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7398 lines
256 KiB
7398 lines
256 KiB
// SPDX-License-Identifier: GPL-2.0-or-later |
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/* |
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* Budget Fair Queueing (BFQ) I/O scheduler. |
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* |
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* Based on ideas and code from CFQ: |
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* Copyright (C) 2003 Jens Axboe <[email protected]> |
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* |
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* Copyright (C) 2008 Fabio Checconi <[email protected]> |
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* Paolo Valente <[email protected]> |
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* |
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* Copyright (C) 2010 Paolo Valente <[email protected]> |
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* Arianna Avanzini <[email protected]> |
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* |
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* Copyright (C) 2017 Paolo Valente <[email protected]> |
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* |
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* BFQ is a proportional-share I/O scheduler, with some extra |
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* low-latency capabilities. BFQ also supports full hierarchical |
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* scheduling through cgroups. Next paragraphs provide an introduction |
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* on BFQ inner workings. Details on BFQ benefits, usage and |
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* limitations can be found in Documentation/block/bfq-iosched.rst. |
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* |
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* BFQ is a proportional-share storage-I/O scheduling algorithm based |
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* on the slice-by-slice service scheme of CFQ. But BFQ assigns |
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* budgets, measured in number of sectors, to processes instead of |
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* time slices. The device is not granted to the in-service process |
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* for a given time slice, but until it has exhausted its assigned |
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* budget. This change from the time to the service domain enables BFQ |
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* to distribute the device throughput among processes as desired, |
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* without any distortion due to throughput fluctuations, or to device |
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* internal queueing. BFQ uses an ad hoc internal scheduler, called |
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* B-WF2Q+, to schedule processes according to their budgets. More |
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* precisely, BFQ schedules queues associated with processes. Each |
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* process/queue is assigned a user-configurable weight, and B-WF2Q+ |
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* guarantees that each queue receives a fraction of the throughput |
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* proportional to its weight. Thanks to the accurate policy of |
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* B-WF2Q+, BFQ can afford to assign high budgets to I/O-bound |
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* processes issuing sequential requests (to boost the throughput), |
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* and yet guarantee a low latency to interactive and soft real-time |
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* applications. |
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* |
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* In particular, to provide these low-latency guarantees, BFQ |
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* explicitly privileges the I/O of two classes of time-sensitive |
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* applications: interactive and soft real-time. In more detail, BFQ |
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* behaves this way if the low_latency parameter is set (default |
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* configuration). This feature enables BFQ to provide applications in |
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* these classes with a very low latency. |
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* |
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* To implement this feature, BFQ constantly tries to detect whether |
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* the I/O requests in a bfq_queue come from an interactive or a soft |
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* real-time application. For brevity, in these cases, the queue is |
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* said to be interactive or soft real-time. In both cases, BFQ |
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* privileges the service of the queue, over that of non-interactive |
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* and non-soft-real-time queues. This privileging is performed, |
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* mainly, by raising the weight of the queue. So, for brevity, we |
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* call just weight-raising periods the time periods during which a |
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* queue is privileged, because deemed interactive or soft real-time. |
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* |
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* The detection of soft real-time queues/applications is described in |
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* detail in the comments on the function |
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* bfq_bfqq_softrt_next_start. On the other hand, the detection of an |
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* interactive queue works as follows: a queue is deemed interactive |
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* if it is constantly non empty only for a limited time interval, |
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* after which it does become empty. The queue may be deemed |
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* interactive again (for a limited time), if it restarts being |
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* constantly non empty, provided that this happens only after the |
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* queue has remained empty for a given minimum idle time. |
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* |
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* By default, BFQ computes automatically the above maximum time |
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* interval, i.e., the time interval after which a constantly |
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* non-empty queue stops being deemed interactive. Since a queue is |
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* weight-raised while it is deemed interactive, this maximum time |
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* interval happens to coincide with the (maximum) duration of the |
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* weight-raising for interactive queues. |
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* |
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* Finally, BFQ also features additional heuristics for |
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* preserving both a low latency and a high throughput on NCQ-capable, |
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* rotational or flash-based devices, and to get the job done quickly |
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* for applications consisting in many I/O-bound processes. |
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* |
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* NOTE: if the main or only goal, with a given device, is to achieve |
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* the maximum-possible throughput at all times, then do switch off |
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* all low-latency heuristics for that device, by setting low_latency |
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* to 0. |
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* |
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* BFQ is described in [1], where also a reference to the initial, |
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* more theoretical paper on BFQ can be found. The interested reader |
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* can find in the latter paper full details on the main algorithm, as |
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* well as formulas of the guarantees and formal proofs of all the |
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* properties. With respect to the version of BFQ presented in these |
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* papers, this implementation adds a few more heuristics, such as the |
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* ones that guarantee a low latency to interactive and soft real-time |
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* applications, and a hierarchical extension based on H-WF2Q+. |
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* |
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* B-WF2Q+ is based on WF2Q+, which is described in [2], together with |
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* H-WF2Q+, while the augmented tree used here to implement B-WF2Q+ |
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* with O(log N) complexity derives from the one introduced with EEVDF |
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* in [3]. |
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* |
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* [1] P. Valente, A. Avanzini, "Evolution of the BFQ Storage I/O |
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* Scheduler", Proceedings of the First Workshop on Mobile System |
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* Technologies (MST-2015), May 2015. |
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* http://algogroup.unimore.it/people/paolo/disk_sched/mst-2015.pdf |
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* |
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* [2] Jon C.R. Bennett and H. Zhang, "Hierarchical Packet Fair Queueing |
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* Algorithms", IEEE/ACM Transactions on Networking, 5(5):675-689, |
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* Oct 1997. |
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* |
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* http://www.cs.cmu.edu/~hzhang/papers/TON-97-Oct.ps.gz |
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* |
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* [3] I. Stoica and H. Abdel-Wahab, "Earliest Eligible Virtual Deadline |
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* First: A Flexible and Accurate Mechanism for Proportional Share |
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* Resource Allocation", technical report. |
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* |
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* http://www.cs.berkeley.edu/~istoica/papers/eevdf-tr-95.pdf |
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*/ |
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#include <linux/module.h> |
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#include <linux/slab.h> |
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#include <linux/blkdev.h> |
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#include <linux/cgroup.h> |
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#include <linux/elevator.h> |
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#include <linux/ktime.h> |
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#include <linux/rbtree.h> |
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#include <linux/ioprio.h> |
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#include <linux/sbitmap.h> |
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#include <linux/delay.h> |
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#include <linux/backing-dev.h> |
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|
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#include <trace/events/block.h> |
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|
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#include "blk.h" |
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#include "blk-mq.h" |
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#include "blk-mq-tag.h" |
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#include "blk-mq-sched.h" |
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#include "bfq-iosched.h" |
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#include "blk-wbt.h" |
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|
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#define BFQ_BFQQ_FNS(name) \ |
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void bfq_mark_bfqq_##name(struct bfq_queue *bfqq) \ |
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{ \ |
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__set_bit(BFQQF_##name, &(bfqq)->flags); \ |
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} \ |
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void bfq_clear_bfqq_##name(struct bfq_queue *bfqq) \ |
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{ \ |
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__clear_bit(BFQQF_##name, &(bfqq)->flags); \ |
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} \ |
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int bfq_bfqq_##name(const struct bfq_queue *bfqq) \ |
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{ \ |
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return test_bit(BFQQF_##name, &(bfqq)->flags); \ |
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} |
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|
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BFQ_BFQQ_FNS(just_created); |
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BFQ_BFQQ_FNS(busy); |
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BFQ_BFQQ_FNS(wait_request); |
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BFQ_BFQQ_FNS(non_blocking_wait_rq); |
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BFQ_BFQQ_FNS(fifo_expire); |
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BFQ_BFQQ_FNS(has_short_ttime); |
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BFQ_BFQQ_FNS(sync); |
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BFQ_BFQQ_FNS(IO_bound); |
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BFQ_BFQQ_FNS(in_large_burst); |
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BFQ_BFQQ_FNS(coop); |
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BFQ_BFQQ_FNS(split_coop); |
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BFQ_BFQQ_FNS(softrt_update); |
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#undef BFQ_BFQQ_FNS \ |
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|
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/* Expiration time of async (0) and sync (1) requests, in ns. */ |
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static const u64 bfq_fifo_expire[2] = { NSEC_PER_SEC / 4, NSEC_PER_SEC / 8 }; |
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/* Maximum backwards seek (magic number lifted from CFQ), in KiB. */ |
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static const int bfq_back_max = 16 * 1024; |
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|
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/* Penalty of a backwards seek, in number of sectors. */ |
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static const int bfq_back_penalty = 2; |
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|
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/* Idling period duration, in ns. */ |
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static u64 bfq_slice_idle = NSEC_PER_SEC / 125; |
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/* Minimum number of assigned budgets for which stats are safe to compute. */ |
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static const int bfq_stats_min_budgets = 194; |
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|
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/* Default maximum budget values, in sectors and number of requests. */ |
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static const int bfq_default_max_budget = 16 * 1024; |
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|
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/* |
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* When a sync request is dispatched, the queue that contains that |
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* request, and all the ancestor entities of that queue, are charged |
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* with the number of sectors of the request. In contrast, if the |
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* request is async, then the queue and its ancestor entities are |
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* charged with the number of sectors of the request, multiplied by |
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* the factor below. This throttles the bandwidth for async I/O, |
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* w.r.t. to sync I/O, and it is done to counter the tendency of async |
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* writes to steal I/O throughput to reads. |
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* |
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* The current value of this parameter is the result of a tuning with |
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* several hardware and software configurations. We tried to find the |
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* lowest value for which writes do not cause noticeable problems to |
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* reads. In fact, the lower this parameter, the stabler I/O control, |
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* in the following respect. The lower this parameter is, the less |
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* the bandwidth enjoyed by a group decreases |
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* - when the group does writes, w.r.t. to when it does reads; |
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* - when other groups do reads, w.r.t. to when they do writes. |
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*/ |
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static const int bfq_async_charge_factor = 3; |
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|
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/* Default timeout values, in jiffies, approximating CFQ defaults. */ |
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const int bfq_timeout = HZ / 8; |
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|
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/* |
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* Time limit for merging (see comments in bfq_setup_cooperator). Set |
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* to the slowest value that, in our tests, proved to be effective in |
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* removing false positives, while not causing true positives to miss |
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* queue merging. |
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* |
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* As can be deduced from the low time limit below, queue merging, if |
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* successful, happens at the very beginning of the I/O of the involved |
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* cooperating processes, as a consequence of the arrival of the very |
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* first requests from each cooperator. After that, there is very |
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* little chance to find cooperators. |
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*/ |
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static const unsigned long bfq_merge_time_limit = HZ/10; |
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static struct kmem_cache *bfq_pool; |
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|
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/* Below this threshold (in ns), we consider thinktime immediate. */ |
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#define BFQ_MIN_TT (2 * NSEC_PER_MSEC) |
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|
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/* hw_tag detection: parallel requests threshold and min samples needed. */ |
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#define BFQ_HW_QUEUE_THRESHOLD 3 |
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#define BFQ_HW_QUEUE_SAMPLES 32 |
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#define BFQQ_SEEK_THR (sector_t)(8 * 100) |
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#define BFQQ_SECT_THR_NONROT (sector_t)(2 * 32) |
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#define BFQ_RQ_SEEKY(bfqd, last_pos, rq) \ |
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(get_sdist(last_pos, rq) > \ |
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BFQQ_SEEK_THR && \ |
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(!blk_queue_nonrot(bfqd->queue) || \ |
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blk_rq_sectors(rq) < BFQQ_SECT_THR_NONROT)) |
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#define BFQQ_CLOSE_THR (sector_t)(8 * 1024) |
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#define BFQQ_SEEKY(bfqq) (hweight32(bfqq->seek_history) > 19) |
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/* |
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* Sync random I/O is likely to be confused with soft real-time I/O, |
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* because it is characterized by limited throughput and apparently |
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* isochronous arrival pattern. To avoid false positives, queues |
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* containing only random (seeky) I/O are prevented from being tagged |
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* as soft real-time. |
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*/ |
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#define BFQQ_TOTALLY_SEEKY(bfqq) (bfqq->seek_history == -1) |
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|
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/* Min number of samples required to perform peak-rate update */ |
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#define BFQ_RATE_MIN_SAMPLES 32 |
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/* Min observation time interval required to perform a peak-rate update (ns) */ |
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#define BFQ_RATE_MIN_INTERVAL (300*NSEC_PER_MSEC) |
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/* Target observation time interval for a peak-rate update (ns) */ |
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#define BFQ_RATE_REF_INTERVAL NSEC_PER_SEC |
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|
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/* |
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* Shift used for peak-rate fixed precision calculations. |
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* With |
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* - the current shift: 16 positions |
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* - the current type used to store rate: u32 |
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* - the current unit of measure for rate: [sectors/usec], or, more precisely, |
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* [(sectors/usec) / 2^BFQ_RATE_SHIFT] to take into account the shift, |
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* the range of rates that can be stored is |
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* [1 / 2^BFQ_RATE_SHIFT, 2^(32 - BFQ_RATE_SHIFT)] sectors/usec = |
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* [1 / 2^16, 2^16] sectors/usec = [15e-6, 65536] sectors/usec = |
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* [15, 65G] sectors/sec |
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* Which, assuming a sector size of 512B, corresponds to a range of |
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* [7.5K, 33T] B/sec |
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*/ |
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#define BFQ_RATE_SHIFT 16 |
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/* |
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* When configured for computing the duration of the weight-raising |
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* for interactive queues automatically (see the comments at the |
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* beginning of this file), BFQ does it using the following formula: |
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* duration = (ref_rate / r) * ref_wr_duration, |
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* where r is the peak rate of the device, and ref_rate and |
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* ref_wr_duration are two reference parameters. In particular, |
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* ref_rate is the peak rate of the reference storage device (see |
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* below), and ref_wr_duration is about the maximum time needed, with |
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* BFQ and while reading two files in parallel, to load typical large |
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* applications on the reference device (see the comments on |
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* max_service_from_wr below, for more details on how ref_wr_duration |
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* is obtained). In practice, the slower/faster the device at hand |
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* is, the more/less it takes to load applications with respect to the |
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* reference device. Accordingly, the longer/shorter BFQ grants |
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* weight raising to interactive applications. |
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* |
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* BFQ uses two different reference pairs (ref_rate, ref_wr_duration), |
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* depending on whether the device is rotational or non-rotational. |
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* |
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* In the following definitions, ref_rate[0] and ref_wr_duration[0] |
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* are the reference values for a rotational device, whereas |
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* ref_rate[1] and ref_wr_duration[1] are the reference values for a |
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* non-rotational device. The reference rates are not the actual peak |
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* rates of the devices used as a reference, but slightly lower |
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* values. The reason for using slightly lower values is that the |
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* peak-rate estimator tends to yield slightly lower values than the |
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* actual peak rate (it can yield the actual peak rate only if there |
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* is only one process doing I/O, and the process does sequential |
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* I/O). |
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* |
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* The reference peak rates are measured in sectors/usec, left-shifted |
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* by BFQ_RATE_SHIFT. |
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*/ |
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static int ref_rate[2] = {14000, 33000}; |
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/* |
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* To improve readability, a conversion function is used to initialize |
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* the following array, which entails that the array can be |
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* initialized only in a function. |
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*/ |
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static int ref_wr_duration[2]; |
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/* |
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* BFQ uses the above-detailed, time-based weight-raising mechanism to |
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* privilege interactive tasks. This mechanism is vulnerable to the |
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* following false positives: I/O-bound applications that will go on |
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* doing I/O for much longer than the duration of weight |
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* raising. These applications have basically no benefit from being |
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* weight-raised at the beginning of their I/O. On the opposite end, |
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* while being weight-raised, these applications |
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* a) unjustly steal throughput to applications that may actually need |
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* low latency; |
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* b) make BFQ uselessly perform device idling; device idling results |
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* in loss of device throughput with most flash-based storage, and may |
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* increase latencies when used purposelessly. |
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* |
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* BFQ tries to reduce these problems, by adopting the following |
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* countermeasure. To introduce this countermeasure, we need first to |
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* finish explaining how the duration of weight-raising for |
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* interactive tasks is computed. |
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* |
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* For a bfq_queue deemed as interactive, the duration of weight |
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* raising is dynamically adjusted, as a function of the estimated |
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* peak rate of the device, so as to be equal to the time needed to |
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* execute the 'largest' interactive task we benchmarked so far. By |
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* largest task, we mean the task for which each involved process has |
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* to do more I/O than for any of the other tasks we benchmarked. This |
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* reference interactive task is the start-up of LibreOffice Writer, |
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* and in this task each process/bfq_queue needs to have at most ~110K |
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* sectors transferred. |
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* |
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* This last piece of information enables BFQ to reduce the actual |
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* duration of weight-raising for at least one class of I/O-bound |
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* applications: those doing sequential or quasi-sequential I/O. An |
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* example is file copy. In fact, once started, the main I/O-bound |
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* processes of these applications usually consume the above 110K |
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* sectors in much less time than the processes of an application that |
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* is starting, because these I/O-bound processes will greedily devote |
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* almost all their CPU cycles only to their target, |
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* throughput-friendly I/O operations. This is even more true if BFQ |
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* happens to be underestimating the device peak rate, and thus |
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* overestimating the duration of weight raising. But, according to |
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* our measurements, once transferred 110K sectors, these processes |
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* have no right to be weight-raised any longer. |
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* |
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* Basing on the last consideration, BFQ ends weight-raising for a |
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* bfq_queue if the latter happens to have received an amount of |
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* service at least equal to the following constant. The constant is |
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* set to slightly more than 110K, to have a minimum safety margin. |
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* |
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* This early ending of weight-raising reduces the amount of time |
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* during which interactive false positives cause the two problems |
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* described at the beginning of these comments. |
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*/ |
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static const unsigned long max_service_from_wr = 120000; |
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|
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/* |
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* Maximum time between the creation of two queues, for stable merge |
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* to be activated (in ms) |
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*/ |
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static const unsigned long bfq_activation_stable_merging = 600; |
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/* |
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* Minimum time to be waited before evaluating delayed stable merge (in ms) |
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*/ |
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static const unsigned long bfq_late_stable_merging = 600; |
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#define RQ_BIC(rq) icq_to_bic((rq)->elv.priv[0]) |
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#define RQ_BFQQ(rq) ((rq)->elv.priv[1]) |
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struct bfq_queue *bic_to_bfqq(struct bfq_io_cq *bic, bool is_sync) |
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{ |
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return bic->bfqq[is_sync]; |
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} |
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static void bfq_put_stable_ref(struct bfq_queue *bfqq); |
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|
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void bic_set_bfqq(struct bfq_io_cq *bic, struct bfq_queue *bfqq, bool is_sync) |
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{ |
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/* |
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* If bfqq != NULL, then a non-stable queue merge between |
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* bic->bfqq and bfqq is happening here. This causes troubles |
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* in the following case: bic->bfqq has also been scheduled |
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* for a possible stable merge with bic->stable_merge_bfqq, |
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* and bic->stable_merge_bfqq == bfqq happens to |
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* hold. Troubles occur because bfqq may then undergo a split, |
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* thereby becoming eligible for a stable merge. Yet, if |
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* bic->stable_merge_bfqq points exactly to bfqq, then bfqq |
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* would be stably merged with itself. To avoid this anomaly, |
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* we cancel the stable merge if |
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* bic->stable_merge_bfqq == bfqq. |
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*/ |
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bic->bfqq[is_sync] = bfqq; |
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|
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if (bfqq && bic->stable_merge_bfqq == bfqq) { |
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/* |
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* Actually, these same instructions are executed also |
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* in bfq_setup_cooperator, in case of abort or actual |
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* execution of a stable merge. We could avoid |
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* repeating these instructions there too, but if we |
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* did so, we would nest even more complexity in this |
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* function. |
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*/ |
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bfq_put_stable_ref(bic->stable_merge_bfqq); |
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|
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bic->stable_merge_bfqq = NULL; |
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} |
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} |
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struct bfq_data *bic_to_bfqd(struct bfq_io_cq *bic) |
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{ |
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return bic->icq.q->elevator->elevator_data; |
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} |
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|
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/** |
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* icq_to_bic - convert iocontext queue structure to bfq_io_cq. |
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* @icq: the iocontext queue. |
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*/ |
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static struct bfq_io_cq *icq_to_bic(struct io_cq *icq) |
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{ |
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/* bic->icq is the first member, %NULL will convert to %NULL */ |
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return container_of(icq, struct bfq_io_cq, icq); |
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} |
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|
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/** |
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* bfq_bic_lookup - search into @ioc a bic associated to @bfqd. |
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* @bfqd: the lookup key. |
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* @ioc: the io_context of the process doing I/O. |
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* @q: the request queue. |
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*/ |
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static struct bfq_io_cq *bfq_bic_lookup(struct bfq_data *bfqd, |
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struct io_context *ioc, |
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struct request_queue *q) |
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{ |
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if (ioc) { |
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unsigned long flags; |
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struct bfq_io_cq *icq; |
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|
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spin_lock_irqsave(&q->queue_lock, flags); |
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icq = icq_to_bic(ioc_lookup_icq(ioc, q)); |
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spin_unlock_irqrestore(&q->queue_lock, flags); |
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return icq; |
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} |
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return NULL; |
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} |
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|
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/* |
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* Scheduler run of queue, if there are requests pending and no one in the |
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* driver that will restart queueing. |
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*/ |
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void bfq_schedule_dispatch(struct bfq_data *bfqd) |
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{ |
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if (bfqd->queued != 0) { |
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bfq_log(bfqd, "schedule dispatch"); |
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blk_mq_run_hw_queues(bfqd->queue, true); |
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} |
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} |
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|
|
#define bfq_class_idle(bfqq) ((bfqq)->ioprio_class == IOPRIO_CLASS_IDLE) |
|
|
|
#define bfq_sample_valid(samples) ((samples) > 80) |
|
|
|
/* |
|
* Lifted from AS - choose which of rq1 and rq2 that is best served now. |
|
* We choose the request that is closer to the head right now. Distance |
|
* behind the head is penalized and only allowed to a certain extent. |
|
*/ |
|
static struct request *bfq_choose_req(struct bfq_data *bfqd, |
|
struct request *rq1, |
|
struct request *rq2, |
|
sector_t last) |
|
{ |
|
sector_t s1, s2, d1 = 0, d2 = 0; |
|
unsigned long back_max; |
|
#define BFQ_RQ1_WRAP 0x01 /* request 1 wraps */ |
|
#define BFQ_RQ2_WRAP 0x02 /* request 2 wraps */ |
|
unsigned int wrap = 0; /* bit mask: requests behind the disk head? */ |
|
|
|
if (!rq1 || rq1 == rq2) |
|
return rq2; |
|
if (!rq2) |
|
return rq1; |
|
|
|
if (rq_is_sync(rq1) && !rq_is_sync(rq2)) |
|
return rq1; |
|
else if (rq_is_sync(rq2) && !rq_is_sync(rq1)) |
|
return rq2; |
|
if ((rq1->cmd_flags & REQ_META) && !(rq2->cmd_flags & REQ_META)) |
|
return rq1; |
|
else if ((rq2->cmd_flags & REQ_META) && !(rq1->cmd_flags & REQ_META)) |
|
return rq2; |
|
|
|
s1 = blk_rq_pos(rq1); |
|
s2 = blk_rq_pos(rq2); |
|
|
|
/* |
|
* By definition, 1KiB is 2 sectors. |
|
*/ |
|
back_max = bfqd->bfq_back_max * 2; |
|
|
|
/* |
|
* Strict one way elevator _except_ in the case where we allow |
|
* short backward seeks which are biased as twice the cost of a |
|
* similar forward seek. |
|
*/ |
|
if (s1 >= last) |
|
d1 = s1 - last; |
|
else if (s1 + back_max >= last) |
|
d1 = (last - s1) * bfqd->bfq_back_penalty; |
|
else |
|
wrap |= BFQ_RQ1_WRAP; |
|
|
|
if (s2 >= last) |
|
d2 = s2 - last; |
|
else if (s2 + back_max >= last) |
|
d2 = (last - s2) * bfqd->bfq_back_penalty; |
|
else |
|
wrap |= BFQ_RQ2_WRAP; |
|
|
|
/* Found required data */ |
|
|
|
/* |
|
* By doing switch() on the bit mask "wrap" we avoid having to |
|
* check two variables for all permutations: --> faster! |
|
*/ |
|
switch (wrap) { |
|
case 0: /* common case for CFQ: rq1 and rq2 not wrapped */ |
|
if (d1 < d2) |
|
return rq1; |
|
else if (d2 < d1) |
|
return rq2; |
|
|
|
if (s1 >= s2) |
|
return rq1; |
|
else |
|
return rq2; |
|
|
|
case BFQ_RQ2_WRAP: |
|
return rq1; |
|
case BFQ_RQ1_WRAP: |
|
return rq2; |
|
case BFQ_RQ1_WRAP|BFQ_RQ2_WRAP: /* both rqs wrapped */ |
|
default: |
|
/* |
|
* Since both rqs are wrapped, |
|
* start with the one that's further behind head |
|
* (--> only *one* back seek required), |
|
* since back seek takes more time than forward. |
|
*/ |
|
if (s1 <= s2) |
|
return rq1; |
|
else |
|
return rq2; |
|
} |
|
} |
|
|
|
/* |
|
* Async I/O can easily starve sync I/O (both sync reads and sync |
|
* writes), by consuming all tags. Similarly, storms of sync writes, |
|
* such as those that sync(2) may trigger, can starve sync reads. |
|
* Limit depths of async I/O and sync writes so as to counter both |
|
* problems. |
|
*/ |
|
static void bfq_limit_depth(unsigned int op, struct blk_mq_alloc_data *data) |
|
{ |
|
struct bfq_data *bfqd = data->q->elevator->elevator_data; |
|
|
|
if (op_is_sync(op) && !op_is_write(op)) |
|
return; |
|
|
|
data->shallow_depth = |
|
bfqd->word_depths[!!bfqd->wr_busy_queues][op_is_sync(op)]; |
|
|
|
bfq_log(bfqd, "[%s] wr_busy %d sync %d depth %u", |
|
__func__, bfqd->wr_busy_queues, op_is_sync(op), |
|
data->shallow_depth); |
|
} |
|
|
|
static struct bfq_queue * |
|
bfq_rq_pos_tree_lookup(struct bfq_data *bfqd, struct rb_root *root, |
|
sector_t sector, struct rb_node **ret_parent, |
|
struct rb_node ***rb_link) |
|
{ |
|
struct rb_node **p, *parent; |
|
struct bfq_queue *bfqq = NULL; |
|
|
|
parent = NULL; |
|
p = &root->rb_node; |
|
while (*p) { |
|
struct rb_node **n; |
|
|
|
parent = *p; |
|
bfqq = rb_entry(parent, struct bfq_queue, pos_node); |
|
|
|
/* |
|
* Sort strictly based on sector. Smallest to the left, |
|
* largest to the right. |
|
*/ |
|
if (sector > blk_rq_pos(bfqq->next_rq)) |
|
n = &(*p)->rb_right; |
|
else if (sector < blk_rq_pos(bfqq->next_rq)) |
|
n = &(*p)->rb_left; |
|
else |
|
break; |
|
p = n; |
|
bfqq = NULL; |
|
} |
|
|
|
*ret_parent = parent; |
|
if (rb_link) |
|
*rb_link = p; |
|
|
|
bfq_log(bfqd, "rq_pos_tree_lookup %llu: returning %d", |
|
(unsigned long long)sector, |
|
bfqq ? bfqq->pid : 0); |
|
|
|
return bfqq; |
|
} |
|
|
|
static bool bfq_too_late_for_merging(struct bfq_queue *bfqq) |
|
{ |
|
return bfqq->service_from_backlogged > 0 && |
|
time_is_before_jiffies(bfqq->first_IO_time + |
|
bfq_merge_time_limit); |
|
} |
|
|
|
/* |
|
* The following function is not marked as __cold because it is |
|
* actually cold, but for the same performance goal described in the |
|
* comments on the likely() at the beginning of |
|
* bfq_setup_cooperator(). Unexpectedly, to reach an even lower |
|
* execution time for the case where this function is not invoked, we |
|
* had to add an unlikely() in each involved if(). |
|
*/ |
|
void __cold |
|
bfq_pos_tree_add_move(struct bfq_data *bfqd, struct bfq_queue *bfqq) |
|
{ |
|
struct rb_node **p, *parent; |
|
struct bfq_queue *__bfqq; |
|
|
|
if (bfqq->pos_root) { |
|
rb_erase(&bfqq->pos_node, bfqq->pos_root); |
|
bfqq->pos_root = NULL; |
|
} |
|
|
|
/* oom_bfqq does not participate in queue merging */ |
|
if (bfqq == &bfqd->oom_bfqq) |
|
return; |
|
|
|
/* |
|
* bfqq cannot be merged any longer (see comments in |
|
* bfq_setup_cooperator): no point in adding bfqq into the |
|
* position tree. |
|
*/ |
|
if (bfq_too_late_for_merging(bfqq)) |
|
return; |
|
|
|
if (bfq_class_idle(bfqq)) |
|
return; |
|
if (!bfqq->next_rq) |
|
return; |
|
|
|
bfqq->pos_root = &bfq_bfqq_to_bfqg(bfqq)->rq_pos_tree; |
|
__bfqq = bfq_rq_pos_tree_lookup(bfqd, bfqq->pos_root, |
|
blk_rq_pos(bfqq->next_rq), &parent, &p); |
|
if (!__bfqq) { |
|
rb_link_node(&bfqq->pos_node, parent, p); |
|
rb_insert_color(&bfqq->pos_node, bfqq->pos_root); |
|
} else |
|
bfqq->pos_root = NULL; |
|
} |
|
|
|
/* |
|
* The following function returns false either if every active queue |
|
* must receive the same share of the throughput (symmetric scenario), |
|
* or, as a special case, if bfqq must receive a share of the |
|
* throughput lower than or equal to the share that every other active |
|
* queue must receive. If bfqq does sync I/O, then these are the only |
|
* two cases where bfqq happens to be guaranteed its share of the |
|
* throughput even if I/O dispatching is not plugged when bfqq remains |
|
* temporarily empty (for more details, see the comments in the |
|
* function bfq_better_to_idle()). For this reason, the return value |
|
* of this function is used to check whether I/O-dispatch plugging can |
|
* be avoided. |
|
* |
|
* The above first case (symmetric scenario) occurs when: |
|
* 1) all active queues have the same weight, |
|
* 2) all active queues belong to the same I/O-priority class, |
|
* 3) all active groups at the same level in the groups tree have the same |
|
* weight, |
|
* 4) all active groups at the same level in the groups tree have the same |
|
* number of children. |
|
* |
|
* Unfortunately, keeping the necessary state for evaluating exactly |
|
* the last two symmetry sub-conditions above would be quite complex |
|
* and time consuming. Therefore this function evaluates, instead, |
|
* only the following stronger three sub-conditions, for which it is |
|
* much easier to maintain the needed state: |
|
* 1) all active queues have the same weight, |
|
* 2) all active queues belong to the same I/O-priority class, |
|
* 3) there are no active groups. |
|
* In particular, the last condition is always true if hierarchical |
|
* support or the cgroups interface are not enabled, thus no state |
|
* needs to be maintained in this case. |
|
*/ |
|
static bool bfq_asymmetric_scenario(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
bool smallest_weight = bfqq && |
|
bfqq->weight_counter && |
|
bfqq->weight_counter == |
|
container_of( |
|
rb_first_cached(&bfqd->queue_weights_tree), |
|
struct bfq_weight_counter, |
|
weights_node); |
|
|
|
/* |
|
* For queue weights to differ, queue_weights_tree must contain |
|
* at least two nodes. |
|
*/ |
|
bool varied_queue_weights = !smallest_weight && |
|
!RB_EMPTY_ROOT(&bfqd->queue_weights_tree.rb_root) && |
|
(bfqd->queue_weights_tree.rb_root.rb_node->rb_left || |
|
bfqd->queue_weights_tree.rb_root.rb_node->rb_right); |
|
|
|
bool multiple_classes_busy = |
|
(bfqd->busy_queues[0] && bfqd->busy_queues[1]) || |
|
(bfqd->busy_queues[0] && bfqd->busy_queues[2]) || |
|
(bfqd->busy_queues[1] && bfqd->busy_queues[2]); |
|
|
|
return varied_queue_weights || multiple_classes_busy |
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED |
|
|| bfqd->num_groups_with_pending_reqs > 0 |
|
#endif |
|
; |
|
} |
|
|
|
/* |
|
* If the weight-counter tree passed as input contains no counter for |
|
* the weight of the input queue, then add that counter; otherwise just |
|
* increment the existing counter. |
|
* |
|
* Note that weight-counter trees contain few nodes in mostly symmetric |
|
* scenarios. For example, if all queues have the same weight, then the |
|
* weight-counter tree for the queues may contain at most one node. |
|
* This holds even if low_latency is on, because weight-raised queues |
|
* are not inserted in the tree. |
|
* In most scenarios, the rate at which nodes are created/destroyed |
|
* should be low too. |
|
*/ |
|
void bfq_weights_tree_add(struct bfq_data *bfqd, struct bfq_queue *bfqq, |
|
struct rb_root_cached *root) |
|
{ |
|
struct bfq_entity *entity = &bfqq->entity; |
|
struct rb_node **new = &(root->rb_root.rb_node), *parent = NULL; |
|
bool leftmost = true; |
|
|
|
/* |
|
* Do not insert if the queue is already associated with a |
|
* counter, which happens if: |
|
* 1) a request arrival has caused the queue to become both |
|
* non-weight-raised, and hence change its weight, and |
|
* backlogged; in this respect, each of the two events |
|
* causes an invocation of this function, |
|
* 2) this is the invocation of this function caused by the |
|
* second event. This second invocation is actually useless, |
|
* and we handle this fact by exiting immediately. More |
|
* efficient or clearer solutions might possibly be adopted. |
|
*/ |
|
if (bfqq->weight_counter) |
|
return; |
|
|
|
while (*new) { |
|
struct bfq_weight_counter *__counter = container_of(*new, |
|
struct bfq_weight_counter, |
|
weights_node); |
|
parent = *new; |
|
|
|
if (entity->weight == __counter->weight) { |
|
bfqq->weight_counter = __counter; |
|
goto inc_counter; |
|
} |
|
if (entity->weight < __counter->weight) |
|
new = &((*new)->rb_left); |
|
else { |
|
new = &((*new)->rb_right); |
|
leftmost = false; |
|
} |
|
} |
|
|
|
bfqq->weight_counter = kzalloc(sizeof(struct bfq_weight_counter), |
|
GFP_ATOMIC); |
|
|
|
/* |
|
* In the unlucky event of an allocation failure, we just |
|
* exit. This will cause the weight of queue to not be |
|
* considered in bfq_asymmetric_scenario, which, in its turn, |
|
* causes the scenario to be deemed wrongly symmetric in case |
|
* bfqq's weight would have been the only weight making the |
|
* scenario asymmetric. On the bright side, no unbalance will |
|
* however occur when bfqq becomes inactive again (the |
|
* invocation of this function is triggered by an activation |
|
* of queue). In fact, bfq_weights_tree_remove does nothing |
|
* if !bfqq->weight_counter. |
|
*/ |
|
if (unlikely(!bfqq->weight_counter)) |
|
return; |
|
|
|
bfqq->weight_counter->weight = entity->weight; |
|
rb_link_node(&bfqq->weight_counter->weights_node, parent, new); |
|
rb_insert_color_cached(&bfqq->weight_counter->weights_node, root, |
|
leftmost); |
|
|
|
inc_counter: |
|
bfqq->weight_counter->num_active++; |
|
bfqq->ref++; |
|
} |
|
|
|
/* |
|
* Decrement the weight counter associated with the queue, and, if the |
|
* counter reaches 0, remove the counter from the tree. |
|
* See the comments to the function bfq_weights_tree_add() for considerations |
|
* about overhead. |
|
*/ |
|
void __bfq_weights_tree_remove(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq, |
|
struct rb_root_cached *root) |
|
{ |
|
if (!bfqq->weight_counter) |
|
return; |
|
|
|
bfqq->weight_counter->num_active--; |
|
if (bfqq->weight_counter->num_active > 0) |
|
goto reset_entity_pointer; |
|
|
|
rb_erase_cached(&bfqq->weight_counter->weights_node, root); |
|
kfree(bfqq->weight_counter); |
|
|
|
reset_entity_pointer: |
|
bfqq->weight_counter = NULL; |
|
bfq_put_queue(bfqq); |
|
} |
|
|
|
/* |
|
* Invoke __bfq_weights_tree_remove on bfqq and decrement the number |
|
* of active groups for each queue's inactive parent entity. |
|
*/ |
|
void bfq_weights_tree_remove(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
struct bfq_entity *entity = bfqq->entity.parent; |
|
|
|
for_each_entity(entity) { |
|
struct bfq_sched_data *sd = entity->my_sched_data; |
|
|
|
if (sd->next_in_service || sd->in_service_entity) { |
|
/* |
|
* entity is still active, because either |
|
* next_in_service or in_service_entity is not |
|
* NULL (see the comments on the definition of |
|
* next_in_service for details on why |
|
* in_service_entity must be checked too). |
|
* |
|
* As a consequence, its parent entities are |
|
* active as well, and thus this loop must |
|
* stop here. |
|
*/ |
|
break; |
|
} |
|
|
|
/* |
|
* The decrement of num_groups_with_pending_reqs is |
|
* not performed immediately upon the deactivation of |
|
* entity, but it is delayed to when it also happens |
|
* that the first leaf descendant bfqq of entity gets |
|
* all its pending requests completed. The following |
|
* instructions perform this delayed decrement, if |
|
* needed. See the comments on |
|
* num_groups_with_pending_reqs for details. |
|
*/ |
|
if (entity->in_groups_with_pending_reqs) { |
|
entity->in_groups_with_pending_reqs = false; |
|
bfqd->num_groups_with_pending_reqs--; |
|
} |
|
} |
|
|
|
/* |
|
* Next function is invoked last, because it causes bfqq to be |
|
* freed if the following holds: bfqq is not in service and |
|
* has no dispatched request. DO NOT use bfqq after the next |
|
* function invocation. |
|
*/ |
|
__bfq_weights_tree_remove(bfqd, bfqq, |
|
&bfqd->queue_weights_tree); |
|
} |
|
|
|
/* |
|
* Return expired entry, or NULL to just start from scratch in rbtree. |
|
*/ |
|
static struct request *bfq_check_fifo(struct bfq_queue *bfqq, |
|
struct request *last) |
|
{ |
|
struct request *rq; |
|
|
|
if (bfq_bfqq_fifo_expire(bfqq)) |
|
return NULL; |
|
|
|
bfq_mark_bfqq_fifo_expire(bfqq); |
|
|
|
rq = rq_entry_fifo(bfqq->fifo.next); |
|
|
|
if (rq == last || ktime_get_ns() < rq->fifo_time) |
|
return NULL; |
|
|
|
bfq_log_bfqq(bfqq->bfqd, bfqq, "check_fifo: returned %p", rq); |
|
return rq; |
|
} |
|
|
|
static struct request *bfq_find_next_rq(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq, |
|
struct request *last) |
|
{ |
|
struct rb_node *rbnext = rb_next(&last->rb_node); |
|
struct rb_node *rbprev = rb_prev(&last->rb_node); |
|
struct request *next, *prev = NULL; |
|
|
|
/* Follow expired path, else get first next available. */ |
|
next = bfq_check_fifo(bfqq, last); |
|
if (next) |
|
return next; |
|
|
|
if (rbprev) |
|
prev = rb_entry_rq(rbprev); |
|
|
|
if (rbnext) |
|
next = rb_entry_rq(rbnext); |
|
else { |
|
rbnext = rb_first(&bfqq->sort_list); |
|
if (rbnext && rbnext != &last->rb_node) |
|
next = rb_entry_rq(rbnext); |
|
} |
|
|
|
return bfq_choose_req(bfqd, next, prev, blk_rq_pos(last)); |
|
} |
|
|
|
/* see the definition of bfq_async_charge_factor for details */ |
|
static unsigned long bfq_serv_to_charge(struct request *rq, |
|
struct bfq_queue *bfqq) |
|
{ |
|
if (bfq_bfqq_sync(bfqq) || bfqq->wr_coeff > 1 || |
|
bfq_asymmetric_scenario(bfqq->bfqd, bfqq)) |
|
return blk_rq_sectors(rq); |
|
|
|
return blk_rq_sectors(rq) * bfq_async_charge_factor; |
|
} |
|
|
|
/** |
|
* bfq_updated_next_req - update the queue after a new next_rq selection. |
|
* @bfqd: the device data the queue belongs to. |
|
* @bfqq: the queue to update. |
|
* |
|
* If the first request of a queue changes we make sure that the queue |
|
* has enough budget to serve at least its first request (if the |
|
* request has grown). We do this because if the queue has not enough |
|
* budget for its first request, it has to go through two dispatch |
|
* rounds to actually get it dispatched. |
|
*/ |
|
static void bfq_updated_next_req(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
struct bfq_entity *entity = &bfqq->entity; |
|
struct request *next_rq = bfqq->next_rq; |
|
unsigned long new_budget; |
|
|
|
if (!next_rq) |
|
return; |
|
|
|
if (bfqq == bfqd->in_service_queue) |
|
/* |
|
* In order not to break guarantees, budgets cannot be |
|
* changed after an entity has been selected. |
|
*/ |
|
return; |
|
|
|
new_budget = max_t(unsigned long, |
|
max_t(unsigned long, bfqq->max_budget, |
|
bfq_serv_to_charge(next_rq, bfqq)), |
|
entity->service); |
|
if (entity->budget != new_budget) { |
|
entity->budget = new_budget; |
|
bfq_log_bfqq(bfqd, bfqq, "updated next rq: new budget %lu", |
|
new_budget); |
|
bfq_requeue_bfqq(bfqd, bfqq, false); |
|
} |
|
} |
|
|
|
static unsigned int bfq_wr_duration(struct bfq_data *bfqd) |
|
{ |
|
u64 dur; |
|
|
|
if (bfqd->bfq_wr_max_time > 0) |
|
return bfqd->bfq_wr_max_time; |
|
|
|
dur = bfqd->rate_dur_prod; |
|
do_div(dur, bfqd->peak_rate); |
|
|
|
/* |
|
* Limit duration between 3 and 25 seconds. The upper limit |
|
* has been conservatively set after the following worst case: |
|
* on a QEMU/KVM virtual machine |
|
* - running in a slow PC |
|
* - with a virtual disk stacked on a slow low-end 5400rpm HDD |
|
* - serving a heavy I/O workload, such as the sequential reading |
|
* of several files |
|
* mplayer took 23 seconds to start, if constantly weight-raised. |
|
* |
|
* As for higher values than that accommodating the above bad |
|
* scenario, tests show that higher values would often yield |
|
* the opposite of the desired result, i.e., would worsen |
|
* responsiveness by allowing non-interactive applications to |
|
* preserve weight raising for too long. |
|
* |
|
* On the other end, lower values than 3 seconds make it |
|
* difficult for most interactive tasks to complete their jobs |
|
* before weight-raising finishes. |
|
*/ |
|
return clamp_val(dur, msecs_to_jiffies(3000), msecs_to_jiffies(25000)); |
|
} |
|
|
|
/* switch back from soft real-time to interactive weight raising */ |
|
static void switch_back_to_interactive_wr(struct bfq_queue *bfqq, |
|
struct bfq_data *bfqd) |
|
{ |
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff; |
|
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd); |
|
bfqq->last_wr_start_finish = bfqq->wr_start_at_switch_to_srt; |
|
} |
|
|
|
static void |
|
bfq_bfqq_resume_state(struct bfq_queue *bfqq, struct bfq_data *bfqd, |
|
struct bfq_io_cq *bic, bool bfq_already_existing) |
|
{ |
|
unsigned int old_wr_coeff = 1; |
|
bool busy = bfq_already_existing && bfq_bfqq_busy(bfqq); |
|
|
|
if (bic->saved_has_short_ttime) |
|
bfq_mark_bfqq_has_short_ttime(bfqq); |
|
else |
|
bfq_clear_bfqq_has_short_ttime(bfqq); |
|
|
|
if (bic->saved_IO_bound) |
|
bfq_mark_bfqq_IO_bound(bfqq); |
|
else |
|
bfq_clear_bfqq_IO_bound(bfqq); |
|
|
|
bfqq->last_serv_time_ns = bic->saved_last_serv_time_ns; |
|
bfqq->inject_limit = bic->saved_inject_limit; |
|
bfqq->decrease_time_jif = bic->saved_decrease_time_jif; |
|
|
|
bfqq->entity.new_weight = bic->saved_weight; |
|
bfqq->ttime = bic->saved_ttime; |
|
bfqq->io_start_time = bic->saved_io_start_time; |
|
bfqq->tot_idle_time = bic->saved_tot_idle_time; |
|
/* |
|
* Restore weight coefficient only if low_latency is on |
|
*/ |
|
if (bfqd->low_latency) { |
|
old_wr_coeff = bfqq->wr_coeff; |
|
bfqq->wr_coeff = bic->saved_wr_coeff; |
|
} |
|
bfqq->service_from_wr = bic->saved_service_from_wr; |
|
bfqq->wr_start_at_switch_to_srt = bic->saved_wr_start_at_switch_to_srt; |
|
bfqq->last_wr_start_finish = bic->saved_last_wr_start_finish; |
|
bfqq->wr_cur_max_time = bic->saved_wr_cur_max_time; |
|
|
|
if (bfqq->wr_coeff > 1 && (bfq_bfqq_in_large_burst(bfqq) || |
|
time_is_before_jiffies(bfqq->last_wr_start_finish + |
|
bfqq->wr_cur_max_time))) { |
|
if (bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time && |
|
!bfq_bfqq_in_large_burst(bfqq) && |
|
time_is_after_eq_jiffies(bfqq->wr_start_at_switch_to_srt + |
|
bfq_wr_duration(bfqd))) { |
|
switch_back_to_interactive_wr(bfqq, bfqd); |
|
} else { |
|
bfqq->wr_coeff = 1; |
|
bfq_log_bfqq(bfqq->bfqd, bfqq, |
|
"resume state: switching off wr"); |
|
} |
|
} |
|
|
|
/* make sure weight will be updated, however we got here */ |
|
bfqq->entity.prio_changed = 1; |
|
|
|
if (likely(!busy)) |
|
return; |
|
|
|
if (old_wr_coeff == 1 && bfqq->wr_coeff > 1) |
|
bfqd->wr_busy_queues++; |
|
else if (old_wr_coeff > 1 && bfqq->wr_coeff == 1) |
|
bfqd->wr_busy_queues--; |
|
} |
|
|
|
static int bfqq_process_refs(struct bfq_queue *bfqq) |
|
{ |
|
return bfqq->ref - bfqq->allocated - bfqq->entity.on_st_or_in_serv - |
|
(bfqq->weight_counter != NULL) - bfqq->stable_ref; |
|
} |
|
|
|
/* Empty burst list and add just bfqq (see comments on bfq_handle_burst) */ |
|
static void bfq_reset_burst_list(struct bfq_data *bfqd, struct bfq_queue *bfqq) |
|
{ |
|
struct bfq_queue *item; |
|
struct hlist_node *n; |
|
|
|
hlist_for_each_entry_safe(item, n, &bfqd->burst_list, burst_list_node) |
|
hlist_del_init(&item->burst_list_node); |
|
|
|
/* |
|
* Start the creation of a new burst list only if there is no |
|
* active queue. See comments on the conditional invocation of |
|
* bfq_handle_burst(). |
|
*/ |
|
if (bfq_tot_busy_queues(bfqd) == 0) { |
|
hlist_add_head(&bfqq->burst_list_node, &bfqd->burst_list); |
|
bfqd->burst_size = 1; |
|
} else |
|
bfqd->burst_size = 0; |
|
|
|
bfqd->burst_parent_entity = bfqq->entity.parent; |
|
} |
|
|
|
/* Add bfqq to the list of queues in current burst (see bfq_handle_burst) */ |
|
static void bfq_add_to_burst(struct bfq_data *bfqd, struct bfq_queue *bfqq) |
|
{ |
|
/* Increment burst size to take into account also bfqq */ |
|
bfqd->burst_size++; |
|
|
|
if (bfqd->burst_size == bfqd->bfq_large_burst_thresh) { |
|
struct bfq_queue *pos, *bfqq_item; |
|
struct hlist_node *n; |
|
|
|
/* |
|
* Enough queues have been activated shortly after each |
|
* other to consider this burst as large. |
|
*/ |
|
bfqd->large_burst = true; |
|
|
|
/* |
|
* We can now mark all queues in the burst list as |
|
* belonging to a large burst. |
|
*/ |
|
hlist_for_each_entry(bfqq_item, &bfqd->burst_list, |
|
burst_list_node) |
|
bfq_mark_bfqq_in_large_burst(bfqq_item); |
|
bfq_mark_bfqq_in_large_burst(bfqq); |
|
|
|
/* |
|
* From now on, and until the current burst finishes, any |
|
* new queue being activated shortly after the last queue |
|
* was inserted in the burst can be immediately marked as |
|
* belonging to a large burst. So the burst list is not |
|
* needed any more. Remove it. |
|
*/ |
|
hlist_for_each_entry_safe(pos, n, &bfqd->burst_list, |
|
burst_list_node) |
|
hlist_del_init(&pos->burst_list_node); |
|
} else /* |
|
* Burst not yet large: add bfqq to the burst list. Do |
|
* not increment the ref counter for bfqq, because bfqq |
|
* is removed from the burst list before freeing bfqq |
|
* in put_queue. |
|
*/ |
|
hlist_add_head(&bfqq->burst_list_node, &bfqd->burst_list); |
|
} |
|
|
|
/* |
|
* If many queues belonging to the same group happen to be created |
|
* shortly after each other, then the processes associated with these |
|
* queues have typically a common goal. In particular, bursts of queue |
|
* creations are usually caused by services or applications that spawn |
|
* many parallel threads/processes. Examples are systemd during boot, |
|
* or git grep. To help these processes get their job done as soon as |
|
* possible, it is usually better to not grant either weight-raising |
|
* or device idling to their queues, unless these queues must be |
|
* protected from the I/O flowing through other active queues. |
|
* |
|
* In this comment we describe, firstly, the reasons why this fact |
|
* holds, and, secondly, the next function, which implements the main |
|
* steps needed to properly mark these queues so that they can then be |
|
* treated in a different way. |
|
* |
|
* The above services or applications benefit mostly from a high |
|
* throughput: the quicker the requests of the activated queues are |
|
* cumulatively served, the sooner the target job of these queues gets |
|
* completed. As a consequence, weight-raising any of these queues, |
|
* which also implies idling the device for it, is almost always |
|
* counterproductive, unless there are other active queues to isolate |
|
* these new queues from. If there no other active queues, then |
|
* weight-raising these new queues just lowers throughput in most |
|
* cases. |
|
* |
|
* On the other hand, a burst of queue creations may be caused also by |
|
* the start of an application that does not consist of a lot of |
|
* parallel I/O-bound threads. In fact, with a complex application, |
|
* several short processes may need to be executed to start-up the |
|
* application. In this respect, to start an application as quickly as |
|
* possible, the best thing to do is in any case to privilege the I/O |
|
* related to the application with respect to all other |
|
* I/O. Therefore, the best strategy to start as quickly as possible |
|
* an application that causes a burst of queue creations is to |
|
* weight-raise all the queues created during the burst. This is the |
|
* exact opposite of the best strategy for the other type of bursts. |
|
* |
|
* In the end, to take the best action for each of the two cases, the |
|
* two types of bursts need to be distinguished. Fortunately, this |
|
* seems relatively easy, by looking at the sizes of the bursts. In |
|
* particular, we found a threshold such that only bursts with a |
|
* larger size than that threshold are apparently caused by |
|
* services or commands such as systemd or git grep. For brevity, |
|
* hereafter we call just 'large' these bursts. BFQ *does not* |
|
* weight-raise queues whose creation occurs in a large burst. In |
|
* addition, for each of these queues BFQ performs or does not perform |
|
* idling depending on which choice boosts the throughput more. The |
|
* exact choice depends on the device and request pattern at |
|
* hand. |
|
* |
|
* Unfortunately, false positives may occur while an interactive task |
|
* is starting (e.g., an application is being started). The |
|
* consequence is that the queues associated with the task do not |
|
* enjoy weight raising as expected. Fortunately these false positives |
|
* are very rare. They typically occur if some service happens to |
|
* start doing I/O exactly when the interactive task starts. |
|
* |
|
* Turning back to the next function, it is invoked only if there are |
|
* no active queues (apart from active queues that would belong to the |
|
* same, possible burst bfqq would belong to), and it implements all |
|
* the steps needed to detect the occurrence of a large burst and to |
|
* properly mark all the queues belonging to it (so that they can then |
|
* be treated in a different way). This goal is achieved by |
|
* maintaining a "burst list" that holds, temporarily, the queues that |
|
* belong to the burst in progress. The list is then used to mark |
|
* these queues as belonging to a large burst if the burst does become |
|
* large. The main steps are the following. |
|
* |
|
* . when the very first queue is created, the queue is inserted into the |
|
* list (as it could be the first queue in a possible burst) |
|
* |
|
* . if the current burst has not yet become large, and a queue Q that does |
|
* not yet belong to the burst is activated shortly after the last time |
|
* at which a new queue entered the burst list, then the function appends |
|
* Q to the burst list |
|
* |
|
* . if, as a consequence of the previous step, the burst size reaches |
|
* the large-burst threshold, then |
|
* |
|
* . all the queues in the burst list are marked as belonging to a |
|
* large burst |
|
* |
|
* . the burst list is deleted; in fact, the burst list already served |
|
* its purpose (keeping temporarily track of the queues in a burst, |
|
* so as to be able to mark them as belonging to a large burst in the |
|
* previous sub-step), and now is not needed any more |
|
* |
|
* . the device enters a large-burst mode |
|
* |
|
* . if a queue Q that does not belong to the burst is created while |
|
* the device is in large-burst mode and shortly after the last time |
|
* at which a queue either entered the burst list or was marked as |
|
* belonging to the current large burst, then Q is immediately marked |
|
* as belonging to a large burst. |
|
* |
|
* . if a queue Q that does not belong to the burst is created a while |
|
* later, i.e., not shortly after, than the last time at which a queue |
|
* either entered the burst list or was marked as belonging to the |
|
* current large burst, then the current burst is deemed as finished and: |
|
* |
|
* . the large-burst mode is reset if set |
|
* |
|
* . the burst list is emptied |
|
* |
|
* . Q is inserted in the burst list, as Q may be the first queue |
|
* in a possible new burst (then the burst list contains just Q |
|
* after this step). |
|
*/ |
|
static void bfq_handle_burst(struct bfq_data *bfqd, struct bfq_queue *bfqq) |
|
{ |
|
/* |
|
* If bfqq is already in the burst list or is part of a large |
|
* burst, or finally has just been split, then there is |
|
* nothing else to do. |
|
*/ |
|
if (!hlist_unhashed(&bfqq->burst_list_node) || |
|
bfq_bfqq_in_large_burst(bfqq) || |
|
time_is_after_eq_jiffies(bfqq->split_time + |
|
msecs_to_jiffies(10))) |
|
return; |
|
|
|
/* |
|
* If bfqq's creation happens late enough, or bfqq belongs to |
|
* a different group than the burst group, then the current |
|
* burst is finished, and related data structures must be |
|
* reset. |
|
* |
|
* In this respect, consider the special case where bfqq is |
|
* the very first queue created after BFQ is selected for this |
|
* device. In this case, last_ins_in_burst and |
|
* burst_parent_entity are not yet significant when we get |
|
* here. But it is easy to verify that, whether or not the |
|
* following condition is true, bfqq will end up being |
|
* inserted into the burst list. In particular the list will |
|
* happen to contain only bfqq. And this is exactly what has |
|
* to happen, as bfqq may be the first queue of the first |
|
* burst. |
|
*/ |
|
if (time_is_before_jiffies(bfqd->last_ins_in_burst + |
|
bfqd->bfq_burst_interval) || |
|
bfqq->entity.parent != bfqd->burst_parent_entity) { |
|
bfqd->large_burst = false; |
|
bfq_reset_burst_list(bfqd, bfqq); |
|
goto end; |
|
} |
|
|
|
/* |
|
* If we get here, then bfqq is being activated shortly after the |
|
* last queue. So, if the current burst is also large, we can mark |
|
* bfqq as belonging to this large burst immediately. |
|
*/ |
|
if (bfqd->large_burst) { |
|
bfq_mark_bfqq_in_large_burst(bfqq); |
|
goto end; |
|
} |
|
|
|
/* |
|
* If we get here, then a large-burst state has not yet been |
|
* reached, but bfqq is being activated shortly after the last |
|
* queue. Then we add bfqq to the burst. |
|
*/ |
|
bfq_add_to_burst(bfqd, bfqq); |
|
end: |
|
/* |
|
* At this point, bfqq either has been added to the current |
|
* burst or has caused the current burst to terminate and a |
|
* possible new burst to start. In particular, in the second |
|
* case, bfqq has become the first queue in the possible new |
|
* burst. In both cases last_ins_in_burst needs to be moved |
|
* forward. |
|
*/ |
|
bfqd->last_ins_in_burst = jiffies; |
|
} |
|
|
|
static int bfq_bfqq_budget_left(struct bfq_queue *bfqq) |
|
{ |
|
struct bfq_entity *entity = &bfqq->entity; |
|
|
|
return entity->budget - entity->service; |
|
} |
|
|
|
/* |
|
* If enough samples have been computed, return the current max budget |
|
* stored in bfqd, which is dynamically updated according to the |
|
* estimated disk peak rate; otherwise return the default max budget |
|
*/ |
|
static int bfq_max_budget(struct bfq_data *bfqd) |
|
{ |
|
if (bfqd->budgets_assigned < bfq_stats_min_budgets) |
|
return bfq_default_max_budget; |
|
else |
|
return bfqd->bfq_max_budget; |
|
} |
|
|
|
/* |
|
* Return min budget, which is a fraction of the current or default |
|
* max budget (trying with 1/32) |
|
*/ |
|
static int bfq_min_budget(struct bfq_data *bfqd) |
|
{ |
|
if (bfqd->budgets_assigned < bfq_stats_min_budgets) |
|
return bfq_default_max_budget / 32; |
|
else |
|
return bfqd->bfq_max_budget / 32; |
|
} |
|
|
|
/* |
|
* The next function, invoked after the input queue bfqq switches from |
|
* idle to busy, updates the budget of bfqq. The function also tells |
|
* whether the in-service queue should be expired, by returning |
|
* true. The purpose of expiring the in-service queue is to give bfqq |
|
* the chance to possibly preempt the in-service queue, and the reason |
|
* for preempting the in-service queue is to achieve one of the two |
|
* goals below. |
|
* |
|
* 1. Guarantee to bfqq its reserved bandwidth even if bfqq has |
|
* expired because it has remained idle. In particular, bfqq may have |
|
* expired for one of the following two reasons: |
|
* |
|
* - BFQQE_NO_MORE_REQUESTS bfqq did not enjoy any device idling |
|
* and did not make it to issue a new request before its last |
|
* request was served; |
|
* |
|
* - BFQQE_TOO_IDLE bfqq did enjoy device idling, but did not issue |
|
* a new request before the expiration of the idling-time. |
|
* |
|
* Even if bfqq has expired for one of the above reasons, the process |
|
* associated with the queue may be however issuing requests greedily, |
|
* and thus be sensitive to the bandwidth it receives (bfqq may have |
|
* remained idle for other reasons: CPU high load, bfqq not enjoying |
|
* idling, I/O throttling somewhere in the path from the process to |
|
* the I/O scheduler, ...). But if, after every expiration for one of |
|
* the above two reasons, bfqq has to wait for the service of at least |
|
* one full budget of another queue before being served again, then |
|
* bfqq is likely to get a much lower bandwidth or resource time than |
|
* its reserved ones. To address this issue, two countermeasures need |
|
* to be taken. |
|
* |
|
* First, the budget and the timestamps of bfqq need to be updated in |
|
* a special way on bfqq reactivation: they need to be updated as if |
|
* bfqq did not remain idle and did not expire. In fact, if they are |
|
* computed as if bfqq expired and remained idle until reactivation, |
|
* then the process associated with bfqq is treated as if, instead of |
|
* being greedy, it stopped issuing requests when bfqq remained idle, |
|
* and restarts issuing requests only on this reactivation. In other |
|
* words, the scheduler does not help the process recover the "service |
|
* hole" between bfqq expiration and reactivation. As a consequence, |
|
* the process receives a lower bandwidth than its reserved one. In |
|
* contrast, to recover this hole, the budget must be updated as if |
|
* bfqq was not expired at all before this reactivation, i.e., it must |
|
* be set to the value of the remaining budget when bfqq was |
|
* expired. Along the same line, timestamps need to be assigned the |
|
* value they had the last time bfqq was selected for service, i.e., |
|
* before last expiration. Thus timestamps need to be back-shifted |
|
* with respect to their normal computation (see [1] for more details |
|
* on this tricky aspect). |
|
* |
|
* Secondly, to allow the process to recover the hole, the in-service |
|
* queue must be expired too, to give bfqq the chance to preempt it |
|
* immediately. In fact, if bfqq has to wait for a full budget of the |
|
* in-service queue to be completed, then it may become impossible to |
|
* let the process recover the hole, even if the back-shifted |
|
* timestamps of bfqq are lower than those of the in-service queue. If |
|
* this happens for most or all of the holes, then the process may not |
|
* receive its reserved bandwidth. In this respect, it is worth noting |
|
* that, being the service of outstanding requests unpreemptible, a |
|
* little fraction of the holes may however be unrecoverable, thereby |
|
* causing a little loss of bandwidth. |
|
* |
|
* The last important point is detecting whether bfqq does need this |
|
* bandwidth recovery. In this respect, the next function deems the |
|
* process associated with bfqq greedy, and thus allows it to recover |
|
* the hole, if: 1) the process is waiting for the arrival of a new |
|
* request (which implies that bfqq expired for one of the above two |
|
* reasons), and 2) such a request has arrived soon. The first |
|
* condition is controlled through the flag non_blocking_wait_rq, |
|
* while the second through the flag arrived_in_time. If both |
|
* conditions hold, then the function computes the budget in the |
|
* above-described special way, and signals that the in-service queue |
|
* should be expired. Timestamp back-shifting is done later in |
|
* __bfq_activate_entity. |
|
* |
|
* 2. Reduce latency. Even if timestamps are not backshifted to let |
|
* the process associated with bfqq recover a service hole, bfqq may |
|
* however happen to have, after being (re)activated, a lower finish |
|
* timestamp than the in-service queue. That is, the next budget of |
|
* bfqq may have to be completed before the one of the in-service |
|
* queue. If this is the case, then preempting the in-service queue |
|
* allows this goal to be achieved, apart from the unpreemptible, |
|
* outstanding requests mentioned above. |
|
* |
|
* Unfortunately, regardless of which of the above two goals one wants |
|
* to achieve, service trees need first to be updated to know whether |
|
* the in-service queue must be preempted. To have service trees |
|
* correctly updated, the in-service queue must be expired and |
|
* rescheduled, and bfqq must be scheduled too. This is one of the |
|
* most costly operations (in future versions, the scheduling |
|
* mechanism may be re-designed in such a way to make it possible to |
|
* know whether preemption is needed without needing to update service |
|
* trees). In addition, queue preemptions almost always cause random |
|
* I/O, which may in turn cause loss of throughput. Finally, there may |
|
* even be no in-service queue when the next function is invoked (so, |
|
* no queue to compare timestamps with). Because of these facts, the |
|
* next function adopts the following simple scheme to avoid costly |
|
* operations, too frequent preemptions and too many dependencies on |
|
* the state of the scheduler: it requests the expiration of the |
|
* in-service queue (unconditionally) only for queues that need to |
|
* recover a hole. Then it delegates to other parts of the code the |
|
* responsibility of handling the above case 2. |
|
*/ |
|
static bool bfq_bfqq_update_budg_for_activation(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq, |
|
bool arrived_in_time) |
|
{ |
|
struct bfq_entity *entity = &bfqq->entity; |
|
|
|
/* |
|
* In the next compound condition, we check also whether there |
|
* is some budget left, because otherwise there is no point in |
|
* trying to go on serving bfqq with this same budget: bfqq |
|
* would be expired immediately after being selected for |
|
* service. This would only cause useless overhead. |
|
*/ |
|
if (bfq_bfqq_non_blocking_wait_rq(bfqq) && arrived_in_time && |
|
bfq_bfqq_budget_left(bfqq) > 0) { |
|
/* |
|
* We do not clear the flag non_blocking_wait_rq here, as |
|
* the latter is used in bfq_activate_bfqq to signal |
|
* that timestamps need to be back-shifted (and is |
|
* cleared right after). |
|
*/ |
|
|
|
/* |
|
* In next assignment we rely on that either |
|
* entity->service or entity->budget are not updated |
|
* on expiration if bfqq is empty (see |
|
* __bfq_bfqq_recalc_budget). Thus both quantities |
|
* remain unchanged after such an expiration, and the |
|
* following statement therefore assigns to |
|
* entity->budget the remaining budget on such an |
|
* expiration. |
|
*/ |
|
entity->budget = min_t(unsigned long, |
|
bfq_bfqq_budget_left(bfqq), |
|
bfqq->max_budget); |
|
|
|
/* |
|
* At this point, we have used entity->service to get |
|
* the budget left (needed for updating |
|
* entity->budget). Thus we finally can, and have to, |
|
* reset entity->service. The latter must be reset |
|
* because bfqq would otherwise be charged again for |
|
* the service it has received during its previous |
|
* service slot(s). |
|
*/ |
|
entity->service = 0; |
|
|
|
return true; |
|
} |
|
|
|
/* |
|
* We can finally complete expiration, by setting service to 0. |
|
*/ |
|
entity->service = 0; |
|
entity->budget = max_t(unsigned long, bfqq->max_budget, |
|
bfq_serv_to_charge(bfqq->next_rq, bfqq)); |
|
bfq_clear_bfqq_non_blocking_wait_rq(bfqq); |
|
return false; |
|
} |
|
|
|
/* |
|
* Return the farthest past time instant according to jiffies |
|
* macros. |
|
*/ |
|
static unsigned long bfq_smallest_from_now(void) |
|
{ |
|
return jiffies - MAX_JIFFY_OFFSET; |
|
} |
|
|
|
static void bfq_update_bfqq_wr_on_rq_arrival(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq, |
|
unsigned int old_wr_coeff, |
|
bool wr_or_deserves_wr, |
|
bool interactive, |
|
bool in_burst, |
|
bool soft_rt) |
|
{ |
|
if (old_wr_coeff == 1 && wr_or_deserves_wr) { |
|
/* start a weight-raising period */ |
|
if (interactive) { |
|
bfqq->service_from_wr = 0; |
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff; |
|
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd); |
|
} else { |
|
/* |
|
* No interactive weight raising in progress |
|
* here: assign minus infinity to |
|
* wr_start_at_switch_to_srt, to make sure |
|
* that, at the end of the soft-real-time |
|
* weight raising periods that is starting |
|
* now, no interactive weight-raising period |
|
* may be wrongly considered as still in |
|
* progress (and thus actually started by |
|
* mistake). |
|
*/ |
|
bfqq->wr_start_at_switch_to_srt = |
|
bfq_smallest_from_now(); |
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff * |
|
BFQ_SOFTRT_WEIGHT_FACTOR; |
|
bfqq->wr_cur_max_time = |
|
bfqd->bfq_wr_rt_max_time; |
|
} |
|
|
|
/* |
|
* If needed, further reduce budget to make sure it is |
|
* close to bfqq's backlog, so as to reduce the |
|
* scheduling-error component due to a too large |
|
* budget. Do not care about throughput consequences, |
|
* but only about latency. Finally, do not assign a |
|
* too small budget either, to avoid increasing |
|
* latency by causing too frequent expirations. |
|
*/ |
|
bfqq->entity.budget = min_t(unsigned long, |
|
bfqq->entity.budget, |
|
2 * bfq_min_budget(bfqd)); |
|
} else if (old_wr_coeff > 1) { |
|
if (interactive) { /* update wr coeff and duration */ |
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff; |
|
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd); |
|
} else if (in_burst) |
|
bfqq->wr_coeff = 1; |
|
else if (soft_rt) { |
|
/* |
|
* The application is now or still meeting the |
|
* requirements for being deemed soft rt. We |
|
* can then correctly and safely (re)charge |
|
* the weight-raising duration for the |
|
* application with the weight-raising |
|
* duration for soft rt applications. |
|
* |
|
* In particular, doing this recharge now, i.e., |
|
* before the weight-raising period for the |
|
* application finishes, reduces the probability |
|
* of the following negative scenario: |
|
* 1) the weight of a soft rt application is |
|
* raised at startup (as for any newly |
|
* created application), |
|
* 2) since the application is not interactive, |
|
* at a certain time weight-raising is |
|
* stopped for the application, |
|
* 3) at that time the application happens to |
|
* still have pending requests, and hence |
|
* is destined to not have a chance to be |
|
* deemed soft rt before these requests are |
|
* completed (see the comments to the |
|
* function bfq_bfqq_softrt_next_start() |
|
* for details on soft rt detection), |
|
* 4) these pending requests experience a high |
|
* latency because the application is not |
|
* weight-raised while they are pending. |
|
*/ |
|
if (bfqq->wr_cur_max_time != |
|
bfqd->bfq_wr_rt_max_time) { |
|
bfqq->wr_start_at_switch_to_srt = |
|
bfqq->last_wr_start_finish; |
|
|
|
bfqq->wr_cur_max_time = |
|
bfqd->bfq_wr_rt_max_time; |
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff * |
|
BFQ_SOFTRT_WEIGHT_FACTOR; |
|
} |
|
bfqq->last_wr_start_finish = jiffies; |
|
} |
|
} |
|
} |
|
|
|
static bool bfq_bfqq_idle_for_long_time(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
return bfqq->dispatched == 0 && |
|
time_is_before_jiffies( |
|
bfqq->budget_timeout + |
|
bfqd->bfq_wr_min_idle_time); |
|
} |
|
|
|
|
|
/* |
|
* Return true if bfqq is in a higher priority class, or has a higher |
|
* weight than the in-service queue. |
|
*/ |
|
static bool bfq_bfqq_higher_class_or_weight(struct bfq_queue *bfqq, |
|
struct bfq_queue *in_serv_bfqq) |
|
{ |
|
int bfqq_weight, in_serv_weight; |
|
|
|
if (bfqq->ioprio_class < in_serv_bfqq->ioprio_class) |
|
return true; |
|
|
|
if (in_serv_bfqq->entity.parent == bfqq->entity.parent) { |
|
bfqq_weight = bfqq->entity.weight; |
|
in_serv_weight = in_serv_bfqq->entity.weight; |
|
} else { |
|
if (bfqq->entity.parent) |
|
bfqq_weight = bfqq->entity.parent->weight; |
|
else |
|
bfqq_weight = bfqq->entity.weight; |
|
if (in_serv_bfqq->entity.parent) |
|
in_serv_weight = in_serv_bfqq->entity.parent->weight; |
|
else |
|
in_serv_weight = in_serv_bfqq->entity.weight; |
|
} |
|
|
|
return bfqq_weight > in_serv_weight; |
|
} |
|
|
|
static bool bfq_better_to_idle(struct bfq_queue *bfqq); |
|
|
|
static void bfq_bfqq_handle_idle_busy_switch(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq, |
|
int old_wr_coeff, |
|
struct request *rq, |
|
bool *interactive) |
|
{ |
|
bool soft_rt, in_burst, wr_or_deserves_wr, |
|
bfqq_wants_to_preempt, |
|
idle_for_long_time = bfq_bfqq_idle_for_long_time(bfqd, bfqq), |
|
/* |
|
* See the comments on |
|
* bfq_bfqq_update_budg_for_activation for |
|
* details on the usage of the next variable. |
|
*/ |
|
arrived_in_time = ktime_get_ns() <= |
|
bfqq->ttime.last_end_request + |
|
bfqd->bfq_slice_idle * 3; |
|
|
|
|
|
/* |
|
* bfqq deserves to be weight-raised if: |
|
* - it is sync, |
|
* - it does not belong to a large burst, |
|
* - it has been idle for enough time or is soft real-time, |
|
* - is linked to a bfq_io_cq (it is not shared in any sense), |
|
* - has a default weight (otherwise we assume the user wanted |
|
* to control its weight explicitly) |
|
*/ |
|
in_burst = bfq_bfqq_in_large_burst(bfqq); |
|
soft_rt = bfqd->bfq_wr_max_softrt_rate > 0 && |
|
!BFQQ_TOTALLY_SEEKY(bfqq) && |
|
!in_burst && |
|
time_is_before_jiffies(bfqq->soft_rt_next_start) && |
|
bfqq->dispatched == 0 && |
|
bfqq->entity.new_weight == 40; |
|
*interactive = !in_burst && idle_for_long_time && |
|
bfqq->entity.new_weight == 40; |
|
/* |
|
* Merged bfq_queues are kept out of weight-raising |
|
* (low-latency) mechanisms. The reason is that these queues |
|
* are usually created for non-interactive and |
|
* non-soft-real-time tasks. Yet this is not the case for |
|
* stably-merged queues. These queues are merged just because |
|
* they are created shortly after each other. So they may |
|
* easily serve the I/O of an interactive or soft-real time |
|
* application, if the application happens to spawn multiple |
|
* processes. So let also stably-merged queued enjoy weight |
|
* raising. |
|
*/ |
|
wr_or_deserves_wr = bfqd->low_latency && |
|
(bfqq->wr_coeff > 1 || |
|
(bfq_bfqq_sync(bfqq) && |
|
(bfqq->bic || RQ_BIC(rq)->stably_merged) && |
|
(*interactive || soft_rt))); |
|
|
|
/* |
|
* Using the last flag, update budget and check whether bfqq |
|
* may want to preempt the in-service queue. |
|
*/ |
|
bfqq_wants_to_preempt = |
|
bfq_bfqq_update_budg_for_activation(bfqd, bfqq, |
|
arrived_in_time); |
|
|
|
/* |
|
* If bfqq happened to be activated in a burst, but has been |
|
* idle for much more than an interactive queue, then we |
|
* assume that, in the overall I/O initiated in the burst, the |
|
* I/O associated with bfqq is finished. So bfqq does not need |
|
* to be treated as a queue belonging to a burst |
|
* anymore. Accordingly, we reset bfqq's in_large_burst flag |
|
* if set, and remove bfqq from the burst list if it's |
|
* there. We do not decrement burst_size, because the fact |
|
* that bfqq does not need to belong to the burst list any |
|
* more does not invalidate the fact that bfqq was created in |
|
* a burst. |
|
*/ |
|
if (likely(!bfq_bfqq_just_created(bfqq)) && |
|
idle_for_long_time && |
|
time_is_before_jiffies( |
|
bfqq->budget_timeout + |
|
msecs_to_jiffies(10000))) { |
|
hlist_del_init(&bfqq->burst_list_node); |
|
bfq_clear_bfqq_in_large_burst(bfqq); |
|
} |
|
|
|
bfq_clear_bfqq_just_created(bfqq); |
|
|
|
if (bfqd->low_latency) { |
|
if (unlikely(time_is_after_jiffies(bfqq->split_time))) |
|
/* wraparound */ |
|
bfqq->split_time = |
|
jiffies - bfqd->bfq_wr_min_idle_time - 1; |
|
|
|
if (time_is_before_jiffies(bfqq->split_time + |
|
bfqd->bfq_wr_min_idle_time)) { |
|
bfq_update_bfqq_wr_on_rq_arrival(bfqd, bfqq, |
|
old_wr_coeff, |
|
wr_or_deserves_wr, |
|
*interactive, |
|
in_burst, |
|
soft_rt); |
|
|
|
if (old_wr_coeff != bfqq->wr_coeff) |
|
bfqq->entity.prio_changed = 1; |
|
} |
|
} |
|
|
|
bfqq->last_idle_bklogged = jiffies; |
|
bfqq->service_from_backlogged = 0; |
|
bfq_clear_bfqq_softrt_update(bfqq); |
|
|
|
bfq_add_bfqq_busy(bfqd, bfqq); |
|
|
|
/* |
|
* Expire in-service queue if preemption may be needed for |
|
* guarantees or throughput. As for guarantees, we care |
|
* explicitly about two cases. The first is that bfqq has to |
|
* recover a service hole, as explained in the comments on |
|
* bfq_bfqq_update_budg_for_activation(), i.e., that |
|
* bfqq_wants_to_preempt is true. However, if bfqq does not |
|
* carry time-critical I/O, then bfqq's bandwidth is less |
|
* important than that of queues that carry time-critical I/O. |
|
* So, as a further constraint, we consider this case only if |
|
* bfqq is at least as weight-raised, i.e., at least as time |
|
* critical, as the in-service queue. |
|
* |
|
* The second case is that bfqq is in a higher priority class, |
|
* or has a higher weight than the in-service queue. If this |
|
* condition does not hold, we don't care because, even if |
|
* bfqq does not start to be served immediately, the resulting |
|
* delay for bfqq's I/O is however lower or much lower than |
|
* the ideal completion time to be guaranteed to bfqq's I/O. |
|
* |
|
* In both cases, preemption is needed only if, according to |
|
* the timestamps of both bfqq and of the in-service queue, |
|
* bfqq actually is the next queue to serve. So, to reduce |
|
* useless preemptions, the return value of |
|
* next_queue_may_preempt() is considered in the next compound |
|
* condition too. Yet next_queue_may_preempt() just checks a |
|
* simple, necessary condition for bfqq to be the next queue |
|
* to serve. In fact, to evaluate a sufficient condition, the |
|
* timestamps of the in-service queue would need to be |
|
* updated, and this operation is quite costly (see the |
|
* comments on bfq_bfqq_update_budg_for_activation()). |
|
* |
|
* As for throughput, we ask bfq_better_to_idle() whether we |
|
* still need to plug I/O dispatching. If bfq_better_to_idle() |
|
* says no, then plugging is not needed any longer, either to |
|
* boost throughput or to perserve service guarantees. Then |
|
* the best option is to stop plugging I/O, as not doing so |
|
* would certainly lower throughput. We may end up in this |
|
* case if: (1) upon a dispatch attempt, we detected that it |
|
* was better to plug I/O dispatch, and to wait for a new |
|
* request to arrive for the currently in-service queue, but |
|
* (2) this switch of bfqq to busy changes the scenario. |
|
*/ |
|
if (bfqd->in_service_queue && |
|
((bfqq_wants_to_preempt && |
|
bfqq->wr_coeff >= bfqd->in_service_queue->wr_coeff) || |
|
bfq_bfqq_higher_class_or_weight(bfqq, bfqd->in_service_queue) || |
|
!bfq_better_to_idle(bfqd->in_service_queue)) && |
|
next_queue_may_preempt(bfqd)) |
|
bfq_bfqq_expire(bfqd, bfqd->in_service_queue, |
|
false, BFQQE_PREEMPTED); |
|
} |
|
|
|
static void bfq_reset_inject_limit(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
/* invalidate baseline total service time */ |
|
bfqq->last_serv_time_ns = 0; |
|
|
|
/* |
|
* Reset pointer in case we are waiting for |
|
* some request completion. |
|
*/ |
|
bfqd->waited_rq = NULL; |
|
|
|
/* |
|
* If bfqq has a short think time, then start by setting the |
|
* inject limit to 0 prudentially, because the service time of |
|
* an injected I/O request may be higher than the think time |
|
* of bfqq, and therefore, if one request was injected when |
|
* bfqq remains empty, this injected request might delay the |
|
* service of the next I/O request for bfqq significantly. In |
|
* case bfqq can actually tolerate some injection, then the |
|
* adaptive update will however raise the limit soon. This |
|
* lucky circumstance holds exactly because bfqq has a short |
|
* think time, and thus, after remaining empty, is likely to |
|
* get new I/O enqueued---and then completed---before being |
|
* expired. This is the very pattern that gives the |
|
* limit-update algorithm the chance to measure the effect of |
|
* injection on request service times, and then to update the |
|
* limit accordingly. |
|
* |
|
* However, in the following special case, the inject limit is |
|
* left to 1 even if the think time is short: bfqq's I/O is |
|
* synchronized with that of some other queue, i.e., bfqq may |
|
* receive new I/O only after the I/O of the other queue is |
|
* completed. Keeping the inject limit to 1 allows the |
|
* blocking I/O to be served while bfqq is in service. And |
|
* this is very convenient both for bfqq and for overall |
|
* throughput, as explained in detail in the comments in |
|
* bfq_update_has_short_ttime(). |
|
* |
|
* On the opposite end, if bfqq has a long think time, then |
|
* start directly by 1, because: |
|
* a) on the bright side, keeping at most one request in |
|
* service in the drive is unlikely to cause any harm to the |
|
* latency of bfqq's requests, as the service time of a single |
|
* request is likely to be lower than the think time of bfqq; |
|
* b) on the downside, after becoming empty, bfqq is likely to |
|
* expire before getting its next request. With this request |
|
* arrival pattern, it is very hard to sample total service |
|
* times and update the inject limit accordingly (see comments |
|
* on bfq_update_inject_limit()). So the limit is likely to be |
|
* never, or at least seldom, updated. As a consequence, by |
|
* setting the limit to 1, we avoid that no injection ever |
|
* occurs with bfqq. On the downside, this proactive step |
|
* further reduces chances to actually compute the baseline |
|
* total service time. Thus it reduces chances to execute the |
|
* limit-update algorithm and possibly raise the limit to more |
|
* than 1. |
|
*/ |
|
if (bfq_bfqq_has_short_ttime(bfqq)) |
|
bfqq->inject_limit = 0; |
|
else |
|
bfqq->inject_limit = 1; |
|
|
|
bfqq->decrease_time_jif = jiffies; |
|
} |
|
|
|
static void bfq_update_io_intensity(struct bfq_queue *bfqq, u64 now_ns) |
|
{ |
|
u64 tot_io_time = now_ns - bfqq->io_start_time; |
|
|
|
if (RB_EMPTY_ROOT(&bfqq->sort_list) && bfqq->dispatched == 0) |
|
bfqq->tot_idle_time += |
|
now_ns - bfqq->ttime.last_end_request; |
|
|
|
if (unlikely(bfq_bfqq_just_created(bfqq))) |
|
return; |
|
|
|
/* |
|
* Must be busy for at least about 80% of the time to be |
|
* considered I/O bound. |
|
*/ |
|
if (bfqq->tot_idle_time * 5 > tot_io_time) |
|
bfq_clear_bfqq_IO_bound(bfqq); |
|
else |
|
bfq_mark_bfqq_IO_bound(bfqq); |
|
|
|
/* |
|
* Keep an observation window of at most 200 ms in the past |
|
* from now. |
|
*/ |
|
if (tot_io_time > 200 * NSEC_PER_MSEC) { |
|
bfqq->io_start_time = now_ns - (tot_io_time>>1); |
|
bfqq->tot_idle_time >>= 1; |
|
} |
|
} |
|
|
|
/* |
|
* Detect whether bfqq's I/O seems synchronized with that of some |
|
* other queue, i.e., whether bfqq, after remaining empty, happens to |
|
* receive new I/O only right after some I/O request of the other |
|
* queue has been completed. We call waker queue the other queue, and |
|
* we assume, for simplicity, that bfqq may have at most one waker |
|
* queue. |
|
* |
|
* A remarkable throughput boost can be reached by unconditionally |
|
* injecting the I/O of the waker queue, every time a new |
|
* bfq_dispatch_request happens to be invoked while I/O is being |
|
* plugged for bfqq. In addition to boosting throughput, this |
|
* unblocks bfqq's I/O, thereby improving bandwidth and latency for |
|
* bfqq. Note that these same results may be achieved with the general |
|
* injection mechanism, but less effectively. For details on this |
|
* aspect, see the comments on the choice of the queue for injection |
|
* in bfq_select_queue(). |
|
* |
|
* Turning back to the detection of a waker queue, a queue Q is deemed |
|
* as a waker queue for bfqq if, for three consecutive times, bfqq |
|
* happens to become non empty right after a request of Q has been |
|
* completed. In this respect, even if bfqq is empty, we do not check |
|
* for a waker if it still has some in-flight I/O. In fact, in this |
|
* case bfqq is actually still being served by the drive, and may |
|
* receive new I/O on the completion of some of the in-flight |
|
* requests. In particular, on the first time, Q is tentatively set as |
|
* a candidate waker queue, while on the third consecutive time that Q |
|
* is detected, the field waker_bfqq is set to Q, to confirm that Q is |
|
* a waker queue for bfqq. These detection steps are performed only if |
|
* bfqq has a long think time, so as to make it more likely that |
|
* bfqq's I/O is actually being blocked by a synchronization. This |
|
* last filter, plus the above three-times requirement, make false |
|
* positives less likely. |
|
* |
|
* NOTE |
|
* |
|
* The sooner a waker queue is detected, the sooner throughput can be |
|
* boosted by injecting I/O from the waker queue. Fortunately, |
|
* detection is likely to be actually fast, for the following |
|
* reasons. While blocked by synchronization, bfqq has a long think |
|
* time. This implies that bfqq's inject limit is at least equal to 1 |
|
* (see the comments in bfq_update_inject_limit()). So, thanks to |
|
* injection, the waker queue is likely to be served during the very |
|
* first I/O-plugging time interval for bfqq. This triggers the first |
|
* step of the detection mechanism. Thanks again to injection, the |
|
* candidate waker queue is then likely to be confirmed no later than |
|
* during the next I/O-plugging interval for bfqq. |
|
* |
|
* ISSUE |
|
* |
|
* On queue merging all waker information is lost. |
|
*/ |
|
static void bfq_check_waker(struct bfq_data *bfqd, struct bfq_queue *bfqq, |
|
u64 now_ns) |
|
{ |
|
if (!bfqd->last_completed_rq_bfqq || |
|
bfqd->last_completed_rq_bfqq == bfqq || |
|
bfq_bfqq_has_short_ttime(bfqq) || |
|
bfqq->dispatched > 0 || |
|
now_ns - bfqd->last_completion >= 4 * NSEC_PER_MSEC || |
|
bfqd->last_completed_rq_bfqq == bfqq->waker_bfqq) |
|
return; |
|
|
|
if (bfqd->last_completed_rq_bfqq != |
|
bfqq->tentative_waker_bfqq) { |
|
/* |
|
* First synchronization detected with a |
|
* candidate waker queue, or with a different |
|
* candidate waker queue from the current one. |
|
*/ |
|
bfqq->tentative_waker_bfqq = |
|
bfqd->last_completed_rq_bfqq; |
|
bfqq->num_waker_detections = 1; |
|
} else /* Same tentative waker queue detected again */ |
|
bfqq->num_waker_detections++; |
|
|
|
if (bfqq->num_waker_detections == 3) { |
|
bfqq->waker_bfqq = bfqd->last_completed_rq_bfqq; |
|
bfqq->tentative_waker_bfqq = NULL; |
|
|
|
/* |
|
* If the waker queue disappears, then |
|
* bfqq->waker_bfqq must be reset. To |
|
* this goal, we maintain in each |
|
* waker queue a list, woken_list, of |
|
* all the queues that reference the |
|
* waker queue through their |
|
* waker_bfqq pointer. When the waker |
|
* queue exits, the waker_bfqq pointer |
|
* of all the queues in the woken_list |
|
* is reset. |
|
* |
|
* In addition, if bfqq is already in |
|
* the woken_list of a waker queue, |
|
* then, before being inserted into |
|
* the woken_list of a new waker |
|
* queue, bfqq must be removed from |
|
* the woken_list of the old waker |
|
* queue. |
|
*/ |
|
if (!hlist_unhashed(&bfqq->woken_list_node)) |
|
hlist_del_init(&bfqq->woken_list_node); |
|
hlist_add_head(&bfqq->woken_list_node, |
|
&bfqd->last_completed_rq_bfqq->woken_list); |
|
} |
|
} |
|
|
|
static void bfq_add_request(struct request *rq) |
|
{ |
|
struct bfq_queue *bfqq = RQ_BFQQ(rq); |
|
struct bfq_data *bfqd = bfqq->bfqd; |
|
struct request *next_rq, *prev; |
|
unsigned int old_wr_coeff = bfqq->wr_coeff; |
|
bool interactive = false; |
|
u64 now_ns = ktime_get_ns(); |
|
|
|
bfq_log_bfqq(bfqd, bfqq, "add_request %d", rq_is_sync(rq)); |
|
bfqq->queued[rq_is_sync(rq)]++; |
|
bfqd->queued++; |
|
|
|
if (RB_EMPTY_ROOT(&bfqq->sort_list) && bfq_bfqq_sync(bfqq)) { |
|
bfq_check_waker(bfqd, bfqq, now_ns); |
|
|
|
/* |
|
* Periodically reset inject limit, to make sure that |
|
* the latter eventually drops in case workload |
|
* changes, see step (3) in the comments on |
|
* bfq_update_inject_limit(). |
|
*/ |
|
if (time_is_before_eq_jiffies(bfqq->decrease_time_jif + |
|
msecs_to_jiffies(1000))) |
|
bfq_reset_inject_limit(bfqd, bfqq); |
|
|
|
/* |
|
* The following conditions must hold to setup a new |
|
* sampling of total service time, and then a new |
|
* update of the inject limit: |
|
* - bfqq is in service, because the total service |
|
* time is evaluated only for the I/O requests of |
|
* the queues in service; |
|
* - this is the right occasion to compute or to |
|
* lower the baseline total service time, because |
|
* there are actually no requests in the drive, |
|
* or |
|
* the baseline total service time is available, and |
|
* this is the right occasion to compute the other |
|
* quantity needed to update the inject limit, i.e., |
|
* the total service time caused by the amount of |
|
* injection allowed by the current value of the |
|
* limit. It is the right occasion because injection |
|
* has actually been performed during the service |
|
* hole, and there are still in-flight requests, |
|
* which are very likely to be exactly the injected |
|
* requests, or part of them; |
|
* - the minimum interval for sampling the total |
|
* service time and updating the inject limit has |
|
* elapsed. |
|
*/ |
|
if (bfqq == bfqd->in_service_queue && |
|
(bfqd->rq_in_driver == 0 || |
|
(bfqq->last_serv_time_ns > 0 && |
|
bfqd->rqs_injected && bfqd->rq_in_driver > 0)) && |
|
time_is_before_eq_jiffies(bfqq->decrease_time_jif + |
|
msecs_to_jiffies(10))) { |
|
bfqd->last_empty_occupied_ns = ktime_get_ns(); |
|
/* |
|
* Start the state machine for measuring the |
|
* total service time of rq: setting |
|
* wait_dispatch will cause bfqd->waited_rq to |
|
* be set when rq will be dispatched. |
|
*/ |
|
bfqd->wait_dispatch = true; |
|
/* |
|
* If there is no I/O in service in the drive, |
|
* then possible injection occurred before the |
|
* arrival of rq will not affect the total |
|
* service time of rq. So the injection limit |
|
* must not be updated as a function of such |
|
* total service time, unless new injection |
|
* occurs before rq is completed. To have the |
|
* injection limit updated only in the latter |
|
* case, reset rqs_injected here (rqs_injected |
|
* will be set in case injection is performed |
|
* on bfqq before rq is completed). |
|
*/ |
|
if (bfqd->rq_in_driver == 0) |
|
bfqd->rqs_injected = false; |
|
} |
|
} |
|
|
|
if (bfq_bfqq_sync(bfqq)) |
|
bfq_update_io_intensity(bfqq, now_ns); |
|
|
|
elv_rb_add(&bfqq->sort_list, rq); |
|
|
|
/* |
|
* Check if this request is a better next-serve candidate. |
|
*/ |
|
prev = bfqq->next_rq; |
|
next_rq = bfq_choose_req(bfqd, bfqq->next_rq, rq, bfqd->last_position); |
|
bfqq->next_rq = next_rq; |
|
|
|
/* |
|
* Adjust priority tree position, if next_rq changes. |
|
* See comments on bfq_pos_tree_add_move() for the unlikely(). |
|
*/ |
|
if (unlikely(!bfqd->nonrot_with_queueing && prev != bfqq->next_rq)) |
|
bfq_pos_tree_add_move(bfqd, bfqq); |
|
|
|
if (!bfq_bfqq_busy(bfqq)) /* switching to busy ... */ |
|
bfq_bfqq_handle_idle_busy_switch(bfqd, bfqq, old_wr_coeff, |
|
rq, &interactive); |
|
else { |
|
if (bfqd->low_latency && old_wr_coeff == 1 && !rq_is_sync(rq) && |
|
time_is_before_jiffies( |
|
bfqq->last_wr_start_finish + |
|
bfqd->bfq_wr_min_inter_arr_async)) { |
|
bfqq->wr_coeff = bfqd->bfq_wr_coeff; |
|
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd); |
|
|
|
bfqd->wr_busy_queues++; |
|
bfqq->entity.prio_changed = 1; |
|
} |
|
if (prev != bfqq->next_rq) |
|
bfq_updated_next_req(bfqd, bfqq); |
|
} |
|
|
|
/* |
|
* Assign jiffies to last_wr_start_finish in the following |
|
* cases: |
|
* |
|
* . if bfqq is not going to be weight-raised, because, for |
|
* non weight-raised queues, last_wr_start_finish stores the |
|
* arrival time of the last request; as of now, this piece |
|
* of information is used only for deciding whether to |
|
* weight-raise async queues |
|
* |
|
* . if bfqq is not weight-raised, because, if bfqq is now |
|
* switching to weight-raised, then last_wr_start_finish |
|
* stores the time when weight-raising starts |
|
* |
|
* . if bfqq is interactive, because, regardless of whether |
|
* bfqq is currently weight-raised, the weight-raising |
|
* period must start or restart (this case is considered |
|
* separately because it is not detected by the above |
|
* conditions, if bfqq is already weight-raised) |
|
* |
|
* last_wr_start_finish has to be updated also if bfqq is soft |
|
* real-time, because the weight-raising period is constantly |
|
* restarted on idle-to-busy transitions for these queues, but |
|
* this is already done in bfq_bfqq_handle_idle_busy_switch if |
|
* needed. |
|
*/ |
|
if (bfqd->low_latency && |
|
(old_wr_coeff == 1 || bfqq->wr_coeff == 1 || interactive)) |
|
bfqq->last_wr_start_finish = jiffies; |
|
} |
|
|
|
static struct request *bfq_find_rq_fmerge(struct bfq_data *bfqd, |
|
struct bio *bio, |
|
struct request_queue *q) |
|
{ |
|
struct bfq_queue *bfqq = bfqd->bio_bfqq; |
|
|
|
|
|
if (bfqq) |
|
return elv_rb_find(&bfqq->sort_list, bio_end_sector(bio)); |
|
|
|
return NULL; |
|
} |
|
|
|
static sector_t get_sdist(sector_t last_pos, struct request *rq) |
|
{ |
|
if (last_pos) |
|
return abs(blk_rq_pos(rq) - last_pos); |
|
|
|
return 0; |
|
} |
|
|
|
#if 0 /* Still not clear if we can do without next two functions */ |
|
static void bfq_activate_request(struct request_queue *q, struct request *rq) |
|
{ |
|
struct bfq_data *bfqd = q->elevator->elevator_data; |
|
|
|
bfqd->rq_in_driver++; |
|
} |
|
|
|
static void bfq_deactivate_request(struct request_queue *q, struct request *rq) |
|
{ |
|
struct bfq_data *bfqd = q->elevator->elevator_data; |
|
|
|
bfqd->rq_in_driver--; |
|
} |
|
#endif |
|
|
|
static void bfq_remove_request(struct request_queue *q, |
|
struct request *rq) |
|
{ |
|
struct bfq_queue *bfqq = RQ_BFQQ(rq); |
|
struct bfq_data *bfqd = bfqq->bfqd; |
|
const int sync = rq_is_sync(rq); |
|
|
|
if (bfqq->next_rq == rq) { |
|
bfqq->next_rq = bfq_find_next_rq(bfqd, bfqq, rq); |
|
bfq_updated_next_req(bfqd, bfqq); |
|
} |
|
|
|
if (rq->queuelist.prev != &rq->queuelist) |
|
list_del_init(&rq->queuelist); |
|
bfqq->queued[sync]--; |
|
bfqd->queued--; |
|
elv_rb_del(&bfqq->sort_list, rq); |
|
|
|
elv_rqhash_del(q, rq); |
|
if (q->last_merge == rq) |
|
q->last_merge = NULL; |
|
|
|
if (RB_EMPTY_ROOT(&bfqq->sort_list)) { |
|
bfqq->next_rq = NULL; |
|
|
|
if (bfq_bfqq_busy(bfqq) && bfqq != bfqd->in_service_queue) { |
|
bfq_del_bfqq_busy(bfqd, bfqq, false); |
|
/* |
|
* bfqq emptied. In normal operation, when |
|
* bfqq is empty, bfqq->entity.service and |
|
* bfqq->entity.budget must contain, |
|
* respectively, the service received and the |
|
* budget used last time bfqq emptied. These |
|
* facts do not hold in this case, as at least |
|
* this last removal occurred while bfqq is |
|
* not in service. To avoid inconsistencies, |
|
* reset both bfqq->entity.service and |
|
* bfqq->entity.budget, if bfqq has still a |
|
* process that may issue I/O requests to it. |
|
*/ |
|
bfqq->entity.budget = bfqq->entity.service = 0; |
|
} |
|
|
|
/* |
|
* Remove queue from request-position tree as it is empty. |
|
*/ |
|
if (bfqq->pos_root) { |
|
rb_erase(&bfqq->pos_node, bfqq->pos_root); |
|
bfqq->pos_root = NULL; |
|
} |
|
} else { |
|
/* see comments on bfq_pos_tree_add_move() for the unlikely() */ |
|
if (unlikely(!bfqd->nonrot_with_queueing)) |
|
bfq_pos_tree_add_move(bfqd, bfqq); |
|
} |
|
|
|
if (rq->cmd_flags & REQ_META) |
|
bfqq->meta_pending--; |
|
|
|
} |
|
|
|
static bool bfq_bio_merge(struct request_queue *q, struct bio *bio, |
|
unsigned int nr_segs) |
|
{ |
|
struct bfq_data *bfqd = q->elevator->elevator_data; |
|
struct request *free = NULL; |
|
/* |
|
* bfq_bic_lookup grabs the queue_lock: invoke it now and |
|
* store its return value for later use, to avoid nesting |
|
* queue_lock inside the bfqd->lock. We assume that the bic |
|
* returned by bfq_bic_lookup does not go away before |
|
* bfqd->lock is taken. |
|
*/ |
|
struct bfq_io_cq *bic = bfq_bic_lookup(bfqd, current->io_context, q); |
|
bool ret; |
|
|
|
spin_lock_irq(&bfqd->lock); |
|
|
|
if (bic) |
|
bfqd->bio_bfqq = bic_to_bfqq(bic, op_is_sync(bio->bi_opf)); |
|
else |
|
bfqd->bio_bfqq = NULL; |
|
bfqd->bio_bic = bic; |
|
|
|
ret = blk_mq_sched_try_merge(q, bio, nr_segs, &free); |
|
|
|
spin_unlock_irq(&bfqd->lock); |
|
if (free) |
|
blk_mq_free_request(free); |
|
|
|
return ret; |
|
} |
|
|
|
static int bfq_request_merge(struct request_queue *q, struct request **req, |
|
struct bio *bio) |
|
{ |
|
struct bfq_data *bfqd = q->elevator->elevator_data; |
|
struct request *__rq; |
|
|
|
__rq = bfq_find_rq_fmerge(bfqd, bio, q); |
|
if (__rq && elv_bio_merge_ok(__rq, bio)) { |
|
*req = __rq; |
|
|
|
if (blk_discard_mergable(__rq)) |
|
return ELEVATOR_DISCARD_MERGE; |
|
return ELEVATOR_FRONT_MERGE; |
|
} |
|
|
|
return ELEVATOR_NO_MERGE; |
|
} |
|
|
|
static struct bfq_queue *bfq_init_rq(struct request *rq); |
|
|
|
static void bfq_request_merged(struct request_queue *q, struct request *req, |
|
enum elv_merge type) |
|
{ |
|
if (type == ELEVATOR_FRONT_MERGE && |
|
rb_prev(&req->rb_node) && |
|
blk_rq_pos(req) < |
|
blk_rq_pos(container_of(rb_prev(&req->rb_node), |
|
struct request, rb_node))) { |
|
struct bfq_queue *bfqq = bfq_init_rq(req); |
|
struct bfq_data *bfqd; |
|
struct request *prev, *next_rq; |
|
|
|
if (!bfqq) |
|
return; |
|
|
|
bfqd = bfqq->bfqd; |
|
|
|
/* Reposition request in its sort_list */ |
|
elv_rb_del(&bfqq->sort_list, req); |
|
elv_rb_add(&bfqq->sort_list, req); |
|
|
|
/* Choose next request to be served for bfqq */ |
|
prev = bfqq->next_rq; |
|
next_rq = bfq_choose_req(bfqd, bfqq->next_rq, req, |
|
bfqd->last_position); |
|
bfqq->next_rq = next_rq; |
|
/* |
|
* If next_rq changes, update both the queue's budget to |
|
* fit the new request and the queue's position in its |
|
* rq_pos_tree. |
|
*/ |
|
if (prev != bfqq->next_rq) { |
|
bfq_updated_next_req(bfqd, bfqq); |
|
/* |
|
* See comments on bfq_pos_tree_add_move() for |
|
* the unlikely(). |
|
*/ |
|
if (unlikely(!bfqd->nonrot_with_queueing)) |
|
bfq_pos_tree_add_move(bfqd, bfqq); |
|
} |
|
} |
|
} |
|
|
|
/* |
|
* This function is called to notify the scheduler that the requests |
|
* rq and 'next' have been merged, with 'next' going away. BFQ |
|
* exploits this hook to address the following issue: if 'next' has a |
|
* fifo_time lower that rq, then the fifo_time of rq must be set to |
|
* the value of 'next', to not forget the greater age of 'next'. |
|
* |
|
* NOTE: in this function we assume that rq is in a bfq_queue, basing |
|
* on that rq is picked from the hash table q->elevator->hash, which, |
|
* in its turn, is filled only with I/O requests present in |
|
* bfq_queues, while BFQ is in use for the request queue q. In fact, |
|
* the function that fills this hash table (elv_rqhash_add) is called |
|
* only by bfq_insert_request. |
|
*/ |
|
static void bfq_requests_merged(struct request_queue *q, struct request *rq, |
|
struct request *next) |
|
{ |
|
struct bfq_queue *bfqq = bfq_init_rq(rq), |
|
*next_bfqq = bfq_init_rq(next); |
|
|
|
if (!bfqq) |
|
goto remove; |
|
|
|
/* |
|
* If next and rq belong to the same bfq_queue and next is older |
|
* than rq, then reposition rq in the fifo (by substituting next |
|
* with rq). Otherwise, if next and rq belong to different |
|
* bfq_queues, never reposition rq: in fact, we would have to |
|
* reposition it with respect to next's position in its own fifo, |
|
* which would most certainly be too expensive with respect to |
|
* the benefits. |
|
*/ |
|
if (bfqq == next_bfqq && |
|
!list_empty(&rq->queuelist) && !list_empty(&next->queuelist) && |
|
next->fifo_time < rq->fifo_time) { |
|
list_del_init(&rq->queuelist); |
|
list_replace_init(&next->queuelist, &rq->queuelist); |
|
rq->fifo_time = next->fifo_time; |
|
} |
|
|
|
if (bfqq->next_rq == next) |
|
bfqq->next_rq = rq; |
|
|
|
bfqg_stats_update_io_merged(bfqq_group(bfqq), next->cmd_flags); |
|
remove: |
|
/* Merged request may be in the IO scheduler. Remove it. */ |
|
if (!RB_EMPTY_NODE(&next->rb_node)) { |
|
bfq_remove_request(next->q, next); |
|
if (next_bfqq) |
|
bfqg_stats_update_io_remove(bfqq_group(next_bfqq), |
|
next->cmd_flags); |
|
} |
|
} |
|
|
|
/* Must be called with bfqq != NULL */ |
|
static void bfq_bfqq_end_wr(struct bfq_queue *bfqq) |
|
{ |
|
/* |
|
* If bfqq has been enjoying interactive weight-raising, then |
|
* reset soft_rt_next_start. We do it for the following |
|
* reason. bfqq may have been conveying the I/O needed to load |
|
* a soft real-time application. Such an application actually |
|
* exhibits a soft real-time I/O pattern after it finishes |
|
* loading, and finally starts doing its job. But, if bfqq has |
|
* been receiving a lot of bandwidth so far (likely to happen |
|
* on a fast device), then soft_rt_next_start now contains a |
|
* high value that. So, without this reset, bfqq would be |
|
* prevented from being possibly considered as soft_rt for a |
|
* very long time. |
|
*/ |
|
|
|
if (bfqq->wr_cur_max_time != |
|
bfqq->bfqd->bfq_wr_rt_max_time) |
|
bfqq->soft_rt_next_start = jiffies; |
|
|
|
if (bfq_bfqq_busy(bfqq)) |
|
bfqq->bfqd->wr_busy_queues--; |
|
bfqq->wr_coeff = 1; |
|
bfqq->wr_cur_max_time = 0; |
|
bfqq->last_wr_start_finish = jiffies; |
|
/* |
|
* Trigger a weight change on the next invocation of |
|
* __bfq_entity_update_weight_prio. |
|
*/ |
|
bfqq->entity.prio_changed = 1; |
|
} |
|
|
|
void bfq_end_wr_async_queues(struct bfq_data *bfqd, |
|
struct bfq_group *bfqg) |
|
{ |
|
int i, j; |
|
|
|
for (i = 0; i < 2; i++) |
|
for (j = 0; j < IOPRIO_NR_LEVELS; j++) |
|
if (bfqg->async_bfqq[i][j]) |
|
bfq_bfqq_end_wr(bfqg->async_bfqq[i][j]); |
|
if (bfqg->async_idle_bfqq) |
|
bfq_bfqq_end_wr(bfqg->async_idle_bfqq); |
|
} |
|
|
|
static void bfq_end_wr(struct bfq_data *bfqd) |
|
{ |
|
struct bfq_queue *bfqq; |
|
|
|
spin_lock_irq(&bfqd->lock); |
|
|
|
list_for_each_entry(bfqq, &bfqd->active_list, bfqq_list) |
|
bfq_bfqq_end_wr(bfqq); |
|
list_for_each_entry(bfqq, &bfqd->idle_list, bfqq_list) |
|
bfq_bfqq_end_wr(bfqq); |
|
bfq_end_wr_async(bfqd); |
|
|
|
spin_unlock_irq(&bfqd->lock); |
|
} |
|
|
|
static sector_t bfq_io_struct_pos(void *io_struct, bool request) |
|
{ |
|
if (request) |
|
return blk_rq_pos(io_struct); |
|
else |
|
return ((struct bio *)io_struct)->bi_iter.bi_sector; |
|
} |
|
|
|
static int bfq_rq_close_to_sector(void *io_struct, bool request, |
|
sector_t sector) |
|
{ |
|
return abs(bfq_io_struct_pos(io_struct, request) - sector) <= |
|
BFQQ_CLOSE_THR; |
|
} |
|
|
|
static struct bfq_queue *bfqq_find_close(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq, |
|
sector_t sector) |
|
{ |
|
struct rb_root *root = &bfq_bfqq_to_bfqg(bfqq)->rq_pos_tree; |
|
struct rb_node *parent, *node; |
|
struct bfq_queue *__bfqq; |
|
|
|
if (RB_EMPTY_ROOT(root)) |
|
return NULL; |
|
|
|
/* |
|
* First, if we find a request starting at the end of the last |
|
* request, choose it. |
|
*/ |
|
__bfqq = bfq_rq_pos_tree_lookup(bfqd, root, sector, &parent, NULL); |
|
if (__bfqq) |
|
return __bfqq; |
|
|
|
/* |
|
* If the exact sector wasn't found, the parent of the NULL leaf |
|
* will contain the closest sector (rq_pos_tree sorted by |
|
* next_request position). |
|
*/ |
|
__bfqq = rb_entry(parent, struct bfq_queue, pos_node); |
|
if (bfq_rq_close_to_sector(__bfqq->next_rq, true, sector)) |
|
return __bfqq; |
|
|
|
if (blk_rq_pos(__bfqq->next_rq) < sector) |
|
node = rb_next(&__bfqq->pos_node); |
|
else |
|
node = rb_prev(&__bfqq->pos_node); |
|
if (!node) |
|
return NULL; |
|
|
|
__bfqq = rb_entry(node, struct bfq_queue, pos_node); |
|
if (bfq_rq_close_to_sector(__bfqq->next_rq, true, sector)) |
|
return __bfqq; |
|
|
|
return NULL; |
|
} |
|
|
|
static struct bfq_queue *bfq_find_close_cooperator(struct bfq_data *bfqd, |
|
struct bfq_queue *cur_bfqq, |
|
sector_t sector) |
|
{ |
|
struct bfq_queue *bfqq; |
|
|
|
/* |
|
* We shall notice if some of the queues are cooperating, |
|
* e.g., working closely on the same area of the device. In |
|
* that case, we can group them together and: 1) don't waste |
|
* time idling, and 2) serve the union of their requests in |
|
* the best possible order for throughput. |
|
*/ |
|
bfqq = bfqq_find_close(bfqd, cur_bfqq, sector); |
|
if (!bfqq || bfqq == cur_bfqq) |
|
return NULL; |
|
|
|
return bfqq; |
|
} |
|
|
|
static struct bfq_queue * |
|
bfq_setup_merge(struct bfq_queue *bfqq, struct bfq_queue *new_bfqq) |
|
{ |
|
int process_refs, new_process_refs; |
|
struct bfq_queue *__bfqq; |
|
|
|
/* |
|
* If there are no process references on the new_bfqq, then it is |
|
* unsafe to follow the ->new_bfqq chain as other bfqq's in the chain |
|
* may have dropped their last reference (not just their last process |
|
* reference). |
|
*/ |
|
if (!bfqq_process_refs(new_bfqq)) |
|
return NULL; |
|
|
|
/* Avoid a circular list and skip interim queue merges. */ |
|
while ((__bfqq = new_bfqq->new_bfqq)) { |
|
if (__bfqq == bfqq) |
|
return NULL; |
|
new_bfqq = __bfqq; |
|
} |
|
|
|
process_refs = bfqq_process_refs(bfqq); |
|
new_process_refs = bfqq_process_refs(new_bfqq); |
|
/* |
|
* If the process for the bfqq has gone away, there is no |
|
* sense in merging the queues. |
|
*/ |
|
if (process_refs == 0 || new_process_refs == 0) |
|
return NULL; |
|
|
|
bfq_log_bfqq(bfqq->bfqd, bfqq, "scheduling merge with queue %d", |
|
new_bfqq->pid); |
|
|
|
/* |
|
* Merging is just a redirection: the requests of the process |
|
* owning one of the two queues are redirected to the other queue. |
|
* The latter queue, in its turn, is set as shared if this is the |
|
* first time that the requests of some process are redirected to |
|
* it. |
|
* |
|
* We redirect bfqq to new_bfqq and not the opposite, because |
|
* we are in the context of the process owning bfqq, thus we |
|
* have the io_cq of this process. So we can immediately |
|
* configure this io_cq to redirect the requests of the |
|
* process to new_bfqq. In contrast, the io_cq of new_bfqq is |
|
* not available any more (new_bfqq->bic == NULL). |
|
* |
|
* Anyway, even in case new_bfqq coincides with the in-service |
|
* queue, redirecting requests the in-service queue is the |
|
* best option, as we feed the in-service queue with new |
|
* requests close to the last request served and, by doing so, |
|
* are likely to increase the throughput. |
|
*/ |
|
bfqq->new_bfqq = new_bfqq; |
|
/* |
|
* The above assignment schedules the following redirections: |
|
* each time some I/O for bfqq arrives, the process that |
|
* generated that I/O is disassociated from bfqq and |
|
* associated with new_bfqq. Here we increases new_bfqq->ref |
|
* in advance, adding the number of processes that are |
|
* expected to be associated with new_bfqq as they happen to |
|
* issue I/O. |
|
*/ |
|
new_bfqq->ref += process_refs; |
|
return new_bfqq; |
|
} |
|
|
|
static bool bfq_may_be_close_cooperator(struct bfq_queue *bfqq, |
|
struct bfq_queue *new_bfqq) |
|
{ |
|
if (bfq_too_late_for_merging(new_bfqq)) |
|
return false; |
|
|
|
if (bfq_class_idle(bfqq) || bfq_class_idle(new_bfqq) || |
|
(bfqq->ioprio_class != new_bfqq->ioprio_class)) |
|
return false; |
|
|
|
/* |
|
* If either of the queues has already been detected as seeky, |
|
* then merging it with the other queue is unlikely to lead to |
|
* sequential I/O. |
|
*/ |
|
if (BFQQ_SEEKY(bfqq) || BFQQ_SEEKY(new_bfqq)) |
|
return false; |
|
|
|
/* |
|
* Interleaved I/O is known to be done by (some) applications |
|
* only for reads, so it does not make sense to merge async |
|
* queues. |
|
*/ |
|
if (!bfq_bfqq_sync(bfqq) || !bfq_bfqq_sync(new_bfqq)) |
|
return false; |
|
|
|
return true; |
|
} |
|
|
|
static bool idling_boosts_thr_without_issues(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq); |
|
|
|
/* |
|
* Attempt to schedule a merge of bfqq with the currently in-service |
|
* queue or with a close queue among the scheduled queues. Return |
|
* NULL if no merge was scheduled, a pointer to the shared bfq_queue |
|
* structure otherwise. |
|
* |
|
* The OOM queue is not allowed to participate to cooperation: in fact, since |
|
* the requests temporarily redirected to the OOM queue could be redirected |
|
* again to dedicated queues at any time, the state needed to correctly |
|
* handle merging with the OOM queue would be quite complex and expensive |
|
* to maintain. Besides, in such a critical condition as an out of memory, |
|
* the benefits of queue merging may be little relevant, or even negligible. |
|
* |
|
* WARNING: queue merging may impair fairness among non-weight raised |
|
* queues, for at least two reasons: 1) the original weight of a |
|
* merged queue may change during the merged state, 2) even being the |
|
* weight the same, a merged queue may be bloated with many more |
|
* requests than the ones produced by its originally-associated |
|
* process. |
|
*/ |
|
static struct bfq_queue * |
|
bfq_setup_cooperator(struct bfq_data *bfqd, struct bfq_queue *bfqq, |
|
void *io_struct, bool request, struct bfq_io_cq *bic) |
|
{ |
|
struct bfq_queue *in_service_bfqq, *new_bfqq; |
|
|
|
/* if a merge has already been setup, then proceed with that first */ |
|
if (bfqq->new_bfqq) |
|
return bfqq->new_bfqq; |
|
|
|
/* |
|
* Check delayed stable merge for rotational or non-queueing |
|
* devs. For this branch to be executed, bfqq must not be |
|
* currently merged with some other queue (i.e., bfqq->bic |
|
* must be non null). If we considered also merged queues, |
|
* then we should also check whether bfqq has already been |
|
* merged with bic->stable_merge_bfqq. But this would be |
|
* costly and complicated. |
|
*/ |
|
if (unlikely(!bfqd->nonrot_with_queueing)) { |
|
/* |
|
* Make sure also that bfqq is sync, because |
|
* bic->stable_merge_bfqq may point to some queue (for |
|
* stable merging) also if bic is associated with a |
|
* sync queue, but this bfqq is async |
|
*/ |
|
if (bfq_bfqq_sync(bfqq) && bic->stable_merge_bfqq && |
|
!bfq_bfqq_just_created(bfqq) && |
|
time_is_before_jiffies(bfqq->split_time + |
|
msecs_to_jiffies(bfq_late_stable_merging)) && |
|
time_is_before_jiffies(bfqq->creation_time + |
|
msecs_to_jiffies(bfq_late_stable_merging))) { |
|
struct bfq_queue *stable_merge_bfqq = |
|
bic->stable_merge_bfqq; |
|
int proc_ref = min(bfqq_process_refs(bfqq), |
|
bfqq_process_refs(stable_merge_bfqq)); |
|
|
|
/* deschedule stable merge, because done or aborted here */ |
|
bfq_put_stable_ref(stable_merge_bfqq); |
|
|
|
bic->stable_merge_bfqq = NULL; |
|
|
|
if (!idling_boosts_thr_without_issues(bfqd, bfqq) && |
|
proc_ref > 0) { |
|
/* next function will take at least one ref */ |
|
struct bfq_queue *new_bfqq = |
|
bfq_setup_merge(bfqq, stable_merge_bfqq); |
|
|
|
bic->stably_merged = true; |
|
if (new_bfqq && new_bfqq->bic) |
|
new_bfqq->bic->stably_merged = true; |
|
return new_bfqq; |
|
} else |
|
return NULL; |
|
} |
|
} |
|
|
|
/* |
|
* Do not perform queue merging if the device is non |
|
* rotational and performs internal queueing. In fact, such a |
|
* device reaches a high speed through internal parallelism |
|
* and pipelining. This means that, to reach a high |
|
* throughput, it must have many requests enqueued at the same |
|
* time. But, in this configuration, the internal scheduling |
|
* algorithm of the device does exactly the job of queue |
|
* merging: it reorders requests so as to obtain as much as |
|
* possible a sequential I/O pattern. As a consequence, with |
|
* the workload generated by processes doing interleaved I/O, |
|
* the throughput reached by the device is likely to be the |
|
* same, with and without queue merging. |
|
* |
|
* Disabling merging also provides a remarkable benefit in |
|
* terms of throughput. Merging tends to make many workloads |
|
* artificially more uneven, because of shared queues |
|
* remaining non empty for incomparably more time than |
|
* non-merged queues. This may accentuate workload |
|
* asymmetries. For example, if one of the queues in a set of |
|
* merged queues has a higher weight than a normal queue, then |
|
* the shared queue may inherit such a high weight and, by |
|
* staying almost always active, may force BFQ to perform I/O |
|
* plugging most of the time. This evidently makes it harder |
|
* for BFQ to let the device reach a high throughput. |
|
* |
|
* Finally, the likely() macro below is not used because one |
|
* of the two branches is more likely than the other, but to |
|
* have the code path after the following if() executed as |
|
* fast as possible for the case of a non rotational device |
|
* with queueing. We want it because this is the fastest kind |
|
* of device. On the opposite end, the likely() may lengthen |
|
* the execution time of BFQ for the case of slower devices |
|
* (rotational or at least without queueing). But in this case |
|
* the execution time of BFQ matters very little, if not at |
|
* all. |
|
*/ |
|
if (likely(bfqd->nonrot_with_queueing)) |
|
return NULL; |
|
|
|
/* |
|
* Prevent bfqq from being merged if it has been created too |
|
* long ago. The idea is that true cooperating processes, and |
|
* thus their associated bfq_queues, are supposed to be |
|
* created shortly after each other. This is the case, e.g., |
|
* for KVM/QEMU and dump I/O threads. Basing on this |
|
* assumption, the following filtering greatly reduces the |
|
* probability that two non-cooperating processes, which just |
|
* happen to do close I/O for some short time interval, have |
|
* their queues merged by mistake. |
|
*/ |
|
if (bfq_too_late_for_merging(bfqq)) |
|
return NULL; |
|
|
|
if (!io_struct || unlikely(bfqq == &bfqd->oom_bfqq)) |
|
return NULL; |
|
|
|
/* If there is only one backlogged queue, don't search. */ |
|
if (bfq_tot_busy_queues(bfqd) == 1) |
|
return NULL; |
|
|
|
in_service_bfqq = bfqd->in_service_queue; |
|
|
|
if (in_service_bfqq && in_service_bfqq != bfqq && |
|
likely(in_service_bfqq != &bfqd->oom_bfqq) && |
|
bfq_rq_close_to_sector(io_struct, request, |
|
bfqd->in_serv_last_pos) && |
|
bfqq->entity.parent == in_service_bfqq->entity.parent && |
|
bfq_may_be_close_cooperator(bfqq, in_service_bfqq)) { |
|
new_bfqq = bfq_setup_merge(bfqq, in_service_bfqq); |
|
if (new_bfqq) |
|
return new_bfqq; |
|
} |
|
/* |
|
* Check whether there is a cooperator among currently scheduled |
|
* queues. The only thing we need is that the bio/request is not |
|
* NULL, as we need it to establish whether a cooperator exists. |
|
*/ |
|
new_bfqq = bfq_find_close_cooperator(bfqd, bfqq, |
|
bfq_io_struct_pos(io_struct, request)); |
|
|
|
if (new_bfqq && likely(new_bfqq != &bfqd->oom_bfqq) && |
|
bfq_may_be_close_cooperator(bfqq, new_bfqq)) |
|
return bfq_setup_merge(bfqq, new_bfqq); |
|
|
|
return NULL; |
|
} |
|
|
|
static void bfq_bfqq_save_state(struct bfq_queue *bfqq) |
|
{ |
|
struct bfq_io_cq *bic = bfqq->bic; |
|
|
|
/* |
|
* If !bfqq->bic, the queue is already shared or its requests |
|
* have already been redirected to a shared queue; both idle window |
|
* and weight raising state have already been saved. Do nothing. |
|
*/ |
|
if (!bic) |
|
return; |
|
|
|
bic->saved_last_serv_time_ns = bfqq->last_serv_time_ns; |
|
bic->saved_inject_limit = bfqq->inject_limit; |
|
bic->saved_decrease_time_jif = bfqq->decrease_time_jif; |
|
|
|
bic->saved_weight = bfqq->entity.orig_weight; |
|
bic->saved_ttime = bfqq->ttime; |
|
bic->saved_has_short_ttime = bfq_bfqq_has_short_ttime(bfqq); |
|
bic->saved_IO_bound = bfq_bfqq_IO_bound(bfqq); |
|
bic->saved_io_start_time = bfqq->io_start_time; |
|
bic->saved_tot_idle_time = bfqq->tot_idle_time; |
|
bic->saved_in_large_burst = bfq_bfqq_in_large_burst(bfqq); |
|
bic->was_in_burst_list = !hlist_unhashed(&bfqq->burst_list_node); |
|
if (unlikely(bfq_bfqq_just_created(bfqq) && |
|
!bfq_bfqq_in_large_burst(bfqq) && |
|
bfqq->bfqd->low_latency)) { |
|
/* |
|
* bfqq being merged right after being created: bfqq |
|
* would have deserved interactive weight raising, but |
|
* did not make it to be set in a weight-raised state, |
|
* because of this early merge. Store directly the |
|
* weight-raising state that would have been assigned |
|
* to bfqq, so that to avoid that bfqq unjustly fails |
|
* to enjoy weight raising if split soon. |
|
*/ |
|
bic->saved_wr_coeff = bfqq->bfqd->bfq_wr_coeff; |
|
bic->saved_wr_start_at_switch_to_srt = bfq_smallest_from_now(); |
|
bic->saved_wr_cur_max_time = bfq_wr_duration(bfqq->bfqd); |
|
bic->saved_last_wr_start_finish = jiffies; |
|
} else { |
|
bic->saved_wr_coeff = bfqq->wr_coeff; |
|
bic->saved_wr_start_at_switch_to_srt = |
|
bfqq->wr_start_at_switch_to_srt; |
|
bic->saved_service_from_wr = bfqq->service_from_wr; |
|
bic->saved_last_wr_start_finish = bfqq->last_wr_start_finish; |
|
bic->saved_wr_cur_max_time = bfqq->wr_cur_max_time; |
|
} |
|
} |
|
|
|
|
|
static void |
|
bfq_reassign_last_bfqq(struct bfq_queue *cur_bfqq, struct bfq_queue *new_bfqq) |
|
{ |
|
if (cur_bfqq->entity.parent && |
|
cur_bfqq->entity.parent->last_bfqq_created == cur_bfqq) |
|
cur_bfqq->entity.parent->last_bfqq_created = new_bfqq; |
|
else if (cur_bfqq->bfqd && cur_bfqq->bfqd->last_bfqq_created == cur_bfqq) |
|
cur_bfqq->bfqd->last_bfqq_created = new_bfqq; |
|
} |
|
|
|
void bfq_release_process_ref(struct bfq_data *bfqd, struct bfq_queue *bfqq) |
|
{ |
|
/* |
|
* To prevent bfqq's service guarantees from being violated, |
|
* bfqq may be left busy, i.e., queued for service, even if |
|
* empty (see comments in __bfq_bfqq_expire() for |
|
* details). But, if no process will send requests to bfqq any |
|
* longer, then there is no point in keeping bfqq queued for |
|
* service. In addition, keeping bfqq queued for service, but |
|
* with no process ref any longer, may have caused bfqq to be |
|
* freed when dequeued from service. But this is assumed to |
|
* never happen. |
|
*/ |
|
if (bfq_bfqq_busy(bfqq) && RB_EMPTY_ROOT(&bfqq->sort_list) && |
|
bfqq != bfqd->in_service_queue) |
|
bfq_del_bfqq_busy(bfqd, bfqq, false); |
|
|
|
bfq_reassign_last_bfqq(bfqq, NULL); |
|
|
|
bfq_put_queue(bfqq); |
|
} |
|
|
|
static void |
|
bfq_merge_bfqqs(struct bfq_data *bfqd, struct bfq_io_cq *bic, |
|
struct bfq_queue *bfqq, struct bfq_queue *new_bfqq) |
|
{ |
|
bfq_log_bfqq(bfqd, bfqq, "merging with queue %lu", |
|
(unsigned long)new_bfqq->pid); |
|
/* Save weight raising and idle window of the merged queues */ |
|
bfq_bfqq_save_state(bfqq); |
|
bfq_bfqq_save_state(new_bfqq); |
|
if (bfq_bfqq_IO_bound(bfqq)) |
|
bfq_mark_bfqq_IO_bound(new_bfqq); |
|
bfq_clear_bfqq_IO_bound(bfqq); |
|
|
|
/* |
|
* The processes associated with bfqq are cooperators of the |
|
* processes associated with new_bfqq. So, if bfqq has a |
|
* waker, then assume that all these processes will be happy |
|
* to let bfqq's waker freely inject I/O when they have no |
|
* I/O. |
|
*/ |
|
if (bfqq->waker_bfqq && !new_bfqq->waker_bfqq && |
|
bfqq->waker_bfqq != new_bfqq) { |
|
new_bfqq->waker_bfqq = bfqq->waker_bfqq; |
|
new_bfqq->tentative_waker_bfqq = NULL; |
|
|
|
/* |
|
* If the waker queue disappears, then |
|
* new_bfqq->waker_bfqq must be reset. So insert |
|
* new_bfqq into the woken_list of the waker. See |
|
* bfq_check_waker for details. |
|
*/ |
|
hlist_add_head(&new_bfqq->woken_list_node, |
|
&new_bfqq->waker_bfqq->woken_list); |
|
|
|
} |
|
|
|
/* |
|
* If bfqq is weight-raised, then let new_bfqq inherit |
|
* weight-raising. To reduce false positives, neglect the case |
|
* where bfqq has just been created, but has not yet made it |
|
* to be weight-raised (which may happen because EQM may merge |
|
* bfqq even before bfq_add_request is executed for the first |
|
* time for bfqq). Handling this case would however be very |
|
* easy, thanks to the flag just_created. |
|
*/ |
|
if (new_bfqq->wr_coeff == 1 && bfqq->wr_coeff > 1) { |
|
new_bfqq->wr_coeff = bfqq->wr_coeff; |
|
new_bfqq->wr_cur_max_time = bfqq->wr_cur_max_time; |
|
new_bfqq->last_wr_start_finish = bfqq->last_wr_start_finish; |
|
new_bfqq->wr_start_at_switch_to_srt = |
|
bfqq->wr_start_at_switch_to_srt; |
|
if (bfq_bfqq_busy(new_bfqq)) |
|
bfqd->wr_busy_queues++; |
|
new_bfqq->entity.prio_changed = 1; |
|
} |
|
|
|
if (bfqq->wr_coeff > 1) { /* bfqq has given its wr to new_bfqq */ |
|
bfqq->wr_coeff = 1; |
|
bfqq->entity.prio_changed = 1; |
|
if (bfq_bfqq_busy(bfqq)) |
|
bfqd->wr_busy_queues--; |
|
} |
|
|
|
bfq_log_bfqq(bfqd, new_bfqq, "merge_bfqqs: wr_busy %d", |
|
bfqd->wr_busy_queues); |
|
|
|
/* |
|
* Merge queues (that is, let bic redirect its requests to new_bfqq) |
|
*/ |
|
bic_set_bfqq(bic, new_bfqq, 1); |
|
bfq_mark_bfqq_coop(new_bfqq); |
|
/* |
|
* new_bfqq now belongs to at least two bics (it is a shared queue): |
|
* set new_bfqq->bic to NULL. bfqq either: |
|
* - does not belong to any bic any more, and hence bfqq->bic must |
|
* be set to NULL, or |
|
* - is a queue whose owning bics have already been redirected to a |
|
* different queue, hence the queue is destined to not belong to |
|
* any bic soon and bfqq->bic is already NULL (therefore the next |
|
* assignment causes no harm). |
|
*/ |
|
new_bfqq->bic = NULL; |
|
/* |
|
* If the queue is shared, the pid is the pid of one of the associated |
|
* processes. Which pid depends on the exact sequence of merge events |
|
* the queue underwent. So printing such a pid is useless and confusing |
|
* because it reports a random pid between those of the associated |
|
* processes. |
|
* We mark such a queue with a pid -1, and then print SHARED instead of |
|
* a pid in logging messages. |
|
*/ |
|
new_bfqq->pid = -1; |
|
bfqq->bic = NULL; |
|
|
|
bfq_reassign_last_bfqq(bfqq, new_bfqq); |
|
|
|
bfq_release_process_ref(bfqd, bfqq); |
|
} |
|
|
|
static bool bfq_allow_bio_merge(struct request_queue *q, struct request *rq, |
|
struct bio *bio) |
|
{ |
|
struct bfq_data *bfqd = q->elevator->elevator_data; |
|
bool is_sync = op_is_sync(bio->bi_opf); |
|
struct bfq_queue *bfqq = bfqd->bio_bfqq, *new_bfqq; |
|
|
|
/* |
|
* Disallow merge of a sync bio into an async request. |
|
*/ |
|
if (is_sync && !rq_is_sync(rq)) |
|
return false; |
|
|
|
/* |
|
* Lookup the bfqq that this bio will be queued with. Allow |
|
* merge only if rq is queued there. |
|
*/ |
|
if (!bfqq) |
|
return false; |
|
|
|
/* |
|
* We take advantage of this function to perform an early merge |
|
* of the queues of possible cooperating processes. |
|
*/ |
|
new_bfqq = bfq_setup_cooperator(bfqd, bfqq, bio, false, bfqd->bio_bic); |
|
if (new_bfqq) { |
|
/* |
|
* bic still points to bfqq, then it has not yet been |
|
* redirected to some other bfq_queue, and a queue |
|
* merge between bfqq and new_bfqq can be safely |
|
* fulfilled, i.e., bic can be redirected to new_bfqq |
|
* and bfqq can be put. |
|
*/ |
|
bfq_merge_bfqqs(bfqd, bfqd->bio_bic, bfqq, |
|
new_bfqq); |
|
/* |
|
* If we get here, bio will be queued into new_queue, |
|
* so use new_bfqq to decide whether bio and rq can be |
|
* merged. |
|
*/ |
|
bfqq = new_bfqq; |
|
|
|
/* |
|
* Change also bqfd->bio_bfqq, as |
|
* bfqd->bio_bic now points to new_bfqq, and |
|
* this function may be invoked again (and then may |
|
* use again bqfd->bio_bfqq). |
|
*/ |
|
bfqd->bio_bfqq = bfqq; |
|
} |
|
|
|
return bfqq == RQ_BFQQ(rq); |
|
} |
|
|
|
/* |
|
* Set the maximum time for the in-service queue to consume its |
|
* budget. This prevents seeky processes from lowering the throughput. |
|
* In practice, a time-slice service scheme is used with seeky |
|
* processes. |
|
*/ |
|
static void bfq_set_budget_timeout(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
unsigned int timeout_coeff; |
|
|
|
if (bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time) |
|
timeout_coeff = 1; |
|
else |
|
timeout_coeff = bfqq->entity.weight / bfqq->entity.orig_weight; |
|
|
|
bfqd->last_budget_start = ktime_get(); |
|
|
|
bfqq->budget_timeout = jiffies + |
|
bfqd->bfq_timeout * timeout_coeff; |
|
} |
|
|
|
static void __bfq_set_in_service_queue(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
if (bfqq) { |
|
bfq_clear_bfqq_fifo_expire(bfqq); |
|
|
|
bfqd->budgets_assigned = (bfqd->budgets_assigned * 7 + 256) / 8; |
|
|
|
if (time_is_before_jiffies(bfqq->last_wr_start_finish) && |
|
bfqq->wr_coeff > 1 && |
|
bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time && |
|
time_is_before_jiffies(bfqq->budget_timeout)) { |
|
/* |
|
* For soft real-time queues, move the start |
|
* of the weight-raising period forward by the |
|
* time the queue has not received any |
|
* service. Otherwise, a relatively long |
|
* service delay is likely to cause the |
|
* weight-raising period of the queue to end, |
|
* because of the short duration of the |
|
* weight-raising period of a soft real-time |
|
* queue. It is worth noting that this move |
|
* is not so dangerous for the other queues, |
|
* because soft real-time queues are not |
|
* greedy. |
|
* |
|
* To not add a further variable, we use the |
|
* overloaded field budget_timeout to |
|
* determine for how long the queue has not |
|
* received service, i.e., how much time has |
|
* elapsed since the queue expired. However, |
|
* this is a little imprecise, because |
|
* budget_timeout is set to jiffies if bfqq |
|
* not only expires, but also remains with no |
|
* request. |
|
*/ |
|
if (time_after(bfqq->budget_timeout, |
|
bfqq->last_wr_start_finish)) |
|
bfqq->last_wr_start_finish += |
|
jiffies - bfqq->budget_timeout; |
|
else |
|
bfqq->last_wr_start_finish = jiffies; |
|
} |
|
|
|
bfq_set_budget_timeout(bfqd, bfqq); |
|
bfq_log_bfqq(bfqd, bfqq, |
|
"set_in_service_queue, cur-budget = %d", |
|
bfqq->entity.budget); |
|
} |
|
|
|
bfqd->in_service_queue = bfqq; |
|
bfqd->in_serv_last_pos = 0; |
|
} |
|
|
|
/* |
|
* Get and set a new queue for service. |
|
*/ |
|
static struct bfq_queue *bfq_set_in_service_queue(struct bfq_data *bfqd) |
|
{ |
|
struct bfq_queue *bfqq = bfq_get_next_queue(bfqd); |
|
|
|
__bfq_set_in_service_queue(bfqd, bfqq); |
|
return bfqq; |
|
} |
|
|
|
static void bfq_arm_slice_timer(struct bfq_data *bfqd) |
|
{ |
|
struct bfq_queue *bfqq = bfqd->in_service_queue; |
|
u32 sl; |
|
|
|
bfq_mark_bfqq_wait_request(bfqq); |
|
|
|
/* |
|
* We don't want to idle for seeks, but we do want to allow |
|
* fair distribution of slice time for a process doing back-to-back |
|
* seeks. So allow a little bit of time for him to submit a new rq. |
|
*/ |
|
sl = bfqd->bfq_slice_idle; |
|
/* |
|
* Unless the queue is being weight-raised or the scenario is |
|
* asymmetric, grant only minimum idle time if the queue |
|
* is seeky. A long idling is preserved for a weight-raised |
|
* queue, or, more in general, in an asymmetric scenario, |
|
* because a long idling is needed for guaranteeing to a queue |
|
* its reserved share of the throughput (in particular, it is |
|
* needed if the queue has a higher weight than some other |
|
* queue). |
|
*/ |
|
if (BFQQ_SEEKY(bfqq) && bfqq->wr_coeff == 1 && |
|
!bfq_asymmetric_scenario(bfqd, bfqq)) |
|
sl = min_t(u64, sl, BFQ_MIN_TT); |
|
else if (bfqq->wr_coeff > 1) |
|
sl = max_t(u32, sl, 20ULL * NSEC_PER_MSEC); |
|
|
|
bfqd->last_idling_start = ktime_get(); |
|
bfqd->last_idling_start_jiffies = jiffies; |
|
|
|
hrtimer_start(&bfqd->idle_slice_timer, ns_to_ktime(sl), |
|
HRTIMER_MODE_REL); |
|
bfqg_stats_set_start_idle_time(bfqq_group(bfqq)); |
|
} |
|
|
|
/* |
|
* In autotuning mode, max_budget is dynamically recomputed as the |
|
* amount of sectors transferred in timeout at the estimated peak |
|
* rate. This enables BFQ to utilize a full timeslice with a full |
|
* budget, even if the in-service queue is served at peak rate. And |
|
* this maximises throughput with sequential workloads. |
|
*/ |
|
static unsigned long bfq_calc_max_budget(struct bfq_data *bfqd) |
|
{ |
|
return (u64)bfqd->peak_rate * USEC_PER_MSEC * |
|
jiffies_to_msecs(bfqd->bfq_timeout)>>BFQ_RATE_SHIFT; |
|
} |
|
|
|
/* |
|
* Update parameters related to throughput and responsiveness, as a |
|
* function of the estimated peak rate. See comments on |
|
* bfq_calc_max_budget(), and on the ref_wr_duration array. |
|
*/ |
|
static void update_thr_responsiveness_params(struct bfq_data *bfqd) |
|
{ |
|
if (bfqd->bfq_user_max_budget == 0) { |
|
bfqd->bfq_max_budget = |
|
bfq_calc_max_budget(bfqd); |
|
bfq_log(bfqd, "new max_budget = %d", bfqd->bfq_max_budget); |
|
} |
|
} |
|
|
|
static void bfq_reset_rate_computation(struct bfq_data *bfqd, |
|
struct request *rq) |
|
{ |
|
if (rq != NULL) { /* new rq dispatch now, reset accordingly */ |
|
bfqd->last_dispatch = bfqd->first_dispatch = ktime_get_ns(); |
|
bfqd->peak_rate_samples = 1; |
|
bfqd->sequential_samples = 0; |
|
bfqd->tot_sectors_dispatched = bfqd->last_rq_max_size = |
|
blk_rq_sectors(rq); |
|
} else /* no new rq dispatched, just reset the number of samples */ |
|
bfqd->peak_rate_samples = 0; /* full re-init on next disp. */ |
|
|
|
bfq_log(bfqd, |
|
"reset_rate_computation at end, sample %u/%u tot_sects %llu", |
|
bfqd->peak_rate_samples, bfqd->sequential_samples, |
|
bfqd->tot_sectors_dispatched); |
|
} |
|
|
|
static void bfq_update_rate_reset(struct bfq_data *bfqd, struct request *rq) |
|
{ |
|
u32 rate, weight, divisor; |
|
|
|
/* |
|
* For the convergence property to hold (see comments on |
|
* bfq_update_peak_rate()) and for the assessment to be |
|
* reliable, a minimum number of samples must be present, and |
|
* a minimum amount of time must have elapsed. If not so, do |
|
* not compute new rate. Just reset parameters, to get ready |
|
* for a new evaluation attempt. |
|
*/ |
|
if (bfqd->peak_rate_samples < BFQ_RATE_MIN_SAMPLES || |
|
bfqd->delta_from_first < BFQ_RATE_MIN_INTERVAL) |
|
goto reset_computation; |
|
|
|
/* |
|
* If a new request completion has occurred after last |
|
* dispatch, then, to approximate the rate at which requests |
|
* have been served by the device, it is more precise to |
|
* extend the observation interval to the last completion. |
|
*/ |
|
bfqd->delta_from_first = |
|
max_t(u64, bfqd->delta_from_first, |
|
bfqd->last_completion - bfqd->first_dispatch); |
|
|
|
/* |
|
* Rate computed in sects/usec, and not sects/nsec, for |
|
* precision issues. |
|
*/ |
|
rate = div64_ul(bfqd->tot_sectors_dispatched<<BFQ_RATE_SHIFT, |
|
div_u64(bfqd->delta_from_first, NSEC_PER_USEC)); |
|
|
|
/* |
|
* Peak rate not updated if: |
|
* - the percentage of sequential dispatches is below 3/4 of the |
|
* total, and rate is below the current estimated peak rate |
|
* - rate is unreasonably high (> 20M sectors/sec) |
|
*/ |
|
if ((bfqd->sequential_samples < (3 * bfqd->peak_rate_samples)>>2 && |
|
rate <= bfqd->peak_rate) || |
|
rate > 20<<BFQ_RATE_SHIFT) |
|
goto reset_computation; |
|
|
|
/* |
|
* We have to update the peak rate, at last! To this purpose, |
|
* we use a low-pass filter. We compute the smoothing constant |
|
* of the filter as a function of the 'weight' of the new |
|
* measured rate. |
|
* |
|
* As can be seen in next formulas, we define this weight as a |
|
* quantity proportional to how sequential the workload is, |
|
* and to how long the observation time interval is. |
|
* |
|
* The weight runs from 0 to 8. The maximum value of the |
|
* weight, 8, yields the minimum value for the smoothing |
|
* constant. At this minimum value for the smoothing constant, |
|
* the measured rate contributes for half of the next value of |
|
* the estimated peak rate. |
|
* |
|
* So, the first step is to compute the weight as a function |
|
* of how sequential the workload is. Note that the weight |
|
* cannot reach 9, because bfqd->sequential_samples cannot |
|
* become equal to bfqd->peak_rate_samples, which, in its |
|
* turn, holds true because bfqd->sequential_samples is not |
|
* incremented for the first sample. |
|
*/ |
|
weight = (9 * bfqd->sequential_samples) / bfqd->peak_rate_samples; |
|
|
|
/* |
|
* Second step: further refine the weight as a function of the |
|
* duration of the observation interval. |
|
*/ |
|
weight = min_t(u32, 8, |
|
div_u64(weight * bfqd->delta_from_first, |
|
BFQ_RATE_REF_INTERVAL)); |
|
|
|
/* |
|
* Divisor ranging from 10, for minimum weight, to 2, for |
|
* maximum weight. |
|
*/ |
|
divisor = 10 - weight; |
|
|
|
/* |
|
* Finally, update peak rate: |
|
* |
|
* peak_rate = peak_rate * (divisor-1) / divisor + rate / divisor |
|
*/ |
|
bfqd->peak_rate *= divisor-1; |
|
bfqd->peak_rate /= divisor; |
|
rate /= divisor; /* smoothing constant alpha = 1/divisor */ |
|
|
|
bfqd->peak_rate += rate; |
|
|
|
/* |
|
* For a very slow device, bfqd->peak_rate can reach 0 (see |
|
* the minimum representable values reported in the comments |
|
* on BFQ_RATE_SHIFT). Push to 1 if this happens, to avoid |
|
* divisions by zero where bfqd->peak_rate is used as a |
|
* divisor. |
|
*/ |
|
bfqd->peak_rate = max_t(u32, 1, bfqd->peak_rate); |
|
|
|
update_thr_responsiveness_params(bfqd); |
|
|
|
reset_computation: |
|
bfq_reset_rate_computation(bfqd, rq); |
|
} |
|
|
|
/* |
|
* Update the read/write peak rate (the main quantity used for |
|
* auto-tuning, see update_thr_responsiveness_params()). |
|
* |
|
* It is not trivial to estimate the peak rate (correctly): because of |
|
* the presence of sw and hw queues between the scheduler and the |
|
* device components that finally serve I/O requests, it is hard to |
|
* say exactly when a given dispatched request is served inside the |
|
* device, and for how long. As a consequence, it is hard to know |
|
* precisely at what rate a given set of requests is actually served |
|
* by the device. |
|
* |
|
* On the opposite end, the dispatch time of any request is trivially |
|
* available, and, from this piece of information, the "dispatch rate" |
|
* of requests can be immediately computed. So, the idea in the next |
|
* function is to use what is known, namely request dispatch times |
|
* (plus, when useful, request completion times), to estimate what is |
|
* unknown, namely in-device request service rate. |
|
* |
|
* The main issue is that, because of the above facts, the rate at |
|
* which a certain set of requests is dispatched over a certain time |
|
* interval can vary greatly with respect to the rate at which the |
|
* same requests are then served. But, since the size of any |
|
* intermediate queue is limited, and the service scheme is lossless |
|
* (no request is silently dropped), the following obvious convergence |
|
* property holds: the number of requests dispatched MUST become |
|
* closer and closer to the number of requests completed as the |
|
* observation interval grows. This is the key property used in |
|
* the next function to estimate the peak service rate as a function |
|
* of the observed dispatch rate. The function assumes to be invoked |
|
* on every request dispatch. |
|
*/ |
|
static void bfq_update_peak_rate(struct bfq_data *bfqd, struct request *rq) |
|
{ |
|
u64 now_ns = ktime_get_ns(); |
|
|
|
if (bfqd->peak_rate_samples == 0) { /* first dispatch */ |
|
bfq_log(bfqd, "update_peak_rate: goto reset, samples %d", |
|
bfqd->peak_rate_samples); |
|
bfq_reset_rate_computation(bfqd, rq); |
|
goto update_last_values; /* will add one sample */ |
|
} |
|
|
|
/* |
|
* Device idle for very long: the observation interval lasting |
|
* up to this dispatch cannot be a valid observation interval |
|
* for computing a new peak rate (similarly to the late- |
|
* completion event in bfq_completed_request()). Go to |
|
* update_rate_and_reset to have the following three steps |
|
* taken: |
|
* - close the observation interval at the last (previous) |
|
* request dispatch or completion |
|
* - compute rate, if possible, for that observation interval |
|
* - start a new observation interval with this dispatch |
|
*/ |
|
if (now_ns - bfqd->last_dispatch > 100*NSEC_PER_MSEC && |
|
bfqd->rq_in_driver == 0) |
|
goto update_rate_and_reset; |
|
|
|
/* Update sampling information */ |
|
bfqd->peak_rate_samples++; |
|
|
|
if ((bfqd->rq_in_driver > 0 || |
|
now_ns - bfqd->last_completion < BFQ_MIN_TT) |
|
&& !BFQ_RQ_SEEKY(bfqd, bfqd->last_position, rq)) |
|
bfqd->sequential_samples++; |
|
|
|
bfqd->tot_sectors_dispatched += blk_rq_sectors(rq); |
|
|
|
/* Reset max observed rq size every 32 dispatches */ |
|
if (likely(bfqd->peak_rate_samples % 32)) |
|
bfqd->last_rq_max_size = |
|
max_t(u32, blk_rq_sectors(rq), bfqd->last_rq_max_size); |
|
else |
|
bfqd->last_rq_max_size = blk_rq_sectors(rq); |
|
|
|
bfqd->delta_from_first = now_ns - bfqd->first_dispatch; |
|
|
|
/* Target observation interval not yet reached, go on sampling */ |
|
if (bfqd->delta_from_first < BFQ_RATE_REF_INTERVAL) |
|
goto update_last_values; |
|
|
|
update_rate_and_reset: |
|
bfq_update_rate_reset(bfqd, rq); |
|
update_last_values: |
|
bfqd->last_position = blk_rq_pos(rq) + blk_rq_sectors(rq); |
|
if (RQ_BFQQ(rq) == bfqd->in_service_queue) |
|
bfqd->in_serv_last_pos = bfqd->last_position; |
|
bfqd->last_dispatch = now_ns; |
|
} |
|
|
|
/* |
|
* Remove request from internal lists. |
|
*/ |
|
static void bfq_dispatch_remove(struct request_queue *q, struct request *rq) |
|
{ |
|
struct bfq_queue *bfqq = RQ_BFQQ(rq); |
|
|
|
/* |
|
* For consistency, the next instruction should have been |
|
* executed after removing the request from the queue and |
|
* dispatching it. We execute instead this instruction before |
|
* bfq_remove_request() (and hence introduce a temporary |
|
* inconsistency), for efficiency. In fact, should this |
|
* dispatch occur for a non in-service bfqq, this anticipated |
|
* increment prevents two counters related to bfqq->dispatched |
|
* from risking to be, first, uselessly decremented, and then |
|
* incremented again when the (new) value of bfqq->dispatched |
|
* happens to be taken into account. |
|
*/ |
|
bfqq->dispatched++; |
|
bfq_update_peak_rate(q->elevator->elevator_data, rq); |
|
|
|
bfq_remove_request(q, rq); |
|
} |
|
|
|
/* |
|
* There is a case where idling does not have to be performed for |
|
* throughput concerns, but to preserve the throughput share of |
|
* the process associated with bfqq. |
|
* |
|
* To introduce this case, we can note that allowing the drive |
|
* to enqueue more than one request at a time, and hence |
|
* delegating de facto final scheduling decisions to the |
|
* drive's internal scheduler, entails loss of control on the |
|
* actual request service order. In particular, the critical |
|
* situation is when requests from different processes happen |
|
* to be present, at the same time, in the internal queue(s) |
|
* of the drive. In such a situation, the drive, by deciding |
|
* the service order of the internally-queued requests, does |
|
* determine also the actual throughput distribution among |
|
* these processes. But the drive typically has no notion or |
|
* concern about per-process throughput distribution, and |
|
* makes its decisions only on a per-request basis. Therefore, |
|
* the service distribution enforced by the drive's internal |
|
* scheduler is likely to coincide with the desired throughput |
|
* distribution only in a completely symmetric, or favorably |
|
* skewed scenario where: |
|
* (i-a) each of these processes must get the same throughput as |
|
* the others, |
|
* (i-b) in case (i-a) does not hold, it holds that the process |
|
* associated with bfqq must receive a lower or equal |
|
* throughput than any of the other processes; |
|
* (ii) the I/O of each process has the same properties, in |
|
* terms of locality (sequential or random), direction |
|
* (reads or writes), request sizes, greediness |
|
* (from I/O-bound to sporadic), and so on; |
|
|
|
* In fact, in such a scenario, the drive tends to treat the requests |
|
* of each process in about the same way as the requests of the |
|
* others, and thus to provide each of these processes with about the |
|
* same throughput. This is exactly the desired throughput |
|
* distribution if (i-a) holds, or, if (i-b) holds instead, this is an |
|
* even more convenient distribution for (the process associated with) |
|
* bfqq. |
|
* |
|
* In contrast, in any asymmetric or unfavorable scenario, device |
|
* idling (I/O-dispatch plugging) is certainly needed to guarantee |
|
* that bfqq receives its assigned fraction of the device throughput |
|
* (see [1] for details). |
|
* |
|
* The problem is that idling may significantly reduce throughput with |
|
* certain combinations of types of I/O and devices. An important |
|
* example is sync random I/O on flash storage with command |
|
* queueing. So, unless bfqq falls in cases where idling also boosts |
|
* throughput, it is important to check conditions (i-a), i(-b) and |
|
* (ii) accurately, so as to avoid idling when not strictly needed for |
|
* service guarantees. |
|
* |
|
* Unfortunately, it is extremely difficult to thoroughly check |
|
* condition (ii). And, in case there are active groups, it becomes |
|
* very difficult to check conditions (i-a) and (i-b) too. In fact, |
|
* if there are active groups, then, for conditions (i-a) or (i-b) to |
|
* become false 'indirectly', it is enough that an active group |
|
* contains more active processes or sub-groups than some other active |
|
* group. More precisely, for conditions (i-a) or (i-b) to become |
|
* false because of such a group, it is not even necessary that the |
|
* group is (still) active: it is sufficient that, even if the group |
|
* has become inactive, some of its descendant processes still have |
|
* some request already dispatched but still waiting for |
|
* completion. In fact, requests have still to be guaranteed their |
|
* share of the throughput even after being dispatched. In this |
|
* respect, it is easy to show that, if a group frequently becomes |
|
* inactive while still having in-flight requests, and if, when this |
|
* happens, the group is not considered in the calculation of whether |
|
* the scenario is asymmetric, then the group may fail to be |
|
* guaranteed its fair share of the throughput (basically because |
|
* idling may not be performed for the descendant processes of the |
|
* group, but it had to be). We address this issue with the following |
|
* bi-modal behavior, implemented in the function |
|
* bfq_asymmetric_scenario(). |
|
* |
|
* If there are groups with requests waiting for completion |
|
* (as commented above, some of these groups may even be |
|
* already inactive), then the scenario is tagged as |
|
* asymmetric, conservatively, without checking any of the |
|
* conditions (i-a), (i-b) or (ii). So the device is idled for bfqq. |
|
* This behavior matches also the fact that groups are created |
|
* exactly if controlling I/O is a primary concern (to |
|
* preserve bandwidth and latency guarantees). |
|
* |
|
* On the opposite end, if there are no groups with requests waiting |
|
* for completion, then only conditions (i-a) and (i-b) are actually |
|
* controlled, i.e., provided that conditions (i-a) or (i-b) holds, |
|
* idling is not performed, regardless of whether condition (ii) |
|
* holds. In other words, only if conditions (i-a) and (i-b) do not |
|
* hold, then idling is allowed, and the device tends to be prevented |
|
* from queueing many requests, possibly of several processes. Since |
|
* there are no groups with requests waiting for completion, then, to |
|
* control conditions (i-a) and (i-b) it is enough to check just |
|
* whether all the queues with requests waiting for completion also |
|
* have the same weight. |
|
* |
|
* Not checking condition (ii) evidently exposes bfqq to the |
|
* risk of getting less throughput than its fair share. |
|
* However, for queues with the same weight, a further |
|
* mechanism, preemption, mitigates or even eliminates this |
|
* problem. And it does so without consequences on overall |
|
* throughput. This mechanism and its benefits are explained |
|
* in the next three paragraphs. |
|
* |
|
* Even if a queue, say Q, is expired when it remains idle, Q |
|
* can still preempt the new in-service queue if the next |
|
* request of Q arrives soon (see the comments on |
|
* bfq_bfqq_update_budg_for_activation). If all queues and |
|
* groups have the same weight, this form of preemption, |
|
* combined with the hole-recovery heuristic described in the |
|
* comments on function bfq_bfqq_update_budg_for_activation, |
|
* are enough to preserve a correct bandwidth distribution in |
|
* the mid term, even without idling. In fact, even if not |
|
* idling allows the internal queues of the device to contain |
|
* many requests, and thus to reorder requests, we can rather |
|
* safely assume that the internal scheduler still preserves a |
|
* minimum of mid-term fairness. |
|
* |
|
* More precisely, this preemption-based, idleless approach |
|
* provides fairness in terms of IOPS, and not sectors per |
|
* second. This can be seen with a simple example. Suppose |
|
* that there are two queues with the same weight, but that |
|
* the first queue receives requests of 8 sectors, while the |
|
* second queue receives requests of 1024 sectors. In |
|
* addition, suppose that each of the two queues contains at |
|
* most one request at a time, which implies that each queue |
|
* always remains idle after it is served. Finally, after |
|
* remaining idle, each queue receives very quickly a new |
|
* request. It follows that the two queues are served |
|
* alternatively, preempting each other if needed. This |
|
* implies that, although both queues have the same weight, |
|
* the queue with large requests receives a service that is |
|
* 1024/8 times as high as the service received by the other |
|
* queue. |
|
* |
|
* The motivation for using preemption instead of idling (for |
|
* queues with the same weight) is that, by not idling, |
|
* service guarantees are preserved (completely or at least in |
|
* part) without minimally sacrificing throughput. And, if |
|
* there is no active group, then the primary expectation for |
|
* this device is probably a high throughput. |
|
* |
|
* We are now left only with explaining the two sub-conditions in the |
|
* additional compound condition that is checked below for deciding |
|
* whether the scenario is asymmetric. To explain the first |
|
* sub-condition, we need to add that the function |
|
* bfq_asymmetric_scenario checks the weights of only |
|
* non-weight-raised queues, for efficiency reasons (see comments on |
|
* bfq_weights_tree_add()). Then the fact that bfqq is weight-raised |
|
* is checked explicitly here. More precisely, the compound condition |
|
* below takes into account also the fact that, even if bfqq is being |
|
* weight-raised, the scenario is still symmetric if all queues with |
|
* requests waiting for completion happen to be |
|
* weight-raised. Actually, we should be even more precise here, and |
|
* differentiate between interactive weight raising and soft real-time |
|
* weight raising. |
|
* |
|
* The second sub-condition checked in the compound condition is |
|
* whether there is a fair amount of already in-flight I/O not |
|
* belonging to bfqq. If so, I/O dispatching is to be plugged, for the |
|
* following reason. The drive may decide to serve in-flight |
|
* non-bfqq's I/O requests before bfqq's ones, thereby delaying the |
|
* arrival of new I/O requests for bfqq (recall that bfqq is sync). If |
|
* I/O-dispatching is not plugged, then, while bfqq remains empty, a |
|
* basically uncontrolled amount of I/O from other queues may be |
|
* dispatched too, possibly causing the service of bfqq's I/O to be |
|
* delayed even longer in the drive. This problem gets more and more |
|
* serious as the speed and the queue depth of the drive grow, |
|
* because, as these two quantities grow, the probability to find no |
|
* queue busy but many requests in flight grows too. By contrast, |
|
* plugging I/O dispatching minimizes the delay induced by already |
|
* in-flight I/O, and enables bfqq to recover the bandwidth it may |
|
* lose because of this delay. |
|
* |
|
* As a side note, it is worth considering that the above |
|
* device-idling countermeasures may however fail in the following |
|
* unlucky scenario: if I/O-dispatch plugging is (correctly) disabled |
|
* in a time period during which all symmetry sub-conditions hold, and |
|
* therefore the device is allowed to enqueue many requests, but at |
|
* some later point in time some sub-condition stops to hold, then it |
|
* may become impossible to make requests be served in the desired |
|
* order until all the requests already queued in the device have been |
|
* served. The last sub-condition commented above somewhat mitigates |
|
* this problem for weight-raised queues. |
|
* |
|
* However, as an additional mitigation for this problem, we preserve |
|
* plugging for a special symmetric case that may suddenly turn into |
|
* asymmetric: the case where only bfqq is busy. In this case, not |
|
* expiring bfqq does not cause any harm to any other queues in terms |
|
* of service guarantees. In contrast, it avoids the following unlucky |
|
* sequence of events: (1) bfqq is expired, (2) a new queue with a |
|
* lower weight than bfqq becomes busy (or more queues), (3) the new |
|
* queue is served until a new request arrives for bfqq, (4) when bfqq |
|
* is finally served, there are so many requests of the new queue in |
|
* the drive that the pending requests for bfqq take a lot of time to |
|
* be served. In particular, event (2) may case even already |
|
* dispatched requests of bfqq to be delayed, inside the drive. So, to |
|
* avoid this series of events, the scenario is preventively declared |
|
* as asymmetric also if bfqq is the only busy queues |
|
*/ |
|
static bool idling_needed_for_service_guarantees(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
int tot_busy_queues = bfq_tot_busy_queues(bfqd); |
|
|
|
/* No point in idling for bfqq if it won't get requests any longer */ |
|
if (unlikely(!bfqq_process_refs(bfqq))) |
|
return false; |
|
|
|
return (bfqq->wr_coeff > 1 && |
|
(bfqd->wr_busy_queues < |
|
tot_busy_queues || |
|
bfqd->rq_in_driver >= |
|
bfqq->dispatched + 4)) || |
|
bfq_asymmetric_scenario(bfqd, bfqq) || |
|
tot_busy_queues == 1; |
|
} |
|
|
|
static bool __bfq_bfqq_expire(struct bfq_data *bfqd, struct bfq_queue *bfqq, |
|
enum bfqq_expiration reason) |
|
{ |
|
/* |
|
* If this bfqq is shared between multiple processes, check |
|
* to make sure that those processes are still issuing I/Os |
|
* within the mean seek distance. If not, it may be time to |
|
* break the queues apart again. |
|
*/ |
|
if (bfq_bfqq_coop(bfqq) && BFQQ_SEEKY(bfqq)) |
|
bfq_mark_bfqq_split_coop(bfqq); |
|
|
|
/* |
|
* Consider queues with a higher finish virtual time than |
|
* bfqq. If idling_needed_for_service_guarantees(bfqq) returns |
|
* true, then bfqq's bandwidth would be violated if an |
|
* uncontrolled amount of I/O from these queues were |
|
* dispatched while bfqq is waiting for its new I/O to |
|
* arrive. This is exactly what may happen if this is a forced |
|
* expiration caused by a preemption attempt, and if bfqq is |
|
* not re-scheduled. To prevent this from happening, re-queue |
|
* bfqq if it needs I/O-dispatch plugging, even if it is |
|
* empty. By doing so, bfqq is granted to be served before the |
|
* above queues (provided that bfqq is of course eligible). |
|
*/ |
|
if (RB_EMPTY_ROOT(&bfqq->sort_list) && |
|
!(reason == BFQQE_PREEMPTED && |
|
idling_needed_for_service_guarantees(bfqd, bfqq))) { |
|
if (bfqq->dispatched == 0) |
|
/* |
|
* Overloading budget_timeout field to store |
|
* the time at which the queue remains with no |
|
* backlog and no outstanding request; used by |
|
* the weight-raising mechanism. |
|
*/ |
|
bfqq->budget_timeout = jiffies; |
|
|
|
bfq_del_bfqq_busy(bfqd, bfqq, true); |
|
} else { |
|
bfq_requeue_bfqq(bfqd, bfqq, true); |
|
/* |
|
* Resort priority tree of potential close cooperators. |
|
* See comments on bfq_pos_tree_add_move() for the unlikely(). |
|
*/ |
|
if (unlikely(!bfqd->nonrot_with_queueing && |
|
!RB_EMPTY_ROOT(&bfqq->sort_list))) |
|
bfq_pos_tree_add_move(bfqd, bfqq); |
|
} |
|
|
|
/* |
|
* All in-service entities must have been properly deactivated |
|
* or requeued before executing the next function, which |
|
* resets all in-service entities as no more in service. This |
|
* may cause bfqq to be freed. If this happens, the next |
|
* function returns true. |
|
*/ |
|
return __bfq_bfqd_reset_in_service(bfqd); |
|
} |
|
|
|
/** |
|
* __bfq_bfqq_recalc_budget - try to adapt the budget to the @bfqq behavior. |
|
* @bfqd: device data. |
|
* @bfqq: queue to update. |
|
* @reason: reason for expiration. |
|
* |
|
* Handle the feedback on @bfqq budget at queue expiration. |
|
* See the body for detailed comments. |
|
*/ |
|
static void __bfq_bfqq_recalc_budget(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq, |
|
enum bfqq_expiration reason) |
|
{ |
|
struct request *next_rq; |
|
int budget, min_budget; |
|
|
|
min_budget = bfq_min_budget(bfqd); |
|
|
|
if (bfqq->wr_coeff == 1) |
|
budget = bfqq->max_budget; |
|
else /* |
|
* Use a constant, low budget for weight-raised queues, |
|
* to help achieve a low latency. Keep it slightly higher |
|
* than the minimum possible budget, to cause a little |
|
* bit fewer expirations. |
|
*/ |
|
budget = 2 * min_budget; |
|
|
|
bfq_log_bfqq(bfqd, bfqq, "recalc_budg: last budg %d, budg left %d", |
|
bfqq->entity.budget, bfq_bfqq_budget_left(bfqq)); |
|
bfq_log_bfqq(bfqd, bfqq, "recalc_budg: last max_budg %d, min budg %d", |
|
budget, bfq_min_budget(bfqd)); |
|
bfq_log_bfqq(bfqd, bfqq, "recalc_budg: sync %d, seeky %d", |
|
bfq_bfqq_sync(bfqq), BFQQ_SEEKY(bfqd->in_service_queue)); |
|
|
|
if (bfq_bfqq_sync(bfqq) && bfqq->wr_coeff == 1) { |
|
switch (reason) { |
|
/* |
|
* Caveat: in all the following cases we trade latency |
|
* for throughput. |
|
*/ |
|
case BFQQE_TOO_IDLE: |
|
/* |
|
* This is the only case where we may reduce |
|
* the budget: if there is no request of the |
|
* process still waiting for completion, then |
|
* we assume (tentatively) that the timer has |
|
* expired because the batch of requests of |
|
* the process could have been served with a |
|
* smaller budget. Hence, betting that |
|
* process will behave in the same way when it |
|
* becomes backlogged again, we reduce its |
|
* next budget. As long as we guess right, |
|
* this budget cut reduces the latency |
|
* experienced by the process. |
|
* |
|
* However, if there are still outstanding |
|
* requests, then the process may have not yet |
|
* issued its next request just because it is |
|
* still waiting for the completion of some of |
|
* the still outstanding ones. So in this |
|
* subcase we do not reduce its budget, on the |
|
* contrary we increase it to possibly boost |
|
* the throughput, as discussed in the |
|
* comments to the BUDGET_TIMEOUT case. |
|
*/ |
|
if (bfqq->dispatched > 0) /* still outstanding reqs */ |
|
budget = min(budget * 2, bfqd->bfq_max_budget); |
|
else { |
|
if (budget > 5 * min_budget) |
|
budget -= 4 * min_budget; |
|
else |
|
budget = min_budget; |
|
} |
|
break; |
|
case BFQQE_BUDGET_TIMEOUT: |
|
/* |
|
* We double the budget here because it gives |
|
* the chance to boost the throughput if this |
|
* is not a seeky process (and has bumped into |
|
* this timeout because of, e.g., ZBR). |
|
*/ |
|
budget = min(budget * 2, bfqd->bfq_max_budget); |
|
break; |
|
case BFQQE_BUDGET_EXHAUSTED: |
|
/* |
|
* The process still has backlog, and did not |
|
* let either the budget timeout or the disk |
|
* idling timeout expire. Hence it is not |
|
* seeky, has a short thinktime and may be |
|
* happy with a higher budget too. So |
|
* definitely increase the budget of this good |
|
* candidate to boost the disk throughput. |
|
*/ |
|
budget = min(budget * 4, bfqd->bfq_max_budget); |
|
break; |
|
case BFQQE_NO_MORE_REQUESTS: |
|
/* |
|
* For queues that expire for this reason, it |
|
* is particularly important to keep the |
|
* budget close to the actual service they |
|
* need. Doing so reduces the timestamp |
|
* misalignment problem described in the |
|
* comments in the body of |
|
* __bfq_activate_entity. In fact, suppose |
|
* that a queue systematically expires for |
|
* BFQQE_NO_MORE_REQUESTS and presents a |
|
* new request in time to enjoy timestamp |
|
* back-shifting. The larger the budget of the |
|
* queue is with respect to the service the |
|
* queue actually requests in each service |
|
* slot, the more times the queue can be |
|
* reactivated with the same virtual finish |
|
* time. It follows that, even if this finish |
|
* time is pushed to the system virtual time |
|
* to reduce the consequent timestamp |
|
* misalignment, the queue unjustly enjoys for |
|
* many re-activations a lower finish time |
|
* than all newly activated queues. |
|
* |
|
* The service needed by bfqq is measured |
|
* quite precisely by bfqq->entity.service. |
|
* Since bfqq does not enjoy device idling, |
|
* bfqq->entity.service is equal to the number |
|
* of sectors that the process associated with |
|
* bfqq requested to read/write before waiting |
|
* for request completions, or blocking for |
|
* other reasons. |
|
*/ |
|
budget = max_t(int, bfqq->entity.service, min_budget); |
|
break; |
|
default: |
|
return; |
|
} |
|
} else if (!bfq_bfqq_sync(bfqq)) { |
|
/* |
|
* Async queues get always the maximum possible |
|
* budget, as for them we do not care about latency |
|
* (in addition, their ability to dispatch is limited |
|
* by the charging factor). |
|
*/ |
|
budget = bfqd->bfq_max_budget; |
|
} |
|
|
|
bfqq->max_budget = budget; |
|
|
|
if (bfqd->budgets_assigned >= bfq_stats_min_budgets && |
|
!bfqd->bfq_user_max_budget) |
|
bfqq->max_budget = min(bfqq->max_budget, bfqd->bfq_max_budget); |
|
|
|
/* |
|
* If there is still backlog, then assign a new budget, making |
|
* sure that it is large enough for the next request. Since |
|
* the finish time of bfqq must be kept in sync with the |
|
* budget, be sure to call __bfq_bfqq_expire() *after* this |
|
* update. |
|
* |
|
* If there is no backlog, then no need to update the budget; |
|
* it will be updated on the arrival of a new request. |
|
*/ |
|
next_rq = bfqq->next_rq; |
|
if (next_rq) |
|
bfqq->entity.budget = max_t(unsigned long, bfqq->max_budget, |
|
bfq_serv_to_charge(next_rq, bfqq)); |
|
|
|
bfq_log_bfqq(bfqd, bfqq, "head sect: %u, new budget %d", |
|
next_rq ? blk_rq_sectors(next_rq) : 0, |
|
bfqq->entity.budget); |
|
} |
|
|
|
/* |
|
* Return true if the process associated with bfqq is "slow". The slow |
|
* flag is used, in addition to the budget timeout, to reduce the |
|
* amount of service provided to seeky processes, and thus reduce |
|
* their chances to lower the throughput. More details in the comments |
|
* on the function bfq_bfqq_expire(). |
|
* |
|
* An important observation is in order: as discussed in the comments |
|
* on the function bfq_update_peak_rate(), with devices with internal |
|
* queues, it is hard if ever possible to know when and for how long |
|
* an I/O request is processed by the device (apart from the trivial |
|
* I/O pattern where a new request is dispatched only after the |
|
* previous one has been completed). This makes it hard to evaluate |
|
* the real rate at which the I/O requests of each bfq_queue are |
|
* served. In fact, for an I/O scheduler like BFQ, serving a |
|
* bfq_queue means just dispatching its requests during its service |
|
* slot (i.e., until the budget of the queue is exhausted, or the |
|
* queue remains idle, or, finally, a timeout fires). But, during the |
|
* service slot of a bfq_queue, around 100 ms at most, the device may |
|
* be even still processing requests of bfq_queues served in previous |
|
* service slots. On the opposite end, the requests of the in-service |
|
* bfq_queue may be completed after the service slot of the queue |
|
* finishes. |
|
* |
|
* Anyway, unless more sophisticated solutions are used |
|
* (where possible), the sum of the sizes of the requests dispatched |
|
* during the service slot of a bfq_queue is probably the only |
|
* approximation available for the service received by the bfq_queue |
|
* during its service slot. And this sum is the quantity used in this |
|
* function to evaluate the I/O speed of a process. |
|
*/ |
|
static bool bfq_bfqq_is_slow(struct bfq_data *bfqd, struct bfq_queue *bfqq, |
|
bool compensate, enum bfqq_expiration reason, |
|
unsigned long *delta_ms) |
|
{ |
|
ktime_t delta_ktime; |
|
u32 delta_usecs; |
|
bool slow = BFQQ_SEEKY(bfqq); /* if delta too short, use seekyness */ |
|
|
|
if (!bfq_bfqq_sync(bfqq)) |
|
return false; |
|
|
|
if (compensate) |
|
delta_ktime = bfqd->last_idling_start; |
|
else |
|
delta_ktime = ktime_get(); |
|
delta_ktime = ktime_sub(delta_ktime, bfqd->last_budget_start); |
|
delta_usecs = ktime_to_us(delta_ktime); |
|
|
|
/* don't use too short time intervals */ |
|
if (delta_usecs < 1000) { |
|
if (blk_queue_nonrot(bfqd->queue)) |
|
/* |
|
* give same worst-case guarantees as idling |
|
* for seeky |
|
*/ |
|
*delta_ms = BFQ_MIN_TT / NSEC_PER_MSEC; |
|
else /* charge at least one seek */ |
|
*delta_ms = bfq_slice_idle / NSEC_PER_MSEC; |
|
|
|
return slow; |
|
} |
|
|
|
*delta_ms = delta_usecs / USEC_PER_MSEC; |
|
|
|
/* |
|
* Use only long (> 20ms) intervals to filter out excessive |
|
* spikes in service rate estimation. |
|
*/ |
|
if (delta_usecs > 20000) { |
|
/* |
|
* Caveat for rotational devices: processes doing I/O |
|
* in the slower disk zones tend to be slow(er) even |
|
* if not seeky. In this respect, the estimated peak |
|
* rate is likely to be an average over the disk |
|
* surface. Accordingly, to not be too harsh with |
|
* unlucky processes, a process is deemed slow only if |
|
* its rate has been lower than half of the estimated |
|
* peak rate. |
|
*/ |
|
slow = bfqq->entity.service < bfqd->bfq_max_budget / 2; |
|
} |
|
|
|
bfq_log_bfqq(bfqd, bfqq, "bfq_bfqq_is_slow: slow %d", slow); |
|
|
|
return slow; |
|
} |
|
|
|
/* |
|
* To be deemed as soft real-time, an application must meet two |
|
* requirements. First, the application must not require an average |
|
* bandwidth higher than the approximate bandwidth required to playback or |
|
* record a compressed high-definition video. |
|
* The next function is invoked on the completion of the last request of a |
|
* batch, to compute the next-start time instant, soft_rt_next_start, such |
|
* that, if the next request of the application does not arrive before |
|
* soft_rt_next_start, then the above requirement on the bandwidth is met. |
|
* |
|
* The second requirement is that the request pattern of the application is |
|
* isochronous, i.e., that, after issuing a request or a batch of requests, |
|
* the application stops issuing new requests until all its pending requests |
|
* have been completed. After that, the application may issue a new batch, |
|
* and so on. |
|
* For this reason the next function is invoked to compute |
|
* soft_rt_next_start only for applications that meet this requirement, |
|
* whereas soft_rt_next_start is set to infinity for applications that do |
|
* not. |
|
* |
|
* Unfortunately, even a greedy (i.e., I/O-bound) application may |
|
* happen to meet, occasionally or systematically, both the above |
|
* bandwidth and isochrony requirements. This may happen at least in |
|
* the following circumstances. First, if the CPU load is high. The |
|
* application may stop issuing requests while the CPUs are busy |
|
* serving other processes, then restart, then stop again for a while, |
|
* and so on. The other circumstances are related to the storage |
|
* device: the storage device is highly loaded or reaches a low-enough |
|
* throughput with the I/O of the application (e.g., because the I/O |
|
* is random and/or the device is slow). In all these cases, the |
|
* I/O of the application may be simply slowed down enough to meet |
|
* the bandwidth and isochrony requirements. To reduce the probability |
|
* that greedy applications are deemed as soft real-time in these |
|
* corner cases, a further rule is used in the computation of |
|
* soft_rt_next_start: the return value of this function is forced to |
|
* be higher than the maximum between the following two quantities. |
|
* |
|
* (a) Current time plus: (1) the maximum time for which the arrival |
|
* of a request is waited for when a sync queue becomes idle, |
|
* namely bfqd->bfq_slice_idle, and (2) a few extra jiffies. We |
|
* postpone for a moment the reason for adding a few extra |
|
* jiffies; we get back to it after next item (b). Lower-bounding |
|
* the return value of this function with the current time plus |
|
* bfqd->bfq_slice_idle tends to filter out greedy applications, |
|
* because the latter issue their next request as soon as possible |
|
* after the last one has been completed. In contrast, a soft |
|
* real-time application spends some time processing data, after a |
|
* batch of its requests has been completed. |
|
* |
|
* (b) Current value of bfqq->soft_rt_next_start. As pointed out |
|
* above, greedy applications may happen to meet both the |
|
* bandwidth and isochrony requirements under heavy CPU or |
|
* storage-device load. In more detail, in these scenarios, these |
|
* applications happen, only for limited time periods, to do I/O |
|
* slowly enough to meet all the requirements described so far, |
|
* including the filtering in above item (a). These slow-speed |
|
* time intervals are usually interspersed between other time |
|
* intervals during which these applications do I/O at a very high |
|
* speed. Fortunately, exactly because of the high speed of the |
|
* I/O in the high-speed intervals, the values returned by this |
|
* function happen to be so high, near the end of any such |
|
* high-speed interval, to be likely to fall *after* the end of |
|
* the low-speed time interval that follows. These high values are |
|
* stored in bfqq->soft_rt_next_start after each invocation of |
|
* this function. As a consequence, if the last value of |
|
* bfqq->soft_rt_next_start is constantly used to lower-bound the |
|
* next value that this function may return, then, from the very |
|
* beginning of a low-speed interval, bfqq->soft_rt_next_start is |
|
* likely to be constantly kept so high that any I/O request |
|
* issued during the low-speed interval is considered as arriving |
|
* to soon for the application to be deemed as soft |
|
* real-time. Then, in the high-speed interval that follows, the |
|
* application will not be deemed as soft real-time, just because |
|
* it will do I/O at a high speed. And so on. |
|
* |
|
* Getting back to the filtering in item (a), in the following two |
|
* cases this filtering might be easily passed by a greedy |
|
* application, if the reference quantity was just |
|
* bfqd->bfq_slice_idle: |
|
* 1) HZ is so low that the duration of a jiffy is comparable to or |
|
* higher than bfqd->bfq_slice_idle. This happens, e.g., on slow |
|
* devices with HZ=100. The time granularity may be so coarse |
|
* that the approximation, in jiffies, of bfqd->bfq_slice_idle |
|
* is rather lower than the exact value. |
|
* 2) jiffies, instead of increasing at a constant rate, may stop increasing |
|
* for a while, then suddenly 'jump' by several units to recover the lost |
|
* increments. This seems to happen, e.g., inside virtual machines. |
|
* To address this issue, in the filtering in (a) we do not use as a |
|
* reference time interval just bfqd->bfq_slice_idle, but |
|
* bfqd->bfq_slice_idle plus a few jiffies. In particular, we add the |
|
* minimum number of jiffies for which the filter seems to be quite |
|
* precise also in embedded systems and KVM/QEMU virtual machines. |
|
*/ |
|
static unsigned long bfq_bfqq_softrt_next_start(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
return max3(bfqq->soft_rt_next_start, |
|
bfqq->last_idle_bklogged + |
|
HZ * bfqq->service_from_backlogged / |
|
bfqd->bfq_wr_max_softrt_rate, |
|
jiffies + nsecs_to_jiffies(bfqq->bfqd->bfq_slice_idle) + 4); |
|
} |
|
|
|
/** |
|
* bfq_bfqq_expire - expire a queue. |
|
* @bfqd: device owning the queue. |
|
* @bfqq: the queue to expire. |
|
* @compensate: if true, compensate for the time spent idling. |
|
* @reason: the reason causing the expiration. |
|
* |
|
* If the process associated with bfqq does slow I/O (e.g., because it |
|
* issues random requests), we charge bfqq with the time it has been |
|
* in service instead of the service it has received (see |
|
* bfq_bfqq_charge_time for details on how this goal is achieved). As |
|
* a consequence, bfqq will typically get higher timestamps upon |
|
* reactivation, and hence it will be rescheduled as if it had |
|
* received more service than what it has actually received. In the |
|
* end, bfqq receives less service in proportion to how slowly its |
|
* associated process consumes its budgets (and hence how seriously it |
|
* tends to lower the throughput). In addition, this time-charging |
|
* strategy guarantees time fairness among slow processes. In |
|
* contrast, if the process associated with bfqq is not slow, we |
|
* charge bfqq exactly with the service it has received. |
|
* |
|
* Charging time to the first type of queues and the exact service to |
|
* the other has the effect of using the WF2Q+ policy to schedule the |
|
* former on a timeslice basis, without violating service domain |
|
* guarantees among the latter. |
|
*/ |
|
void bfq_bfqq_expire(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq, |
|
bool compensate, |
|
enum bfqq_expiration reason) |
|
{ |
|
bool slow; |
|
unsigned long delta = 0; |
|
struct bfq_entity *entity = &bfqq->entity; |
|
|
|
/* |
|
* Check whether the process is slow (see bfq_bfqq_is_slow). |
|
*/ |
|
slow = bfq_bfqq_is_slow(bfqd, bfqq, compensate, reason, &delta); |
|
|
|
/* |
|
* As above explained, charge slow (typically seeky) and |
|
* timed-out queues with the time and not the service |
|
* received, to favor sequential workloads. |
|
* |
|
* Processes doing I/O in the slower disk zones will tend to |
|
* be slow(er) even if not seeky. Therefore, since the |
|
* estimated peak rate is actually an average over the disk |
|
* surface, these processes may timeout just for bad luck. To |
|
* avoid punishing them, do not charge time to processes that |
|
* succeeded in consuming at least 2/3 of their budget. This |
|
* allows BFQ to preserve enough elasticity to still perform |
|
* bandwidth, and not time, distribution with little unlucky |
|
* or quasi-sequential processes. |
|
*/ |
|
if (bfqq->wr_coeff == 1 && |
|
(slow || |
|
(reason == BFQQE_BUDGET_TIMEOUT && |
|
bfq_bfqq_budget_left(bfqq) >= entity->budget / 3))) |
|
bfq_bfqq_charge_time(bfqd, bfqq, delta); |
|
|
|
if (bfqd->low_latency && bfqq->wr_coeff == 1) |
|
bfqq->last_wr_start_finish = jiffies; |
|
|
|
if (bfqd->low_latency && bfqd->bfq_wr_max_softrt_rate > 0 && |
|
RB_EMPTY_ROOT(&bfqq->sort_list)) { |
|
/* |
|
* If we get here, and there are no outstanding |
|
* requests, then the request pattern is isochronous |
|
* (see the comments on the function |
|
* bfq_bfqq_softrt_next_start()). Therefore we can |
|
* compute soft_rt_next_start. |
|
* |
|
* If, instead, the queue still has outstanding |
|
* requests, then we have to wait for the completion |
|
* of all the outstanding requests to discover whether |
|
* the request pattern is actually isochronous. |
|
*/ |
|
if (bfqq->dispatched == 0) |
|
bfqq->soft_rt_next_start = |
|
bfq_bfqq_softrt_next_start(bfqd, bfqq); |
|
else if (bfqq->dispatched > 0) { |
|
/* |
|
* Schedule an update of soft_rt_next_start to when |
|
* the task may be discovered to be isochronous. |
|
*/ |
|
bfq_mark_bfqq_softrt_update(bfqq); |
|
} |
|
} |
|
|
|
bfq_log_bfqq(bfqd, bfqq, |
|
"expire (%d, slow %d, num_disp %d, short_ttime %d)", reason, |
|
slow, bfqq->dispatched, bfq_bfqq_has_short_ttime(bfqq)); |
|
|
|
/* |
|
* bfqq expired, so no total service time needs to be computed |
|
* any longer: reset state machine for measuring total service |
|
* times. |
|
*/ |
|
bfqd->rqs_injected = bfqd->wait_dispatch = false; |
|
bfqd->waited_rq = NULL; |
|
|
|
/* |
|
* Increase, decrease or leave budget unchanged according to |
|
* reason. |
|
*/ |
|
__bfq_bfqq_recalc_budget(bfqd, bfqq, reason); |
|
if (__bfq_bfqq_expire(bfqd, bfqq, reason)) |
|
/* bfqq is gone, no more actions on it */ |
|
return; |
|
|
|
/* mark bfqq as waiting a request only if a bic still points to it */ |
|
if (!bfq_bfqq_busy(bfqq) && |
|
reason != BFQQE_BUDGET_TIMEOUT && |
|
reason != BFQQE_BUDGET_EXHAUSTED) { |
|
bfq_mark_bfqq_non_blocking_wait_rq(bfqq); |
|
/* |
|
* Not setting service to 0, because, if the next rq |
|
* arrives in time, the queue will go on receiving |
|
* service with this same budget (as if it never expired) |
|
*/ |
|
} else |
|
entity->service = 0; |
|
|
|
/* |
|
* Reset the received-service counter for every parent entity. |
|
* Differently from what happens with bfqq->entity.service, |
|
* the resetting of this counter never needs to be postponed |
|
* for parent entities. In fact, in case bfqq may have a |
|
* chance to go on being served using the last, partially |
|
* consumed budget, bfqq->entity.service needs to be kept, |
|
* because if bfqq then actually goes on being served using |
|
* the same budget, the last value of bfqq->entity.service is |
|
* needed to properly decrement bfqq->entity.budget by the |
|
* portion already consumed. In contrast, it is not necessary |
|
* to keep entity->service for parent entities too, because |
|
* the bubble up of the new value of bfqq->entity.budget will |
|
* make sure that the budgets of parent entities are correct, |
|
* even in case bfqq and thus parent entities go on receiving |
|
* service with the same budget. |
|
*/ |
|
entity = entity->parent; |
|
for_each_entity(entity) |
|
entity->service = 0; |
|
} |
|
|
|
/* |
|
* Budget timeout is not implemented through a dedicated timer, but |
|
* just checked on request arrivals and completions, as well as on |
|
* idle timer expirations. |
|
*/ |
|
static bool bfq_bfqq_budget_timeout(struct bfq_queue *bfqq) |
|
{ |
|
return time_is_before_eq_jiffies(bfqq->budget_timeout); |
|
} |
|
|
|
/* |
|
* If we expire a queue that is actively waiting (i.e., with the |
|
* device idled) for the arrival of a new request, then we may incur |
|
* the timestamp misalignment problem described in the body of the |
|
* function __bfq_activate_entity. Hence we return true only if this |
|
* condition does not hold, or if the queue is slow enough to deserve |
|
* only to be kicked off for preserving a high throughput. |
|
*/ |
|
static bool bfq_may_expire_for_budg_timeout(struct bfq_queue *bfqq) |
|
{ |
|
bfq_log_bfqq(bfqq->bfqd, bfqq, |
|
"may_budget_timeout: wait_request %d left %d timeout %d", |
|
bfq_bfqq_wait_request(bfqq), |
|
bfq_bfqq_budget_left(bfqq) >= bfqq->entity.budget / 3, |
|
bfq_bfqq_budget_timeout(bfqq)); |
|
|
|
return (!bfq_bfqq_wait_request(bfqq) || |
|
bfq_bfqq_budget_left(bfqq) >= bfqq->entity.budget / 3) |
|
&& |
|
bfq_bfqq_budget_timeout(bfqq); |
|
} |
|
|
|
static bool idling_boosts_thr_without_issues(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
bool rot_without_queueing = |
|
!blk_queue_nonrot(bfqd->queue) && !bfqd->hw_tag, |
|
bfqq_sequential_and_IO_bound, |
|
idling_boosts_thr; |
|
|
|
/* No point in idling for bfqq if it won't get requests any longer */ |
|
if (unlikely(!bfqq_process_refs(bfqq))) |
|
return false; |
|
|
|
bfqq_sequential_and_IO_bound = !BFQQ_SEEKY(bfqq) && |
|
bfq_bfqq_IO_bound(bfqq) && bfq_bfqq_has_short_ttime(bfqq); |
|
|
|
/* |
|
* The next variable takes into account the cases where idling |
|
* boosts the throughput. |
|
* |
|
* The value of the variable is computed considering, first, that |
|
* idling is virtually always beneficial for the throughput if: |
|
* (a) the device is not NCQ-capable and rotational, or |
|
* (b) regardless of the presence of NCQ, the device is rotational and |
|
* the request pattern for bfqq is I/O-bound and sequential, or |
|
* (c) regardless of whether it is rotational, the device is |
|
* not NCQ-capable and the request pattern for bfqq is |
|
* I/O-bound and sequential. |
|
* |
|
* Secondly, and in contrast to the above item (b), idling an |
|
* NCQ-capable flash-based device would not boost the |
|
* throughput even with sequential I/O; rather it would lower |
|
* the throughput in proportion to how fast the device |
|
* is. Accordingly, the next variable is true if any of the |
|
* above conditions (a), (b) or (c) is true, and, in |
|
* particular, happens to be false if bfqd is an NCQ-capable |
|
* flash-based device. |
|
*/ |
|
idling_boosts_thr = rot_without_queueing || |
|
((!blk_queue_nonrot(bfqd->queue) || !bfqd->hw_tag) && |
|
bfqq_sequential_and_IO_bound); |
|
|
|
/* |
|
* The return value of this function is equal to that of |
|
* idling_boosts_thr, unless a special case holds. In this |
|
* special case, described below, idling may cause problems to |
|
* weight-raised queues. |
|
* |
|
* When the request pool is saturated (e.g., in the presence |
|
* of write hogs), if the processes associated with |
|
* non-weight-raised queues ask for requests at a lower rate, |
|
* then processes associated with weight-raised queues have a |
|
* higher probability to get a request from the pool |
|
* immediately (or at least soon) when they need one. Thus |
|
* they have a higher probability to actually get a fraction |
|
* of the device throughput proportional to their high |
|
* weight. This is especially true with NCQ-capable drives, |
|
* which enqueue several requests in advance, and further |
|
* reorder internally-queued requests. |
|
* |
|
* For this reason, we force to false the return value if |
|
* there are weight-raised busy queues. In this case, and if |
|
* bfqq is not weight-raised, this guarantees that the device |
|
* is not idled for bfqq (if, instead, bfqq is weight-raised, |
|
* then idling will be guaranteed by another variable, see |
|
* below). Combined with the timestamping rules of BFQ (see |
|
* [1] for details), this behavior causes bfqq, and hence any |
|
* sync non-weight-raised queue, to get a lower number of |
|
* requests served, and thus to ask for a lower number of |
|
* requests from the request pool, before the busy |
|
* weight-raised queues get served again. This often mitigates |
|
* starvation problems in the presence of heavy write |
|
* workloads and NCQ, thereby guaranteeing a higher |
|
* application and system responsiveness in these hostile |
|
* scenarios. |
|
*/ |
|
return idling_boosts_thr && |
|
bfqd->wr_busy_queues == 0; |
|
} |
|
|
|
/* |
|
* For a queue that becomes empty, device idling is allowed only if |
|
* this function returns true for that queue. As a consequence, since |
|
* device idling plays a critical role for both throughput boosting |
|
* and service guarantees, the return value of this function plays a |
|
* critical role as well. |
|
* |
|
* In a nutshell, this function returns true only if idling is |
|
* beneficial for throughput or, even if detrimental for throughput, |
|
* idling is however necessary to preserve service guarantees (low |
|
* latency, desired throughput distribution, ...). In particular, on |
|
* NCQ-capable devices, this function tries to return false, so as to |
|
* help keep the drives' internal queues full, whenever this helps the |
|
* device boost the throughput without causing any service-guarantee |
|
* issue. |
|
* |
|
* Most of the issues taken into account to get the return value of |
|
* this function are not trivial. We discuss these issues in the two |
|
* functions providing the main pieces of information needed by this |
|
* function. |
|
*/ |
|
static bool bfq_better_to_idle(struct bfq_queue *bfqq) |
|
{ |
|
struct bfq_data *bfqd = bfqq->bfqd; |
|
bool idling_boosts_thr_with_no_issue, idling_needed_for_service_guar; |
|
|
|
/* No point in idling for bfqq if it won't get requests any longer */ |
|
if (unlikely(!bfqq_process_refs(bfqq))) |
|
return false; |
|
|
|
if (unlikely(bfqd->strict_guarantees)) |
|
return true; |
|
|
|
/* |
|
* Idling is performed only if slice_idle > 0. In addition, we |
|
* do not idle if |
|
* (a) bfqq is async |
|
* (b) bfqq is in the idle io prio class: in this case we do |
|
* not idle because we want to minimize the bandwidth that |
|
* queues in this class can steal to higher-priority queues |
|
*/ |
|
if (bfqd->bfq_slice_idle == 0 || !bfq_bfqq_sync(bfqq) || |
|
bfq_class_idle(bfqq)) |
|
return false; |
|
|
|
idling_boosts_thr_with_no_issue = |
|
idling_boosts_thr_without_issues(bfqd, bfqq); |
|
|
|
idling_needed_for_service_guar = |
|
idling_needed_for_service_guarantees(bfqd, bfqq); |
|
|
|
/* |
|
* We have now the two components we need to compute the |
|
* return value of the function, which is true only if idling |
|
* either boosts the throughput (without issues), or is |
|
* necessary to preserve service guarantees. |
|
*/ |
|
return idling_boosts_thr_with_no_issue || |
|
idling_needed_for_service_guar; |
|
} |
|
|
|
/* |
|
* If the in-service queue is empty but the function bfq_better_to_idle |
|
* returns true, then: |
|
* 1) the queue must remain in service and cannot be expired, and |
|
* 2) the device must be idled to wait for the possible arrival of a new |
|
* request for the queue. |
|
* See the comments on the function bfq_better_to_idle for the reasons |
|
* why performing device idling is the best choice to boost the throughput |
|
* and preserve service guarantees when bfq_better_to_idle itself |
|
* returns true. |
|
*/ |
|
static bool bfq_bfqq_must_idle(struct bfq_queue *bfqq) |
|
{ |
|
return RB_EMPTY_ROOT(&bfqq->sort_list) && bfq_better_to_idle(bfqq); |
|
} |
|
|
|
/* |
|
* This function chooses the queue from which to pick the next extra |
|
* I/O request to inject, if it finds a compatible queue. See the |
|
* comments on bfq_update_inject_limit() for details on the injection |
|
* mechanism, and for the definitions of the quantities mentioned |
|
* below. |
|
*/ |
|
static struct bfq_queue * |
|
bfq_choose_bfqq_for_injection(struct bfq_data *bfqd) |
|
{ |
|
struct bfq_queue *bfqq, *in_serv_bfqq = bfqd->in_service_queue; |
|
unsigned int limit = in_serv_bfqq->inject_limit; |
|
/* |
|
* If |
|
* - bfqq is not weight-raised and therefore does not carry |
|
* time-critical I/O, |
|
* or |
|
* - regardless of whether bfqq is weight-raised, bfqq has |
|
* however a long think time, during which it can absorb the |
|
* effect of an appropriate number of extra I/O requests |
|
* from other queues (see bfq_update_inject_limit for |
|
* details on the computation of this number); |
|
* then injection can be performed without restrictions. |
|
*/ |
|
bool in_serv_always_inject = in_serv_bfqq->wr_coeff == 1 || |
|
!bfq_bfqq_has_short_ttime(in_serv_bfqq); |
|
|
|
/* |
|
* If |
|
* - the baseline total service time could not be sampled yet, |
|
* so the inject limit happens to be still 0, and |
|
* - a lot of time has elapsed since the plugging of I/O |
|
* dispatching started, so drive speed is being wasted |
|
* significantly; |
|
* then temporarily raise inject limit to one request. |
|
*/ |
|
if (limit == 0 && in_serv_bfqq->last_serv_time_ns == 0 && |
|
bfq_bfqq_wait_request(in_serv_bfqq) && |
|
time_is_before_eq_jiffies(bfqd->last_idling_start_jiffies + |
|
bfqd->bfq_slice_idle) |
|
) |
|
limit = 1; |
|
|
|
if (bfqd->rq_in_driver >= limit) |
|
return NULL; |
|
|
|
/* |
|
* Linear search of the source queue for injection; but, with |
|
* a high probability, very few steps are needed to find a |
|
* candidate queue, i.e., a queue with enough budget left for |
|
* its next request. In fact: |
|
* - BFQ dynamically updates the budget of every queue so as |
|
* to accommodate the expected backlog of the queue; |
|
* - if a queue gets all its requests dispatched as injected |
|
* service, then the queue is removed from the active list |
|
* (and re-added only if it gets new requests, but then it |
|
* is assigned again enough budget for its new backlog). |
|
*/ |
|
list_for_each_entry(bfqq, &bfqd->active_list, bfqq_list) |
|
if (!RB_EMPTY_ROOT(&bfqq->sort_list) && |
|
(in_serv_always_inject || bfqq->wr_coeff > 1) && |
|
bfq_serv_to_charge(bfqq->next_rq, bfqq) <= |
|
bfq_bfqq_budget_left(bfqq)) { |
|
/* |
|
* Allow for only one large in-flight request |
|
* on non-rotational devices, for the |
|
* following reason. On non-rotationl drives, |
|
* large requests take much longer than |
|
* smaller requests to be served. In addition, |
|
* the drive prefers to serve large requests |
|
* w.r.t. to small ones, if it can choose. So, |
|
* having more than one large requests queued |
|
* in the drive may easily make the next first |
|
* request of the in-service queue wait for so |
|
* long to break bfqq's service guarantees. On |
|
* the bright side, large requests let the |
|
* drive reach a very high throughput, even if |
|
* there is only one in-flight large request |
|
* at a time. |
|
*/ |
|
if (blk_queue_nonrot(bfqd->queue) && |
|
blk_rq_sectors(bfqq->next_rq) >= |
|
BFQQ_SECT_THR_NONROT) |
|
limit = min_t(unsigned int, 1, limit); |
|
else |
|
limit = in_serv_bfqq->inject_limit; |
|
|
|
if (bfqd->rq_in_driver < limit) { |
|
bfqd->rqs_injected = true; |
|
return bfqq; |
|
} |
|
} |
|
|
|
return NULL; |
|
} |
|
|
|
/* |
|
* Select a queue for service. If we have a current queue in service, |
|
* check whether to continue servicing it, or retrieve and set a new one. |
|
*/ |
|
static struct bfq_queue *bfq_select_queue(struct bfq_data *bfqd) |
|
{ |
|
struct bfq_queue *bfqq; |
|
struct request *next_rq; |
|
enum bfqq_expiration reason = BFQQE_BUDGET_TIMEOUT; |
|
|
|
bfqq = bfqd->in_service_queue; |
|
if (!bfqq) |
|
goto new_queue; |
|
|
|
bfq_log_bfqq(bfqd, bfqq, "select_queue: already in-service queue"); |
|
|
|
/* |
|
* Do not expire bfqq for budget timeout if bfqq may be about |
|
* to enjoy device idling. The reason why, in this case, we |
|
* prevent bfqq from expiring is the same as in the comments |
|
* on the case where bfq_bfqq_must_idle() returns true, in |
|
* bfq_completed_request(). |
|
*/ |
|
if (bfq_may_expire_for_budg_timeout(bfqq) && |
|
!bfq_bfqq_must_idle(bfqq)) |
|
goto expire; |
|
|
|
check_queue: |
|
/* |
|
* This loop is rarely executed more than once. Even when it |
|
* happens, it is much more convenient to re-execute this loop |
|
* than to return NULL and trigger a new dispatch to get a |
|
* request served. |
|
*/ |
|
next_rq = bfqq->next_rq; |
|
/* |
|
* If bfqq has requests queued and it has enough budget left to |
|
* serve them, keep the queue, otherwise expire it. |
|
*/ |
|
if (next_rq) { |
|
if (bfq_serv_to_charge(next_rq, bfqq) > |
|
bfq_bfqq_budget_left(bfqq)) { |
|
/* |
|
* Expire the queue for budget exhaustion, |
|
* which makes sure that the next budget is |
|
* enough to serve the next request, even if |
|
* it comes from the fifo expired path. |
|
*/ |
|
reason = BFQQE_BUDGET_EXHAUSTED; |
|
goto expire; |
|
} else { |
|
/* |
|
* The idle timer may be pending because we may |
|
* not disable disk idling even when a new request |
|
* arrives. |
|
*/ |
|
if (bfq_bfqq_wait_request(bfqq)) { |
|
/* |
|
* If we get here: 1) at least a new request |
|
* has arrived but we have not disabled the |
|
* timer because the request was too small, |
|
* 2) then the block layer has unplugged |
|
* the device, causing the dispatch to be |
|
* invoked. |
|
* |
|
* Since the device is unplugged, now the |
|
* requests are probably large enough to |
|
* provide a reasonable throughput. |
|
* So we disable idling. |
|
*/ |
|
bfq_clear_bfqq_wait_request(bfqq); |
|
hrtimer_try_to_cancel(&bfqd->idle_slice_timer); |
|
} |
|
goto keep_queue; |
|
} |
|
} |
|
|
|
/* |
|
* No requests pending. However, if the in-service queue is idling |
|
* for a new request, or has requests waiting for a completion and |
|
* may idle after their completion, then keep it anyway. |
|
* |
|
* Yet, inject service from other queues if it boosts |
|
* throughput and is possible. |
|
*/ |
|
if (bfq_bfqq_wait_request(bfqq) || |
|
(bfqq->dispatched != 0 && bfq_better_to_idle(bfqq))) { |
|
struct bfq_queue *async_bfqq = |
|
bfqq->bic && bfqq->bic->bfqq[0] && |
|
bfq_bfqq_busy(bfqq->bic->bfqq[0]) && |
|
bfqq->bic->bfqq[0]->next_rq ? |
|
bfqq->bic->bfqq[0] : NULL; |
|
struct bfq_queue *blocked_bfqq = |
|
!hlist_empty(&bfqq->woken_list) ? |
|
container_of(bfqq->woken_list.first, |
|
struct bfq_queue, |
|
woken_list_node) |
|
: NULL; |
|
|
|
/* |
|
* The next four mutually-exclusive ifs decide |
|
* whether to try injection, and choose the queue to |
|
* pick an I/O request from. |
|
* |
|
* The first if checks whether the process associated |
|
* with bfqq has also async I/O pending. If so, it |
|
* injects such I/O unconditionally. Injecting async |
|
* I/O from the same process can cause no harm to the |
|
* process. On the contrary, it can only increase |
|
* bandwidth and reduce latency for the process. |
|
* |
|
* The second if checks whether there happens to be a |
|
* non-empty waker queue for bfqq, i.e., a queue whose |
|
* I/O needs to be completed for bfqq to receive new |
|
* I/O. This happens, e.g., if bfqq is associated with |
|
* a process that does some sync. A sync generates |
|
* extra blocking I/O, which must be completed before |
|
* the process associated with bfqq can go on with its |
|
* I/O. If the I/O of the waker queue is not served, |
|
* then bfqq remains empty, and no I/O is dispatched, |
|
* until the idle timeout fires for bfqq. This is |
|
* likely to result in lower bandwidth and higher |
|
* latencies for bfqq, and in a severe loss of total |
|
* throughput. The best action to take is therefore to |
|
* serve the waker queue as soon as possible. So do it |
|
* (without relying on the third alternative below for |
|
* eventually serving waker_bfqq's I/O; see the last |
|
* paragraph for further details). This systematic |
|
* injection of I/O from the waker queue does not |
|
* cause any delay to bfqq's I/O. On the contrary, |
|
* next bfqq's I/O is brought forward dramatically, |
|
* for it is not blocked for milliseconds. |
|
* |
|
* The third if checks whether there is a queue woken |
|
* by bfqq, and currently with pending I/O. Such a |
|
* woken queue does not steal bandwidth from bfqq, |
|
* because it remains soon without I/O if bfqq is not |
|
* served. So there is virtually no risk of loss of |
|
* bandwidth for bfqq if this woken queue has I/O |
|
* dispatched while bfqq is waiting for new I/O. |
|
* |
|
* The fourth if checks whether bfqq is a queue for |
|
* which it is better to avoid injection. It is so if |
|
* bfqq delivers more throughput when served without |
|
* any further I/O from other queues in the middle, or |
|
* if the service times of bfqq's I/O requests both |
|
* count more than overall throughput, and may be |
|
* easily increased by injection (this happens if bfqq |
|
* has a short think time). If none of these |
|
* conditions holds, then a candidate queue for |
|
* injection is looked for through |
|
* bfq_choose_bfqq_for_injection(). Note that the |
|
* latter may return NULL (for example if the inject |
|
* limit for bfqq is currently 0). |
|
* |
|
* NOTE: motivation for the second alternative |
|
* |
|
* Thanks to the way the inject limit is updated in |
|
* bfq_update_has_short_ttime(), it is rather likely |
|
* that, if I/O is being plugged for bfqq and the |
|
* waker queue has pending I/O requests that are |
|
* blocking bfqq's I/O, then the fourth alternative |
|
* above lets the waker queue get served before the |
|
* I/O-plugging timeout fires. So one may deem the |
|
* second alternative superfluous. It is not, because |
|
* the fourth alternative may be way less effective in |
|
* case of a synchronization. For two main |
|
* reasons. First, throughput may be low because the |
|
* inject limit may be too low to guarantee the same |
|
* amount of injected I/O, from the waker queue or |
|
* other queues, that the second alternative |
|
* guarantees (the second alternative unconditionally |
|
* injects a pending I/O request of the waker queue |
|
* for each bfq_dispatch_request()). Second, with the |
|
* fourth alternative, the duration of the plugging, |
|
* i.e., the time before bfqq finally receives new I/O, |
|
* may not be minimized, because the waker queue may |
|
* happen to be served only after other queues. |
|
*/ |
|
if (async_bfqq && |
|
icq_to_bic(async_bfqq->next_rq->elv.icq) == bfqq->bic && |
|
bfq_serv_to_charge(async_bfqq->next_rq, async_bfqq) <= |
|
bfq_bfqq_budget_left(async_bfqq)) |
|
bfqq = bfqq->bic->bfqq[0]; |
|
else if (bfqq->waker_bfqq && |
|
bfq_bfqq_busy(bfqq->waker_bfqq) && |
|
bfqq->waker_bfqq->next_rq && |
|
bfq_serv_to_charge(bfqq->waker_bfqq->next_rq, |
|
bfqq->waker_bfqq) <= |
|
bfq_bfqq_budget_left(bfqq->waker_bfqq) |
|
) |
|
bfqq = bfqq->waker_bfqq; |
|
else if (blocked_bfqq && |
|
bfq_bfqq_busy(blocked_bfqq) && |
|
blocked_bfqq->next_rq && |
|
bfq_serv_to_charge(blocked_bfqq->next_rq, |
|
blocked_bfqq) <= |
|
bfq_bfqq_budget_left(blocked_bfqq) |
|
) |
|
bfqq = blocked_bfqq; |
|
else if (!idling_boosts_thr_without_issues(bfqd, bfqq) && |
|
(bfqq->wr_coeff == 1 || bfqd->wr_busy_queues > 1 || |
|
!bfq_bfqq_has_short_ttime(bfqq))) |
|
bfqq = bfq_choose_bfqq_for_injection(bfqd); |
|
else |
|
bfqq = NULL; |
|
|
|
goto keep_queue; |
|
} |
|
|
|
reason = BFQQE_NO_MORE_REQUESTS; |
|
expire: |
|
bfq_bfqq_expire(bfqd, bfqq, false, reason); |
|
new_queue: |
|
bfqq = bfq_set_in_service_queue(bfqd); |
|
if (bfqq) { |
|
bfq_log_bfqq(bfqd, bfqq, "select_queue: checking new queue"); |
|
goto check_queue; |
|
} |
|
keep_queue: |
|
if (bfqq) |
|
bfq_log_bfqq(bfqd, bfqq, "select_queue: returned this queue"); |
|
else |
|
bfq_log(bfqd, "select_queue: no queue returned"); |
|
|
|
return bfqq; |
|
} |
|
|
|
static void bfq_update_wr_data(struct bfq_data *bfqd, struct bfq_queue *bfqq) |
|
{ |
|
struct bfq_entity *entity = &bfqq->entity; |
|
|
|
if (bfqq->wr_coeff > 1) { /* queue is being weight-raised */ |
|
bfq_log_bfqq(bfqd, bfqq, |
|
"raising period dur %u/%u msec, old coeff %u, w %d(%d)", |
|
jiffies_to_msecs(jiffies - bfqq->last_wr_start_finish), |
|
jiffies_to_msecs(bfqq->wr_cur_max_time), |
|
bfqq->wr_coeff, |
|
bfqq->entity.weight, bfqq->entity.orig_weight); |
|
|
|
if (entity->prio_changed) |
|
bfq_log_bfqq(bfqd, bfqq, "WARN: pending prio change"); |
|
|
|
/* |
|
* If the queue was activated in a burst, or too much |
|
* time has elapsed from the beginning of this |
|
* weight-raising period, then end weight raising. |
|
*/ |
|
if (bfq_bfqq_in_large_burst(bfqq)) |
|
bfq_bfqq_end_wr(bfqq); |
|
else if (time_is_before_jiffies(bfqq->last_wr_start_finish + |
|
bfqq->wr_cur_max_time)) { |
|
if (bfqq->wr_cur_max_time != bfqd->bfq_wr_rt_max_time || |
|
time_is_before_jiffies(bfqq->wr_start_at_switch_to_srt + |
|
bfq_wr_duration(bfqd))) { |
|
/* |
|
* Either in interactive weight |
|
* raising, or in soft_rt weight |
|
* raising with the |
|
* interactive-weight-raising period |
|
* elapsed (so no switch back to |
|
* interactive weight raising). |
|
*/ |
|
bfq_bfqq_end_wr(bfqq); |
|
} else { /* |
|
* soft_rt finishing while still in |
|
* interactive period, switch back to |
|
* interactive weight raising |
|
*/ |
|
switch_back_to_interactive_wr(bfqq, bfqd); |
|
bfqq->entity.prio_changed = 1; |
|
} |
|
} |
|
if (bfqq->wr_coeff > 1 && |
|
bfqq->wr_cur_max_time != bfqd->bfq_wr_rt_max_time && |
|
bfqq->service_from_wr > max_service_from_wr) { |
|
/* see comments on max_service_from_wr */ |
|
bfq_bfqq_end_wr(bfqq); |
|
} |
|
} |
|
/* |
|
* To improve latency (for this or other queues), immediately |
|
* update weight both if it must be raised and if it must be |
|
* lowered. Since, entity may be on some active tree here, and |
|
* might have a pending change of its ioprio class, invoke |
|
* next function with the last parameter unset (see the |
|
* comments on the function). |
|
*/ |
|
if ((entity->weight > entity->orig_weight) != (bfqq->wr_coeff > 1)) |
|
__bfq_entity_update_weight_prio(bfq_entity_service_tree(entity), |
|
entity, false); |
|
} |
|
|
|
/* |
|
* Dispatch next request from bfqq. |
|
*/ |
|
static struct request *bfq_dispatch_rq_from_bfqq(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
struct request *rq = bfqq->next_rq; |
|
unsigned long service_to_charge; |
|
|
|
service_to_charge = bfq_serv_to_charge(rq, bfqq); |
|
|
|
bfq_bfqq_served(bfqq, service_to_charge); |
|
|
|
if (bfqq == bfqd->in_service_queue && bfqd->wait_dispatch) { |
|
bfqd->wait_dispatch = false; |
|
bfqd->waited_rq = rq; |
|
} |
|
|
|
bfq_dispatch_remove(bfqd->queue, rq); |
|
|
|
if (bfqq != bfqd->in_service_queue) |
|
goto return_rq; |
|
|
|
/* |
|
* If weight raising has to terminate for bfqq, then next |
|
* function causes an immediate update of bfqq's weight, |
|
* without waiting for next activation. As a consequence, on |
|
* expiration, bfqq will be timestamped as if has never been |
|
* weight-raised during this service slot, even if it has |
|
* received part or even most of the service as a |
|
* weight-raised queue. This inflates bfqq's timestamps, which |
|
* is beneficial, as bfqq is then more willing to leave the |
|
* device immediately to possible other weight-raised queues. |
|
*/ |
|
bfq_update_wr_data(bfqd, bfqq); |
|
|
|
/* |
|
* Expire bfqq, pretending that its budget expired, if bfqq |
|
* belongs to CLASS_IDLE and other queues are waiting for |
|
* service. |
|
*/ |
|
if (!(bfq_tot_busy_queues(bfqd) > 1 && bfq_class_idle(bfqq))) |
|
goto return_rq; |
|
|
|
bfq_bfqq_expire(bfqd, bfqq, false, BFQQE_BUDGET_EXHAUSTED); |
|
|
|
return_rq: |
|
return rq; |
|
} |
|
|
|
static bool bfq_has_work(struct blk_mq_hw_ctx *hctx) |
|
{ |
|
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data; |
|
|
|
/* |
|
* Avoiding lock: a race on bfqd->busy_queues should cause at |
|
* most a call to dispatch for nothing |
|
*/ |
|
return !list_empty_careful(&bfqd->dispatch) || |
|
bfq_tot_busy_queues(bfqd) > 0; |
|
} |
|
|
|
static struct request *__bfq_dispatch_request(struct blk_mq_hw_ctx *hctx) |
|
{ |
|
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data; |
|
struct request *rq = NULL; |
|
struct bfq_queue *bfqq = NULL; |
|
|
|
if (!list_empty(&bfqd->dispatch)) { |
|
rq = list_first_entry(&bfqd->dispatch, struct request, |
|
queuelist); |
|
list_del_init(&rq->queuelist); |
|
|
|
bfqq = RQ_BFQQ(rq); |
|
|
|
if (bfqq) { |
|
/* |
|
* Increment counters here, because this |
|
* dispatch does not follow the standard |
|
* dispatch flow (where counters are |
|
* incremented) |
|
*/ |
|
bfqq->dispatched++; |
|
|
|
goto inc_in_driver_start_rq; |
|
} |
|
|
|
/* |
|
* We exploit the bfq_finish_requeue_request hook to |
|
* decrement rq_in_driver, but |
|
* bfq_finish_requeue_request will not be invoked on |
|
* this request. So, to avoid unbalance, just start |
|
* this request, without incrementing rq_in_driver. As |
|
* a negative consequence, rq_in_driver is deceptively |
|
* lower than it should be while this request is in |
|
* service. This may cause bfq_schedule_dispatch to be |
|
* invoked uselessly. |
|
* |
|
* As for implementing an exact solution, the |
|
* bfq_finish_requeue_request hook, if defined, is |
|
* probably invoked also on this request. So, by |
|
* exploiting this hook, we could 1) increment |
|
* rq_in_driver here, and 2) decrement it in |
|
* bfq_finish_requeue_request. Such a solution would |
|
* let the value of the counter be always accurate, |
|
* but it would entail using an extra interface |
|
* function. This cost seems higher than the benefit, |
|
* being the frequency of non-elevator-private |
|
* requests very low. |
|
*/ |
|
goto start_rq; |
|
} |
|
|
|
bfq_log(bfqd, "dispatch requests: %d busy queues", |
|
bfq_tot_busy_queues(bfqd)); |
|
|
|
if (bfq_tot_busy_queues(bfqd) == 0) |
|
goto exit; |
|
|
|
/* |
|
* Force device to serve one request at a time if |
|
* strict_guarantees is true. Forcing this service scheme is |
|
* currently the ONLY way to guarantee that the request |
|
* service order enforced by the scheduler is respected by a |
|
* queueing device. Otherwise the device is free even to make |
|
* some unlucky request wait for as long as the device |
|
* wishes. |
|
* |
|
* Of course, serving one request at a time may cause loss of |
|
* throughput. |
|
*/ |
|
if (bfqd->strict_guarantees && bfqd->rq_in_driver > 0) |
|
goto exit; |
|
|
|
bfqq = bfq_select_queue(bfqd); |
|
if (!bfqq) |
|
goto exit; |
|
|
|
rq = bfq_dispatch_rq_from_bfqq(bfqd, bfqq); |
|
|
|
if (rq) { |
|
inc_in_driver_start_rq: |
|
bfqd->rq_in_driver++; |
|
start_rq: |
|
rq->rq_flags |= RQF_STARTED; |
|
} |
|
exit: |
|
return rq; |
|
} |
|
|
|
#ifdef CONFIG_BFQ_CGROUP_DEBUG |
|
static void bfq_update_dispatch_stats(struct request_queue *q, |
|
struct request *rq, |
|
struct bfq_queue *in_serv_queue, |
|
bool idle_timer_disabled) |
|
{ |
|
struct bfq_queue *bfqq = rq ? RQ_BFQQ(rq) : NULL; |
|
|
|
if (!idle_timer_disabled && !bfqq) |
|
return; |
|
|
|
/* |
|
* rq and bfqq are guaranteed to exist until this function |
|
* ends, for the following reasons. First, rq can be |
|
* dispatched to the device, and then can be completed and |
|
* freed, only after this function ends. Second, rq cannot be |
|
* merged (and thus freed because of a merge) any longer, |
|
* because it has already started. Thus rq cannot be freed |
|
* before this function ends, and, since rq has a reference to |
|
* bfqq, the same guarantee holds for bfqq too. |
|
* |
|
* In addition, the following queue lock guarantees that |
|
* bfqq_group(bfqq) exists as well. |
|
*/ |
|
spin_lock_irq(&q->queue_lock); |
|
if (idle_timer_disabled) |
|
/* |
|
* Since the idle timer has been disabled, |
|
* in_serv_queue contained some request when |
|
* __bfq_dispatch_request was invoked above, which |
|
* implies that rq was picked exactly from |
|
* in_serv_queue. Thus in_serv_queue == bfqq, and is |
|
* therefore guaranteed to exist because of the above |
|
* arguments. |
|
*/ |
|
bfqg_stats_update_idle_time(bfqq_group(in_serv_queue)); |
|
if (bfqq) { |
|
struct bfq_group *bfqg = bfqq_group(bfqq); |
|
|
|
bfqg_stats_update_avg_queue_size(bfqg); |
|
bfqg_stats_set_start_empty_time(bfqg); |
|
bfqg_stats_update_io_remove(bfqg, rq->cmd_flags); |
|
} |
|
spin_unlock_irq(&q->queue_lock); |
|
} |
|
#else |
|
static inline void bfq_update_dispatch_stats(struct request_queue *q, |
|
struct request *rq, |
|
struct bfq_queue *in_serv_queue, |
|
bool idle_timer_disabled) {} |
|
#endif /* CONFIG_BFQ_CGROUP_DEBUG */ |
|
|
|
static struct request *bfq_dispatch_request(struct blk_mq_hw_ctx *hctx) |
|
{ |
|
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data; |
|
struct request *rq; |
|
struct bfq_queue *in_serv_queue; |
|
bool waiting_rq, idle_timer_disabled; |
|
|
|
spin_lock_irq(&bfqd->lock); |
|
|
|
in_serv_queue = bfqd->in_service_queue; |
|
waiting_rq = in_serv_queue && bfq_bfqq_wait_request(in_serv_queue); |
|
|
|
rq = __bfq_dispatch_request(hctx); |
|
|
|
idle_timer_disabled = |
|
waiting_rq && !bfq_bfqq_wait_request(in_serv_queue); |
|
|
|
spin_unlock_irq(&bfqd->lock); |
|
|
|
bfq_update_dispatch_stats(hctx->queue, rq, in_serv_queue, |
|
idle_timer_disabled); |
|
|
|
return rq; |
|
} |
|
|
|
/* |
|
* Task holds one reference to the queue, dropped when task exits. Each rq |
|
* in-flight on this queue also holds a reference, dropped when rq is freed. |
|
* |
|
* Scheduler lock must be held here. Recall not to use bfqq after calling |
|
* this function on it. |
|
*/ |
|
void bfq_put_queue(struct bfq_queue *bfqq) |
|
{ |
|
struct bfq_queue *item; |
|
struct hlist_node *n; |
|
struct bfq_group *bfqg = bfqq_group(bfqq); |
|
|
|
if (bfqq->bfqd) |
|
bfq_log_bfqq(bfqq->bfqd, bfqq, "put_queue: %p %d", |
|
bfqq, bfqq->ref); |
|
|
|
bfqq->ref--; |
|
if (bfqq->ref) |
|
return; |
|
|
|
if (!hlist_unhashed(&bfqq->burst_list_node)) { |
|
hlist_del_init(&bfqq->burst_list_node); |
|
/* |
|
* Decrement also burst size after the removal, if the |
|
* process associated with bfqq is exiting, and thus |
|
* does not contribute to the burst any longer. This |
|
* decrement helps filter out false positives of large |
|
* bursts, when some short-lived process (often due to |
|
* the execution of commands by some service) happens |
|
* to start and exit while a complex application is |
|
* starting, and thus spawning several processes that |
|
* do I/O (and that *must not* be treated as a large |
|
* burst, see comments on bfq_handle_burst). |
|
* |
|
* In particular, the decrement is performed only if: |
|
* 1) bfqq is not a merged queue, because, if it is, |
|
* then this free of bfqq is not triggered by the exit |
|
* of the process bfqq is associated with, but exactly |
|
* by the fact that bfqq has just been merged. |
|
* 2) burst_size is greater than 0, to handle |
|
* unbalanced decrements. Unbalanced decrements may |
|
* happen in te following case: bfqq is inserted into |
|
* the current burst list--without incrementing |
|
* bust_size--because of a split, but the current |
|
* burst list is not the burst list bfqq belonged to |
|
* (see comments on the case of a split in |
|
* bfq_set_request). |
|
*/ |
|
if (bfqq->bic && bfqq->bfqd->burst_size > 0) |
|
bfqq->bfqd->burst_size--; |
|
} |
|
|
|
/* |
|
* bfqq does not exist any longer, so it cannot be woken by |
|
* any other queue, and cannot wake any other queue. Then bfqq |
|
* must be removed from the woken list of its possible waker |
|
* queue, and all queues in the woken list of bfqq must stop |
|
* having a waker queue. Strictly speaking, these updates |
|
* should be performed when bfqq remains with no I/O source |
|
* attached to it, which happens before bfqq gets freed. In |
|
* particular, this happens when the last process associated |
|
* with bfqq exits or gets associated with a different |
|
* queue. However, both events lead to bfqq being freed soon, |
|
* and dangling references would come out only after bfqq gets |
|
* freed. So these updates are done here, as a simple and safe |
|
* way to handle all cases. |
|
*/ |
|
/* remove bfqq from woken list */ |
|
if (!hlist_unhashed(&bfqq->woken_list_node)) |
|
hlist_del_init(&bfqq->woken_list_node); |
|
|
|
/* reset waker for all queues in woken list */ |
|
hlist_for_each_entry_safe(item, n, &bfqq->woken_list, |
|
woken_list_node) { |
|
item->waker_bfqq = NULL; |
|
hlist_del_init(&item->woken_list_node); |
|
} |
|
|
|
if (bfqq->bfqd && bfqq->bfqd->last_completed_rq_bfqq == bfqq) |
|
bfqq->bfqd->last_completed_rq_bfqq = NULL; |
|
|
|
kmem_cache_free(bfq_pool, bfqq); |
|
bfqg_and_blkg_put(bfqg); |
|
} |
|
|
|
static void bfq_put_stable_ref(struct bfq_queue *bfqq) |
|
{ |
|
bfqq->stable_ref--; |
|
bfq_put_queue(bfqq); |
|
} |
|
|
|
static void bfq_put_cooperator(struct bfq_queue *bfqq) |
|
{ |
|
struct bfq_queue *__bfqq, *next; |
|
|
|
/* |
|
* If this queue was scheduled to merge with another queue, be |
|
* sure to drop the reference taken on that queue (and others in |
|
* the merge chain). See bfq_setup_merge and bfq_merge_bfqqs. |
|
*/ |
|
__bfqq = bfqq->new_bfqq; |
|
while (__bfqq) { |
|
if (__bfqq == bfqq) |
|
break; |
|
next = __bfqq->new_bfqq; |
|
bfq_put_queue(__bfqq); |
|
__bfqq = next; |
|
} |
|
} |
|
|
|
static void bfq_exit_bfqq(struct bfq_data *bfqd, struct bfq_queue *bfqq) |
|
{ |
|
if (bfqq == bfqd->in_service_queue) { |
|
__bfq_bfqq_expire(bfqd, bfqq, BFQQE_BUDGET_TIMEOUT); |
|
bfq_schedule_dispatch(bfqd); |
|
} |
|
|
|
bfq_log_bfqq(bfqd, bfqq, "exit_bfqq: %p, %d", bfqq, bfqq->ref); |
|
|
|
bfq_put_cooperator(bfqq); |
|
|
|
bfq_release_process_ref(bfqd, bfqq); |
|
} |
|
|
|
static void bfq_exit_icq_bfqq(struct bfq_io_cq *bic, bool is_sync) |
|
{ |
|
struct bfq_queue *bfqq = bic_to_bfqq(bic, is_sync); |
|
struct bfq_data *bfqd; |
|
|
|
if (bfqq) |
|
bfqd = bfqq->bfqd; /* NULL if scheduler already exited */ |
|
|
|
if (bfqq && bfqd) { |
|
unsigned long flags; |
|
|
|
spin_lock_irqsave(&bfqd->lock, flags); |
|
bfqq->bic = NULL; |
|
bfq_exit_bfqq(bfqd, bfqq); |
|
bic_set_bfqq(bic, NULL, is_sync); |
|
spin_unlock_irqrestore(&bfqd->lock, flags); |
|
} |
|
} |
|
|
|
static void bfq_exit_icq(struct io_cq *icq) |
|
{ |
|
struct bfq_io_cq *bic = icq_to_bic(icq); |
|
|
|
if (bic->stable_merge_bfqq) { |
|
struct bfq_data *bfqd = bic->stable_merge_bfqq->bfqd; |
|
|
|
/* |
|
* bfqd is NULL if scheduler already exited, and in |
|
* that case this is the last time bfqq is accessed. |
|
*/ |
|
if (bfqd) { |
|
unsigned long flags; |
|
|
|
spin_lock_irqsave(&bfqd->lock, flags); |
|
bfq_put_stable_ref(bic->stable_merge_bfqq); |
|
spin_unlock_irqrestore(&bfqd->lock, flags); |
|
} else { |
|
bfq_put_stable_ref(bic->stable_merge_bfqq); |
|
} |
|
} |
|
|
|
bfq_exit_icq_bfqq(bic, true); |
|
bfq_exit_icq_bfqq(bic, false); |
|
} |
|
|
|
/* |
|
* Update the entity prio values; note that the new values will not |
|
* be used until the next (re)activation. |
|
*/ |
|
static void |
|
bfq_set_next_ioprio_data(struct bfq_queue *bfqq, struct bfq_io_cq *bic) |
|
{ |
|
struct task_struct *tsk = current; |
|
int ioprio_class; |
|
struct bfq_data *bfqd = bfqq->bfqd; |
|
|
|
if (!bfqd) |
|
return; |
|
|
|
ioprio_class = IOPRIO_PRIO_CLASS(bic->ioprio); |
|
switch (ioprio_class) { |
|
default: |
|
pr_err("bdi %s: bfq: bad prio class %d\n", |
|
bdi_dev_name(bfqq->bfqd->queue->disk->bdi), |
|
ioprio_class); |
|
fallthrough; |
|
case IOPRIO_CLASS_NONE: |
|
/* |
|
* No prio set, inherit CPU scheduling settings. |
|
*/ |
|
bfqq->new_ioprio = task_nice_ioprio(tsk); |
|
bfqq->new_ioprio_class = task_nice_ioclass(tsk); |
|
break; |
|
case IOPRIO_CLASS_RT: |
|
bfqq->new_ioprio = IOPRIO_PRIO_DATA(bic->ioprio); |
|
bfqq->new_ioprio_class = IOPRIO_CLASS_RT; |
|
break; |
|
case IOPRIO_CLASS_BE: |
|
bfqq->new_ioprio = IOPRIO_PRIO_DATA(bic->ioprio); |
|
bfqq->new_ioprio_class = IOPRIO_CLASS_BE; |
|
break; |
|
case IOPRIO_CLASS_IDLE: |
|
bfqq->new_ioprio_class = IOPRIO_CLASS_IDLE; |
|
bfqq->new_ioprio = 7; |
|
break; |
|
} |
|
|
|
if (bfqq->new_ioprio >= IOPRIO_NR_LEVELS) { |
|
pr_crit("bfq_set_next_ioprio_data: new_ioprio %d\n", |
|
bfqq->new_ioprio); |
|
bfqq->new_ioprio = IOPRIO_NR_LEVELS - 1; |
|
} |
|
|
|
bfqq->entity.new_weight = bfq_ioprio_to_weight(bfqq->new_ioprio); |
|
bfq_log_bfqq(bfqd, bfqq, "new_ioprio %d new_weight %d", |
|
bfqq->new_ioprio, bfqq->entity.new_weight); |
|
bfqq->entity.prio_changed = 1; |
|
} |
|
|
|
static struct bfq_queue *bfq_get_queue(struct bfq_data *bfqd, |
|
struct bio *bio, bool is_sync, |
|
struct bfq_io_cq *bic, |
|
bool respawn); |
|
|
|
static void bfq_check_ioprio_change(struct bfq_io_cq *bic, struct bio *bio) |
|
{ |
|
struct bfq_data *bfqd = bic_to_bfqd(bic); |
|
struct bfq_queue *bfqq; |
|
int ioprio = bic->icq.ioc->ioprio; |
|
|
|
/* |
|
* This condition may trigger on a newly created bic, be sure to |
|
* drop the lock before returning. |
|
*/ |
|
if (unlikely(!bfqd) || likely(bic->ioprio == ioprio)) |
|
return; |
|
|
|
bic->ioprio = ioprio; |
|
|
|
bfqq = bic_to_bfqq(bic, false); |
|
if (bfqq) { |
|
bfq_release_process_ref(bfqd, bfqq); |
|
bfqq = bfq_get_queue(bfqd, bio, BLK_RW_ASYNC, bic, true); |
|
bic_set_bfqq(bic, bfqq, false); |
|
} |
|
|
|
bfqq = bic_to_bfqq(bic, true); |
|
if (bfqq) |
|
bfq_set_next_ioprio_data(bfqq, bic); |
|
} |
|
|
|
static void bfq_init_bfqq(struct bfq_data *bfqd, struct bfq_queue *bfqq, |
|
struct bfq_io_cq *bic, pid_t pid, int is_sync) |
|
{ |
|
u64 now_ns = ktime_get_ns(); |
|
|
|
RB_CLEAR_NODE(&bfqq->entity.rb_node); |
|
INIT_LIST_HEAD(&bfqq->fifo); |
|
INIT_HLIST_NODE(&bfqq->burst_list_node); |
|
INIT_HLIST_NODE(&bfqq->woken_list_node); |
|
INIT_HLIST_HEAD(&bfqq->woken_list); |
|
|
|
bfqq->ref = 0; |
|
bfqq->bfqd = bfqd; |
|
|
|
if (bic) |
|
bfq_set_next_ioprio_data(bfqq, bic); |
|
|
|
if (is_sync) { |
|
/* |
|
* No need to mark as has_short_ttime if in |
|
* idle_class, because no device idling is performed |
|
* for queues in idle class |
|
*/ |
|
if (!bfq_class_idle(bfqq)) |
|
/* tentatively mark as has_short_ttime */ |
|
bfq_mark_bfqq_has_short_ttime(bfqq); |
|
bfq_mark_bfqq_sync(bfqq); |
|
bfq_mark_bfqq_just_created(bfqq); |
|
} else |
|
bfq_clear_bfqq_sync(bfqq); |
|
|
|
/* set end request to minus infinity from now */ |
|
bfqq->ttime.last_end_request = now_ns + 1; |
|
|
|
bfqq->creation_time = jiffies; |
|
|
|
bfqq->io_start_time = now_ns; |
|
|
|
bfq_mark_bfqq_IO_bound(bfqq); |
|
|
|
bfqq->pid = pid; |
|
|
|
/* Tentative initial value to trade off between thr and lat */ |
|
bfqq->max_budget = (2 * bfq_max_budget(bfqd)) / 3; |
|
bfqq->budget_timeout = bfq_smallest_from_now(); |
|
|
|
bfqq->wr_coeff = 1; |
|
bfqq->last_wr_start_finish = jiffies; |
|
bfqq->wr_start_at_switch_to_srt = bfq_smallest_from_now(); |
|
bfqq->split_time = bfq_smallest_from_now(); |
|
|
|
/* |
|
* To not forget the possibly high bandwidth consumed by a |
|
* process/queue in the recent past, |
|
* bfq_bfqq_softrt_next_start() returns a value at least equal |
|
* to the current value of bfqq->soft_rt_next_start (see |
|
* comments on bfq_bfqq_softrt_next_start). Set |
|
* soft_rt_next_start to now, to mean that bfqq has consumed |
|
* no bandwidth so far. |
|
*/ |
|
bfqq->soft_rt_next_start = jiffies; |
|
|
|
/* first request is almost certainly seeky */ |
|
bfqq->seek_history = 1; |
|
} |
|
|
|
static struct bfq_queue **bfq_async_queue_prio(struct bfq_data *bfqd, |
|
struct bfq_group *bfqg, |
|
int ioprio_class, int ioprio) |
|
{ |
|
switch (ioprio_class) { |
|
case IOPRIO_CLASS_RT: |
|
return &bfqg->async_bfqq[0][ioprio]; |
|
case IOPRIO_CLASS_NONE: |
|
ioprio = IOPRIO_BE_NORM; |
|
fallthrough; |
|
case IOPRIO_CLASS_BE: |
|
return &bfqg->async_bfqq[1][ioprio]; |
|
case IOPRIO_CLASS_IDLE: |
|
return &bfqg->async_idle_bfqq; |
|
default: |
|
return NULL; |
|
} |
|
} |
|
|
|
static struct bfq_queue * |
|
bfq_do_early_stable_merge(struct bfq_data *bfqd, struct bfq_queue *bfqq, |
|
struct bfq_io_cq *bic, |
|
struct bfq_queue *last_bfqq_created) |
|
{ |
|
struct bfq_queue *new_bfqq = |
|
bfq_setup_merge(bfqq, last_bfqq_created); |
|
|
|
if (!new_bfqq) |
|
return bfqq; |
|
|
|
if (new_bfqq->bic) |
|
new_bfqq->bic->stably_merged = true; |
|
bic->stably_merged = true; |
|
|
|
/* |
|
* Reusing merge functions. This implies that |
|
* bfqq->bic must be set too, for |
|
* bfq_merge_bfqqs to correctly save bfqq's |
|
* state before killing it. |
|
*/ |
|
bfqq->bic = bic; |
|
bfq_merge_bfqqs(bfqd, bic, bfqq, new_bfqq); |
|
|
|
return new_bfqq; |
|
} |
|
|
|
/* |
|
* Many throughput-sensitive workloads are made of several parallel |
|
* I/O flows, with all flows generated by the same application, or |
|
* more generically by the same task (e.g., system boot). The most |
|
* counterproductive action with these workloads is plugging I/O |
|
* dispatch when one of the bfq_queues associated with these flows |
|
* remains temporarily empty. |
|
* |
|
* To avoid this plugging, BFQ has been using a burst-handling |
|
* mechanism for years now. This mechanism has proven effective for |
|
* throughput, and not detrimental for service guarantees. The |
|
* following function pushes this mechanism a little bit further, |
|
* basing on the following two facts. |
|
* |
|
* First, all the I/O flows of a the same application or task |
|
* contribute to the execution/completion of that common application |
|
* or task. So the performance figures that matter are total |
|
* throughput of the flows and task-wide I/O latency. In particular, |
|
* these flows do not need to be protected from each other, in terms |
|
* of individual bandwidth or latency. |
|
* |
|
* Second, the above fact holds regardless of the number of flows. |
|
* |
|
* Putting these two facts together, this commits merges stably the |
|
* bfq_queues associated with these I/O flows, i.e., with the |
|
* processes that generate these IO/ flows, regardless of how many the |
|
* involved processes are. |
|
* |
|
* To decide whether a set of bfq_queues is actually associated with |
|
* the I/O flows of a common application or task, and to merge these |
|
* queues stably, this function operates as follows: given a bfq_queue, |
|
* say Q2, currently being created, and the last bfq_queue, say Q1, |
|
* created before Q2, Q2 is merged stably with Q1 if |
|
* - very little time has elapsed since when Q1 was created |
|
* - Q2 has the same ioprio as Q1 |
|
* - Q2 belongs to the same group as Q1 |
|
* |
|
* Merging bfq_queues also reduces scheduling overhead. A fio test |
|
* with ten random readers on /dev/nullb shows a throughput boost of |
|
* 40%, with a quadcore. Since BFQ's execution time amounts to ~50% of |
|
* the total per-request processing time, the above throughput boost |
|
* implies that BFQ's overhead is reduced by more than 50%. |
|
* |
|
* This new mechanism most certainly obsoletes the current |
|
* burst-handling heuristics. We keep those heuristics for the moment. |
|
*/ |
|
static struct bfq_queue *bfq_do_or_sched_stable_merge(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq, |
|
struct bfq_io_cq *bic) |
|
{ |
|
struct bfq_queue **source_bfqq = bfqq->entity.parent ? |
|
&bfqq->entity.parent->last_bfqq_created : |
|
&bfqd->last_bfqq_created; |
|
|
|
struct bfq_queue *last_bfqq_created = *source_bfqq; |
|
|
|
/* |
|
* If last_bfqq_created has not been set yet, then init it. If |
|
* it has been set already, but too long ago, then move it |
|
* forward to bfqq. Finally, move also if bfqq belongs to a |
|
* different group than last_bfqq_created, or if bfqq has a |
|
* different ioprio or ioprio_class. If none of these |
|
* conditions holds true, then try an early stable merge or |
|
* schedule a delayed stable merge. |
|
* |
|
* A delayed merge is scheduled (instead of performing an |
|
* early merge), in case bfqq might soon prove to be more |
|
* throughput-beneficial if not merged. Currently this is |
|
* possible only if bfqd is rotational with no queueing. For |
|
* such a drive, not merging bfqq is better for throughput if |
|
* bfqq happens to contain sequential I/O. So, we wait a |
|
* little bit for enough I/O to flow through bfqq. After that, |
|
* if such an I/O is sequential, then the merge is |
|
* canceled. Otherwise the merge is finally performed. |
|
*/ |
|
if (!last_bfqq_created || |
|
time_before(last_bfqq_created->creation_time + |
|
msecs_to_jiffies(bfq_activation_stable_merging), |
|
bfqq->creation_time) || |
|
bfqq->entity.parent != last_bfqq_created->entity.parent || |
|
bfqq->ioprio != last_bfqq_created->ioprio || |
|
bfqq->ioprio_class != last_bfqq_created->ioprio_class) |
|
*source_bfqq = bfqq; |
|
else if (time_after_eq(last_bfqq_created->creation_time + |
|
bfqd->bfq_burst_interval, |
|
bfqq->creation_time)) { |
|
if (likely(bfqd->nonrot_with_queueing)) |
|
/* |
|
* With this type of drive, leaving |
|
* bfqq alone may provide no |
|
* throughput benefits compared with |
|
* merging bfqq. So merge bfqq now. |
|
*/ |
|
bfqq = bfq_do_early_stable_merge(bfqd, bfqq, |
|
bic, |
|
last_bfqq_created); |
|
else { /* schedule tentative stable merge */ |
|
/* |
|
* get reference on last_bfqq_created, |
|
* to prevent it from being freed, |
|
* until we decide whether to merge |
|
*/ |
|
last_bfqq_created->ref++; |
|
/* |
|
* need to keep track of stable refs, to |
|
* compute process refs correctly |
|
*/ |
|
last_bfqq_created->stable_ref++; |
|
/* |
|
* Record the bfqq to merge to. |
|
*/ |
|
bic->stable_merge_bfqq = last_bfqq_created; |
|
} |
|
} |
|
|
|
return bfqq; |
|
} |
|
|
|
|
|
static struct bfq_queue *bfq_get_queue(struct bfq_data *bfqd, |
|
struct bio *bio, bool is_sync, |
|
struct bfq_io_cq *bic, |
|
bool respawn) |
|
{ |
|
const int ioprio = IOPRIO_PRIO_DATA(bic->ioprio); |
|
const int ioprio_class = IOPRIO_PRIO_CLASS(bic->ioprio); |
|
struct bfq_queue **async_bfqq = NULL; |
|
struct bfq_queue *bfqq; |
|
struct bfq_group *bfqg; |
|
|
|
rcu_read_lock(); |
|
|
|
bfqg = bfq_find_set_group(bfqd, __bio_blkcg(bio)); |
|
if (!bfqg) { |
|
bfqq = &bfqd->oom_bfqq; |
|
goto out; |
|
} |
|
|
|
if (!is_sync) { |
|
async_bfqq = bfq_async_queue_prio(bfqd, bfqg, ioprio_class, |
|
ioprio); |
|
bfqq = *async_bfqq; |
|
if (bfqq) |
|
goto out; |
|
} |
|
|
|
bfqq = kmem_cache_alloc_node(bfq_pool, |
|
GFP_NOWAIT | __GFP_ZERO | __GFP_NOWARN, |
|
bfqd->queue->node); |
|
|
|
if (bfqq) { |
|
bfq_init_bfqq(bfqd, bfqq, bic, current->pid, |
|
is_sync); |
|
bfq_init_entity(&bfqq->entity, bfqg); |
|
bfq_log_bfqq(bfqd, bfqq, "allocated"); |
|
} else { |
|
bfqq = &bfqd->oom_bfqq; |
|
bfq_log_bfqq(bfqd, bfqq, "using oom bfqq"); |
|
goto out; |
|
} |
|
|
|
/* |
|
* Pin the queue now that it's allocated, scheduler exit will |
|
* prune it. |
|
*/ |
|
if (async_bfqq) { |
|
bfqq->ref++; /* |
|
* Extra group reference, w.r.t. sync |
|
* queue. This extra reference is removed |
|
* only if bfqq->bfqg disappears, to |
|
* guarantee that this queue is not freed |
|
* until its group goes away. |
|
*/ |
|
bfq_log_bfqq(bfqd, bfqq, "get_queue, bfqq not in async: %p, %d", |
|
bfqq, bfqq->ref); |
|
*async_bfqq = bfqq; |
|
} |
|
|
|
out: |
|
bfqq->ref++; /* get a process reference to this queue */ |
|
|
|
if (bfqq != &bfqd->oom_bfqq && is_sync && !respawn) |
|
bfqq = bfq_do_or_sched_stable_merge(bfqd, bfqq, bic); |
|
|
|
rcu_read_unlock(); |
|
return bfqq; |
|
} |
|
|
|
static void bfq_update_io_thinktime(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
struct bfq_ttime *ttime = &bfqq->ttime; |
|
u64 elapsed; |
|
|
|
/* |
|
* We are really interested in how long it takes for the queue to |
|
* become busy when there is no outstanding IO for this queue. So |
|
* ignore cases when the bfq queue has already IO queued. |
|
*/ |
|
if (bfqq->dispatched || bfq_bfqq_busy(bfqq)) |
|
return; |
|
elapsed = ktime_get_ns() - bfqq->ttime.last_end_request; |
|
elapsed = min_t(u64, elapsed, 2ULL * bfqd->bfq_slice_idle); |
|
|
|
ttime->ttime_samples = (7*ttime->ttime_samples + 256) / 8; |
|
ttime->ttime_total = div_u64(7*ttime->ttime_total + 256*elapsed, 8); |
|
ttime->ttime_mean = div64_ul(ttime->ttime_total + 128, |
|
ttime->ttime_samples); |
|
} |
|
|
|
static void |
|
bfq_update_io_seektime(struct bfq_data *bfqd, struct bfq_queue *bfqq, |
|
struct request *rq) |
|
{ |
|
bfqq->seek_history <<= 1; |
|
bfqq->seek_history |= BFQ_RQ_SEEKY(bfqd, bfqq->last_request_pos, rq); |
|
|
|
if (bfqq->wr_coeff > 1 && |
|
bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time && |
|
BFQQ_TOTALLY_SEEKY(bfqq)) { |
|
if (time_is_before_jiffies(bfqq->wr_start_at_switch_to_srt + |
|
bfq_wr_duration(bfqd))) { |
|
/* |
|
* In soft_rt weight raising with the |
|
* interactive-weight-raising period |
|
* elapsed (so no switch back to |
|
* interactive weight raising). |
|
*/ |
|
bfq_bfqq_end_wr(bfqq); |
|
} else { /* |
|
* stopping soft_rt weight raising |
|
* while still in interactive period, |
|
* switch back to interactive weight |
|
* raising |
|
*/ |
|
switch_back_to_interactive_wr(bfqq, bfqd); |
|
bfqq->entity.prio_changed = 1; |
|
} |
|
} |
|
} |
|
|
|
static void bfq_update_has_short_ttime(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq, |
|
struct bfq_io_cq *bic) |
|
{ |
|
bool has_short_ttime = true, state_changed; |
|
|
|
/* |
|
* No need to update has_short_ttime if bfqq is async or in |
|
* idle io prio class, or if bfq_slice_idle is zero, because |
|
* no device idling is performed for bfqq in this case. |
|
*/ |
|
if (!bfq_bfqq_sync(bfqq) || bfq_class_idle(bfqq) || |
|
bfqd->bfq_slice_idle == 0) |
|
return; |
|
|
|
/* Idle window just restored, statistics are meaningless. */ |
|
if (time_is_after_eq_jiffies(bfqq->split_time + |
|
bfqd->bfq_wr_min_idle_time)) |
|
return; |
|
|
|
/* Think time is infinite if no process is linked to |
|
* bfqq. Otherwise check average think time to decide whether |
|
* to mark as has_short_ttime. To this goal, compare average |
|
* think time with half the I/O-plugging timeout. |
|
*/ |
|
if (atomic_read(&bic->icq.ioc->active_ref) == 0 || |
|
(bfq_sample_valid(bfqq->ttime.ttime_samples) && |
|
bfqq->ttime.ttime_mean > bfqd->bfq_slice_idle>>1)) |
|
has_short_ttime = false; |
|
|
|
state_changed = has_short_ttime != bfq_bfqq_has_short_ttime(bfqq); |
|
|
|
if (has_short_ttime) |
|
bfq_mark_bfqq_has_short_ttime(bfqq); |
|
else |
|
bfq_clear_bfqq_has_short_ttime(bfqq); |
|
|
|
/* |
|
* Until the base value for the total service time gets |
|
* finally computed for bfqq, the inject limit does depend on |
|
* the think-time state (short|long). In particular, the limit |
|
* is 0 or 1 if the think time is deemed, respectively, as |
|
* short or long (details in the comments in |
|
* bfq_update_inject_limit()). Accordingly, the next |
|
* instructions reset the inject limit if the think-time state |
|
* has changed and the above base value is still to be |
|
* computed. |
|
* |
|
* However, the reset is performed only if more than 100 ms |
|
* have elapsed since the last update of the inject limit, or |
|
* (inclusive) if the change is from short to long think |
|
* time. The reason for this waiting is as follows. |
|
* |
|
* bfqq may have a long think time because of a |
|
* synchronization with some other queue, i.e., because the |
|
* I/O of some other queue may need to be completed for bfqq |
|
* to receive new I/O. Details in the comments on the choice |
|
* of the queue for injection in bfq_select_queue(). |
|
* |
|
* As stressed in those comments, if such a synchronization is |
|
* actually in place, then, without injection on bfqq, the |
|
* blocking I/O cannot happen to served while bfqq is in |
|
* service. As a consequence, if bfqq is granted |
|
* I/O-dispatch-plugging, then bfqq remains empty, and no I/O |
|
* is dispatched, until the idle timeout fires. This is likely |
|
* to result in lower bandwidth and higher latencies for bfqq, |
|
* and in a severe loss of total throughput. |
|
* |
|
* On the opposite end, a non-zero inject limit may allow the |
|
* I/O that blocks bfqq to be executed soon, and therefore |
|
* bfqq to receive new I/O soon. |
|
* |
|
* But, if the blocking gets actually eliminated, then the |
|
* next think-time sample for bfqq may be very low. This in |
|
* turn may cause bfqq's think time to be deemed |
|
* short. Without the 100 ms barrier, this new state change |
|
* would cause the body of the next if to be executed |
|
* immediately. But this would set to 0 the inject |
|
* limit. Without injection, the blocking I/O would cause the |
|
* think time of bfqq to become long again, and therefore the |
|
* inject limit to be raised again, and so on. The only effect |
|
* of such a steady oscillation between the two think-time |
|
* states would be to prevent effective injection on bfqq. |
|
* |
|
* In contrast, if the inject limit is not reset during such a |
|
* long time interval as 100 ms, then the number of short |
|
* think time samples can grow significantly before the reset |
|
* is performed. As a consequence, the think time state can |
|
* become stable before the reset. Therefore there will be no |
|
* state change when the 100 ms elapse, and no reset of the |
|
* inject limit. The inject limit remains steadily equal to 1 |
|
* both during and after the 100 ms. So injection can be |
|
* performed at all times, and throughput gets boosted. |
|
* |
|
* An inject limit equal to 1 is however in conflict, in |
|
* general, with the fact that the think time of bfqq is |
|
* short, because injection may be likely to delay bfqq's I/O |
|
* (as explained in the comments in |
|
* bfq_update_inject_limit()). But this does not happen in |
|
* this special case, because bfqq's low think time is due to |
|
* an effective handling of a synchronization, through |
|
* injection. In this special case, bfqq's I/O does not get |
|
* delayed by injection; on the contrary, bfqq's I/O is |
|
* brought forward, because it is not blocked for |
|
* milliseconds. |
|
* |
|
* In addition, serving the blocking I/O much sooner, and much |
|
* more frequently than once per I/O-plugging timeout, makes |
|
* it much quicker to detect a waker queue (the concept of |
|
* waker queue is defined in the comments in |
|
* bfq_add_request()). This makes it possible to start sooner |
|
* to boost throughput more effectively, by injecting the I/O |
|
* of the waker queue unconditionally on every |
|
* bfq_dispatch_request(). |
|
* |
|
* One last, important benefit of not resetting the inject |
|
* limit before 100 ms is that, during this time interval, the |
|
* base value for the total service time is likely to get |
|
* finally computed for bfqq, freeing the inject limit from |
|
* its relation with the think time. |
|
*/ |
|
if (state_changed && bfqq->last_serv_time_ns == 0 && |
|
(time_is_before_eq_jiffies(bfqq->decrease_time_jif + |
|
msecs_to_jiffies(100)) || |
|
!has_short_ttime)) |
|
bfq_reset_inject_limit(bfqd, bfqq); |
|
} |
|
|
|
/* |
|
* Called when a new fs request (rq) is added to bfqq. Check if there's |
|
* something we should do about it. |
|
*/ |
|
static void bfq_rq_enqueued(struct bfq_data *bfqd, struct bfq_queue *bfqq, |
|
struct request *rq) |
|
{ |
|
if (rq->cmd_flags & REQ_META) |
|
bfqq->meta_pending++; |
|
|
|
bfqq->last_request_pos = blk_rq_pos(rq) + blk_rq_sectors(rq); |
|
|
|
if (bfqq == bfqd->in_service_queue && bfq_bfqq_wait_request(bfqq)) { |
|
bool small_req = bfqq->queued[rq_is_sync(rq)] == 1 && |
|
blk_rq_sectors(rq) < 32; |
|
bool budget_timeout = bfq_bfqq_budget_timeout(bfqq); |
|
|
|
/* |
|
* There is just this request queued: if |
|
* - the request is small, and |
|
* - we are idling to boost throughput, and |
|
* - the queue is not to be expired, |
|
* then just exit. |
|
* |
|
* In this way, if the device is being idled to wait |
|
* for a new request from the in-service queue, we |
|
* avoid unplugging the device and committing the |
|
* device to serve just a small request. In contrast |
|
* we wait for the block layer to decide when to |
|
* unplug the device: hopefully, new requests will be |
|
* merged to this one quickly, then the device will be |
|
* unplugged and larger requests will be dispatched. |
|
*/ |
|
if (small_req && idling_boosts_thr_without_issues(bfqd, bfqq) && |
|
!budget_timeout) |
|
return; |
|
|
|
/* |
|
* A large enough request arrived, or idling is being |
|
* performed to preserve service guarantees, or |
|
* finally the queue is to be expired: in all these |
|
* cases disk idling is to be stopped, so clear |
|
* wait_request flag and reset timer. |
|
*/ |
|
bfq_clear_bfqq_wait_request(bfqq); |
|
hrtimer_try_to_cancel(&bfqd->idle_slice_timer); |
|
|
|
/* |
|
* The queue is not empty, because a new request just |
|
* arrived. Hence we can safely expire the queue, in |
|
* case of budget timeout, without risking that the |
|
* timestamps of the queue are not updated correctly. |
|
* See [1] for more details. |
|
*/ |
|
if (budget_timeout) |
|
bfq_bfqq_expire(bfqd, bfqq, false, |
|
BFQQE_BUDGET_TIMEOUT); |
|
} |
|
} |
|
|
|
/* returns true if it causes the idle timer to be disabled */ |
|
static bool __bfq_insert_request(struct bfq_data *bfqd, struct request *rq) |
|
{ |
|
struct bfq_queue *bfqq = RQ_BFQQ(rq), |
|
*new_bfqq = bfq_setup_cooperator(bfqd, bfqq, rq, true, |
|
RQ_BIC(rq)); |
|
bool waiting, idle_timer_disabled = false; |
|
|
|
if (new_bfqq) { |
|
/* |
|
* Release the request's reference to the old bfqq |
|
* and make sure one is taken to the shared queue. |
|
*/ |
|
new_bfqq->allocated++; |
|
bfqq->allocated--; |
|
new_bfqq->ref++; |
|
/* |
|
* If the bic associated with the process |
|
* issuing this request still points to bfqq |
|
* (and thus has not been already redirected |
|
* to new_bfqq or even some other bfq_queue), |
|
* then complete the merge and redirect it to |
|
* new_bfqq. |
|
*/ |
|
if (bic_to_bfqq(RQ_BIC(rq), 1) == bfqq) |
|
bfq_merge_bfqqs(bfqd, RQ_BIC(rq), |
|
bfqq, new_bfqq); |
|
|
|
bfq_clear_bfqq_just_created(bfqq); |
|
/* |
|
* rq is about to be enqueued into new_bfqq, |
|
* release rq reference on bfqq |
|
*/ |
|
bfq_put_queue(bfqq); |
|
rq->elv.priv[1] = new_bfqq; |
|
bfqq = new_bfqq; |
|
} |
|
|
|
bfq_update_io_thinktime(bfqd, bfqq); |
|
bfq_update_has_short_ttime(bfqd, bfqq, RQ_BIC(rq)); |
|
bfq_update_io_seektime(bfqd, bfqq, rq); |
|
|
|
waiting = bfqq && bfq_bfqq_wait_request(bfqq); |
|
bfq_add_request(rq); |
|
idle_timer_disabled = waiting && !bfq_bfqq_wait_request(bfqq); |
|
|
|
rq->fifo_time = ktime_get_ns() + bfqd->bfq_fifo_expire[rq_is_sync(rq)]; |
|
list_add_tail(&rq->queuelist, &bfqq->fifo); |
|
|
|
bfq_rq_enqueued(bfqd, bfqq, rq); |
|
|
|
return idle_timer_disabled; |
|
} |
|
|
|
#ifdef CONFIG_BFQ_CGROUP_DEBUG |
|
static void bfq_update_insert_stats(struct request_queue *q, |
|
struct bfq_queue *bfqq, |
|
bool idle_timer_disabled, |
|
unsigned int cmd_flags) |
|
{ |
|
if (!bfqq) |
|
return; |
|
|
|
/* |
|
* bfqq still exists, because it can disappear only after |
|
* either it is merged with another queue, or the process it |
|
* is associated with exits. But both actions must be taken by |
|
* the same process currently executing this flow of |
|
* instructions. |
|
* |
|
* In addition, the following queue lock guarantees that |
|
* bfqq_group(bfqq) exists as well. |
|
*/ |
|
spin_lock_irq(&q->queue_lock); |
|
bfqg_stats_update_io_add(bfqq_group(bfqq), bfqq, cmd_flags); |
|
if (idle_timer_disabled) |
|
bfqg_stats_update_idle_time(bfqq_group(bfqq)); |
|
spin_unlock_irq(&q->queue_lock); |
|
} |
|
#else |
|
static inline void bfq_update_insert_stats(struct request_queue *q, |
|
struct bfq_queue *bfqq, |
|
bool idle_timer_disabled, |
|
unsigned int cmd_flags) {} |
|
#endif /* CONFIG_BFQ_CGROUP_DEBUG */ |
|
|
|
static void bfq_insert_request(struct blk_mq_hw_ctx *hctx, struct request *rq, |
|
bool at_head) |
|
{ |
|
struct request_queue *q = hctx->queue; |
|
struct bfq_data *bfqd = q->elevator->elevator_data; |
|
struct bfq_queue *bfqq; |
|
bool idle_timer_disabled = false; |
|
unsigned int cmd_flags; |
|
LIST_HEAD(free); |
|
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED |
|
if (!cgroup_subsys_on_dfl(io_cgrp_subsys) && rq->bio) |
|
bfqg_stats_update_legacy_io(q, rq); |
|
#endif |
|
spin_lock_irq(&bfqd->lock); |
|
if (blk_mq_sched_try_insert_merge(q, rq, &free)) { |
|
spin_unlock_irq(&bfqd->lock); |
|
blk_mq_free_requests(&free); |
|
return; |
|
} |
|
|
|
spin_unlock_irq(&bfqd->lock); |
|
|
|
trace_block_rq_insert(rq); |
|
|
|
spin_lock_irq(&bfqd->lock); |
|
bfqq = bfq_init_rq(rq); |
|
|
|
/* |
|
* Reqs with at_head or passthrough flags set are to be put |
|
* directly into dispatch list. Additional case for putting rq |
|
* directly into the dispatch queue: the only active |
|
* bfq_queues are bfqq and either its waker bfq_queue or one |
|
* of its woken bfq_queues. The rationale behind this |
|
* additional condition is as follows: |
|
* - consider a bfq_queue, say Q1, detected as a waker of |
|
* another bfq_queue, say Q2 |
|
* - by definition of a waker, Q1 blocks the I/O of Q2, i.e., |
|
* some I/O of Q1 needs to be completed for new I/O of Q2 |
|
* to arrive. A notable example of waker is journald |
|
* - so, Q1 and Q2 are in any respect the queues of two |
|
* cooperating processes (or of two cooperating sets of |
|
* processes): the goal of Q1's I/O is doing what needs to |
|
* be done so that new Q2's I/O can finally be |
|
* issued. Therefore, if the service of Q1's I/O is delayed, |
|
* then Q2's I/O is delayed too. Conversely, if Q2's I/O is |
|
* delayed, the goal of Q1's I/O is hindered. |
|
* - as a consequence, if some I/O of Q1/Q2 arrives while |
|
* Q2/Q1 is the only queue in service, there is absolutely |
|
* no point in delaying the service of such an I/O. The |
|
* only possible result is a throughput loss |
|
* - so, when the above condition holds, the best option is to |
|
* have the new I/O dispatched as soon as possible |
|
* - the most effective and efficient way to attain the above |
|
* goal is to put the new I/O directly in the dispatch |
|
* list |
|
* - as an additional restriction, Q1 and Q2 must be the only |
|
* busy queues for this commit to put the I/O of Q2/Q1 in |
|
* the dispatch list. This is necessary, because, if also |
|
* other queues are waiting for service, then putting new |
|
* I/O directly in the dispatch list may evidently cause a |
|
* violation of service guarantees for the other queues |
|
*/ |
|
if (!bfqq || |
|
(bfqq != bfqd->in_service_queue && |
|
bfqd->in_service_queue != NULL && |
|
bfq_tot_busy_queues(bfqd) == 1 + bfq_bfqq_busy(bfqq) && |
|
(bfqq->waker_bfqq == bfqd->in_service_queue || |
|
bfqd->in_service_queue->waker_bfqq == bfqq)) || at_head) { |
|
if (at_head) |
|
list_add(&rq->queuelist, &bfqd->dispatch); |
|
else |
|
list_add_tail(&rq->queuelist, &bfqd->dispatch); |
|
} else { |
|
idle_timer_disabled = __bfq_insert_request(bfqd, rq); |
|
/* |
|
* Update bfqq, because, if a queue merge has occurred |
|
* in __bfq_insert_request, then rq has been |
|
* redirected into a new queue. |
|
*/ |
|
bfqq = RQ_BFQQ(rq); |
|
|
|
if (rq_mergeable(rq)) { |
|
elv_rqhash_add(q, rq); |
|
if (!q->last_merge) |
|
q->last_merge = rq; |
|
} |
|
} |
|
|
|
/* |
|
* Cache cmd_flags before releasing scheduler lock, because rq |
|
* may disappear afterwards (for example, because of a request |
|
* merge). |
|
*/ |
|
cmd_flags = rq->cmd_flags; |
|
|
|
spin_unlock_irq(&bfqd->lock); |
|
|
|
bfq_update_insert_stats(q, bfqq, idle_timer_disabled, |
|
cmd_flags); |
|
} |
|
|
|
static void bfq_insert_requests(struct blk_mq_hw_ctx *hctx, |
|
struct list_head *list, bool at_head) |
|
{ |
|
while (!list_empty(list)) { |
|
struct request *rq; |
|
|
|
rq = list_first_entry(list, struct request, queuelist); |
|
list_del_init(&rq->queuelist); |
|
bfq_insert_request(hctx, rq, at_head); |
|
} |
|
} |
|
|
|
static void bfq_update_hw_tag(struct bfq_data *bfqd) |
|
{ |
|
struct bfq_queue *bfqq = bfqd->in_service_queue; |
|
|
|
bfqd->max_rq_in_driver = max_t(int, bfqd->max_rq_in_driver, |
|
bfqd->rq_in_driver); |
|
|
|
if (bfqd->hw_tag == 1) |
|
return; |
|
|
|
/* |
|
* This sample is valid if the number of outstanding requests |
|
* is large enough to allow a queueing behavior. Note that the |
|
* sum is not exact, as it's not taking into account deactivated |
|
* requests. |
|
*/ |
|
if (bfqd->rq_in_driver + bfqd->queued <= BFQ_HW_QUEUE_THRESHOLD) |
|
return; |
|
|
|
/* |
|
* If active queue hasn't enough requests and can idle, bfq might not |
|
* dispatch sufficient requests to hardware. Don't zero hw_tag in this |
|
* case |
|
*/ |
|
if (bfqq && bfq_bfqq_has_short_ttime(bfqq) && |
|
bfqq->dispatched + bfqq->queued[0] + bfqq->queued[1] < |
|
BFQ_HW_QUEUE_THRESHOLD && |
|
bfqd->rq_in_driver < BFQ_HW_QUEUE_THRESHOLD) |
|
return; |
|
|
|
if (bfqd->hw_tag_samples++ < BFQ_HW_QUEUE_SAMPLES) |
|
return; |
|
|
|
bfqd->hw_tag = bfqd->max_rq_in_driver > BFQ_HW_QUEUE_THRESHOLD; |
|
bfqd->max_rq_in_driver = 0; |
|
bfqd->hw_tag_samples = 0; |
|
|
|
bfqd->nonrot_with_queueing = |
|
blk_queue_nonrot(bfqd->queue) && bfqd->hw_tag; |
|
} |
|
|
|
static void bfq_completed_request(struct bfq_queue *bfqq, struct bfq_data *bfqd) |
|
{ |
|
u64 now_ns; |
|
u32 delta_us; |
|
|
|
bfq_update_hw_tag(bfqd); |
|
|
|
bfqd->rq_in_driver--; |
|
bfqq->dispatched--; |
|
|
|
if (!bfqq->dispatched && !bfq_bfqq_busy(bfqq)) { |
|
/* |
|
* Set budget_timeout (which we overload to store the |
|
* time at which the queue remains with no backlog and |
|
* no outstanding request; used by the weight-raising |
|
* mechanism). |
|
*/ |
|
bfqq->budget_timeout = jiffies; |
|
|
|
bfq_weights_tree_remove(bfqd, bfqq); |
|
} |
|
|
|
now_ns = ktime_get_ns(); |
|
|
|
bfqq->ttime.last_end_request = now_ns; |
|
|
|
/* |
|
* Using us instead of ns, to get a reasonable precision in |
|
* computing rate in next check. |
|
*/ |
|
delta_us = div_u64(now_ns - bfqd->last_completion, NSEC_PER_USEC); |
|
|
|
/* |
|
* If the request took rather long to complete, and, according |
|
* to the maximum request size recorded, this completion latency |
|
* implies that the request was certainly served at a very low |
|
* rate (less than 1M sectors/sec), then the whole observation |
|
* interval that lasts up to this time instant cannot be a |
|
* valid time interval for computing a new peak rate. Invoke |
|
* bfq_update_rate_reset to have the following three steps |
|
* taken: |
|
* - close the observation interval at the last (previous) |
|
* request dispatch or completion |
|
* - compute rate, if possible, for that observation interval |
|
* - reset to zero samples, which will trigger a proper |
|
* re-initialization of the observation interval on next |
|
* dispatch |
|
*/ |
|
if (delta_us > BFQ_MIN_TT/NSEC_PER_USEC && |
|
(bfqd->last_rq_max_size<<BFQ_RATE_SHIFT)/delta_us < |
|
1UL<<(BFQ_RATE_SHIFT - 10)) |
|
bfq_update_rate_reset(bfqd, NULL); |
|
bfqd->last_completion = now_ns; |
|
/* |
|
* Shared queues are likely to receive I/O at a high |
|
* rate. This may deceptively let them be considered as wakers |
|
* of other queues. But a false waker will unjustly steal |
|
* bandwidth to its supposedly woken queue. So considering |
|
* also shared queues in the waking mechanism may cause more |
|
* control troubles than throughput benefits. Then reset |
|
* last_completed_rq_bfqq if bfqq is a shared queue. |
|
*/ |
|
if (!bfq_bfqq_coop(bfqq)) |
|
bfqd->last_completed_rq_bfqq = bfqq; |
|
else |
|
bfqd->last_completed_rq_bfqq = NULL; |
|
|
|
/* |
|
* If we are waiting to discover whether the request pattern |
|
* of the task associated with the queue is actually |
|
* isochronous, and both requisites for this condition to hold |
|
* are now satisfied, then compute soft_rt_next_start (see the |
|
* comments on the function bfq_bfqq_softrt_next_start()). We |
|
* do not compute soft_rt_next_start if bfqq is in interactive |
|
* weight raising (see the comments in bfq_bfqq_expire() for |
|
* an explanation). We schedule this delayed update when bfqq |
|
* expires, if it still has in-flight requests. |
|
*/ |
|
if (bfq_bfqq_softrt_update(bfqq) && bfqq->dispatched == 0 && |
|
RB_EMPTY_ROOT(&bfqq->sort_list) && |
|
bfqq->wr_coeff != bfqd->bfq_wr_coeff) |
|
bfqq->soft_rt_next_start = |
|
bfq_bfqq_softrt_next_start(bfqd, bfqq); |
|
|
|
/* |
|
* If this is the in-service queue, check if it needs to be expired, |
|
* or if we want to idle in case it has no pending requests. |
|
*/ |
|
if (bfqd->in_service_queue == bfqq) { |
|
if (bfq_bfqq_must_idle(bfqq)) { |
|
if (bfqq->dispatched == 0) |
|
bfq_arm_slice_timer(bfqd); |
|
/* |
|
* If we get here, we do not expire bfqq, even |
|
* if bfqq was in budget timeout or had no |
|
* more requests (as controlled in the next |
|
* conditional instructions). The reason for |
|
* not expiring bfqq is as follows. |
|
* |
|
* Here bfqq->dispatched > 0 holds, but |
|
* bfq_bfqq_must_idle() returned true. This |
|
* implies that, even if no request arrives |
|
* for bfqq before bfqq->dispatched reaches 0, |
|
* bfqq will, however, not be expired on the |
|
* completion event that causes bfqq->dispatch |
|
* to reach zero. In contrast, on this event, |
|
* bfqq will start enjoying device idling |
|
* (I/O-dispatch plugging). |
|
* |
|
* But, if we expired bfqq here, bfqq would |
|
* not have the chance to enjoy device idling |
|
* when bfqq->dispatched finally reaches |
|
* zero. This would expose bfqq to violation |
|
* of its reserved service guarantees. |
|
*/ |
|
return; |
|
} else if (bfq_may_expire_for_budg_timeout(bfqq)) |
|
bfq_bfqq_expire(bfqd, bfqq, false, |
|
BFQQE_BUDGET_TIMEOUT); |
|
else if (RB_EMPTY_ROOT(&bfqq->sort_list) && |
|
(bfqq->dispatched == 0 || |
|
!bfq_better_to_idle(bfqq))) |
|
bfq_bfqq_expire(bfqd, bfqq, false, |
|
BFQQE_NO_MORE_REQUESTS); |
|
} |
|
|
|
if (!bfqd->rq_in_driver) |
|
bfq_schedule_dispatch(bfqd); |
|
} |
|
|
|
static void bfq_finish_requeue_request_body(struct bfq_queue *bfqq) |
|
{ |
|
bfqq->allocated--; |
|
|
|
bfq_put_queue(bfqq); |
|
} |
|
|
|
/* |
|
* The processes associated with bfqq may happen to generate their |
|
* cumulative I/O at a lower rate than the rate at which the device |
|
* could serve the same I/O. This is rather probable, e.g., if only |
|
* one process is associated with bfqq and the device is an SSD. It |
|
* results in bfqq becoming often empty while in service. In this |
|
* respect, if BFQ is allowed to switch to another queue when bfqq |
|
* remains empty, then the device goes on being fed with I/O requests, |
|
* and the throughput is not affected. In contrast, if BFQ is not |
|
* allowed to switch to another queue---because bfqq is sync and |
|
* I/O-dispatch needs to be plugged while bfqq is temporarily |
|
* empty---then, during the service of bfqq, there will be frequent |
|
* "service holes", i.e., time intervals during which bfqq gets empty |
|
* and the device can only consume the I/O already queued in its |
|
* hardware queues. During service holes, the device may even get to |
|
* remaining idle. In the end, during the service of bfqq, the device |
|
* is driven at a lower speed than the one it can reach with the kind |
|
* of I/O flowing through bfqq. |
|
* |
|
* To counter this loss of throughput, BFQ implements a "request |
|
* injection mechanism", which tries to fill the above service holes |
|
* with I/O requests taken from other queues. The hard part in this |
|
* mechanism is finding the right amount of I/O to inject, so as to |
|
* both boost throughput and not break bfqq's bandwidth and latency |
|
* guarantees. In this respect, the mechanism maintains a per-queue |
|
* inject limit, computed as below. While bfqq is empty, the injection |
|
* mechanism dispatches extra I/O requests only until the total number |
|
* of I/O requests in flight---i.e., already dispatched but not yet |
|
* completed---remains lower than this limit. |
|
* |
|
* A first definition comes in handy to introduce the algorithm by |
|
* which the inject limit is computed. We define as first request for |
|
* bfqq, an I/O request for bfqq that arrives while bfqq is in |
|
* service, and causes bfqq to switch from empty to non-empty. The |
|
* algorithm updates the limit as a function of the effect of |
|
* injection on the service times of only the first requests of |
|
* bfqq. The reason for this restriction is that these are the |
|
* requests whose service time is affected most, because they are the |
|
* first to arrive after injection possibly occurred. |
|
* |
|
* To evaluate the effect of injection, the algorithm measures the |
|
* "total service time" of first requests. We define as total service |
|
* time of an I/O request, the time that elapses since when the |
|
* request is enqueued into bfqq, to when it is completed. This |
|
* quantity allows the whole effect of injection to be measured. It is |
|
* easy to see why. Suppose that some requests of other queues are |
|
* actually injected while bfqq is empty, and that a new request R |
|
* then arrives for bfqq. If the device does start to serve all or |
|
* part of the injected requests during the service hole, then, |
|
* because of this extra service, it may delay the next invocation of |
|
* the dispatch hook of BFQ. Then, even after R gets eventually |
|
* dispatched, the device may delay the actual service of R if it is |
|
* still busy serving the extra requests, or if it decides to serve, |
|
* before R, some extra request still present in its queues. As a |
|
* conclusion, the cumulative extra delay caused by injection can be |
|
* easily evaluated by just comparing the total service time of first |
|
* requests with and without injection. |
|
* |
|
* The limit-update algorithm works as follows. On the arrival of a |
|
* first request of bfqq, the algorithm measures the total time of the |
|
* request only if one of the three cases below holds, and, for each |
|
* case, it updates the limit as described below: |
|
* |
|
* (1) If there is no in-flight request. This gives a baseline for the |
|
* total service time of the requests of bfqq. If the baseline has |
|
* not been computed yet, then, after computing it, the limit is |
|
* set to 1, to start boosting throughput, and to prepare the |
|
* ground for the next case. If the baseline has already been |
|
* computed, then it is updated, in case it results to be lower |
|
* than the previous value. |
|
* |
|
* (2) If the limit is higher than 0 and there are in-flight |
|
* requests. By comparing the total service time in this case with |
|
* the above baseline, it is possible to know at which extent the |
|
* current value of the limit is inflating the total service |
|
* time. If the inflation is below a certain threshold, then bfqq |
|
* is assumed to be suffering from no perceivable loss of its |
|
* service guarantees, and the limit is even tentatively |
|
* increased. If the inflation is above the threshold, then the |
|
* limit is decreased. Due to the lack of any hysteresis, this |
|
* logic makes the limit oscillate even in steady workload |
|
* conditions. Yet we opted for it, because it is fast in reaching |
|
* the best value for the limit, as a function of the current I/O |
|
* workload. To reduce oscillations, this step is disabled for a |
|
* short time interval after the limit happens to be decreased. |
|
* |
|
* (3) Periodically, after resetting the limit, to make sure that the |
|
* limit eventually drops in case the workload changes. This is |
|
* needed because, after the limit has gone safely up for a |
|
* certain workload, it is impossible to guess whether the |
|
* baseline total service time may have changed, without measuring |
|
* it again without injection. A more effective version of this |
|
* step might be to just sample the baseline, by interrupting |
|
* injection only once, and then to reset/lower the limit only if |
|
* the total service time with the current limit does happen to be |
|
* too large. |
|
* |
|
* More details on each step are provided in the comments on the |
|
* pieces of code that implement these steps: the branch handling the |
|
* transition from empty to non empty in bfq_add_request(), the branch |
|
* handling injection in bfq_select_queue(), and the function |
|
* bfq_choose_bfqq_for_injection(). These comments also explain some |
|
* exceptions, made by the injection mechanism in some special cases. |
|
*/ |
|
static void bfq_update_inject_limit(struct bfq_data *bfqd, |
|
struct bfq_queue *bfqq) |
|
{ |
|
u64 tot_time_ns = ktime_get_ns() - bfqd->last_empty_occupied_ns; |
|
unsigned int old_limit = bfqq->inject_limit; |
|
|
|
if (bfqq->last_serv_time_ns > 0 && bfqd->rqs_injected) { |
|
u64 threshold = (bfqq->last_serv_time_ns * 3)>>1; |
|
|
|
if (tot_time_ns >= threshold && old_limit > 0) { |
|
bfqq->inject_limit--; |
|
bfqq->decrease_time_jif = jiffies; |
|
} else if (tot_time_ns < threshold && |
|
old_limit <= bfqd->max_rq_in_driver) |
|
bfqq->inject_limit++; |
|
} |
|
|
|
/* |
|
* Either we still have to compute the base value for the |
|
* total service time, and there seem to be the right |
|
* conditions to do it, or we can lower the last base value |
|
* computed. |
|
* |
|
* NOTE: (bfqd->rq_in_driver == 1) means that there is no I/O |
|
* request in flight, because this function is in the code |
|
* path that handles the completion of a request of bfqq, and, |
|
* in particular, this function is executed before |
|
* bfqd->rq_in_driver is decremented in such a code path. |
|
*/ |
|
if ((bfqq->last_serv_time_ns == 0 && bfqd->rq_in_driver == 1) || |
|
tot_time_ns < bfqq->last_serv_time_ns) { |
|
if (bfqq->last_serv_time_ns == 0) { |
|
/* |
|
* Now we certainly have a base value: make sure we |
|
* start trying injection. |
|
*/ |
|
bfqq->inject_limit = max_t(unsigned int, 1, old_limit); |
|
} |
|
bfqq->last_serv_time_ns = tot_time_ns; |
|
} else if (!bfqd->rqs_injected && bfqd->rq_in_driver == 1) |
|
/* |
|
* No I/O injected and no request still in service in |
|
* the drive: these are the exact conditions for |
|
* computing the base value of the total service time |
|
* for bfqq. So let's update this value, because it is |
|
* rather variable. For example, it varies if the size |
|
* or the spatial locality of the I/O requests in bfqq |
|
* change. |
|
*/ |
|
bfqq->last_serv_time_ns = tot_time_ns; |
|
|
|
|
|
/* update complete, not waiting for any request completion any longer */ |
|
bfqd->waited_rq = NULL; |
|
bfqd->rqs_injected = false; |
|
} |
|
|
|
/* |
|
* Handle either a requeue or a finish for rq. The things to do are |
|
* the same in both cases: all references to rq are to be dropped. In |
|
* particular, rq is considered completed from the point of view of |
|
* the scheduler. |
|
*/ |
|
static void bfq_finish_requeue_request(struct request *rq) |
|
{ |
|
struct bfq_queue *bfqq = RQ_BFQQ(rq); |
|
struct bfq_data *bfqd; |
|
unsigned long flags; |
|
|
|
/* |
|
* rq either is not associated with any icq, or is an already |
|
* requeued request that has not (yet) been re-inserted into |
|
* a bfq_queue. |
|
*/ |
|
if (!rq->elv.icq || !bfqq) |
|
return; |
|
|
|
bfqd = bfqq->bfqd; |
|
|
|
if (rq->rq_flags & RQF_STARTED) |
|
bfqg_stats_update_completion(bfqq_group(bfqq), |
|
rq->start_time_ns, |
|
rq->io_start_time_ns, |
|
rq->cmd_flags); |
|
|
|
spin_lock_irqsave(&bfqd->lock, flags); |
|
if (likely(rq->rq_flags & RQF_STARTED)) { |
|
if (rq == bfqd->waited_rq) |
|
bfq_update_inject_limit(bfqd, bfqq); |
|
|
|
bfq_completed_request(bfqq, bfqd); |
|
} |
|
bfq_finish_requeue_request_body(bfqq); |
|
spin_unlock_irqrestore(&bfqd->lock, flags); |
|
|
|
/* |
|
* Reset private fields. In case of a requeue, this allows |
|
* this function to correctly do nothing if it is spuriously |
|
* invoked again on this same request (see the check at the |
|
* beginning of the function). Probably, a better general |
|
* design would be to prevent blk-mq from invoking the requeue |
|
* or finish hooks of an elevator, for a request that is not |
|
* referred by that elevator. |
|
* |
|
* Resetting the following fields would break the |
|
* request-insertion logic if rq is re-inserted into a bfq |
|
* internal queue, without a re-preparation. Here we assume |
|
* that re-insertions of requeued requests, without |
|
* re-preparation, can happen only for pass_through or at_head |
|
* requests (which are not re-inserted into bfq internal |
|
* queues). |
|
*/ |
|
rq->elv.priv[0] = NULL; |
|
rq->elv.priv[1] = NULL; |
|
} |
|
|
|
/* |
|
* Removes the association between the current task and bfqq, assuming |
|
* that bic points to the bfq iocontext of the task. |
|
* Returns NULL if a new bfqq should be allocated, or the old bfqq if this |
|
* was the last process referring to that bfqq. |
|
*/ |
|
static struct bfq_queue * |
|
bfq_split_bfqq(struct bfq_io_cq *bic, struct bfq_queue *bfqq) |
|
{ |
|
bfq_log_bfqq(bfqq->bfqd, bfqq, "splitting queue"); |
|
|
|
if (bfqq_process_refs(bfqq) == 1) { |
|
bfqq->pid = current->pid; |
|
bfq_clear_bfqq_coop(bfqq); |
|
bfq_clear_bfqq_split_coop(bfqq); |
|
return bfqq; |
|
} |
|
|
|
bic_set_bfqq(bic, NULL, 1); |
|
|
|
bfq_put_cooperator(bfqq); |
|
|
|
bfq_release_process_ref(bfqq->bfqd, bfqq); |
|
return NULL; |
|
} |
|
|
|
static struct bfq_queue *bfq_get_bfqq_handle_split(struct bfq_data *bfqd, |
|
struct bfq_io_cq *bic, |
|
struct bio *bio, |
|
bool split, bool is_sync, |
|
bool *new_queue) |
|
{ |
|
struct bfq_queue *bfqq = bic_to_bfqq(bic, is_sync); |
|
|
|
if (likely(bfqq && bfqq != &bfqd->oom_bfqq)) |
|
return bfqq; |
|
|
|
if (new_queue) |
|
*new_queue = true; |
|
|
|
if (bfqq) |
|
bfq_put_queue(bfqq); |
|
bfqq = bfq_get_queue(bfqd, bio, is_sync, bic, split); |
|
|
|
bic_set_bfqq(bic, bfqq, is_sync); |
|
if (split && is_sync) { |
|
if ((bic->was_in_burst_list && bfqd->large_burst) || |
|
bic->saved_in_large_burst) |
|
bfq_mark_bfqq_in_large_burst(bfqq); |
|
else { |
|
bfq_clear_bfqq_in_large_burst(bfqq); |
|
if (bic->was_in_burst_list) |
|
/* |
|
* If bfqq was in the current |
|
* burst list before being |
|
* merged, then we have to add |
|
* it back. And we do not need |
|
* to increase burst_size, as |
|
* we did not decrement |
|
* burst_size when we removed |
|
* bfqq from the burst list as |
|
* a consequence of a merge |
|
* (see comments in |
|
* bfq_put_queue). In this |
|
* respect, it would be rather |
|
* costly to know whether the |
|
* current burst list is still |
|
* the same burst list from |
|
* which bfqq was removed on |
|
* the merge. To avoid this |
|
* cost, if bfqq was in a |
|
* burst list, then we add |
|
* bfqq to the current burst |
|
* list without any further |
|
* check. This can cause |
|
* inappropriate insertions, |
|
* but rarely enough to not |
|
* harm the detection of large |
|
* bursts significantly. |
|
*/ |
|
hlist_add_head(&bfqq->burst_list_node, |
|
&bfqd->burst_list); |
|
} |
|
bfqq->split_time = jiffies; |
|
} |
|
|
|
return bfqq; |
|
} |
|
|
|
/* |
|
* Only reset private fields. The actual request preparation will be |
|
* performed by bfq_init_rq, when rq is either inserted or merged. See |
|
* comments on bfq_init_rq for the reason behind this delayed |
|
* preparation. |
|
*/ |
|
static void bfq_prepare_request(struct request *rq) |
|
{ |
|
/* |
|
* Regardless of whether we have an icq attached, we have to |
|
* clear the scheduler pointers, as they might point to |
|
* previously allocated bic/bfqq structs. |
|
*/ |
|
rq->elv.priv[0] = rq->elv.priv[1] = NULL; |
|
} |
|
|
|
/* |
|
* If needed, init rq, allocate bfq data structures associated with |
|
* rq, and increment reference counters in the destination bfq_queue |
|
* for rq. Return the destination bfq_queue for rq, or NULL is rq is |
|
* not associated with any bfq_queue. |
|
* |
|
* This function is invoked by the functions that perform rq insertion |
|
* or merging. One may have expected the above preparation operations |
|
* to be performed in bfq_prepare_request, and not delayed to when rq |
|
* is inserted or merged. The rationale behind this delayed |
|
* preparation is that, after the prepare_request hook is invoked for |
|
* rq, rq may still be transformed into a request with no icq, i.e., a |
|
* request not associated with any queue. No bfq hook is invoked to |
|
* signal this transformation. As a consequence, should these |
|
* preparation operations be performed when the prepare_request hook |
|
* is invoked, and should rq be transformed one moment later, bfq |
|
* would end up in an inconsistent state, because it would have |
|
* incremented some queue counters for an rq destined to |
|
* transformation, without any chance to correctly lower these |
|
* counters back. In contrast, no transformation can still happen for |
|
* rq after rq has been inserted or merged. So, it is safe to execute |
|
* these preparation operations when rq is finally inserted or merged. |
|
*/ |
|
static struct bfq_queue *bfq_init_rq(struct request *rq) |
|
{ |
|
struct request_queue *q = rq->q; |
|
struct bio *bio = rq->bio; |
|
struct bfq_data *bfqd = q->elevator->elevator_data; |
|
struct bfq_io_cq *bic; |
|
const int is_sync = rq_is_sync(rq); |
|
struct bfq_queue *bfqq; |
|
bool new_queue = false; |
|
bool bfqq_already_existing = false, split = false; |
|
|
|
if (unlikely(!rq->elv.icq)) |
|
return NULL; |
|
|
|
/* |
|
* Assuming that elv.priv[1] is set only if everything is set |
|
* for this rq. This holds true, because this function is |
|
* invoked only for insertion or merging, and, after such |
|
* events, a request cannot be manipulated any longer before |
|
* being removed from bfq. |
|
*/ |
|
if (rq->elv.priv[1]) |
|
return rq->elv.priv[1]; |
|
|
|
bic = icq_to_bic(rq->elv.icq); |
|
|
|
bfq_check_ioprio_change(bic, bio); |
|
|
|
bfq_bic_update_cgroup(bic, bio); |
|
|
|
bfqq = bfq_get_bfqq_handle_split(bfqd, bic, bio, false, is_sync, |
|
&new_queue); |
|
|
|
if (likely(!new_queue)) { |
|
/* If the queue was seeky for too long, break it apart. */ |
|
if (bfq_bfqq_coop(bfqq) && bfq_bfqq_split_coop(bfqq) && |
|
!bic->stably_merged) { |
|
struct bfq_queue *old_bfqq = bfqq; |
|
|
|
/* Update bic before losing reference to bfqq */ |
|
if (bfq_bfqq_in_large_burst(bfqq)) |
|
bic->saved_in_large_burst = true; |
|
|
|
bfqq = bfq_split_bfqq(bic, bfqq); |
|
split = true; |
|
|
|
if (!bfqq) { |
|
bfqq = bfq_get_bfqq_handle_split(bfqd, bic, bio, |
|
true, is_sync, |
|
NULL); |
|
bfqq->waker_bfqq = old_bfqq->waker_bfqq; |
|
bfqq->tentative_waker_bfqq = NULL; |
|
|
|
/* |
|
* If the waker queue disappears, then |
|
* new_bfqq->waker_bfqq must be |
|
* reset. So insert new_bfqq into the |
|
* woken_list of the waker. See |
|
* bfq_check_waker for details. |
|
*/ |
|
if (bfqq->waker_bfqq) |
|
hlist_add_head(&bfqq->woken_list_node, |
|
&bfqq->waker_bfqq->woken_list); |
|
} else |
|
bfqq_already_existing = true; |
|
} |
|
} |
|
|
|
bfqq->allocated++; |
|
bfqq->ref++; |
|
bfq_log_bfqq(bfqd, bfqq, "get_request %p: bfqq %p, %d", |
|
rq, bfqq, bfqq->ref); |
|
|
|
rq->elv.priv[0] = bic; |
|
rq->elv.priv[1] = bfqq; |
|
|
|
/* |
|
* If a bfq_queue has only one process reference, it is owned |
|
* by only this bic: we can then set bfqq->bic = bic. in |
|
* addition, if the queue has also just been split, we have to |
|
* resume its state. |
|
*/ |
|
if (likely(bfqq != &bfqd->oom_bfqq) && bfqq_process_refs(bfqq) == 1) { |
|
bfqq->bic = bic; |
|
if (split) { |
|
/* |
|
* The queue has just been split from a shared |
|
* queue: restore the idle window and the |
|
* possible weight raising period. |
|
*/ |
|
bfq_bfqq_resume_state(bfqq, bfqd, bic, |
|
bfqq_already_existing); |
|
} |
|
} |
|
|
|
/* |
|
* Consider bfqq as possibly belonging to a burst of newly |
|
* created queues only if: |
|
* 1) A burst is actually happening (bfqd->burst_size > 0) |
|
* or |
|
* 2) There is no other active queue. In fact, if, in |
|
* contrast, there are active queues not belonging to the |
|
* possible burst bfqq may belong to, then there is no gain |
|
* in considering bfqq as belonging to a burst, and |
|
* therefore in not weight-raising bfqq. See comments on |
|
* bfq_handle_burst(). |
|
* |
|
* This filtering also helps eliminating false positives, |
|
* occurring when bfqq does not belong to an actual large |
|
* burst, but some background task (e.g., a service) happens |
|
* to trigger the creation of new queues very close to when |
|
* bfqq and its possible companion queues are created. See |
|
* comments on bfq_handle_burst() for further details also on |
|
* this issue. |
|
*/ |
|
if (unlikely(bfq_bfqq_just_created(bfqq) && |
|
(bfqd->burst_size > 0 || |
|
bfq_tot_busy_queues(bfqd) == 0))) |
|
bfq_handle_burst(bfqd, bfqq); |
|
|
|
return bfqq; |
|
} |
|
|
|
static void |
|
bfq_idle_slice_timer_body(struct bfq_data *bfqd, struct bfq_queue *bfqq) |
|
{ |
|
enum bfqq_expiration reason; |
|
unsigned long flags; |
|
|
|
spin_lock_irqsave(&bfqd->lock, flags); |
|
|
|
/* |
|
* Considering that bfqq may be in race, we should firstly check |
|
* whether bfqq is in service before doing something on it. If |
|
* the bfqq in race is not in service, it has already been expired |
|
* through __bfq_bfqq_expire func and its wait_request flags has |
|
* been cleared in __bfq_bfqd_reset_in_service func. |
|
*/ |
|
if (bfqq != bfqd->in_service_queue) { |
|
spin_unlock_irqrestore(&bfqd->lock, flags); |
|
return; |
|
} |
|
|
|
bfq_clear_bfqq_wait_request(bfqq); |
|
|
|
if (bfq_bfqq_budget_timeout(bfqq)) |
|
/* |
|
* Also here the queue can be safely expired |
|
* for budget timeout without wasting |
|
* guarantees |
|
*/ |
|
reason = BFQQE_BUDGET_TIMEOUT; |
|
else if (bfqq->queued[0] == 0 && bfqq->queued[1] == 0) |
|
/* |
|
* The queue may not be empty upon timer expiration, |
|
* because we may not disable the timer when the |
|
* first request of the in-service queue arrives |
|
* during disk idling. |
|
*/ |
|
reason = BFQQE_TOO_IDLE; |
|
else |
|
goto schedule_dispatch; |
|
|
|
bfq_bfqq_expire(bfqd, bfqq, true, reason); |
|
|
|
schedule_dispatch: |
|
spin_unlock_irqrestore(&bfqd->lock, flags); |
|
bfq_schedule_dispatch(bfqd); |
|
} |
|
|
|
/* |
|
* Handler of the expiration of the timer running if the in-service queue |
|
* is idling inside its time slice. |
|
*/ |
|
static enum hrtimer_restart bfq_idle_slice_timer(struct hrtimer *timer) |
|
{ |
|
struct bfq_data *bfqd = container_of(timer, struct bfq_data, |
|
idle_slice_timer); |
|
struct bfq_queue *bfqq = bfqd->in_service_queue; |
|
|
|
/* |
|
* Theoretical race here: the in-service queue can be NULL or |
|
* different from the queue that was idling if a new request |
|
* arrives for the current queue and there is a full dispatch |
|
* cycle that changes the in-service queue. This can hardly |
|
* happen, but in the worst case we just expire a queue too |
|
* early. |
|
*/ |
|
if (bfqq) |
|
bfq_idle_slice_timer_body(bfqd, bfqq); |
|
|
|
return HRTIMER_NORESTART; |
|
} |
|
|
|
static void __bfq_put_async_bfqq(struct bfq_data *bfqd, |
|
struct bfq_queue **bfqq_ptr) |
|
{ |
|
struct bfq_queue *bfqq = *bfqq_ptr; |
|
|
|
bfq_log(bfqd, "put_async_bfqq: %p", bfqq); |
|
if (bfqq) { |
|
bfq_bfqq_move(bfqd, bfqq, bfqd->root_group); |
|
|
|
bfq_log_bfqq(bfqd, bfqq, "put_async_bfqq: putting %p, %d", |
|
bfqq, bfqq->ref); |
|
bfq_put_queue(bfqq); |
|
*bfqq_ptr = NULL; |
|
} |
|
} |
|
|
|
/* |
|
* Release all the bfqg references to its async queues. If we are |
|
* deallocating the group these queues may still contain requests, so |
|
* we reparent them to the root cgroup (i.e., the only one that will |
|
* exist for sure until all the requests on a device are gone). |
|
*/ |
|
void bfq_put_async_queues(struct bfq_data *bfqd, struct bfq_group *bfqg) |
|
{ |
|
int i, j; |
|
|
|
for (i = 0; i < 2; i++) |
|
for (j = 0; j < IOPRIO_NR_LEVELS; j++) |
|
__bfq_put_async_bfqq(bfqd, &bfqg->async_bfqq[i][j]); |
|
|
|
__bfq_put_async_bfqq(bfqd, &bfqg->async_idle_bfqq); |
|
} |
|
|
|
/* |
|
* See the comments on bfq_limit_depth for the purpose of |
|
* the depths set in the function. Return minimum shallow depth we'll use. |
|
*/ |
|
static unsigned int bfq_update_depths(struct bfq_data *bfqd, |
|
struct sbitmap_queue *bt) |
|
{ |
|
unsigned int i, j, min_shallow = UINT_MAX; |
|
|
|
/* |
|
* In-word depths if no bfq_queue is being weight-raised: |
|
* leaving 25% of tags only for sync reads. |
|
* |
|
* In next formulas, right-shift the value |
|
* (1U<<bt->sb.shift), instead of computing directly |
|
* (1U<<(bt->sb.shift - something)), to be robust against |
|
* any possible value of bt->sb.shift, without having to |
|
* limit 'something'. |
|
*/ |
|
/* no more than 50% of tags for async I/O */ |
|
bfqd->word_depths[0][0] = max((1U << bt->sb.shift) >> 1, 1U); |
|
/* |
|
* no more than 75% of tags for sync writes (25% extra tags |
|
* w.r.t. async I/O, to prevent async I/O from starving sync |
|
* writes) |
|
*/ |
|
bfqd->word_depths[0][1] = max(((1U << bt->sb.shift) * 3) >> 2, 1U); |
|
|
|
/* |
|
* In-word depths in case some bfq_queue is being weight- |
|
* raised: leaving ~63% of tags for sync reads. This is the |
|
* highest percentage for which, in our tests, application |
|
* start-up times didn't suffer from any regression due to tag |
|
* shortage. |
|
*/ |
|
/* no more than ~18% of tags for async I/O */ |
|
bfqd->word_depths[1][0] = max(((1U << bt->sb.shift) * 3) >> 4, 1U); |
|
/* no more than ~37% of tags for sync writes (~20% extra tags) */ |
|
bfqd->word_depths[1][1] = max(((1U << bt->sb.shift) * 6) >> 4, 1U); |
|
|
|
for (i = 0; i < 2; i++) |
|
for (j = 0; j < 2; j++) |
|
min_shallow = min(min_shallow, bfqd->word_depths[i][j]); |
|
|
|
return min_shallow; |
|
} |
|
|
|
static void bfq_depth_updated(struct blk_mq_hw_ctx *hctx) |
|
{ |
|
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data; |
|
struct blk_mq_tags *tags = hctx->sched_tags; |
|
unsigned int min_shallow; |
|
|
|
min_shallow = bfq_update_depths(bfqd, tags->bitmap_tags); |
|
sbitmap_queue_min_shallow_depth(tags->bitmap_tags, min_shallow); |
|
} |
|
|
|
static int bfq_init_hctx(struct blk_mq_hw_ctx *hctx, unsigned int index) |
|
{ |
|
bfq_depth_updated(hctx); |
|
return 0; |
|
} |
|
|
|
static void bfq_exit_queue(struct elevator_queue *e) |
|
{ |
|
struct bfq_data *bfqd = e->elevator_data; |
|
struct bfq_queue *bfqq, *n; |
|
|
|
hrtimer_cancel(&bfqd->idle_slice_timer); |
|
|
|
spin_lock_irq(&bfqd->lock); |
|
list_for_each_entry_safe(bfqq, n, &bfqd->idle_list, bfqq_list) |
|
bfq_deactivate_bfqq(bfqd, bfqq, false, false); |
|
spin_unlock_irq(&bfqd->lock); |
|
|
|
hrtimer_cancel(&bfqd->idle_slice_timer); |
|
|
|
/* release oom-queue reference to root group */ |
|
bfqg_and_blkg_put(bfqd->root_group); |
|
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED |
|
blkcg_deactivate_policy(bfqd->queue, &blkcg_policy_bfq); |
|
#else |
|
spin_lock_irq(&bfqd->lock); |
|
bfq_put_async_queues(bfqd, bfqd->root_group); |
|
kfree(bfqd->root_group); |
|
spin_unlock_irq(&bfqd->lock); |
|
#endif |
|
|
|
kfree(bfqd); |
|
} |
|
|
|
static void bfq_init_root_group(struct bfq_group *root_group, |
|
struct bfq_data *bfqd) |
|
{ |
|
int i; |
|
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED |
|
root_group->entity.parent = NULL; |
|
root_group->my_entity = NULL; |
|
root_group->bfqd = bfqd; |
|
#endif |
|
root_group->rq_pos_tree = RB_ROOT; |
|
for (i = 0; i < BFQ_IOPRIO_CLASSES; i++) |
|
root_group->sched_data.service_tree[i] = BFQ_SERVICE_TREE_INIT; |
|
root_group->sched_data.bfq_class_idle_last_service = jiffies; |
|
} |
|
|
|
static int bfq_init_queue(struct request_queue *q, struct elevator_type *e) |
|
{ |
|
struct bfq_data *bfqd; |
|
struct elevator_queue *eq; |
|
|
|
eq = elevator_alloc(q, e); |
|
if (!eq) |
|
return -ENOMEM; |
|
|
|
bfqd = kzalloc_node(sizeof(*bfqd), GFP_KERNEL, q->node); |
|
if (!bfqd) { |
|
kobject_put(&eq->kobj); |
|
return -ENOMEM; |
|
} |
|
eq->elevator_data = bfqd; |
|
|
|
spin_lock_irq(&q->queue_lock); |
|
q->elevator = eq; |
|
spin_unlock_irq(&q->queue_lock); |
|
|
|
/* |
|
* Our fallback bfqq if bfq_find_alloc_queue() runs into OOM issues. |
|
* Grab a permanent reference to it, so that the normal code flow |
|
* will not attempt to free it. |
|
*/ |
|
bfq_init_bfqq(bfqd, &bfqd->oom_bfqq, NULL, 1, 0); |
|
bfqd->oom_bfqq.ref++; |
|
bfqd->oom_bfqq.new_ioprio = BFQ_DEFAULT_QUEUE_IOPRIO; |
|
bfqd->oom_bfqq.new_ioprio_class = IOPRIO_CLASS_BE; |
|
bfqd->oom_bfqq.entity.new_weight = |
|
bfq_ioprio_to_weight(bfqd->oom_bfqq.new_ioprio); |
|
|
|
/* oom_bfqq does not participate to bursts */ |
|
bfq_clear_bfqq_just_created(&bfqd->oom_bfqq); |
|
|
|
/* |
|
* Trigger weight initialization, according to ioprio, at the |
|
* oom_bfqq's first activation. The oom_bfqq's ioprio and ioprio |
|
* class won't be changed any more. |
|
*/ |
|
bfqd->oom_bfqq.entity.prio_changed = 1; |
|
|
|
bfqd->queue = q; |
|
|
|
INIT_LIST_HEAD(&bfqd->dispatch); |
|
|
|
hrtimer_init(&bfqd->idle_slice_timer, CLOCK_MONOTONIC, |
|
HRTIMER_MODE_REL); |
|
bfqd->idle_slice_timer.function = bfq_idle_slice_timer; |
|
|
|
bfqd->queue_weights_tree = RB_ROOT_CACHED; |
|
bfqd->num_groups_with_pending_reqs = 0; |
|
|
|
INIT_LIST_HEAD(&bfqd->active_list); |
|
INIT_LIST_HEAD(&bfqd->idle_list); |
|
INIT_HLIST_HEAD(&bfqd->burst_list); |
|
|
|
bfqd->hw_tag = -1; |
|
bfqd->nonrot_with_queueing = blk_queue_nonrot(bfqd->queue); |
|
|
|
bfqd->bfq_max_budget = bfq_default_max_budget; |
|
|
|
bfqd->bfq_fifo_expire[0] = bfq_fifo_expire[0]; |
|
bfqd->bfq_fifo_expire[1] = bfq_fifo_expire[1]; |
|
bfqd->bfq_back_max = bfq_back_max; |
|
bfqd->bfq_back_penalty = bfq_back_penalty; |
|
bfqd->bfq_slice_idle = bfq_slice_idle; |
|
bfqd->bfq_timeout = bfq_timeout; |
|
|
|
bfqd->bfq_large_burst_thresh = 8; |
|
bfqd->bfq_burst_interval = msecs_to_jiffies(180); |
|
|
|
bfqd->low_latency = true; |
|
|
|
/* |
|
* Trade-off between responsiveness and fairness. |
|
*/ |
|
bfqd->bfq_wr_coeff = 30; |
|
bfqd->bfq_wr_rt_max_time = msecs_to_jiffies(300); |
|
bfqd->bfq_wr_max_time = 0; |
|
bfqd->bfq_wr_min_idle_time = msecs_to_jiffies(2000); |
|
bfqd->bfq_wr_min_inter_arr_async = msecs_to_jiffies(500); |
|
bfqd->bfq_wr_max_softrt_rate = 7000; /* |
|
* Approximate rate required |
|
* to playback or record a |
|
* high-definition compressed |
|
* video. |
|
*/ |
|
bfqd->wr_busy_queues = 0; |
|
|
|
/* |
|
* Begin by assuming, optimistically, that the device peak |
|
* rate is equal to 2/3 of the highest reference rate. |
|
*/ |
|
bfqd->rate_dur_prod = ref_rate[blk_queue_nonrot(bfqd->queue)] * |
|
ref_wr_duration[blk_queue_nonrot(bfqd->queue)]; |
|
bfqd->peak_rate = ref_rate[blk_queue_nonrot(bfqd->queue)] * 2 / 3; |
|
|
|
spin_lock_init(&bfqd->lock); |
|
|
|
/* |
|
* The invocation of the next bfq_create_group_hierarchy |
|
* function is the head of a chain of function calls |
|
* (bfq_create_group_hierarchy->blkcg_activate_policy-> |
|
* blk_mq_freeze_queue) that may lead to the invocation of the |
|
* has_work hook function. For this reason, |
|
* bfq_create_group_hierarchy is invoked only after all |
|
* scheduler data has been initialized, apart from the fields |
|
* that can be initialized only after invoking |
|
* bfq_create_group_hierarchy. This, in particular, enables |
|
* has_work to correctly return false. Of course, to avoid |
|
* other inconsistencies, the blk-mq stack must then refrain |
|
* from invoking further scheduler hooks before this init |
|
* function is finished. |
|
*/ |
|
bfqd->root_group = bfq_create_group_hierarchy(bfqd, q->node); |
|
if (!bfqd->root_group) |
|
goto out_free; |
|
bfq_init_root_group(bfqd->root_group, bfqd); |
|
bfq_init_entity(&bfqd->oom_bfqq.entity, bfqd->root_group); |
|
|
|
wbt_disable_default(q); |
|
return 0; |
|
|
|
out_free: |
|
kfree(bfqd); |
|
kobject_put(&eq->kobj); |
|
return -ENOMEM; |
|
} |
|
|
|
static void bfq_slab_kill(void) |
|
{ |
|
kmem_cache_destroy(bfq_pool); |
|
} |
|
|
|
static int __init bfq_slab_setup(void) |
|
{ |
|
bfq_pool = KMEM_CACHE(bfq_queue, 0); |
|
if (!bfq_pool) |
|
return -ENOMEM; |
|
return 0; |
|
} |
|
|
|
static ssize_t bfq_var_show(unsigned int var, char *page) |
|
{ |
|
return sprintf(page, "%u\n", var); |
|
} |
|
|
|
static int bfq_var_store(unsigned long *var, const char *page) |
|
{ |
|
unsigned long new_val; |
|
int ret = kstrtoul(page, 10, &new_val); |
|
|
|
if (ret) |
|
return ret; |
|
*var = new_val; |
|
return 0; |
|
} |
|
|
|
#define SHOW_FUNCTION(__FUNC, __VAR, __CONV) \ |
|
static ssize_t __FUNC(struct elevator_queue *e, char *page) \ |
|
{ \ |
|
struct bfq_data *bfqd = e->elevator_data; \ |
|
u64 __data = __VAR; \ |
|
if (__CONV == 1) \ |
|
__data = jiffies_to_msecs(__data); \ |
|
else if (__CONV == 2) \ |
|
__data = div_u64(__data, NSEC_PER_MSEC); \ |
|
return bfq_var_show(__data, (page)); \ |
|
} |
|
SHOW_FUNCTION(bfq_fifo_expire_sync_show, bfqd->bfq_fifo_expire[1], 2); |
|
SHOW_FUNCTION(bfq_fifo_expire_async_show, bfqd->bfq_fifo_expire[0], 2); |
|
SHOW_FUNCTION(bfq_back_seek_max_show, bfqd->bfq_back_max, 0); |
|
SHOW_FUNCTION(bfq_back_seek_penalty_show, bfqd->bfq_back_penalty, 0); |
|
SHOW_FUNCTION(bfq_slice_idle_show, bfqd->bfq_slice_idle, 2); |
|
SHOW_FUNCTION(bfq_max_budget_show, bfqd->bfq_user_max_budget, 0); |
|
SHOW_FUNCTION(bfq_timeout_sync_show, bfqd->bfq_timeout, 1); |
|
SHOW_FUNCTION(bfq_strict_guarantees_show, bfqd->strict_guarantees, 0); |
|
SHOW_FUNCTION(bfq_low_latency_show, bfqd->low_latency, 0); |
|
#undef SHOW_FUNCTION |
|
|
|
#define USEC_SHOW_FUNCTION(__FUNC, __VAR) \ |
|
static ssize_t __FUNC(struct elevator_queue *e, char *page) \ |
|
{ \ |
|
struct bfq_data *bfqd = e->elevator_data; \ |
|
u64 __data = __VAR; \ |
|
__data = div_u64(__data, NSEC_PER_USEC); \ |
|
return bfq_var_show(__data, (page)); \ |
|
} |
|
USEC_SHOW_FUNCTION(bfq_slice_idle_us_show, bfqd->bfq_slice_idle); |
|
#undef USEC_SHOW_FUNCTION |
|
|
|
#define STORE_FUNCTION(__FUNC, __PTR, MIN, MAX, __CONV) \ |
|
static ssize_t \ |
|
__FUNC(struct elevator_queue *e, const char *page, size_t count) \ |
|
{ \ |
|
struct bfq_data *bfqd = e->elevator_data; \ |
|
unsigned long __data, __min = (MIN), __max = (MAX); \ |
|
int ret; \ |
|
\ |
|
ret = bfq_var_store(&__data, (page)); \ |
|
if (ret) \ |
|
return ret; \ |
|
if (__data < __min) \ |
|
__data = __min; \ |
|
else if (__data > __max) \ |
|
__data = __max; \ |
|
if (__CONV == 1) \ |
|
*(__PTR) = msecs_to_jiffies(__data); \ |
|
else if (__CONV == 2) \ |
|
*(__PTR) = (u64)__data * NSEC_PER_MSEC; \ |
|
else \ |
|
*(__PTR) = __data; \ |
|
return count; \ |
|
} |
|
STORE_FUNCTION(bfq_fifo_expire_sync_store, &bfqd->bfq_fifo_expire[1], 1, |
|
INT_MAX, 2); |
|
STORE_FUNCTION(bfq_fifo_expire_async_store, &bfqd->bfq_fifo_expire[0], 1, |
|
INT_MAX, 2); |
|
STORE_FUNCTION(bfq_back_seek_max_store, &bfqd->bfq_back_max, 0, INT_MAX, 0); |
|
STORE_FUNCTION(bfq_back_seek_penalty_store, &bfqd->bfq_back_penalty, 1, |
|
INT_MAX, 0); |
|
STORE_FUNCTION(bfq_slice_idle_store, &bfqd->bfq_slice_idle, 0, INT_MAX, 2); |
|
#undef STORE_FUNCTION |
|
|
|
#define USEC_STORE_FUNCTION(__FUNC, __PTR, MIN, MAX) \ |
|
static ssize_t __FUNC(struct elevator_queue *e, const char *page, size_t count)\ |
|
{ \ |
|
struct bfq_data *bfqd = e->elevator_data; \ |
|
unsigned long __data, __min = (MIN), __max = (MAX); \ |
|
int ret; \ |
|
\ |
|
ret = bfq_var_store(&__data, (page)); \ |
|
if (ret) \ |
|
return ret; \ |
|
if (__data < __min) \ |
|
__data = __min; \ |
|
else if (__data > __max) \ |
|
__data = __max; \ |
|
*(__PTR) = (u64)__data * NSEC_PER_USEC; \ |
|
return count; \ |
|
} |
|
USEC_STORE_FUNCTION(bfq_slice_idle_us_store, &bfqd->bfq_slice_idle, 0, |
|
UINT_MAX); |
|
#undef USEC_STORE_FUNCTION |
|
|
|
static ssize_t bfq_max_budget_store(struct elevator_queue *e, |
|
const char *page, size_t count) |
|
{ |
|
struct bfq_data *bfqd = e->elevator_data; |
|
unsigned long __data; |
|
int ret; |
|
|
|
ret = bfq_var_store(&__data, (page)); |
|
if (ret) |
|
return ret; |
|
|
|
if (__data == 0) |
|
bfqd->bfq_max_budget = bfq_calc_max_budget(bfqd); |
|
else { |
|
if (__data > INT_MAX) |
|
__data = INT_MAX; |
|
bfqd->bfq_max_budget = __data; |
|
} |
|
|
|
bfqd->bfq_user_max_budget = __data; |
|
|
|
return count; |
|
} |
|
|
|
/* |
|
* Leaving this name to preserve name compatibility with cfq |
|
* parameters, but this timeout is used for both sync and async. |
|
*/ |
|
static ssize_t bfq_timeout_sync_store(struct elevator_queue *e, |
|
const char *page, size_t count) |
|
{ |
|
struct bfq_data *bfqd = e->elevator_data; |
|
unsigned long __data; |
|
int ret; |
|
|
|
ret = bfq_var_store(&__data, (page)); |
|
if (ret) |
|
return ret; |
|
|
|
if (__data < 1) |
|
__data = 1; |
|
else if (__data > INT_MAX) |
|
__data = INT_MAX; |
|
|
|
bfqd->bfq_timeout = msecs_to_jiffies(__data); |
|
if (bfqd->bfq_user_max_budget == 0) |
|
bfqd->bfq_max_budget = bfq_calc_max_budget(bfqd); |
|
|
|
return count; |
|
} |
|
|
|
static ssize_t bfq_strict_guarantees_store(struct elevator_queue *e, |
|
const char *page, size_t count) |
|
{ |
|
struct bfq_data *bfqd = e->elevator_data; |
|
unsigned long __data; |
|
int ret; |
|
|
|
ret = bfq_var_store(&__data, (page)); |
|
if (ret) |
|
return ret; |
|
|
|
if (__data > 1) |
|
__data = 1; |
|
if (!bfqd->strict_guarantees && __data == 1 |
|
&& bfqd->bfq_slice_idle < 8 * NSEC_PER_MSEC) |
|
bfqd->bfq_slice_idle = 8 * NSEC_PER_MSEC; |
|
|
|
bfqd->strict_guarantees = __data; |
|
|
|
return count; |
|
} |
|
|
|
static ssize_t bfq_low_latency_store(struct elevator_queue *e, |
|
const char *page, size_t count) |
|
{ |
|
struct bfq_data *bfqd = e->elevator_data; |
|
unsigned long __data; |
|
int ret; |
|
|
|
ret = bfq_var_store(&__data, (page)); |
|
if (ret) |
|
return ret; |
|
|
|
if (__data > 1) |
|
__data = 1; |
|
if (__data == 0 && bfqd->low_latency != 0) |
|
bfq_end_wr(bfqd); |
|
bfqd->low_latency = __data; |
|
|
|
return count; |
|
} |
|
|
|
#define BFQ_ATTR(name) \ |
|
__ATTR(name, 0644, bfq_##name##_show, bfq_##name##_store) |
|
|
|
static struct elv_fs_entry bfq_attrs[] = { |
|
BFQ_ATTR(fifo_expire_sync), |
|
BFQ_ATTR(fifo_expire_async), |
|
BFQ_ATTR(back_seek_max), |
|
BFQ_ATTR(back_seek_penalty), |
|
BFQ_ATTR(slice_idle), |
|
BFQ_ATTR(slice_idle_us), |
|
BFQ_ATTR(max_budget), |
|
BFQ_ATTR(timeout_sync), |
|
BFQ_ATTR(strict_guarantees), |
|
BFQ_ATTR(low_latency), |
|
__ATTR_NULL |
|
}; |
|
|
|
static struct elevator_type iosched_bfq_mq = { |
|
.ops = { |
|
.limit_depth = bfq_limit_depth, |
|
.prepare_request = bfq_prepare_request, |
|
.requeue_request = bfq_finish_requeue_request, |
|
.finish_request = bfq_finish_requeue_request, |
|
.exit_icq = bfq_exit_icq, |
|
.insert_requests = bfq_insert_requests, |
|
.dispatch_request = bfq_dispatch_request, |
|
.next_request = elv_rb_latter_request, |
|
.former_request = elv_rb_former_request, |
|
.allow_merge = bfq_allow_bio_merge, |
|
.bio_merge = bfq_bio_merge, |
|
.request_merge = bfq_request_merge, |
|
.requests_merged = bfq_requests_merged, |
|
.request_merged = bfq_request_merged, |
|
.has_work = bfq_has_work, |
|
.depth_updated = bfq_depth_updated, |
|
.init_hctx = bfq_init_hctx, |
|
.init_sched = bfq_init_queue, |
|
.exit_sched = bfq_exit_queue, |
|
}, |
|
|
|
.icq_size = sizeof(struct bfq_io_cq), |
|
.icq_align = __alignof__(struct bfq_io_cq), |
|
.elevator_attrs = bfq_attrs, |
|
.elevator_name = "bfq", |
|
.elevator_owner = THIS_MODULE, |
|
}; |
|
MODULE_ALIAS("bfq-iosched"); |
|
|
|
static int __init bfq_init(void) |
|
{ |
|
int ret; |
|
|
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED |
|
ret = blkcg_policy_register(&blkcg_policy_bfq); |
|
if (ret) |
|
return ret; |
|
#endif |
|
|
|
ret = -ENOMEM; |
|
if (bfq_slab_setup()) |
|
goto err_pol_unreg; |
|
|
|
/* |
|
* Times to load large popular applications for the typical |
|
* systems installed on the reference devices (see the |
|
* comments before the definition of the next |
|
* array). Actually, we use slightly lower values, as the |
|
* estimated peak rate tends to be smaller than the actual |
|
* peak rate. The reason for this last fact is that estimates |
|
* are computed over much shorter time intervals than the long |
|
* intervals typically used for benchmarking. Why? First, to |
|
* adapt more quickly to variations. Second, because an I/O |
|
* scheduler cannot rely on a peak-rate-evaluation workload to |
|
* be run for a long time. |
|
*/ |
|
ref_wr_duration[0] = msecs_to_jiffies(7000); /* actually 8 sec */ |
|
ref_wr_duration[1] = msecs_to_jiffies(2500); /* actually 3 sec */ |
|
|
|
ret = elv_register(&iosched_bfq_mq); |
|
if (ret) |
|
goto slab_kill; |
|
|
|
return 0; |
|
|
|
slab_kill: |
|
bfq_slab_kill(); |
|
err_pol_unreg: |
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED |
|
blkcg_policy_unregister(&blkcg_policy_bfq); |
|
#endif |
|
return ret; |
|
} |
|
|
|
static void __exit bfq_exit(void) |
|
{ |
|
elv_unregister(&iosched_bfq_mq); |
|
#ifdef CONFIG_BFQ_GROUP_IOSCHED |
|
blkcg_policy_unregister(&blkcg_policy_bfq); |
|
#endif |
|
bfq_slab_kill(); |
|
} |
|
|
|
module_init(bfq_init); |
|
module_exit(bfq_exit); |
|
|
|
MODULE_AUTHOR("Paolo Valente"); |
|
MODULE_LICENSE("GPL"); |
|
MODULE_DESCRIPTION("MQ Budget Fair Queueing I/O Scheduler");
|
|
|