| // SPDX-License-Identifier: GPL-2.0-or-later |
| /* |
| * Budget Fair Queueing (BFQ) I/O scheduler. |
| * |
| * Based on ideas and code from CFQ: |
| * Copyright (C) 2003 Jens Axboe <axboe@kernel.dk> |
| * |
| * Copyright (C) 2008 Fabio Checconi <fabio@gandalf.sssup.it> |
| * Paolo Valente <paolo.valente@unimore.it> |
| * |
| * Copyright (C) 2010 Paolo Valente <paolo.valente@unimore.it> |
| * Arianna Avanzini <avanzini@google.com> |
| * |
| * Copyright (C) 2017 Paolo Valente <paolo.valente@linaro.org> |
| * |
| * BFQ is a proportional-share I/O scheduler, with some extra |
| * low-latency capabilities. BFQ also supports full hierarchical |
| * scheduling through cgroups. Next paragraphs provide an introduction |
| * on BFQ inner workings. Details on BFQ benefits, usage and |
| * limitations can be found in Documentation/block/bfq-iosched.rst. |
| * |
| * BFQ is a proportional-share storage-I/O scheduling algorithm based |
| * on the slice-by-slice service scheme of CFQ. But BFQ assigns |
| * budgets, measured in number of sectors, to processes instead of |
| * time slices. The device is not granted to the in-service process |
| * for a given time slice, but until it has exhausted its assigned |
| * budget. This change from the time to the service domain enables BFQ |
| * to distribute the device throughput among processes as desired, |
| * without any distortion due to throughput fluctuations, or to device |
| * internal queueing. BFQ uses an ad hoc internal scheduler, called |
| * B-WF2Q+, to schedule processes according to their budgets. More |
| * precisely, BFQ schedules queues associated with processes. Each |
| * process/queue is assigned a user-configurable weight, and B-WF2Q+ |
| * guarantees that each queue receives a fraction of the throughput |
| * proportional to its weight. Thanks to the accurate policy of |
| * B-WF2Q+, BFQ can afford to assign high budgets to I/O-bound |
| * processes issuing sequential requests (to boost the throughput), |
| * and yet guarantee a low latency to interactive and soft real-time |
| * applications. |
| * |
| * In particular, to provide these low-latency guarantees, BFQ |
| * explicitly privileges the I/O of two classes of time-sensitive |
| * applications: interactive and soft real-time. In more detail, BFQ |
| * behaves this way if the low_latency parameter is set (default |
| * configuration). This feature enables BFQ to provide applications in |
| * these classes with a very low latency. |
| * |
| * To implement this feature, BFQ constantly tries to detect whether |
| * the I/O requests in a bfq_queue come from an interactive or a soft |
| * real-time application. For brevity, in these cases, the queue is |
| * said to be interactive or soft real-time. In both cases, BFQ |
| * privileges the service of the queue, over that of non-interactive |
| * and non-soft-real-time queues. This privileging is performed, |
| * mainly, by raising the weight of the queue. So, for brevity, we |
| * call just weight-raising periods the time periods during which a |
| * queue is privileged, because deemed interactive or soft real-time. |
| * |
| * The detection of soft real-time queues/applications is described in |
| * detail in the comments on the function |
| * bfq_bfqq_softrt_next_start. On the other hand, the detection of an |
| * interactive queue works as follows: a queue is deemed interactive |
| * if it is constantly non empty only for a limited time interval, |
| * after which it does become empty. The queue may be deemed |
| * interactive again (for a limited time), if it restarts being |
| * constantly non empty, provided that this happens only after the |
| * queue has remained empty for a given minimum idle time. |
| * |
| * By default, BFQ computes automatically the above maximum time |
| * interval, i.e., the time interval after which a constantly |
| * non-empty queue stops being deemed interactive. Since a queue is |
| * weight-raised while it is deemed interactive, this maximum time |
| * interval happens to coincide with the (maximum) duration of the |
| * weight-raising for interactive queues. |
| * |
| * Finally, BFQ also features additional heuristics for |
| * preserving both a low latency and a high throughput on NCQ-capable, |
| * rotational or flash-based devices, and to get the job done quickly |
| * for applications consisting in many I/O-bound processes. |
| * |
| * NOTE: if the main or only goal, with a given device, is to achieve |
| * the maximum-possible throughput at all times, then do switch off |
| * all low-latency heuristics for that device, by setting low_latency |
| * to 0. |
| * |
| * BFQ is described in [1], where also a reference to the initial, |
| * more theoretical paper on BFQ can be found. The interested reader |
| * can find in the latter paper full details on the main algorithm, as |
| * well as formulas of the guarantees and formal proofs of all the |
| * properties. With respect to the version of BFQ presented in these |
| * papers, this implementation adds a few more heuristics, such as the |
| * ones that guarantee a low latency to interactive and soft real-time |
| * applications, and a hierarchical extension based on H-WF2Q+. |
| * |
| * B-WF2Q+ is based on WF2Q+, which is described in [2], together with |
| * H-WF2Q+, while the augmented tree used here to implement B-WF2Q+ |
| * with O(log N) complexity derives from the one introduced with EEVDF |
| * in [3]. |
| * |
| * [1] P. Valente, A. Avanzini, "Evolution of the BFQ Storage I/O |
| * Scheduler", Proceedings of the First Workshop on Mobile System |
| * Technologies (MST-2015), May 2015. |
| * http://algogroup.unimore.it/people/paolo/disk_sched/mst-2015.pdf |
| * |
| * [2] Jon C.R. Bennett and H. Zhang, "Hierarchical Packet Fair Queueing |
| * Algorithms", IEEE/ACM Transactions on Networking, 5(5):675-689, |
| * Oct 1997. |
| * |
| * http://www.cs.cmu.edu/~hzhang/papers/TON-97-Oct.ps.gz |
| * |
| * [3] I. Stoica and H. Abdel-Wahab, "Earliest Eligible Virtual Deadline |
| * First: A Flexible and Accurate Mechanism for Proportional Share |
| * Resource Allocation", technical report. |
| * |
| * http://www.cs.berkeley.edu/~istoica/papers/eevdf-tr-95.pdf |
| */ |
| #include <linux/module.h> |
| #include <linux/slab.h> |
| #include <linux/blkdev.h> |
| #include <linux/cgroup.h> |
| #include <linux/elevator.h> |
| #include <linux/ktime.h> |
| #include <linux/rbtree.h> |
| #include <linux/ioprio.h> |
| #include <linux/sbitmap.h> |
| #include <linux/delay.h> |
| |
| #include "blk.h" |
| #include "blk-mq.h" |
| #include "blk-mq-tag.h" |
| #include "blk-mq-sched.h" |
| #include "bfq-iosched.h" |
| #include "blk-wbt.h" |
| |
| #define BFQ_BFQQ_FNS(name) \ |
| void bfq_mark_bfqq_##name(struct bfq_queue *bfqq) \ |
| { \ |
| __set_bit(BFQQF_##name, &(bfqq)->flags); \ |
| } \ |
| void bfq_clear_bfqq_##name(struct bfq_queue *bfqq) \ |
| { \ |
| __clear_bit(BFQQF_##name, &(bfqq)->flags); \ |
| } \ |
| int bfq_bfqq_##name(const struct bfq_queue *bfqq) \ |
| { \ |
| return test_bit(BFQQF_##name, &(bfqq)->flags); \ |
| } |
| |
| BFQ_BFQQ_FNS(just_created); |
| BFQ_BFQQ_FNS(busy); |
| BFQ_BFQQ_FNS(wait_request); |
| BFQ_BFQQ_FNS(non_blocking_wait_rq); |
| BFQ_BFQQ_FNS(fifo_expire); |
| BFQ_BFQQ_FNS(has_short_ttime); |
| BFQ_BFQQ_FNS(sync); |
| BFQ_BFQQ_FNS(IO_bound); |
| BFQ_BFQQ_FNS(in_large_burst); |
| BFQ_BFQQ_FNS(coop); |
| BFQ_BFQQ_FNS(split_coop); |
| BFQ_BFQQ_FNS(softrt_update); |
| BFQ_BFQQ_FNS(has_waker); |
| #undef BFQ_BFQQ_FNS \ |
| |
| /* Expiration time of sync (0) and async (1) requests, in ns. */ |
| static const u64 bfq_fifo_expire[2] = { NSEC_PER_SEC / 4, NSEC_PER_SEC / 8 }; |
| |
| /* Maximum backwards seek (magic number lifted from CFQ), in KiB. */ |
| static const int bfq_back_max = 16 * 1024; |
| |
| /* Penalty of a backwards seek, in number of sectors. */ |
| static const int bfq_back_penalty = 2; |
| |
| /* Idling period duration, in ns. */ |
| static u64 bfq_slice_idle = NSEC_PER_SEC / 125; |
| |
| /* Minimum number of assigned budgets for which stats are safe to compute. */ |
| static const int bfq_stats_min_budgets = 194; |
| |
| /* Default maximum budget values, in sectors and number of requests. */ |
| static const int bfq_default_max_budget = 16 * 1024; |
| |
| /* |
| * When a sync request is dispatched, the queue that contains that |
| * request, and all the ancestor entities of that queue, are charged |
| * with the number of sectors of the request. In contrast, if the |
| * request is async, then the queue and its ancestor entities are |
| * charged with the number of sectors of the request, multiplied by |
| * the factor below. This throttles the bandwidth for async I/O, |
| * w.r.t. to sync I/O, and it is done to counter the tendency of async |
| * writes to steal I/O throughput to reads. |
| * |
| * The current value of this parameter is the result of a tuning with |
| * several hardware and software configurations. We tried to find the |
| * lowest value for which writes do not cause noticeable problems to |
| * reads. In fact, the lower this parameter, the stabler I/O control, |
| * in the following respect. The lower this parameter is, the less |
| * the bandwidth enjoyed by a group decreases |
| * - when the group does writes, w.r.t. to when it does reads; |
| * - when other groups do reads, w.r.t. to when they do writes. |
| */ |
| static const int bfq_async_charge_factor = 3; |
| |
| /* Default timeout values, in jiffies, approximating CFQ defaults. */ |
| const int bfq_timeout = HZ / 8; |
| |
| /* |
| * Time limit for merging (see comments in bfq_setup_cooperator). Set |
| * to the slowest value that, in our tests, proved to be effective in |
| * removing false positives, while not causing true positives to miss |
| * queue merging. |
| * |
| * As can be deduced from the low time limit below, queue merging, if |
| * successful, happens at the very beginning of the I/O of the involved |
| * cooperating processes, as a consequence of the arrival of the very |
| * first requests from each cooperator. After that, there is very |
| * little chance to find cooperators. |
| */ |
| static const unsigned long bfq_merge_time_limit = HZ/10; |
| |
| static struct kmem_cache *bfq_pool; |
| |
| /* Below this threshold (in ns), we consider thinktime immediate. */ |
| #define BFQ_MIN_TT (2 * NSEC_PER_MSEC) |
| |
| /* hw_tag detection: parallel requests threshold and min samples needed. */ |
| #define BFQ_HW_QUEUE_THRESHOLD 3 |
| #define BFQ_HW_QUEUE_SAMPLES 32 |
| |
| #define BFQQ_SEEK_THR (sector_t)(8 * 100) |
| #define BFQQ_SECT_THR_NONROT (sector_t)(2 * 32) |
| #define BFQ_RQ_SEEKY(bfqd, last_pos, rq) \ |
| (get_sdist(last_pos, rq) > \ |
| BFQQ_SEEK_THR && \ |
| (!blk_queue_nonrot(bfqd->queue) || \ |
| blk_rq_sectors(rq) < BFQQ_SECT_THR_NONROT)) |
| #define BFQQ_CLOSE_THR (sector_t)(8 * 1024) |
| #define BFQQ_SEEKY(bfqq) (hweight32(bfqq->seek_history) > 19) |
| /* |
| * Sync random I/O is likely to be confused with soft real-time I/O, |
| * because it is characterized by limited throughput and apparently |
| * isochronous arrival pattern. To avoid false positives, queues |
| * containing only random (seeky) I/O are prevented from being tagged |
| * as soft real-time. |
| */ |
| #define BFQQ_TOTALLY_SEEKY(bfqq) (bfqq->seek_history == -1) |
| |
| /* Min number of samples required to perform peak-rate update */ |
| #define BFQ_RATE_MIN_SAMPLES 32 |
| /* Min observation time interval required to perform a peak-rate update (ns) */ |
| #define BFQ_RATE_MIN_INTERVAL (300*NSEC_PER_MSEC) |
| /* Target observation time interval for a peak-rate update (ns) */ |
| #define BFQ_RATE_REF_INTERVAL NSEC_PER_SEC |
| |
| /* |
| * Shift used for peak-rate fixed precision calculations. |
| * With |
| * - the current shift: 16 positions |
| * - the current type used to store rate: u32 |
| * - the current unit of measure for rate: [sectors/usec], or, more precisely, |
| * [(sectors/usec) / 2^BFQ_RATE_SHIFT] to take into account the shift, |
| * the range of rates that can be stored is |
| * [1 / 2^BFQ_RATE_SHIFT, 2^(32 - BFQ_RATE_SHIFT)] sectors/usec = |
| * [1 / 2^16, 2^16] sectors/usec = [15e-6, 65536] sectors/usec = |
| * [15, 65G] sectors/sec |
| * Which, assuming a sector size of 512B, corresponds to a range of |
| * [7.5K, 33T] B/sec |
| */ |
| #define BFQ_RATE_SHIFT 16 |
| |
| /* |
| * When configured for computing the duration of the weight-raising |
| * for interactive queues automatically (see the comments at the |
| * beginning of this file), BFQ does it using the following formula: |
| * duration = (ref_rate / r) * ref_wr_duration, |
| * where r is the peak rate of the device, and ref_rate and |
| * ref_wr_duration are two reference parameters. In particular, |
| * ref_rate is the peak rate of the reference storage device (see |
| * below), and ref_wr_duration is about the maximum time needed, with |
| * BFQ and while reading two files in parallel, to load typical large |
| * applications on the reference device (see the comments on |
| * max_service_from_wr below, for more details on how ref_wr_duration |
| * is obtained). In practice, the slower/faster the device at hand |
| * is, the more/less it takes to load applications with respect to the |
| * reference device. Accordingly, the longer/shorter BFQ grants |
| * weight raising to interactive applications. |
| * |
| * BFQ uses two different reference pairs (ref_rate, ref_wr_duration), |
| * depending on whether the device is rotational or non-rotational. |
| * |
| * In the following definitions, ref_rate[0] and ref_wr_duration[0] |
| * are the reference values for a rotational device, whereas |
| * ref_rate[1] and ref_wr_duration[1] are the reference values for a |
| * non-rotational device. The reference rates are not the actual peak |
| * rates of the devices used as a reference, but slightly lower |
| * values. The reason for using slightly lower values is that the |
| * peak-rate estimator tends to yield slightly lower values than the |
| * actual peak rate (it can yield the actual peak rate only if there |
| * is only one process doing I/O, and the process does sequential |
| * I/O). |
| * |
| * The reference peak rates are measured in sectors/usec, left-shifted |
| * by BFQ_RATE_SHIFT. |
| */ |
| static int ref_rate[2] = {14000, 33000}; |
| /* |
| * To improve readability, a conversion function is used to initialize |
| * the following array, which entails that the array can be |
| * initialized only in a function. |
| */ |
| static int ref_wr_duration[2]; |
| |
| /* |
| * BFQ uses the above-detailed, time-based weight-raising mechanism to |
| * privilege interactive tasks. This mechanism is vulnerable to the |
| * following false positives: I/O-bound applications that will go on |
| * doing I/O for much longer than the duration of weight |
| * raising. These applications have basically no benefit from being |
| * weight-raised at the beginning of their I/O. On the opposite end, |
| * while being weight-raised, these applications |
| * a) unjustly steal throughput to applications that may actually need |
| * low latency; |
| * b) make BFQ uselessly perform device idling; device idling results |
| * in loss of device throughput with most flash-based storage, and may |
| * increase latencies when used purposelessly. |
| * |
| * BFQ tries to reduce these problems, by adopting the following |
| * countermeasure. To introduce this countermeasure, we need first to |
| * finish explaining how the duration of weight-raising for |
| * interactive tasks is computed. |
| * |
| * For a bfq_queue deemed as interactive, the duration of weight |
| * raising is dynamically adjusted, as a function of the estimated |
| * peak rate of the device, so as to be equal to the time needed to |
| * execute the 'largest' interactive task we benchmarked so far. By |
| * largest task, we mean the task for which each involved process has |
| * to do more I/O than for any of the other tasks we benchmarked. This |
| * reference interactive task is the start-up of LibreOffice Writer, |
| * and in this task each process/bfq_queue needs to have at most ~110K |
| * sectors transferred. |
| * |
| * This last piece of information enables BFQ to reduce the actual |
| * duration of weight-raising for at least one class of I/O-bound |
| * applications: those doing sequential or quasi-sequential I/O. An |
| * example is file copy. In fact, once started, the main I/O-bound |
| * processes of these applications usually consume the above 110K |
| * sectors in much less time than the processes of an application that |
| * is starting, because these I/O-bound processes will greedily devote |
| * almost all their CPU cycles only to their target, |
| * throughput-friendly I/O operations. This is even more true if BFQ |
| * happens to be underestimating the device peak rate, and thus |
| * overestimating the duration of weight raising. But, according to |
| * our measurements, once transferred 110K sectors, these processes |
| * have no right to be weight-raised any longer. |
| * |
| * Basing on the last consideration, BFQ ends weight-raising for a |
| * bfq_queue if the latter happens to have received an amount of |
| * service at least equal to the following constant. The constant is |
| * set to slightly more than 110K, to have a minimum safety margin. |
| * |
| * This early ending of weight-raising reduces the amount of time |
| * during which interactive false positives cause the two problems |
| * described at the beginning of these comments. |
| */ |
| static const unsigned long max_service_from_wr = 120000; |
| |
| #define RQ_BIC(rq) icq_to_bic((rq)->elv.priv[0]) |
| #define RQ_BFQQ(rq) ((rq)->elv.priv[1]) |
| |
| struct bfq_queue *bic_to_bfqq(struct bfq_io_cq *bic, bool is_sync) |
| { |
| return bic->bfqq[is_sync]; |
| } |
| |
| void bic_set_bfqq(struct bfq_io_cq *bic, struct bfq_queue *bfqq, bool is_sync) |
| { |
| bic->bfqq[is_sync] = bfqq; |
| } |
| |
| struct bfq_data *bic_to_bfqd(struct bfq_io_cq *bic) |
| { |
| return bic->icq.q->elevator->elevator_data; |
| } |
| |
| /** |
| * icq_to_bic - convert iocontext queue structure to bfq_io_cq. |
| * @icq: the iocontext queue. |
| */ |
| static struct bfq_io_cq *icq_to_bic(struct io_cq *icq) |
| { |
| /* bic->icq is the first member, %NULL will convert to %NULL */ |
| return container_of(icq, struct bfq_io_cq, icq); |
| } |
| |
| /** |
| * bfq_bic_lookup - search into @ioc a bic associated to @bfqd. |
| * @bfqd: the lookup key. |
| * @ioc: the io_context of the process doing I/O. |
| * @q: the request queue. |
| */ |
| static struct bfq_io_cq *bfq_bic_lookup(struct bfq_data *bfqd, |
| struct io_context *ioc, |
| struct request_queue *q) |
| { |
| if (ioc) { |
| unsigned long flags; |
| struct bfq_io_cq *icq; |
| |
| spin_lock_irqsave(&q->queue_lock, flags); |
| icq = icq_to_bic(ioc_lookup_icq(ioc, q)); |
| spin_unlock_irqrestore(&q->queue_lock, flags); |
| |
| return icq; |
| } |
| |
| return NULL; |
| } |
| |
| /* |
| * Scheduler run of queue, if there are requests pending and no one in the |
| * driver that will restart queueing. |
| */ |
| void bfq_schedule_dispatch(struct bfq_data *bfqd) |
| { |
| if (bfqd->queued != 0) { |
| bfq_log(bfqd, "schedule dispatch"); |
| blk_mq_run_hw_queues(bfqd->queue, true); |
| } |
| } |
| |
| #define bfq_class_idle(bfqq) ((bfqq)->ioprio_class == IOPRIO_CLASS_IDLE) |
| #define bfq_class_rt(bfqq) ((bfqq)->ioprio_class == IOPRIO_CLASS_RT) |
| |
| #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; |
| } |
| |
| /* |
| * 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 = bfqq->wr_coeff; |
| 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->entity.new_weight = bic->saved_weight; |
| bfqq->ttime = bic->saved_ttime; |
| bfqq->wr_coeff = bic->saved_wr_coeff; |
| 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 - |
| (bfqq->weight_counter != NULL); |
| } |
| |
| /* 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 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). |
| */ |
| 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; |
| *interactive = !in_burst && idle_for_long_time; |
| wr_or_deserves_wr = bfqd->low_latency && |
| (bfqq->wr_coeff > 1 || |
| (bfq_bfqq_sync(bfqq) && |
| bfqq->bic && (*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 (!bfq_bfqq_IO_bound(bfqq)) { |
| if (arrived_in_time) { |
| bfqq->requests_within_timer++; |
| if (bfqq->requests_within_timer >= |
| bfqd->bfq_requests_within_timer) |
| bfq_mark_bfqq_IO_bound(bfqq); |
| } else |
| bfqq->requests_within_timer = 0; |
| } |
| |
| 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 only if preemption may be needed |
| * for guarantees. In particular, we care only 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()). |
| */ |
| 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)) && |
| 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_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; |
| |
| 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)) { |
| /* |
| * 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 |
| * two consecutive times, bfqq happens to become non |
| * empty right after a request of Q has been |
| * completed. In particular, on the first time, Q is |
| * tentatively set as a candidate waker queue, while |
| * on the second time, the flag |
| * bfq_bfqq_has_waker(bfqq) is set 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 two-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. |
| */ |
| if (bfqd->last_completed_rq_bfqq && |
| !bfq_bfqq_has_short_ttime(bfqq) && |
| ktime_get_ns() - bfqd->last_completion < |
| 200 * NSEC_PER_USEC) { |
| if (bfqd->last_completed_rq_bfqq != bfqq && |
| bfqd->last_completed_rq_bfqq != |
| bfqq->waker_bfqq) { |
| /* |
| * First synchronization detected with |
| * a candidate waker queue, or with a |
| * different candidate waker queue |
| * from the current one. |
| */ |
| bfqq->waker_bfqq = bfqd->last_completed_rq_bfqq; |
| |
| /* |
| * 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); |
| |
| bfq_clear_bfqq_has_waker(bfqq); |
| } else if (bfqd->last_completed_rq_bfqq == |
| bfqq->waker_bfqq && |
| !bfq_bfqq_has_waker(bfqq)) { |
| /* |
| * synchronization with waker_bfqq |
| * seen for the second time |
| */ |
| bfq_mark_bfqq_has_waker(bfqq); |
| } |
| } |
| |
| /* |
| * 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(100))) { |
| 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; |
| bfqd->rqs_injected = false; |
| } |
| } |
| |
| 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 blk_mq_hw_ctx *hctx, struct bio *bio, |
| unsigned int nr_segs) |
| { |
| struct request_queue *q = hctx->queue; |
| 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); |
| |
| if (free) |
| blk_mq_free_request(free); |
| spin_unlock_irq(&bfqd->lock); |
| |
| 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; |
| 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) |
| return; |
| |
| /* |
| * 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); |
| } |
| |
| /* Must be called with bfqq != NULL */ |
| static void bfq_bfqq_end_wr(struct bfq_queue *bfqq) |
| { |
| 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_BE_NR; 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; |
| 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; |
| } |
| |
| /* |
| * 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_queue *in_service_bfqq, *new_bfqq; |
| |
| /* |
| * 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 (bfqq->new_bfqq) |
| return bfqq->new_bfqq; |
| |
| 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_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_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_last_wr_start_finish = bfqq->last_wr_start_finish; |
| bic->saved_wr_cur_max_time = bfqq->wr_cur_max_time; |
| } |
| } |
| |
| 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); |
| |
| /* |
| * 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; |
| /* release process reference to bfqq */ |
| bfq_put_queue(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); |
| 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; |
| } |
| |
| /* |
| * 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. |
| */ |
| static bool idling_needed_for_service_guarantees(struct bfq_data *bfqd, |
| struct bfq_queue *bfqq) |
| { |
| return (bfqq->wr_coeff > 1 && |
| (bfqd->wr_busy_queues < |
| bfq_tot_busy_queues(bfqd) || |
| bfqd->rq_in_driver >= |
| bfqq->dispatched + 4)) || |
| bfq_asymmetric_scenario(bfqd, bfqq); |
| } |
| |
| 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 (reason == BFQQE_TOO_IDLE && |
| entity->service <= 2 * entity->budget / 10) |
| bfq_clear_bfqq_IO_bound(bfqq); |
| |
| 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()). Thus we can compute |
| * soft_rt_next_start. And we do it, unless bfqq is in |
| * interactive weight raising. We do not do it in the |
| * latter subcase, for the following reason. bfqq may |
| * be conveying the I/O needed to load a soft |
| * real-time application. Such an application will |
| * actually exhibit a soft real-time I/O pattern after |
| * it finally starts doing its job. But, if |
| * soft_rt_next_start is computed here for an |
| * interactive bfqq, and bfqq had received a lot of |
| * service before remaining with no outstanding |
| * request (likely to happen on a fast device), then |
| * soft_rt_next_start would be assigned such a high |
| * value that, for a very long time, bfqq would be |
| * prevented from being possibly considered as soft |
| * real time. |
| * |
| * 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->wr_coeff != bfqd->bfq_wr_coeff) |
| 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; |
| |
| 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; |
| |
| 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; |
| |
| /* |
| * The next three 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 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 third 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 third 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 |
| * third 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 (bfq_bfqq_has_waker(bfqq) && |
| bfq_bfqq_busy(bfqq->waker_bfqq) && |
| 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 (!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))) |
| bfq_bfqq_end_wr(bfqq); |
| else { |
| 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 at 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; |
| #ifdef CONFIG_BFQ_GROUP_IOSCHED |
| struct bfq_group *bfqg = bfqq_group(bfqq); |
| #endif |
| |
| 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; |
| bfq_clear_bfqq_has_waker(item); |
| 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); |
| #ifdef CONFIG_BFQ_GROUP_IOSCHED |
| bfqg_and_blkg_put(bfqg); |
| #endif |
| } |
| |
| 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_put_queue(bfqq); /* release process reference */ |
| } |
| |
| 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); |
| |
| 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: |
| dev_err(bfqq->bfqd->queue->backing_dev_info->dev, |
| "bfq: bad prio class %d\n", ioprio_class); |
| /* fall through */ |
| 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_BE_NR) { |
| pr_crit("bfq_set_next_ioprio_data: new_ioprio %d\n", |
| bfqq->new_ioprio); |
| bfqq->new_ioprio = IOPRIO_BE_NR; |
| } |
| |
| bfqq->entity.new_weight = bfq_ioprio_to_weight(bfqq->new_ioprio); |
| 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); |
| |
| 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) { |
| /* release process reference on this queue */ |
| bfq_put_queue(bfqq); |
| bfqq = bfq_get_queue(bfqd, bio, BLK_RW_ASYNC, bic); |
| 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) |
| { |
| 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 = ktime_get_ns() + 1; |
| |
| 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_NORM; |
| /* fall through */ |
| 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_get_queue(struct bfq_data *bfqd, |
| struct bio *bio, bool is_sync, |
| struct bfq_io_cq *bic) |
| { |
| 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 */ |
| bfq_log_bfqq(bfqd, bfqq, "get_queue, at end: %p, %d", bfqq, bfqq->ref); |
| 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 = ktime_get_ns() - bfqq->ttime.last_end_request; |
| |
| elapsed = min_t(u64, elapsed, 2ULL * bfqd->bfq_slice_idle); |
| |
| ttime->ttime_samples = (7*bfqq->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)) |
| bfq_bfqq_end_wr(bfqq); |
| } |
| |
| 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 |
| */ |
| if (atomic_read(&bic->icq.ioc->active_ref) == 0 || |
| (bfq_sample_valid(bfqq->ttime.ttime_samples) && |
| bfqq->ttime.ttime_mean > bfqd->bfq_slice_idle)) |
| 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); |
| 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; |
| |
| spin_lock_irq(&bfqd->lock); |
| if (blk_mq_sched_try_insert_merge(q, rq)) { |
| spin_unlock_irq(&bfqd->lock); |
| return; |
| } |
| |
| spin_unlock_irq(&bfqd->lock); |
| |
| blk_mq_sched_request_inserted(rq); |
| |
| spin_lock_irq(&bfqd->lock); |
| bfqq = bfq_init_rq(rq); |
| if (!bfqq || at_head || blk_rq_is_passthrough(rq)) { |
| 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; |
| bfqd->last_completed_rq_bfqq = bfqq; |
| |
| /* |
| * 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) { |
| 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<<1) |
| 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) { |
| bfqq->last_serv_time_ns = tot_time_ns; |
| /* |
| * Now we certainly have a base value: make sure we |
| * start trying injection. |
| */ |
| bfqq->inject_limit = max_t(unsigned int, 1, old_limit); |
| } 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; |
| } |
| |
| /* |
| * 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; |
| |
| /* |
| * Requeue and finish hooks are invoked in blk-mq without |
| * checking whether the involved request is actually still |
| * referenced in the scheduler. To handle this fact, the |
| * following two checks make this function exit in case of |
| * spurious invocations, for which there is nothing to do. |
| * |
| * First, check whether rq has nothing to do with an elevator. |
| */ |
| if (unlikely(!(rq->rq_flags & RQF_ELVPRIV))) |
| return; |
| |
| /* |
| * 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); |
| |
| if (likely(rq->rq_flags & RQF_STARTED)) { |
| unsigned long flags; |
| |
| spin_lock_irqsave(&bfqd->lock, flags); |
| |
| 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); |
| } else { |
| /* |
| * Request rq may be still/already in the scheduler, |
| * in which case we need to remove it (this should |
| * never happen in case of requeue). And we cannot |
| * defer such a check and removal, to avoid |
| * inconsistencies in the time interval from the end |
| * of this function to the start of the deferred work. |
| * This situation seems to occur only in process |
| * context, as a consequence of a merge. In the |
| * current version of the code, this implies that the |
| * lock is held. |
| */ |
| |
| if (!RB_EMPTY_NODE(&rq->rb_node)) { |
| bfq_remove_request(rq->q, rq); |
| bfqg_stats_update_io_remove(bfqq_group(bfqq), |
| rq->cmd_flags); |
| } |
| bfq_finish_requeue_request_body(bfqq); |
| } |
| |
| /* |
| * 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; |
| } |
| |
| /* |
| * 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_put_queue(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); |
| |
| 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, struct bio *bio) |
| { |
| /* |
| * 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)) { |
| bfq_log_bfqq(bfqd, bfqq, "breaking apart 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); |
| 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_queue *bfqq) |
| { |
| struct bfq_data *bfqd = bfqq->bfqd; |
| enum bfqq_expiration reason; |
| unsigned long flags; |
| |
| spin_lock_irqsave(&bfqd->lock, flags); |
| bfq_clear_bfqq_wait_request(bfqq); |
| |
| if (bfqq != bfqd->in_service_queue) { |
| spin_unlock_irqrestore(&bfqd->lock, flags); |
| return; |
| } |
| |
| 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(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_BE_NR; 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); |
| |
| #ifdef CONFIG_BFQ_GROUP_IOSCHED |
| /* release oom-queue reference to root group */ |
| bfqg_and_blkg_put(bfqd->root_group); |
| |
| 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_requests_within_timer = 120; |
| |
| 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"); |