| // 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; |
| } |
| |
| /* oom_bfqq does not participate in queue merging */ |
| if (bfqq == &bfqd->oom_bfqq) |
| return; |
| |
| /* |
| * bfqq cannot be merged any longer (see comments in |
| * bfq_setup_cooperator): no point in adding bfqq into the |
| * position tree. |
| */ |
| if (bfq_too_late_for_merging(bfqq)) |
| return; |
| |
| if (bfq_class_idle(bfqq)) |
| return; |
| if (!bfqq->next_rq) |
| return; |
| |
| bfqq->pos_root = &bfq_bfqq_to_bfqg(bfqq)->rq_pos_tree; |
| __bfqq = bfq_rq_pos_tree_lookup(bfqd, bfqq->pos_root, |
| blk_rq_pos(bfqq->next_rq), &parent, &p); |
| if (!__bfqq) { |
| rb_link_node(&bfqq->pos_node, parent, p); |
| rb_insert_color(&bfqq->pos_node, bfqq->pos_root); |
| } else |
| bfqq->pos_root = NULL; |
| } |
| |
| /* |
| * The following function returns false either if every active queue |
| * must receive the same share of the throughput (symmetric scenario), |
| * or, as a special case, if bfqq must receive a share of the |
| * throughput lower than or equal to the share that every other active |
| * queue must receive. If bfqq does sync I/O, then these are the only |
| * two cases where bfqq happens to be guaranteed its share of the |
| * throughput even if I/O dispatching is not plugged when bfqq remains |
| * temporarily empty (for more details, see the comments in the |
| * function bfq_better_to_idle()). For this reason, the return value |
| * of this function is used to check whether I/O-dispatch plugging can |
| * be avoided. |
| * |
| * The above first case (symmetric scenario) occurs when: |
| * 1) all active queues have the same weight, |
| * 2) all active queues belong to the same I/O-priority class, |
| * 3) all active groups at the same level in the groups tree have the same |
| * weight, |
| * 4) all active groups at the same level in the groups tree have the same |
| * number of children. |
| * |
| * Unfortunately, keeping the necessary state for evaluating exactly |
| * the last two symmetry sub-conditions above would be quite complex |
| * and time consuming. Therefore this function evaluates, instead, |
| * only the following stronger three sub-conditions, for which it is |
| * much easier to maintain the needed state: |
| * 1) all active queues have the same weight, |
| * 2) all active queues belong to the same I/O-priority class, |
| * 3) there are no active groups. |
| * In particular, the last condition is always true if hierarchical |
| * support or the cgroups interface are not enabled, thus no state |
| * needs to be maintained in this case. |
| */ |
| static bool bfq_asymmetric_scenario(struct bfq_data *bfqd, |
| struct bfq_queue *bfqq) |
| { |
| bool smallest_weight = bfqq && |
| bfqq->weight_counter && |
| bfqq->weight_counter == |
| container_of( |
| rb_first_cached(&bfqd->queue_weights_tree), |
| struct bfq_weight_counter, |
| weights_node); |
| |
| /* |
| * For queue weights to differ, queue_weights_tree must contain |
| * at least two nodes. |
| */ |
| bool varied_queue_weights = !smallest_weight && |
| !RB_EMPTY_ROOT(&bfqd->queue_weights_tree.rb_root) && |
| (bfqd->queue_weights_tree.rb_root.rb_node->rb_left || |
| bfqd->queue_weights_tree.rb_root.rb_node->rb_right); |
| |
| bool multiple_classes_busy = |
| (bfqd->busy_queues[0] && bfqd->busy_queues[1]) || |
| (bfqd->busy_queues[0] && bfqd->busy_queues[2]) || |
| (bfqd->busy_queues[1] && bfqd->busy_queues[2]); |
| |
| return varied_queue_weights || multiple_classes_busy |
| #ifdef CONFIG_BFQ_GROUP_IOSCHED |
| || bfqd->num_groups_with_pending_reqs > 0 |
| #endif |
| ; |
| } |
| |
| /* |
| * If the weight-counter tree passed as input contains no counter for |
| * the weight of the input queue, then add that counter; otherwise just |
| * increment the existing counter. |
| * |
| * Note that weight-counter trees contain few nodes in mostly symmetric |
| * scenarios. For example, if all queues have the same weight, then the |
| * weight-counter tree for the queues may contain at most one node. |
| * This holds even if low_latency is on, because weight-raised queues |
| * are not inserted in the tree. |
| * In most scenarios, the rate at which nodes are created/destroyed |
| * should be low too. |
| */ |
| void bfq_weights_tree_add(struct bfq_data *bfqd, struct bfq_queue *bfqq, |
| struct rb_root_cached *root) |
| { |
| struct bfq_entity *entity = &bfqq->entity; |
| struct rb_node **new = &(root->rb_root.rb_node), *parent = NULL; |
| bool leftmost = true; |
| |
| /* |
| * Do not insert if the queue is already associated with a |
| * counter, which happens if: |
| * 1) a request arrival has caused the queue to become both |
| * non-weight-raised, and hence change its weight, and |
| * backlogged; in this respect, each of the two events |
| * causes an invocation of this function, |
| * 2) this is the invocation of this function caused by the |
| * second event. This second invocation is actually useless, |
| * and we handle this fact by exiting immediately. More |
| * efficient or clearer solutions might possibly be adopted. |
| */ |
| if (bfqq->weight_counter) |
| return; |
| |
| while (*new) { |
| struct bfq_weight_counter *__counter = container_of(*new, |
| struct bfq_weight_counter, |
| weights_node); |
| parent = *new; |
| |
| if (entity->weight == __counter->weight) { |
| bfqq->weight_counter = __counter; |
| goto inc_counter; |
| } |
| if (entity->weight < __counter->weight) |
| new = &((*new)->rb_left); |
| else { |
| new = &((*new)->rb_right); |
| leftmost = false; |
| } |
| } |
| |
| bfqq->weight_counter = kzalloc(sizeof(struct bfq_weight_counter), |
| GFP_ATOMIC); |
| |
| /* |
| * In the unlucky event of an allocation failure, we just |
| * exit. This will cause the weight of queue to not be |
| * considered in bfq_asymmetric_scenario, which, in its turn, |
| * causes the scenario to be deemed wrongly symmetric in case |
| * bfqq's weight would have been the only weight making the |
| * scenario asymmetric. On the bright side, no unbalance will |
| * however occur when bfqq becomes inactive again (the |
| * invocation of this function is triggered by an activation |
| * of queue). In fact, bfq_weights_tree_remove does nothing |
| * if !bfqq->weight_counter. |
| */ |
| if (unlikely(!bfqq->weight_counter)) |
| return; |
| |
| bfqq->weight_counter->weight = entity->weight; |
| rb_link_node(&bfqq->weight_counter->weights_node, parent, new); |
| rb_insert_color_cached(&bfqq->weight_counter->weights_node, root, |
| leftmost); |
| |
| inc_counter: |
| bfqq->weight_counter->num_active++; |
| bfqq->ref++; |
| } |
| |
| /* |
| * Decrement the weight counter associated with the queue, and, if the |
| * counter reaches 0, remove the counter from the tree. |
| * See the comments to the function bfq_weights_tree_add() for considerations |
| * about overhead. |
| */ |
| void __bfq_weights_tree_remove(struct bfq_data *bfqd, |
| struct bfq_queue *bfqq, |
| struct rb_root_cached *root) |
| { |
| if (!bfqq->weight_counter) |
| return; |
| |
| bfqq->weight_counter->num_active--; |
| if (bfqq->weight_counter->num_active > 0) |
| goto reset_entity_pointer; |
| |
| rb_erase_cached(&bfqq->weight_counter->weights_node, root); |
| kfree(bfqq->weight_counter); |
| |
| reset_entity_pointer: |
| bfqq->weight_counter = NULL; |
| bfq_put_queue(bfqq); |
| } |
| |
| /* |
| * Invoke __bfq_weights_tree_remove on bfqq and decrement the number |
| * of active groups for each queue's inactive parent entity. |
| */ |
| void bfq_weights_tree_remove(struct bfq_data *bfqd, |
| struct bfq_queue *bfqq) |
| { |
| struct bfq_entity *entity = bfqq->entity.parent; |
| |
| for_each_entity(entity) { |
| struct bfq_sched_data *sd = entity->my_sched_data; |
| |
| if (sd->next_in_service || sd->in_service_entity) { |
| /* |
| * entity is still active, because either |
| * next_in_service or in_service_entity is not |
| * NULL (see the comments on the definition of |
| * next_in_service for details on why |
| * in_service_entity must be checked too). |
| * |
| * As a consequence, its parent entities are |
| * active as well, and thus this loop must |
| * stop here. |
| */ |
| break; |
| } |
| |
| /* |
| * The decrement of num_groups_with_pending_reqs is |
| * not performed immediately upon the deactivation of |
| * entity, but it is delayed to when it also happens |
| * that the first leaf descendant bfqq of entity gets |
| * all its pending requests completed. The following |
| * instructions perform this delayed decrement, if |
| * needed. See the comments on |
| * num_groups_with_pending_reqs for details. |
| */ |
| if (entity->in_groups_with_pending_reqs) { |
| entity->in_groups_with_pending_reqs = false; |
| bfqd->num_groups_with_pending_reqs--; |
| } |
| } |
| |
| /* |
| * Next function is invoked last, because it causes bfqq to be |
| * freed if the following holds: bfqq is not in service and |
| * has no dispatched request. DO NOT use bfqq after the next |
| * function invocation. |
| */ |
| __bfq_weights_tree_remove(bfqd, bfqq, |
| &bfqd->queue_weights_tree); |
| } |
| |
| /* |
| * Return expired entry, or NULL to just start from scratch in rbtree. |
| */ |
| static struct request *bfq_check_fifo(struct bfq_queue *bfqq, |
| struct request *last) |
| { |
| struct request *rq; |
| |
| if (bfq_bfqq_fifo_expire(bfqq)) |
| return NULL; |
| |
| bfq_mark_bfqq_fifo_expire(bfqq); |
| |
| rq = rq_entry_fifo(bfqq->fifo.next); |
| |
| if (rq == last || ktime_get_ns() < rq->fifo_time) |
| return NULL; |
| |
| bfq_log_bfqq(bfqq->bfqd, bfqq, "check_fifo: returned %p", rq); |
| return rq; |
| } |
| |
| static struct request *bfq_find_next_rq(struct bfq_data *bfqd, |
| struct bfq_queue *bfqq, |
| struct request *last) |
| { |
| struct rb_node *rbnext = rb_next(&last->rb_node); |
| struct rb_node *rbprev = rb_prev(&last->rb_node); |
| struct request *next, *prev = NULL; |
| |
| /* Follow expired path, else get first next available. */ |
| next = bfq_check_fifo(bfqq, last); |
| if (next) |
| return next; |
| |
| if (rbprev) |
| prev = rb_entry_rq(rbprev); |
| |
| if (rbnext) |
| next = rb_entry_rq(rbnext); |
| else { |
| rbnext = rb_first(&bfqq->sort_list); |
| if (rbnext && rbnext != &last->rb_node) |
| next = rb_entry_rq(rbnext); |
| } |
| |
| return bfq_choose_req(bfqd, next, prev, blk_rq_pos(last)); |
| } |
| |
| /* see the definition of bfq_async_charge_factor for details */ |
| static unsigned long bfq_serv_to_charge(struct request *rq, |
| struct bfq_queue *bfqq) |
| { |
| if (bfq_bfqq_sync(bfqq) || bfqq->wr_coeff > 1 || |
| bfq_asymmetric_scenario(bfqq->bfqd, bfqq)) |
| return blk_rq_sectors(rq); |
| |
| return blk_rq_sectors(rq) * bfq_async_charge_factor; |
| } |
| |
| /** |
| * bfq_updated_next_req - update the queue after a new next_rq selection. |
| * @bfqd: the device data the queue belongs to. |
| * @bfqq: the queue to update. |
| * |
| * If the first request of a queue changes we make sure that the queue |
| * has enough budget to serve at least its first request (if the |
| * request has grown). We do this because if the queue has not enough |
| * budget for its first request, it has to go through two dispatch |
| * rounds to actually get it dispatched. |
| */ |
| static void bfq_updated_next_req(struct bfq_data *bfqd, |
| struct bfq_queue *bfqq) |
| { |
| struct bfq_entity *entity = &bfqq->entity; |
| struct request *next_rq = bfqq->next_rq; |
| unsigned long new_budget; |
| |
| if (!next_rq) |
| return; |
| |
| if (bfqq == bfqd->in_service_queue) |
| /* |
| * In order not to break guarantees, budgets cannot be |
| * changed after an entity has been selected. |
| */ |
| return; |
| |
| new_budget = max_t(unsigned long, |
| max_t(unsigned long, bfqq->max_budget, |
| bfq_serv_to_charge(next_rq, bfqq)), |
| entity->service); |
| if (entity->budget != new_budget) { |
| entity->budget = new_budget; |
| bfq_log_bfqq(bfqd, bfqq, "updated next rq: new budget %lu", |
| new_budget); |
| bfq_requeue_bfqq(bfqd, bfqq, false); |
| } |
| } |
| |
| static unsigned int bfq_wr_duration(struct bfq_data *bfqd) |
| { |
| u64 dur; |
| |
| if (bfqd->bfq_wr_max_time > 0) |
| return bfqd->bfq_wr_max_time; |
| |
| dur = bfqd->rate_dur_prod; |
| do_div(dur, bfqd->peak_rate); |
| |
| /* |
| * Limit duration between 3 and 25 seconds. The upper limit |
| * has been conservatively set after the following worst case: |
| * on a QEMU/KVM virtual machine |
| * - running in a slow PC |
| * - with a virtual disk stacked on a slow low-end 5400rpm HDD |
| * - serving a heavy I/O workload, such as the sequential reading |
| * of several files |
| * mplayer took 23 seconds to start, if constantly weight-raised. |
| * |
| * As for higher values than that accommodating the above bad |
| * scenario, tests show that higher values would often yield |
| * the opposite of the desired result, i.e., would worsen |
| * responsiveness by allowing non-interactive applications to |
| * preserve weight raising for too long. |
| * |
| * On the other end, lower values than 3 seconds make it |
| * difficult for most interactive tasks to complete their jobs |
| * before weight-raising finishes. |
| */ |
| return clamp_val(dur, msecs_to_jiffies(3000), msecs_to_jiffies(25000)); |
| } |
| |
| /* switch back from soft real-time to interactive weight raising */ |
| static void switch_back_to_interactive_wr(struct bfq_queue *bfqq, |
| struct bfq_data *bfqd) |
| { |
| bfqq->wr_coeff = bfqd->bfq_wr_coeff; |
| bfqq->wr_cur_max_time = bfq_wr_duration(bfqd); |
| bfqq->last_wr_start_finish = bfqq->wr_start_at_switch_to_srt; |
| } |
| |
| static void |
| bfq_bfqq_resume_state(struct bfq_queue *bfqq, struct bfq_data *bfqd, |
| struct bfq_io_cq *bic, bool bfq_already_existing) |
| { |
| unsigned int old_wr_coeff = 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(10))) { |
| bfqd->last_empty_occupied_ns = ktime_get_ns(); |
| /* |
| * Start the state machine for measuring the |
| * total service time of rq: setting |
| * wait_dispatch will cause bfqd->waited_rq to |
| * be set when rq will be dispatched. |
| */ |
| bfqd->wait_dispatch = true; |
| /* |
| * If there is no I/O in service in the drive, |
| * then possible injection occurred before the |
| * arrival of rq will not affect the total |
| * service time of rq. So the injection limit |
| * must not be updated as a function of such |
| * total service time, unless new injection |
| * occurs before rq is completed. To have the |
| * injection limit updated only in the latter |
| * case, reset rqs_injected here (rqs_injected |
| * will be set in case injection is performed |
| * on bfqq before rq is completed). |
| */ |
| if (bfqd->rq_in_driver == 0) |
| bfqd->rqs_injected = false; |
| } |
| } |
| |
| elv_rb_add(&bfqq->sort_list, rq); |
| |
| /* |
| * Check if this request is a better next-serve candidate. |
| */ |
| prev = bfqq->next_rq; |
| next_rq = bfq_choose_req(bfqd, bfqq->next_rq, rq, bfqd->last_position); |
| bfqq->next_rq = next_rq; |
| |
| /* |
| * Adjust priority tree position, if next_rq changes. |
| * See comments on bfq_pos_tree_add_move() for the unlikely(). |
| */ |
| if (unlikely(!bfqd->nonrot_with_queueing && prev != bfqq->next_rq)) |
| bfq_pos_tree_add_move(bfqd, bfqq); |
| |
| if (!bfq_bfqq_busy(bfqq)) /* switching to busy ... */ |
| bfq_bfqq_handle_idle_busy_switch(bfqd, bfqq, old_wr_coeff, |
| rq, &interactive); |
| else { |
| if (bfqd->low_latency && old_wr_coeff == 1 && !rq_is_sync(rq) && |
| time_is_before_jiffies( |
| bfqq->last_wr_start_finish + |
| bfqd->bfq_wr_min_inter_arr_async)) { |
| bfqq->wr_coeff = bfqd->bfq_wr_coeff; |
| bfqq->wr_cur_max_time = bfq_wr_duration(bfqd); |
| |
| bfqd->wr_busy_queues++; |
| bfqq->entity.prio_changed = 1; |
| } |
| if (prev != bfqq->next_rq) |
| bfq_updated_next_req(bfqd, bfqq); |
| } |
| |
| /* |
| * Assign jiffies to last_wr_start_finish in the following |
| * cases: |
| * |
| * . if bfqq is not going to be weight-raised, because, for |
| * non weight-raised queues, last_wr_start_finish stores the |
| * arrival time of the last request; as of now, this piece |
| * of information is used only for deciding whether to |
| * weight-raise async queues |
| * |
| * . if bfqq is not weight-raised, because, if bfqq is now |
| * switching to weight-raised, then last_wr_start_finish |
| * stores the time when weight-raising starts |
| * |
| * . if bfqq is interactive, because, regardless of whether |
| * bfqq is currently weight-raised, the weight-raising |
| * period must start or restart (this case is considered |
| * separately because it is not detected by the above |
| * conditions, if bfqq is already weight-raised) |
| * |
| * last_wr_start_finish has to be updated also if bfqq is soft |
| * real-time, because the weight-raising period is constantly |
| * restarted on idle-to-busy transitions for these queues, but |
| * this is already done in bfq_bfqq_handle_idle_busy_switch if |
| * needed. |
| */ |
| if (bfqd->low_latency && |
| (old_wr_coeff == 1 || bfqq->wr_coeff == 1 || interactive)) |
| bfqq->last_wr_start_finish = jiffies; |
| } |
| |
| static struct request *bfq_find_rq_fmerge(struct bfq_data *bfqd, |
| struct bio *bio, |
| struct request_queue *q) |
| { |
| struct bfq_queue *bfqq = bfqd->bio_bfqq; |
| |
| |
| if (bfqq) |
| return elv_rb_find(&bfqq->sort_list, bio_end_sector(bio)); |
| |
| return NULL; |
| } |
| |
| static sector_t get_sdist(sector_t last_pos, struct request *rq) |
| { |
| if (last_pos) |
| return abs(blk_rq_pos(rq) - last_pos); |
| |
| return 0; |
| } |
| |
| #if 0 /* Still not clear if we can do without next two functions */ |
| static void bfq_activate_request(struct request_queue *q, struct request *rq) |
| { |
| struct bfq_data *bfqd = q->elevator->elevator_data; |
| |
| bfqd->rq_in_driver++; |
| } |
| |
| static void bfq_deactivate_request(struct request_queue *q, struct request *rq) |
| { |
| struct bfq_data *bfqd = q->elevator->elevator_data; |
| |
| bfqd->rq_in_driver--; |
| } |
| #endif |
| |
| static void bfq_remove_request(struct request_queue *q, |
| struct request *rq) |
| { |
| struct bfq_queue *bfqq = RQ_BFQQ(rq); |
| struct bfq_data *bfqd = bfqq->bfqd; |
| const int sync = rq_is_sync(rq); |
| |
| if (bfqq->next_rq == rq) { |
| bfqq->next_rq = bfq_find_next_rq(bfqd, bfqq, rq); |
| bfq_updated_next_req(bfqd, bfqq); |
| } |
| |
| if (rq->queuelist.prev != &rq->queuelist) |
| list_del_init(&rq->queuelist); |
| bfqq->queued[sync]--; |
| bfqd->queued--; |
| elv_rb_del(&bfqq->sort_list, rq); |
| |
| elv_rqhash_del(q, rq); |
| if (q->last_merge == rq) |
| q->last_merge = NULL; |
| |
| if (RB_EMPTY_ROOT(&bfqq->sort_list)) { |
| bfqq->next_rq = NULL; |
| |
| if (bfq_bfqq_busy(bfqq) && bfqq != bfqd->in_service_queue) { |
| bfq_del_bfqq_busy(bfqd, bfqq, false); |
| /* |
| * bfqq emptied. In normal operation, when |
| * bfqq is empty, bfqq->entity.service and |
| * bfqq->entity.budget must contain, |
| * respectively, the service received and the |
| * budget used last time bfqq emptied. These |
| * facts do not hold in this case, as at least |
| * this last removal occurred while bfqq is |
| * not in service. To avoid inconsistencies, |
| * reset both bfqq->entity.service and |
| * bfqq->entity.budget, if bfqq has still a |
| * process that may issue I/O requests to it. |
| */ |
| bfqq->entity.budget = bfqq->entity.service = 0; |
| } |
| |
| /* |
| * Remove queue from request-position tree as it is empty. |
| */ |
| if (bfqq->pos_root) { |
| rb_erase(&bfqq->pos_node, bfqq->pos_root); |
| bfqq->pos_root = NULL; |
| } |
| } else { |
| /* see comments on bfq_pos_tree_add_move() for the unlikely() */ |
| if (unlikely(!bfqd->nonrot_with_queueing)) |
| bfq_pos_tree_add_move(bfqd, bfqq); |
| } |
| |
| if (rq->cmd_flags & REQ_META) |
| bfqq->meta_pending--; |
| |
| } |
| |
| static bool bfq_bio_merge(struct request_queue *q, struct bio *bio, |
| unsigned int nr_segs) |
| { |
| struct bfq_data *bfqd = q->elevator->elevator_data; |
| struct request *free = NULL; |
| /* |
| * bfq_bic_lookup grabs the queue_lock: invoke it now and |
| * store its return value for later use, to avoid nesting |
| * queue_lock inside the bfqd->lock. We assume that the bic |
| * returned by bfq_bic_lookup does not go away before |
| * bfqd->lock is taken. |
| */ |
| struct bfq_io_cq *bic = bfq_bic_lookup(bfqd, current->io_context, q); |
| bool ret; |
| |
| spin_lock_irq(&bfqd->lock); |
| |
| if (bic) |
| bfqd->bio_bfqq = bic_to_bfqq(bic, op_is_sync(bio->bi_opf)); |
| else |
| bfqd->bio_bfqq = NULL; |
| bfqd->bio_bic = bic; |
| |
| ret = blk_mq_sched_try_merge(q, bio, nr_segs, &free); |
| |
| 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; |
| /* |
| * The above assignment schedules the following redirections: |
| * each time some I/O for bfqq arrives, the process that |
| * generated that I/O is disassociated from bfqq and |
| * associated with new_bfqq. Here we increases new_bfqq->ref |
| * in advance, adding the number of processes that are |
| * expected to be associated with new_bfqq as they happen to |
| * issue I/O. |
| */ |
| new_bfqq->ref += process_refs; |
| return new_bfqq; |
| } |
| |
| static bool bfq_may_be_close_cooperator(struct bfq_queue *bfqq, |
| struct bfq_queue *new_bfqq) |
| { |
| if (bfq_too_late_for_merging(new_bfqq)) |
| return false; |
| |
| if (bfq_class_idle(bfqq) || bfq_class_idle(new_bfqq) || |
| (bfqq->ioprio_class != new_bfqq->ioprio_class)) |
| return false; |
| |
| /* |
| * If either of the queues has already been detected as seeky, |
| * then merging it with the other queue is unlikely to lead to |
| * sequential I/O. |
| */ |
| if (BFQQ_SEEKY(bfqq) || BFQQ_SEEKY(new_bfqq)) |
| return false; |
| |
| /* |
| * Interleaved I/O is known to be done by (some) applications |
| * only for reads, so it does not make sense to merge async |
| * queues. |
| */ |
| if (!bfq_bfqq_sync(bfqq) || !bfq_bfqq_sync(new_bfqq)) |
| return false; |
| |
| return true; |
| } |
| |
| /* |
| * 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; |
| |
| /* if a merge has already been setup, then proceed with that first */ |
| if (bfqq->new_bfqq) |
| return 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 (!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; |
| } |
| } |
| |
| void bfq_release_process_ref(struct bfq_data *bfqd, struct bfq_queue *bfqq) |
| { |
| /* |
| * To prevent bfqq's service guarantees from being violated, |
| * bfqq may be left busy, i.e., queued for service, even if |
| * empty (see comments in __bfq_bfqq_expire() for |
| * details). But, if no process will send requests to bfqq any |
| * longer, then there is no point in keeping bfqq queued for |
| * service. In addition, keeping bfqq queued for service, but |
| * with no process ref any longer, may have caused bfqq to be |
| * freed when dequeued from service. But this is assumed to |
| * never happen. |
| */ |
| if (bfq_bfqq_busy(bfqq) && RB_EMPTY_ROOT(&bfqq->sort_list) && |
| bfqq != bfqd->in_service_queue) |
| bfq_del_bfqq_busy(bfqd, bfqq, false); |
| |
| bfq_put_queue(bfqq); |
| } |
| |
| static void |
| bfq_merge_bfqqs(struct bfq_data *bfqd, struct bfq_io_cq *bic, |
| struct bfq_queue *bfqq, struct bfq_queue *new_bfqq) |
| { |
| bfq_log_bfqq(bfqd, bfqq, "merging with queue %lu", |
| (unsigned long)new_bfqq->pid); |
| /* Save weight raising and idle window of the merged queues */ |
| bfq_bfqq_save_state(bfqq); |
| bfq_bfqq_save_state(new_bfqq); |
| if (bfq_bfqq_IO_bound(bfqq)) |
| bfq_mark_bfqq_IO_bound(new_bfqq); |
| bfq_clear_bfqq_IO_bound(bfqq); |
| |
| /* |
| * If bfqq is weight-raised, then let new_bfqq inherit |
| * weight-raising. To reduce false positives, neglect the case |
| * where bfqq has just been created, but has not yet made it |
| * to be weight-raised (which may happen because EQM may merge |
| * bfqq even before bfq_add_request is executed for the first |
| * time for bfqq). Handling this case would however be very |
| * easy, thanks to the flag just_created. |
| */ |
| if (new_bfqq->wr_coeff == 1 && bfqq->wr_coeff > 1) { |
| new_bfqq->wr_coeff = bfqq->wr_coeff; |
| new_bfqq->wr_cur_max_time = bfqq->wr_cur_max_time; |
| new_bfqq->last_wr_start_finish = bfqq->last_wr_start_finish; |
| new_bfqq->wr_start_at_switch_to_srt = |
| bfqq->wr_start_at_switch_to_srt; |
| if (bfq_bfqq_busy(new_bfqq)) |
| bfqd->wr_busy_queues++; |
| new_bfqq->entity.prio_changed = 1; |
| } |
| |
| if (bfqq->wr_coeff > 1) { /* bfqq has given its wr to new_bfqq */ |
| bfqq->wr_coeff = 1; |
| bfqq->entity.prio_changed = 1; |
| if (bfq_bfqq_busy(bfqq)) |
| bfqd->wr_busy_queues--; |
| } |
| |
| bfq_log_bfqq(bfqd, new_bfqq, "merge_bfqqs: wr_busy %d", |
| bfqd->wr_busy_queues); |
| |
| /* |
| * Merge queues (that is, let bic redirect its requests to new_bfqq) |
| */ |
| bic_set_bfqq(bic, new_bfqq, 1); |
| bfq_mark_bfqq_coop(new_bfqq); |
| /* |
| * new_bfqq now belongs to at least two bics (it is a shared queue): |
| * set new_bfqq->bic to NULL. bfqq either: |
| * - does not belong to any bic any more, and hence bfqq->bic must |
| * be set to NULL, or |
| * - is a queue whose owning bics have already been redirected to a |
| * different queue, hence the queue is destined to not belong to |
| * any bic soon and bfqq->bic is already NULL (therefore the next |
| * assignment causes no harm). |
| */ |
| new_bfqq->bic = NULL; |
| /* |
| * If the queue is shared, the pid is the pid of one of the associated |
| * processes. Which pid depends on the exact sequence of merge events |
| * the queue underwent. So printing such a pid is useless and confusing |
| * because it reports a random pid between those of the associated |
| * processes. |
| * We mark such a queue with a pid -1, and then print SHARED instead of |
| * a pid in logging messages. |
| */ |
| new_bfqq->pid = -1; |
| bfqq->bic = NULL; |
| bfq_release_process_ref(bfqd, bfqq); |
| } |
| |
| static bool bfq_allow_bio_merge(struct request_queue *q, struct request *rq, |
| struct bio *bio) |
| { |
| struct bfq_data *bfqd = q->elevator->elevator_data; |
| bool is_sync = op_is_sync(bio->bi_opf); |
| struct bfq_queue *bfqq = bfqd->bio_bfqq, *new_bfqq; |
| |
| /* |
| * Disallow merge of a sync bio into an async request. |
| */ |
| if (is_sync && !rq_is_sync(rq)) |
| return false; |
| |
| /* |
| * Lookup the bfqq that this bio will be queued with. Allow |
| * merge only if rq is queued there. |
| */ |
| if (!bfqq) |
| return false; |
| |
| /* |
| * We take advantage of this function to perform an early merge |
| * of the queues of possible cooperating processes. |
| */ |
| new_bfqq = bfq_setup_cooperator(bfqd, bfqq, bio, false); |
| if (new_bfqq) { |
| /* |
| * bic still points to bfqq, then it has not yet been |
| * redirected to some other bfq_queue, and a queue |
| * merge between bfqq and new_bfqq can be safely |
| * fulfilled, i.e., bic can be redirected to new_bfqq |
| * and bfqq can be put. |
| */ |
| bfq_merge_bfqqs(bfqd, bfqd->bio_bic, bfqq, |
| new_bfqq); |
| /* |
| * If we get here, bio will be queued into new_queue, |
| * so use new_bfqq to decide whether bio and rq can be |
| * merged. |
| */ |
| bfqq = new_bfqq; |
| |
| /* |
| * Change also bqfd->bio_bfqq, as |
| * bfqd->bio_bic now points to new_bfqq, and |
| * this function may be invoked again (and then may |
| * use again bqfd->bio_bfqq). |
| */ |
| bfqd->bio_bfqq = bfqq; |
| } |
| |
| return bfqq == RQ_BFQQ(rq); |
| } |
| |
| /* |
| * Set the maximum time for the in-service queue to consume its |
| * budget. This prevents seeky processes from lowering the throughput. |
| * In practice, a time-slice service scheme is used with seeky |
| * processes. |
| */ |
| static void bfq_set_budget_timeout(struct bfq_data *bfqd, |
| struct bfq_queue *bfqq) |
| { |
| unsigned int timeout_coeff; |
| |
| if (bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time) |
| timeout_coeff = 1; |
| else |
| timeout_coeff = bfqq->entity.weight / bfqq->entity.orig_weight; |
| |
| bfqd->last_budget_start = ktime_get(); |
| |
| bfqq->budget_timeout = jiffies + |
| bfqd->bfq_timeout * timeout_coeff; |
| } |
| |
| static void __bfq_set_in_service_queue(struct bfq_data *bfqd, |
| struct bfq_queue *bfqq) |
| { |
| if (bfqq) { |
| bfq_clear_bfqq_fifo_expire(bfqq); |
| |
| bfqd->budgets_assigned = (bfqd->budgets_assigned * 7 + 256) / 8; |
| |
| if (time_is_before_jiffies(bfqq->last_wr_start_finish) && |
| bfqq->wr_coeff > 1 && |
| bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time && |
| time_is_before_jiffies(bfqq->budget_timeout)) { |
| /* |
| * For soft real-time queues, move the start |
| * of the weight-raising period forward by the |
| * time the queue has not received any |
| * service. Otherwise, a relatively long |
| * service delay is likely to cause the |
| * weight-raising period of the queue to end, |
| * because of the short duration of the |
| * weight-raising period of a soft real-time |
| * queue. It is worth noting that this move |
| * is not so dangerous for the other queues, |
| * because soft real-time queues are not |
| * greedy. |
| * |
| * To not add a further variable, we use the |
| * overloaded field budget_timeout to |
| * determine for how long the queue has not |
| * received service, i.e., how much time has |
| * elapsed since the queue expired. However, |
| * this is a little imprecise, because |
| * budget_timeout is set to jiffies if bfqq |
| * not only expires, but also remains with no |
| * request. |
| */ |
| if (time_after(bfqq->budget_timeout, |
| bfqq->last_wr_start_finish)) |
| bfqq->last_wr_start_finish += |
| jiffies - bfqq->budget_timeout; |
| else |
| bfqq->last_wr_start_finish = jiffies; |
| } |
| |
| bfq_set_budget_timeout(bfqd, bfqq); |
| bfq_log_bfqq(bfqd, bfqq, |
| "set_in_service_queue, cur-budget = %d", |
| bfqq->entity.budget); |
| } |
| |
| bfqd->in_service_queue = bfqq; |
| bfqd->in_serv_last_pos = 0; |
| } |
| |
| /* |
| * Get and set a new queue for service. |
| */ |
| static struct bfq_queue *bfq_set_in_service_queue(struct bfq_data *bfqd) |
| { |
| struct bfq_queue *bfqq = bfq_get_next_queue(bfqd); |
| |
| __bfq_set_in_service_queue(bfqd, bfqq); |
| return bfqq; |
| } |
| |
| static void bfq_arm_slice_timer(struct bfq_data *bfqd) |
| { |
| struct bfq_queue *bfqq = bfqd->in_service_queue; |
| u32 sl; |
| |
| bfq_mark_bfqq_wait_request(bfqq); |
| |
| /* |
| * We don't want to idle for seeks, but we do want to allow |
| * fair distribution of slice time for a process doing back-to-back |
| * seeks. So allow a little bit of time for him to submit a new rq. |
| */ |
| sl = bfqd->bfq_slice_idle; |
| /* |
| * Unless the queue is being weight-raised or the scenario is |
| * asymmetric, grant only minimum idle time if the queue |
| * is seeky. A long idling is preserved for a weight-raised |
| * queue, or, more in general, in an asymmetric scenario, |
| * because a long idling is needed for guaranteeing to a queue |
| * its reserved share of the throughput (in particular, it is |
| * needed if the queue has a higher weight than some other |
| * queue). |
| */ |
| if (BFQQ_SEEKY(bfqq) && bfqq->wr_coeff == 1 && |
| !bfq_asymmetric_scenario(bfqd, bfqq)) |
| sl = min_t(u64, sl, BFQ_MIN_TT); |
| else if (bfqq->wr_coeff > 1) |
| sl = max_t(u32, sl, 20ULL * NSEC_PER_MSEC); |
| |
| bfqd->last_idling_start = ktime_get(); |
| bfqd->last_idling_start_jiffies = jiffies; |
| |
| hrtimer_start(&bfqd->idle_slice_timer, ns_to_ktime(sl), |
| HRTIMER_MODE_REL); |
| bfqg_stats_set_start_idle_time(bfqq_group(bfqq)); |
| } |
| |
| /* |
| * In autotuning mode, max_budget is dynamically recomputed as the |
| * amount of sectors transferred in timeout at the estimated peak |
| * rate. This enables BFQ to utilize a full timeslice with a full |
| * budget, even if the in-service queue is served at peak rate. And |
| * this maximises throughput with sequential workloads. |
| */ |
| static unsigned long bfq_calc_max_budget(struct bfq_data *bfqd) |
| { |
| return (u64)bfqd->peak_rate * USEC_PER_MSEC * |
| jiffies_to_msecs(bfqd->bfq_timeout)>>BFQ_RATE_SHIFT; |
| } |
| |
| /* |
| * Update parameters related to throughput and responsiveness, as a |
| * function of the estimated peak rate. See comments on |
| * bfq_calc_max_budget(), and on the ref_wr_duration array. |
| */ |
| static void update_thr_responsiveness_params(struct bfq_data *bfqd) |
| { |
| if (bfqd->bfq_user_max_budget == 0) { |
| bfqd->bfq_max_budget = |
| bfq_calc_max_budget(bfqd); |
| bfq_log(bfqd, "new max_budget = %d", bfqd->bfq_max_budget); |
| } |
| } |
| |
| static void bfq_reset_rate_computation(struct bfq_data *bfqd, |
| struct request *rq) |
| { |
| if (rq != NULL) { /* new rq dispatch now, reset accordingly */ |
| bfqd->last_dispatch = bfqd->first_dispatch = ktime_get_ns(); |
| bfqd->peak_rate_samples = 1; |
| bfqd->sequential_samples = 0; |
| bfqd->tot_sectors_dispatched = bfqd->last_rq_max_size = |
| blk_rq_sectors(rq); |
| } else /* no new rq dispatched, just reset the number of samples */ |
| bfqd->peak_rate_samples = 0; /* full re-init on next disp. */ |
| |
| bfq_log(bfqd, |
| "reset_rate_computation at end, sample %u/%u tot_sects %llu", |
| bfqd->peak_rate_samples, bfqd->sequential_samples, |
| bfqd->tot_sectors_dispatched); |
| } |
| |
| static void bfq_update_rate_reset(struct bfq_data *bfqd, struct request *rq) |
| { |
| u32 rate, weight, divisor; |
| |
| /* |
| * For the convergence property to hold (see comments on |
| * bfq_update_peak_rate()) and for the assessment to be |
| * reliable, a minimum number of samples must be present, and |
| * a minimum amount of time must have elapsed. If not so, do |
| * not compute new rate. Just reset parameters, to get ready |
| * for a new evaluation attempt. |
| */ |
| if (bfqd->peak_rate_samples < BFQ_RATE_MIN_SAMPLES || |
| bfqd->delta_from_first < BFQ_RATE_MIN_INTERVAL) |
| goto reset_computation; |
| |
| /* |
| * If a new request completion has occurred after last |
| * dispatch, then, to approximate the rate at which requests |
| * have been served by the device, it is more precise to |
| * extend the observation interval to the last completion. |
| */ |
| bfqd->delta_from_first = |
| max_t(u64, bfqd->delta_from_first, |
| bfqd->last_completion - bfqd->first_dispatch); |
| |
| /* |
| * Rate computed in sects/usec, and not sects/nsec, for |
| * precision issues. |
| */ |
| rate = div64_ul(bfqd->tot_sectors_dispatched<<BFQ_RATE_SHIFT, |
| div_u64(bfqd->delta_from_first, NSEC_PER_USEC)); |
| |
| /* |
| * Peak rate not updated if: |
| * - the percentage of sequential dispatches is below 3/4 of the |
| * total, and rate is below the current estimated peak rate |
| * - rate is unreasonably high (> 20M sectors/sec) |
| */ |
| if ((bfqd->sequential_samples < (3 * bfqd->peak_rate_samples)>>2 && |
| rate <= bfqd->peak_rate) || |
| rate > 20<<BFQ_RATE_SHIFT) |
| goto reset_computation; |
| |
| /* |
| * We have to update the peak rate, at last! To this purpose, |
| * we use a low-pass filter. We compute the smoothing constant |
| * of the filter as a function of the 'weight' of the new |
| * measured rate. |
| * |
| * As can be seen in next formulas, we define this weight as a |
| * quantity proportional to how sequential the workload is, |
| * and to how long the observation time interval is. |
| * |
| * The weight runs from 0 to 8. The maximum value of the |
| * weight, 8, yields the minimum value for the smoothing |
| * constant. At this minimum value for the smoothing constant, |
| * the measured rate contributes for half of the next value of |
| * the estimated peak rate. |
| * |
| * So, the first step is to compute the weight as a function |
| * of how sequential the workload is. Note that the weight |
| * cannot reach 9, because bfqd->sequential_samples cannot |
| * become equal to bfqd->peak_rate_samples, which, in its |
| * turn, holds true because bfqd->sequential_samples is not |
| * incremented for the first sample. |
| */ |
| weight = (9 * bfqd->sequential_samples) / bfqd->peak_rate_samples; |
| |
| /* |
| * Second step: further refine the weight as a function of the |
| * duration of the observation interval. |
| */ |
| weight = min_t(u32, 8, |
| div_u64(weight * bfqd->delta_from_first, |
| BFQ_RATE_REF_INTERVAL)); |
| |
| /* |
| * Divisor ranging from 10, for minimum weight, to 2, for |
| * maximum weight. |
| */ |
| divisor = 10 - weight; |
| |
| /* |
| * Finally, update peak rate: |
| * |
| * peak_rate = peak_rate * (divisor-1) / divisor + rate / divisor |
| */ |
| bfqd->peak_rate *= divisor-1; |
| bfqd->peak_rate /= divisor; |
| rate /= divisor; /* smoothing constant alpha = 1/divisor */ |
| |
| bfqd->peak_rate += rate; |
| |
| /* |
| * For a very slow device, bfqd->peak_rate can reach 0 (see |
| * the minimum representable values reported in the comments |
| * on BFQ_RATE_SHIFT). Push to 1 if this happens, to avoid |
| * divisions by zero where bfqd->peak_rate is used as a |
| * divisor. |
| */ |
| bfqd->peak_rate = max_t(u32, 1, bfqd->peak_rate); |
| |
| update_thr_responsiveness_params(bfqd); |
| |
| reset_computation: |
| bfq_reset_rate_computation(bfqd, rq); |
| } |
| |
| /* |
| * Update the read/write peak rate (the main quantity used for |
| * auto-tuning, see update_thr_responsiveness_params()). |
| * |
| * It is not trivial to estimate the peak rate (correctly): because of |
| * the presence of sw and hw queues between the scheduler and the |
| * device components that finally serve I/O requests, it is hard to |
| * say exactly when a given dispatched request is served inside the |
| * device, and for how long. As a consequence, it is hard to know |
| * precisely at what rate a given set of requests is actually served |
| * by the device. |
| * |
| * On the opposite end, the dispatch time of any request is trivially |
| * available, and, from this piece of information, the "dispatch rate" |
| * of requests can be immediately computed. So, the idea in the next |
| * function is to use what is known, namely request dispatch times |
| * (plus, when useful, request completion times), to estimate what is |
| * unknown, namely in-device request service rate. |
| * |
| * The main issue is that, because of the above facts, the rate at |
| * which a certain set of requests is dispatched over a certain time |
| * interval can vary greatly with respect to the rate at which the |
| * same requests are then served. But, since the size of any |
| * intermediate queue is limited, and the service scheme is lossless |
| * (no request is silently dropped), the following obvious convergence |
| * property holds: the number of requests dispatched MUST become |
| * closer and closer to the number of requests completed as the |
| * observation interval grows. This is the key property used in |
| * the next function to estimate the peak service rate as a function |
| * of the observed dispatch rate. The function assumes to be invoked |
| * on every request dispatch. |
| */ |
| static void bfq_update_peak_rate(struct bfq_data *bfqd, struct request *rq) |
| { |
| u64 now_ns = ktime_get_ns(); |
| |
| if (bfqd->peak_rate_samples == 0) { /* first dispatch */ |
| bfq_log(bfqd, "update_peak_rate: goto reset, samples %d", |
| bfqd->peak_rate_samples); |
| bfq_reset_rate_computation(bfqd, rq); |
| goto update_last_values; /* will add one sample */ |
| } |
| |
| /* |
| * Device idle for very long: the observation interval lasting |
| * up to this dispatch cannot be a valid observation interval |
| * for computing a new peak rate (similarly to the late- |
| * completion event in bfq_completed_request()). Go to |
| * update_rate_and_reset to have the following three steps |
| * taken: |
| * - close the observation interval at the last (previous) |
| * request dispatch or completion |
| * - compute rate, if possible, for that observation interval |
| * - start a new observation interval with this dispatch |
| */ |
| if (now_ns - bfqd->last_dispatch > 100*NSEC_PER_MSEC && |
| bfqd->rq_in_driver == 0) |
| goto update_rate_and_reset; |
| |
| /* Update sampling information */ |
| bfqd->peak_rate_samples++; |
| |
| if ((bfqd->rq_in_driver > 0 || |
| now_ns - bfqd->last_completion < BFQ_MIN_TT) |
| && !BFQ_RQ_SEEKY(bfqd, bfqd->last_position, rq)) |
| bfqd->sequential_samples++; |
| |
| bfqd->tot_sectors_dispatched += blk_rq_sectors(rq); |
| |
| /* Reset max observed rq size every 32 dispatches */ |
| if (likely(bfqd->peak_rate_samples % 32)) |
| bfqd->last_rq_max_size = |
| max_t(u32, blk_rq_sectors(rq), bfqd->last_rq_max_size); |
| else |
| bfqd->last_rq_max_size = blk_rq_sectors(rq); |
| |
| bfqd->delta_from_first = now_ns - bfqd->first_dispatch; |
| |
| /* Target observation interval not yet reached, go on sampling */ |
| if (bfqd->delta_from_first < BFQ_RATE_REF_INTERVAL) |
| goto update_last_values; |
| |
| update_rate_and_reset: |
| bfq_update_rate_reset(bfqd, rq); |
| update_last_values: |
| bfqd->last_position = blk_rq_pos(rq) + blk_rq_sectors(rq); |
| if (RQ_BFQQ(rq) == bfqd->in_service_queue) |
| bfqd->in_serv_last_pos = bfqd->last_position; |
| bfqd->last_dispatch = now_ns; |
| } |
| |
| /* |
| * Remove request from internal lists. |
| */ |
| static void bfq_dispatch_remove(struct request_queue *q, struct request *rq) |
| { |
| struct bfq_queue *bfqq = RQ_BFQQ(rq); |
| |
| /* |
| * For consistency, the next instruction should have been |
| * executed after removing the request from the queue and |
| * dispatching it. We execute instead this instruction before |
| * bfq_remove_request() (and hence introduce a temporary |
| * inconsistency), for efficiency. In fact, should this |
| * dispatch occur for a non in-service bfqq, this anticipated |
| * increment prevents two counters related to bfqq->dispatched |
| * from risking to be, first, uselessly decremented, and then |
| * incremented again when the (new) value of bfqq->dispatched |
| * happens to be taken into account. |
| */ |
| bfqq->dispatched++; |
| bfq_update_peak_rate(q->elevator->elevator_data, rq); |
| |
| bfq_remove_request(q, rq); |
| } |
| |
| /* |
| * There is a case where idling does not have to be performed for |
| * throughput concerns, but to preserve the throughput share of |
| * the process associated with bfqq. |
| * |
| * To introduce this case, we can note that allowing the drive |
| * to enqueue more than one request at a time, and hence |
| * delegating de facto final scheduling decisions to the |
| * drive's internal scheduler, entails loss of control on the |
| * actual request service order. In particular, the critical |
| * situation is when requests from different processes happen |
| * to be present, at the same time, in the internal queue(s) |
| * of the drive. In such a situation, the drive, by deciding |
| * the service order of the internally-queued requests, does |
| * determine also the actual throughput distribution among |
| * these processes. But the drive typically has no notion or |
| * concern about per-process throughput distribution, and |
| * makes its decisions only on a per-request basis. Therefore, |
| * the service distribution enforced by the drive's internal |
| * scheduler is likely to coincide with the desired throughput |
| * distribution only in a completely symmetric, or favorably |
| * skewed scenario where: |
| * (i-a) each of these processes must get the same throughput as |
| * the others, |
| * (i-b) in case (i-a) does not hold, it holds that the process |
| * associated with bfqq must receive a lower or equal |
| * throughput than any of the other processes; |
| * (ii) the I/O of each process has the same properties, in |
| * terms of locality (sequential or random), direction |
| * (reads or writes), request sizes, greediness |
| * (from I/O-bound to sporadic), and so on; |
| |
| * In fact, in such a scenario, the drive tends to treat the requests |
| * of each process in about the same way as the requests of the |
| * others, and thus to provide each of these processes with about the |
| * same throughput. This is exactly the desired throughput |
| * distribution if (i-a) holds, or, if (i-b) holds instead, this is an |
| * even more convenient distribution for (the process associated with) |
| * bfqq. |
| * |
| * In contrast, in any asymmetric or unfavorable scenario, device |
| * idling (I/O-dispatch plugging) is certainly needed to guarantee |
| * that bfqq receives its assigned fraction of the device throughput |
| * (see [1] for details). |
| * |
| * The problem is that idling may significantly reduce throughput with |
| * certain combinations of types of I/O and devices. An important |
| * example is sync random I/O on flash storage with command |
| * queueing. So, unless bfqq falls in cases where idling also boosts |
| * throughput, it is important to check conditions (i-a), i(-b) and |
| * (ii) accurately, so as to avoid idling when not strictly needed for |
| * service guarantees. |
| * |
| * Unfortunately, it is extremely difficult to thoroughly check |
| * condition (ii). And, in case there are active groups, it becomes |
| * very difficult to check conditions (i-a) and (i-b) too. In fact, |
| * if there are active groups, then, for conditions (i-a) or (i-b) to |
| * become false 'indirectly', it is enough that an active group |
| * contains more active processes or sub-groups than some other active |
| * group. More precisely, for conditions (i-a) or (i-b) to become |
| * false because of such a group, it is not even necessary that the |
| * group is (still) active: it is sufficient that, even if the group |
| * has become inactive, some of its descendant processes still have |
| * some request already dispatched but still waiting for |
| * completion. In fact, requests have still to be guaranteed their |
| * share of the throughput even after being dispatched. In this |
| * respect, it is easy to show that, if a group frequently becomes |
| * inactive while still having in-flight requests, and if, when this |
| * happens, the group is not considered in the calculation of whether |
| * the scenario is asymmetric, then the group may fail to be |
| * guaranteed its fair share of the throughput (basically because |
| * idling may not be performed for the descendant processes of the |
| * group, but it had to be). We address this issue with the following |
| * bi-modal behavior, implemented in the function |
| * bfq_asymmetric_scenario(). |
| * |
| * If there are groups with requests waiting for completion |
| * (as commented above, some of these groups may even be |
| * already inactive), then the scenario is tagged as |
| * asymmetric, conservatively, without checking any of the |
| * conditions (i-a), (i-b) or (ii). So the device is idled for bfqq. |
| * This behavior matches also the fact that groups are created |
| * exactly if controlling I/O is a primary concern (to |
| * preserve bandwidth and latency guarantees). |
| * |
| * On the opposite end, if there are no groups with requests waiting |
| * for completion, then only conditions (i-a) and (i-b) are actually |
| * controlled, i.e., provided that conditions (i-a) or (i-b) holds, |
| * idling is not performed, regardless of whether condition (ii) |
| * holds. In other words, only if conditions (i-a) and (i-b) do not |
| * hold, then idling is allowed, and the device tends to be prevented |
| * from queueing many requests, possibly of several processes. Since |
| * there are no groups with requests waiting for completion, then, to |
| * control conditions (i-a) and (i-b) it is enough to check just |
| * whether all the queues with requests waiting for completion also |
| * have the same weight. |
| * |
| * Not checking condition (ii) evidently exposes bfqq to the |
| * risk of getting less throughput than its fair share. |
| * However, for queues with the same weight, a further |
| * mechanism, preemption, mitigates or even eliminates this |
| * problem. And it does so without consequences on overall |
| * throughput. This mechanism and its benefits are explained |
| * in the next three paragraphs. |
| * |
| * Even if a queue, say Q, is expired when it remains idle, Q |
| * can still preempt the new in-service queue if the next |
| * request of Q arrives soon (see the comments on |
| * bfq_bfqq_update_budg_for_activation). If all queues and |
| * groups have the same weight, this form of preemption, |
| * combined with the hole-recovery heuristic described in the |
| * comments on function bfq_bfqq_update_budg_for_activation, |
| * are enough to preserve a correct bandwidth distribution in |
| * the mid term, even without idling. In fact, even if not |
| * idling allows the internal queues of the device to contain |
| * many requests, and thus to reorder requests, we can rather |
| * safely assume that the internal scheduler still preserves a |
| * minimum of mid-term fairness. |
| * |
| * More precisely, this preemption-based, idleless approach |
| * provides fairness in terms of IOPS, and not sectors per |
| * second. This can be seen with a simple example. Suppose |
| * that there are two queues with the same weight, but that |
| * the first queue receives requests of 8 sectors, while the |
| * second queue receives requests of 1024 sectors. In |
| * addition, suppose that each of the two queues contains at |
| * most one request at a time, which implies that each queue |
| * always remains idle after it is served. Finally, after |
| * remaining idle, each queue receives very quickly a new |
| * request. It follows that the two queues are served |
| * alternatively, preempting each other if needed. This |
| * implies that, although both queues have the same weight, |
| * the queue with large requests receives a service that is |
| * 1024/8 times as high as the service received by the other |
| * queue. |
| * |
| * The motivation for using preemption instead of idling (for |
| * queues with the same weight) is that, by not idling, |
| * service guarantees are preserved (completely or at least in |
| * part) without minimally sacrificing throughput. And, if |
| * there is no active group, then the primary expectation for |
| * this device is probably a high throughput. |
| * |
| * We are now left only with explaining the two sub-conditions in the |
| * additional compound condition that is checked below for deciding |
| * whether the scenario is asymmetric. To explain the first |
| * sub-condition, we need to add that the function |
| * bfq_asymmetric_scenario checks the weights of only |
| * non-weight-raised queues, for efficiency reasons (see comments on |
| * bfq_weights_tree_add()). Then the fact that bfqq is weight-raised |
| * is checked explicitly here. More precisely, the compound condition |
| * below takes into account also the fact that, even if bfqq is being |
| * weight-raised, the scenario is still symmetric if all queues with |
| * requests waiting for completion happen to be |
| * weight-raised. Actually, we should be even more precise here, and |
| * differentiate between interactive weight raising and soft real-time |
| * weight raising. |
| * |
| * The second sub-condition checked in the compound condition is |
| * whether there is a fair amount of already in-flight I/O not |
| * belonging to bfqq. If so, I/O dispatching is to be plugged, for the |
| * following reason. The drive may decide to serve in-flight |
| * non-bfqq's I/O requests before bfqq's ones, thereby delaying the |
| * arrival of new I/O requests for bfqq (recall that bfqq is sync). If |
| * I/O-dispatching is not plugged, then, while bfqq remains empty, a |
| * basically uncontrolled amount of I/O from other queues may be |
| * dispatched too, possibly causing the service of bfqq's I/O to be |
| * delayed even longer in the drive. This problem gets more and more |
| * serious as the speed and the queue depth of the drive grow, |
| * because, as these two quantities grow, the probability to find no |
| * queue busy but many requests in flight grows too. By contrast, |
| * plugging I/O dispatching minimizes the delay induced by already |
| * in-flight I/O, and enables bfqq to recover the bandwidth it may |
| * lose because of this delay. |
| * |
| * As a side note, it is worth considering that the above |
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