| Author | Tokens | Token Proportion | Commits | Commit Proportion |
|---|---|---|---|---|
| Paolo Valente | 14172 | 69.46% | 108 | 45.76% |
| Arianna Avanzini | 3926 | 19.24% | 4 | 1.69% |
| Jan Kara | 813 | 3.98% | 18 | 7.63% |
| Davide Zini | 321 | 1.57% | 3 | 1.27% |
| Federico Gavioli | 237 | 1.16% | 1 | 0.42% |
| Christoph Hellwig | 157 | 0.77% | 17 | 7.20% |
| yu kuai | 153 | 0.75% | 13 | 5.51% |
| Bart Van Assche | 145 | 0.71% | 6 | 2.54% |
| Jens Axboe | 98 | 0.48% | 11 | 4.66% |
| Khazhismel Kumykov | 44 | 0.22% | 2 | 0.85% |
| Federico Motta | 29 | 0.14% | 1 | 0.42% |
| Francesco Pollicino | 26 | 0.13% | 2 | 0.85% |
| Tejun Heo | 25 | 0.12% | 2 | 0.85% |
| Pietro Pedroni | 24 | 0.12% | 1 | 0.42% |
| Luca Miccio | 24 | 0.12% | 2 | 0.85% |
| KeMeng Shi | 22 | 0.11% | 5 | 2.12% |
| Angelo Ruocco | 19 | 0.09% | 3 | 1.27% |
| Ming Lei | 18 | 0.09% | 2 | 0.85% |
| Zhang Wensheng | 16 | 0.08% | 1 | 0.42% |
| Colin Ian King | 14 | 0.07% | 3 | 1.27% |
| weiping zhang | 14 | 0.07% | 2 | 0.85% |
| Damien Le Moal | 13 | 0.06% | 4 | 1.69% |
| Zhiqiang Liu | 13 | 0.06% | 1 | 0.42% |
| Davide Sapienza | 12 | 0.06% | 2 | 0.85% |
| Hou Tao | 11 | 0.05% | 1 | 0.42% |
| qiulaibin | 7 | 0.03% | 1 | 0.42% |
| Stephen Bates | 6 | 0.03% | 1 | 0.42% |
| Omar Sandoval | 6 | 0.03% | 2 | 0.85% |
| Yufen Yu | 6 | 0.03% | 1 | 0.42% |
| Yuwei Guan | 5 | 0.02% | 1 | 0.42% |
| Ben Hutchings | 5 | 0.02% | 1 | 0.42% |
| Gustavo A. R. Silva | 4 | 0.02% | 1 | 0.42% |
| Andrew Morton | 3 | 0.01% | 2 | 0.85% |
| John Garry | 2 | 0.01% | 1 | 0.42% |
| Doug Anderson | 2 | 0.01% | 1 | 0.42% |
| Jia Cheng Hu | 2 | 0.01% | 1 | 0.42% |
| Joseph Qi | 1 | 0.00% | 1 | 0.42% |
| Mauro Carvalho Chehab | 1 | 0.00% | 1 | 0.42% |
| Randy Dunlap | 1 | 0.00% | 1 | 0.42% |
| Chaitanya Kulkarni | 1 | 0.00% | 1 | 0.42% |
| Konstantin Khlebnikov | 1 | 0.00% | 1 | 0.42% |
| Neil Brown | 1 | 0.00% | 1 | 0.42% |
| Filippo Muzzini | 1 | 0.00% | 1 | 0.42% |
| Luca Mariotti | 1 | 0.00% | 1 | 0.42% |
| Total | 20402 | 236 |
// 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/ktime.h> #include <linux/rbtree.h> #include <linux/ioprio.h> #include <linux/sbitmap.h> #include <linux/delay.h> #include <linux/backing-dev.h> #include <trace/events/block.h> #include "elevator.h" #include "blk.h" #include "blk-mq.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); #undef BFQ_BFQQ_FNS \ /* Expiration time of async (0) and sync (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; /* * Maximum time between the creation of two queues, for stable merge * to be activated (in ms) */ static const unsigned long bfq_activation_stable_merging = 600; /* * Minimum time to be waited before evaluating delayed stable merge (in ms) */ static const unsigned long bfq_late_stable_merging = 600; #define RQ_BIC(rq) ((struct bfq_io_cq *)((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, unsigned int actuator_idx) { if (is_sync) return bic->bfqq[1][actuator_idx]; return bic->bfqq[0][actuator_idx]; } static void bfq_put_stable_ref(struct bfq_queue *bfqq); void bic_set_bfqq(struct bfq_io_cq *bic, struct bfq_queue *bfqq, bool is_sync, unsigned int actuator_idx) { struct bfq_queue *old_bfqq = bic->bfqq[is_sync][actuator_idx]; /* * If bfqq != NULL, then a non-stable queue merge between * bic->bfqq and bfqq is happening here. This causes troubles * in the following case: bic->bfqq has also been scheduled * for a possible stable merge with bic->stable_merge_bfqq, * and bic->stable_merge_bfqq == bfqq happens to * hold. Troubles occur because bfqq may then undergo a split, * thereby becoming eligible for a stable merge. Yet, if * bic->stable_merge_bfqq points exactly to bfqq, then bfqq * would be stably merged with itself. To avoid this anomaly, * we cancel the stable merge if * bic->stable_merge_bfqq == bfqq. */ struct bfq_iocq_bfqq_data *bfqq_data = &bic->bfqq_data[actuator_idx]; /* Clear bic pointer if bfqq is detached from this bic */ if (old_bfqq && old_bfqq->bic == bic) old_bfqq->bic = NULL; if (is_sync) bic->bfqq[1][actuator_idx] = bfqq; else bic->bfqq[0][actuator_idx] = bfqq; if (bfqq && bfqq_data->stable_merge_bfqq == bfqq) { /* * Actually, these same instructions are executed also * in bfq_setup_cooperator, in case of abort or actual * execution of a stable merge. We could avoid * repeating these instructions there too, but if we * did so, we would nest even more complexity in this * function. */ bfq_put_stable_ref(bfqq_data->stable_merge_bfqq); bfqq_data->stable_merge_bfqq = NULL; } } 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. * @q: the request queue. */ static struct bfq_io_cq *bfq_bic_lookup(struct request_queue *q) { struct bfq_io_cq *icq; unsigned long flags; if (!current->io_context) return NULL; spin_lock_irqsave(&q->queue_lock, flags); icq = icq_to_bic(ioc_lookup_icq(q)); spin_unlock_irqrestore(&q->queue_lock, flags); return icq; } /* * 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) { lockdep_assert_held(&bfqd->lock); 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_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; } } #define BFQ_LIMIT_INLINE_DEPTH 16 #ifdef CONFIG_BFQ_GROUP_IOSCHED static bool bfqq_request_over_limit(struct bfq_queue *bfqq, int limit) { struct bfq_data *bfqd = bfqq->bfqd; struct bfq_entity *entity = &bfqq->entity; struct bfq_entity *inline_entities[BFQ_LIMIT_INLINE_DEPTH]; struct bfq_entity **entities = inline_entities; int depth, level, alloc_depth = BFQ_LIMIT_INLINE_DEPTH; int class_idx = bfqq->ioprio_class - 1; struct bfq_sched_data *sched_data; unsigned long wsum; bool ret = false; if (!entity->on_st_or_in_serv) return false; retry: spin_lock_irq(&bfqd->lock); /* +1 for bfqq entity, root cgroup not included */ depth = bfqg_to_blkg(bfqq_group(bfqq))->blkcg->css.cgroup->level + 1; if (depth > alloc_depth) { spin_unlock_irq(&bfqd->lock); if (entities != inline_entities) kfree(entities); entities = kmalloc_array(depth, sizeof(*entities), GFP_NOIO); if (!entities) return false; alloc_depth = depth; goto retry; } sched_data = entity->sched_data; /* Gather our ancestors as we need to traverse them in reverse order */ level = 0; for_each_entity(entity) { /* * If at some level entity is not even active, allow request * queueing so that BFQ knows there's work to do and activate * entities. */ if (!entity->on_st_or_in_serv) goto out; /* Uh, more parents than cgroup subsystem thinks? */ if (WARN_ON_ONCE(level >= depth)) break; entities[level++] = entity; } WARN_ON_ONCE(level != depth); for (level--; level >= 0; level--) { entity = entities[level]; if (level > 0) { wsum = bfq_entity_service_tree(entity)->wsum; } else { int i; /* * For bfqq itself we take into account service trees * of all higher priority classes and multiply their * weights so that low prio queue from higher class * gets more requests than high prio queue from lower * class. */ wsum = 0; for (i = 0; i <= class_idx; i++) { wsum = wsum * IOPRIO_BE_NR + sched_data->service_tree[i].wsum; } } if (!wsum) continue; limit = DIV_ROUND_CLOSEST(limit * entity->weight, wsum); if (entity->allocated >= limit) { bfq_log_bfqq(bfqq->bfqd, bfqq, "too many requests: allocated %d limit %d level %d", entity->allocated, limit, level); ret = true; break; } } out: spin_unlock_irq(&bfqd->lock); if (entities != inline_entities) kfree(entities); return ret; } #else static bool bfqq_request_over_limit(struct bfq_queue *bfqq, int limit) { return false; } #endif /* * 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. * * Also if a bfq queue or its parent cgroup consume more tags than would be * appropriate for their weight, we trim the available tag depth to 1. This * avoids a situation where one cgroup can starve another cgroup from tags and * thus block service differentiation among cgroups. Note that because the * queue / cgroup already has many requests allocated and queued, this does not * significantly affect service guarantees coming from the BFQ scheduling * algorithm. */ static void bfq_limit_depth(blk_opf_t opf, struct blk_mq_alloc_data *data) { struct bfq_data *bfqd = data->q->elevator->elevator_data; struct bfq_io_cq *bic = bfq_bic_lookup(data->q); int depth; unsigned limit = data->q->nr_requests; unsigned int act_idx; /* Sync reads have full depth available */ if (op_is_sync(opf) && !op_is_write(opf)) { depth = 0; } else { depth = bfqd->word_depths[!!bfqd->wr_busy_queues][op_is_sync(opf)]; limit = (limit * depth) >> bfqd->full_depth_shift; } for (act_idx = 0; bic && act_idx < bfqd->num_actuators; act_idx++) { struct bfq_queue *bfqq = bic_to_bfqq(bic, op_is_sync(opf), act_idx); /* * Does queue (or any parent entity) exceed number of * requests that should be available to it? Heavily * limit depth so that it cannot consume more * available requests and thus starve other entities. */ if (bfqq && bfqq_request_over_limit(bfqq, limit)) { depth = 1; break; } } bfq_log(bfqd, "[%s] wr_busy %d sync %d depth %u", __func__, bfqd->wr_busy_queues, op_is_sync(opf), depth); if (depth) data->shallow_depth = 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 = &bfqq_group(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 is at most one active group. * 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 > 1 #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_queue *bfqq) { struct rb_root_cached *root = &bfqq->bfqd->queue_weights_tree; 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_queue *bfqq) { struct rb_root_cached *root; if (!bfqq->weight_counter) return; root = &bfqq->bfqd->queue_weights_tree; 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); } /* * 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 || blk_time_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; dur = bfqd->rate_dur_prod; do_div(dur, bfqd->peak_rate); /* * Limit duration between 3 and 25 seconds. The upper limit * has been conservatively set after the following worst case: * on a QEMU/KVM virtual machine * - running in a slow PC * - with a virtual disk stacked on a slow low-end 5400rpm HDD * - serving a heavy I/O workload, such as the sequential reading * of several files * mplayer took 23 seconds to start, if constantly weight-raised. * * As for higher values than that accommodating the above bad * scenario, tests show that higher values would often yield * the opposite of the desired result, i.e., would worsen * responsiveness by allowing non-interactive applications to * preserve weight raising for too long. * * On the other end, lower values than 3 seconds make it * difficult for most interactive tasks to complete their jobs * before weight-raising finishes. */ return clamp_val(dur, msecs_to_jiffies(3000), msecs_to_jiffies(25000)); } /* switch back from soft real-time to interactive weight raising */ static void switch_back_to_interactive_wr(struct bfq_queue *bfqq, struct bfq_data *bfqd) { bfqq->wr_coeff = bfqd->bfq_wr_coeff; bfqq->wr_cur_max_time = bfq_wr_duration(bfqd); bfqq->last_wr_start_finish = bfqq->wr_start_at_switch_to_srt; } static void bfq_bfqq_resume_state(struct bfq_queue *bfqq, struct bfq_data *bfqd, struct bfq_io_cq *bic, bool bfq_already_existing) { unsigned int old_wr_coeff = 1; bool busy = bfq_already_existing && bfq_bfqq_busy(bfqq); unsigned int a_idx = bfqq->actuator_idx; struct bfq_iocq_bfqq_data *bfqq_data = &bic->bfqq_data[a_idx]; if (bfqq_data->saved_has_short_ttime) bfq_mark_bfqq_has_short_ttime(bfqq); else bfq_clear_bfqq_has_short_ttime(bfqq); if (bfqq_data->saved_IO_bound) bfq_mark_bfqq_IO_bound(bfqq); else bfq_clear_bfqq_IO_bound(bfqq); bfqq->last_serv_time_ns = bfqq_data->saved_last_serv_time_ns; bfqq->inject_limit = bfqq_data->saved_inject_limit; bfqq->decrease_time_jif = bfqq_data->saved_decrease_time_jif; bfqq->entity.new_weight = bfqq_data->saved_weight; bfqq->ttime = bfqq_data->saved_ttime; bfqq->io_start_time = bfqq_data->saved_io_start_time; bfqq->tot_idle_time = bfqq_data->saved_tot_idle_time; /* * Restore weight coefficient only if low_latency is on */ if (bfqd->low_latency) { old_wr_coeff = bfqq->wr_coeff; bfqq->wr_coeff = bfqq_data->saved_wr_coeff; } bfqq->service_from_wr = bfqq_data->saved_service_from_wr; bfqq->wr_start_at_switch_to_srt = bfqq_data->saved_wr_start_at_switch_to_srt; bfqq->last_wr_start_finish = bfqq_data->saved_last_wr_start_finish; bfqq->wr_cur_max_time = bfqq_data->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->entity.allocated - bfqq->entity.on_st_or_in_serv - (bfqq->weight_counter != NULL) - bfqq->stable_ref; } /* Empty burst list and add just bfqq (see comments on bfq_handle_burst) */ static void bfq_reset_burst_list(struct bfq_data *bfqd, struct bfq_queue *bfqq) { struct bfq_queue *item; struct hlist_node *n; hlist_for_each_entry_safe(item, n, &bfqd->burst_list, burst_list_node) hlist_del_init(&item->burst_list_node); /* * Start the creation of a new burst list only if there is no * active queue. See comments on the conditional invocation of * bfq_handle_burst(). */ if (bfq_tot_busy_queues(bfqd) == 0) { hlist_add_head(&bfqq->burst_list_node, &bfqd->burst_list); bfqd->burst_size = 1; } else bfqd->burst_size = 0; bfqd->burst_parent_entity = bfqq->entity.parent; } /* Add bfqq to the list of queues in current burst (see bfq_handle_burst) */ static void bfq_add_to_burst(struct bfq_data *bfqd, struct bfq_queue *bfqq) { /* Increment burst size to take into account also bfqq */ bfqd->burst_size++; if (bfqd->burst_size == bfqd->bfq_large_burst_thresh) { struct bfq_queue *pos, *bfqq_item; struct hlist_node *n; /* * Enough queues have been activated shortly after each * other to consider this burst as large. */ bfqd->large_burst = true; /* * We can now mark all queues in the burst list as * belonging to a large burst. */ hlist_for_each_entry(bfqq_item, &bfqd->burst_list, burst_list_node) bfq_mark_bfqq_in_large_burst(bfqq_item); bfq_mark_bfqq_in_large_burst(bfqq); /* * From now on, and until the current burst finishes, any * new queue being activated shortly after the last queue * was inserted in the burst can be immediately marked as * belonging to a large burst. So the burst list is not * needed any more. Remove it. */ hlist_for_each_entry_safe(pos, n, &bfqd->burst_list, burst_list_node) hlist_del_init(&pos->burst_list_node); } else /* * Burst not yet large: add bfqq to the burst list. Do * not increment the ref counter for bfqq, because bfqq * is removed from the burst list before freeing bfqq * in put_queue. */ hlist_add_head(&bfqq->burst_list_node, &bfqd->burst_list); } /* * If many queues belonging to the same group happen to be created * shortly after each other, then the processes associated with these * queues have typically a common goal. In particular, bursts of queue * creations are usually caused by services or applications that spawn * many parallel threads/processes. Examples are systemd during boot, * or git grep. To help these processes get their job done as soon as * possible, it is usually better to not grant either weight-raising * or device idling to their queues, unless these queues must be * protected from the I/O flowing through other active queues. * * In this comment we describe, firstly, the reasons why this fact * holds, and, secondly, the next function, which implements the main * steps needed to properly mark these queues so that they can then be * treated in a different way. * * The above services or applications benefit mostly from a high * throughput: the quicker the requests of the activated queues are * cumulatively served, the sooner the target job of these queues gets * completed. As a consequence, weight-raising any of these queues, * which also implies idling the device for it, is almost always * counterproductive, unless there are other active queues to isolate * these new queues from. If there no other active queues, then * weight-raising these new queues just lowers throughput in most * cases. * * On the other hand, a burst of queue creations may be caused also by * the start of an application that does not consist of a lot of * parallel I/O-bound threads. In fact, with a complex application, * several short processes may need to be executed to start-up the * application. In this respect, to start an application as quickly as * possible, the best thing to do is in any case to privilege the I/O * related to the application with respect to all other * I/O. Therefore, the best strategy to start as quickly as possible * an application that causes a burst of queue creations is to * weight-raise all the queues created during the burst. This is the * exact opposite of the best strategy for the other type of bursts. * * In the end, to take the best action for each of the two cases, the * two types of bursts need to be distinguished. Fortunately, this * seems relatively easy, by looking at the sizes of the bursts. In * particular, we found a threshold such that only bursts with a * larger size than that threshold are apparently caused by * services or commands such as systemd or git grep. For brevity, * hereafter we call just 'large' these bursts. BFQ *does not* * weight-raise queues whose creation occurs in a large burst. In * addition, for each of these queues BFQ performs or does not perform * idling depending on which choice boosts the throughput more. The * exact choice depends on the device and request pattern at * hand. * * Unfortunately, false positives may occur while an interactive task * is starting (e.g., an application is being started). The * consequence is that the queues associated with the task do not * enjoy weight raising as expected. Fortunately these false positives * are very rare. They typically occur if some service happens to * start doing I/O exactly when the interactive task starts. * * Turning back to the next function, it is invoked only if there are * no active queues (apart from active queues that would belong to the * same, possible burst bfqq would belong to), and it implements all * the steps needed to detect the occurrence of a large burst and to * properly mark all the queues belonging to it (so that they can then * be treated in a different way). This goal is achieved by * maintaining a "burst list" that holds, temporarily, the queues that * belong to the burst in progress. The list is then used to mark * these queues as belonging to a large burst if the burst does become * large. The main steps are the following. * * . when the very first queue is created, the queue is inserted into the * list (as it could be the first queue in a possible burst) * * . if the current burst has not yet become large, and a queue Q that does * not yet belong to the burst is activated shortly after the last time * at which a new queue entered the burst list, then the function appends * Q to the burst list * * . if, as a consequence of the previous step, the burst size reaches * the large-burst threshold, then * * . all the queues in the burst list are marked as belonging to a * large burst * * . the burst list is deleted; in fact, the burst list already served * its purpose (keeping temporarily track of the queues in a burst, * so as to be able to mark them as belonging to a large burst in the * previous sub-step), and now is not needed any more * * . the device enters a large-burst mode * * . if a queue Q that does not belong to the burst is created while * the device is in large-burst mode and shortly after the last time * at which a queue either entered the burst list or was marked as * belonging to the current large burst, then Q is immediately marked * as belonging to a large burst. * * . if a queue Q that does not belong to the burst is created a while * later, i.e., not shortly after, than the last time at which a queue * either entered the burst list or was marked as belonging to the * current large burst, then the current burst is deemed as finished and: * * . the large-burst mode is reset if set * * . the burst list is emptied * * . Q is inserted in the burst list, as Q may be the first queue * in a possible new burst (then the burst list contains just Q * after this step). */ static void bfq_handle_burst(struct bfq_data *bfqd, struct bfq_queue *bfqq) { /* * If bfqq is already in the burst list or is part of a large * burst, or finally has just been split, then there is * nothing else to do. */ if (!hlist_unhashed(&bfqq->burst_list_node) || bfq_bfqq_in_large_burst(bfqq) || time_is_after_eq_jiffies(bfqq->split_time + msecs_to_jiffies(10))) return; /* * If bfqq's creation happens late enough, or bfqq belongs to * a different group than the burst group, then the current * burst is finished, and related data structures must be * reset. * * In this respect, consider the special case where bfqq is * the very first queue created after BFQ is selected for this * device. In this case, last_ins_in_burst and * burst_parent_entity are not yet significant when we get * here. But it is easy to verify that, whether or not the * following condition is true, bfqq will end up being * inserted into the burst list. In particular the list will * happen to contain only bfqq. And this is exactly what has * to happen, as bfqq may be the first queue of the first * burst. */ if (time_is_before_jiffies(bfqd->last_ins_in_burst + bfqd->bfq_burst_interval) || bfqq->entity.parent != bfqd->burst_parent_entity) { bfqd->large_burst = false; bfq_reset_burst_list(bfqd, bfqq); goto end; } /* * If we get here, then bfqq is being activated shortly after the * last queue. So, if the current burst is also large, we can mark * bfqq as belonging to this large burst immediately. */ if (bfqd->large_burst) { bfq_mark_bfqq_in_large_burst(bfqq); goto end; } /* * If we get here, then a large-burst state has not yet been * reached, but bfqq is being activated shortly after the last * queue. Then we add bfqq to the burst. */ bfq_add_to_burst(bfqd, bfqq); end: /* * At this point, bfqq either has been added to the current * burst or has caused the current burst to terminate and a * possible new burst to start. In particular, in the second * case, bfqq has become the first queue in the possible new * burst. In both cases last_ins_in_burst needs to be moved * forward. */ bfqd->last_ins_in_burst = jiffies; } static int bfq_bfqq_budget_left(struct bfq_queue *bfqq) { struct bfq_entity *entity = &bfqq->entity; return entity->budget - entity->service; } /* * If enough samples have been computed, return the current max budget * stored in bfqd, which is dynamically updated according to the * estimated disk peak rate; otherwise return the default max budget */ static int bfq_max_budget(struct bfq_data *bfqd) { if (bfqd->budgets_assigned < bfq_stats_min_budgets) return bfq_default_max_budget; else return bfqd->bfq_max_budget; } /* * Return min budget, which is a fraction of the current or default * max budget (trying with 1/32) */ static int bfq_min_budget(struct bfq_data *bfqd) { if (bfqd->budgets_assigned < bfq_stats_min_budgets) return bfq_default_max_budget / 32; else return bfqd->bfq_max_budget / 32; } /* * The next function, invoked after the input queue bfqq switches from * idle to busy, updates the budget of bfqq. The function also tells * whether the in-service queue should be expired, by returning * true. The purpose of expiring the in-service queue is to give bfqq * the chance to possibly preempt the in-service queue, and the reason * for preempting the in-service queue is to achieve one of the two * goals below. * * 1. Guarantee to bfqq its reserved bandwidth even if bfqq has * expired because it has remained idle. In particular, bfqq may have * expired for one of the following two reasons: * * - BFQQE_NO_MORE_REQUESTS bfqq did not enjoy any device idling * and did not make it to issue a new request before its last * request was served; * * - BFQQE_TOO_IDLE bfqq did enjoy device idling, but did not issue * a new request before the expiration of the idling-time. * * Even if bfqq has expired for one of the above reasons, the process * associated with the queue may be however issuing requests greedily, * and thus be sensitive to the bandwidth it receives (bfqq may have * remained idle for other reasons: CPU high load, bfqq not enjoying * idling, I/O throttling somewhere in the path from the process to * the I/O scheduler, ...). But if, after every expiration for one of * the above two reasons, bfqq has to wait for the service of at least * one full budget of another queue before being served again, then * bfqq is likely to get a much lower bandwidth or resource time than * its reserved ones. To address this issue, two countermeasures need * to be taken. * * First, the budget and the timestamps of bfqq need to be updated in * a special way on bfqq reactivation: they need to be updated as if * bfqq did not remain idle and did not expire. In fact, if they are * computed as if bfqq expired and remained idle until reactivation, * then the process associated with bfqq is treated as if, instead of * being greedy, it stopped issuing requests when bfqq remained idle, * and restarts issuing requests only on this reactivation. In other * words, the scheduler does not help the process recover the "service * hole" between bfqq expiration and reactivation. As a consequence, * the process receives a lower bandwidth than its reserved one. In * contrast, to recover this hole, the budget must be updated as if * bfqq was not expired at all before this reactivation, i.e., it must * be set to the value of the remaining budget when bfqq was * expired. Along the same line, timestamps need to be assigned the * value they had the last time bfqq was selected for service, i.e., * before last expiration. Thus timestamps need to be back-shifted * with respect to their normal computation (see [1] for more details * on this tricky aspect). * * Secondly, to allow the process to recover the hole, the in-service * queue must be expired too, to give bfqq the chance to preempt it * immediately. In fact, if bfqq has to wait for a full budget of the * in-service queue to be completed, then it may become impossible to * let the process recover the hole, even if the back-shifted * timestamps of bfqq are lower than those of the in-service queue. If * this happens for most or all of the holes, then the process may not * receive its reserved bandwidth. In this respect, it is worth noting * that, being the service of outstanding requests unpreemptible, a * little fraction of the holes may however be unrecoverable, thereby * causing a little loss of bandwidth. * * The last important point is detecting whether bfqq does need this * bandwidth recovery. In this respect, the next function deems the * process associated with bfqq greedy, and thus allows it to recover * the hole, if: 1) the process is waiting for the arrival of a new * request (which implies that bfqq expired for one of the above two * reasons), and 2) such a request has arrived soon. The first * condition is controlled through the flag non_blocking_wait_rq, * while the second through the flag arrived_in_time. If both * conditions hold, then the function computes the budget in the * above-described special way, and signals that the in-service queue * should be expired. Timestamp back-shifting is done later in * __bfq_activate_entity. * * 2. Reduce latency. Even if timestamps are not backshifted to let * the process associated with bfqq recover a service hole, bfqq may * however happen to have, after being (re)activated, a lower finish * timestamp than the in-service queue. That is, the next budget of * bfqq may have to be completed before the one of the in-service * queue. If this is the case, then preempting the in-service queue * allows this goal to be achieved, apart from the unpreemptible, * outstanding requests mentioned above. * * Unfortunately, regardless of which of the above two goals one wants * to achieve, service trees need first to be updated to know whether * the in-service queue must be preempted. To have service trees * correctly updated, the in-service queue must be expired and * rescheduled, and bfqq must be scheduled too. This is one of the * most costly operations (in future versions, the scheduling * mechanism may be re-designed in such a way to make it possible to * know whether preemption is needed without needing to update service * trees). In addition, queue preemptions almost always cause random * I/O, which may in turn cause loss of throughput. Finally, there may * even be no in-service queue when the next function is invoked (so, * no queue to compare timestamps with). Because of these facts, the * next function adopts the following simple scheme to avoid costly * operations, too frequent preemptions and too many dependencies on * the state of the scheduler: it requests the expiration of the * in-service queue (unconditionally) only for queues that need to * recover a hole. Then it delegates to other parts of the code the * responsibility of handling the above case 2. */ static bool bfq_bfqq_update_budg_for_activation(struct bfq_data *bfqd, struct bfq_queue *bfqq, bool arrived_in_time) { struct bfq_entity *entity = &bfqq->entity; /* * In the next compound condition, we check also whether there * is some budget left, because otherwise there is no point in * trying to go on serving bfqq with this same budget: bfqq * would be expired immediately after being selected for * service. This would only cause useless overhead. */ if (bfq_bfqq_non_blocking_wait_rq(bfqq) && arrived_in_time && bfq_bfqq_budget_left(bfqq) > 0) { /* * We do not clear the flag non_blocking_wait_rq here, as * the latter is used in bfq_activate_bfqq to signal * that timestamps need to be back-shifted (and is * cleared right after). */ /* * In next assignment we rely on that either * entity->service or entity->budget are not updated * on expiration if bfqq is empty (see * __bfq_bfqq_recalc_budget). Thus both quantities * remain unchanged after such an expiration, and the * following statement therefore assigns to * entity->budget the remaining budget on such an * expiration. */ entity->budget = min_t(unsigned long, bfq_bfqq_budget_left(bfqq), bfqq->max_budget); /* * At this point, we have used entity->service to get * the budget left (needed for updating * entity->budget). Thus we finally can, and have to, * reset entity->service. The latter must be reset * because bfqq would otherwise be charged again for * the service it has received during its previous * service slot(s). */ entity->service = 0; return true; } /* * We can finally complete expiration, by setting service to 0. */ entity->service = 0; entity->budget = max_t(unsigned long, bfqq->max_budget, bfq_serv_to_charge(bfqq->next_rq, bfqq)); bfq_clear_bfqq_non_blocking_wait_rq(bfqq); return false; } /* * Return the farthest past time instant according to jiffies * macros. */ static unsigned long bfq_smallest_from_now(void) { return jiffies - MAX_JIFFY_OFFSET; } static void bfq_update_bfqq_wr_on_rq_arrival(struct bfq_data *bfqd, struct bfq_queue *bfqq, unsigned int old_wr_coeff, bool wr_or_deserves_wr, bool interactive, bool in_burst, bool soft_rt) { if (old_wr_coeff == 1 && wr_or_deserves_wr) { /* start a weight-raising period */ if (interactive) { bfqq->service_from_wr = 0; bfqq->wr_coeff = bfqd->bfq_wr_coeff; bfqq->wr_cur_max_time = bfq_wr_duration(bfqd); } else { /* * No interactive weight raising in progress * here: assign minus infinity to * wr_start_at_switch_to_srt, to make sure * that, at the end of the soft-real-time * weight raising periods that is starting * now, no interactive weight-raising period * may be wrongly considered as still in * progress (and thus actually started by * mistake). */ bfqq->wr_start_at_switch_to_srt = bfq_smallest_from_now(); bfqq->wr_coeff = bfqd->bfq_wr_coeff * BFQ_SOFTRT_WEIGHT_FACTOR; bfqq->wr_cur_max_time = bfqd->bfq_wr_rt_max_time; } /* * If needed, further reduce budget to make sure it is * close to bfqq's backlog, so as to reduce the * scheduling-error component due to a too large * budget. Do not care about throughput consequences, * but only about latency. Finally, do not assign a * too small budget either, to avoid increasing * latency by causing too frequent expirations. */ bfqq->entity.budget = min_t(unsigned long, bfqq->entity.budget, 2 * bfq_min_budget(bfqd)); } else if (old_wr_coeff > 1) { if (interactive) { /* update wr coeff and duration */ bfqq->wr_coeff = bfqd->bfq_wr_coeff; bfqq->wr_cur_max_time = bfq_wr_duration(bfqd); } else if (in_burst) bfqq->wr_coeff = 1; else if (soft_rt) { /* * The application is now or still meeting the * requirements for being deemed soft rt. We * can then correctly and safely (re)charge * the weight-raising duration for the * application with the weight-raising * duration for soft rt applications. * * In particular, doing this recharge now, i.e., * before the weight-raising period for the * application finishes, reduces the probability * of the following negative scenario: * 1) the weight of a soft rt application is * raised at startup (as for any newly * created application), * 2) since the application is not interactive, * at a certain time weight-raising is * stopped for the application, * 3) at that time the application happens to * still have pending requests, and hence * is destined to not have a chance to be * deemed soft rt before these requests are * completed (see the comments to the * function bfq_bfqq_softrt_next_start() * for details on soft rt detection), * 4) these pending requests experience a high * latency because the application is not * weight-raised while they are pending. */ if (bfqq->wr_cur_max_time != bfqd->bfq_wr_rt_max_time) { bfqq->wr_start_at_switch_to_srt = bfqq->last_wr_start_finish; bfqq->wr_cur_max_time = bfqd->bfq_wr_rt_max_time; bfqq->wr_coeff = bfqd->bfq_wr_coeff * BFQ_SOFTRT_WEIGHT_FACTOR; } bfqq->last_wr_start_finish = jiffies; } } } static bool bfq_bfqq_idle_for_long_time(struct bfq_data *bfqd, struct bfq_queue *bfqq) { return bfqq->dispatched == 0 && time_is_before_jiffies( bfqq->budget_timeout + bfqd->bfq_wr_min_idle_time); } /* * Return true if bfqq is in a higher priority class, or has a higher * weight than the in-service queue. */ static bool bfq_bfqq_higher_class_or_weight(struct bfq_queue *bfqq, struct bfq_queue *in_serv_bfqq) { int bfqq_weight, in_serv_weight; if (bfqq->ioprio_class < in_serv_bfqq->ioprio_class) return true; if (in_serv_bfqq->entity.parent == bfqq->entity.parent) { bfqq_weight = bfqq->entity.weight; in_serv_weight = in_serv_bfqq->entity.weight; } else { if (bfqq->entity.parent) bfqq_weight = bfqq->entity.parent->weight; else bfqq_weight = bfqq->entity.weight; if (in_serv_bfqq->entity.parent) in_serv_weight = in_serv_bfqq->entity.parent->weight; else in_serv_weight = in_serv_bfqq->entity.weight; } return bfqq_weight > in_serv_weight; } /* * Get the index of the actuator that will serve bio. */ static unsigned int bfq_actuator_index(struct bfq_data *bfqd, struct bio *bio) { unsigned int i; sector_t end; /* no search needed if one or zero ranges present */ if (bfqd->num_actuators == 1) return 0; /* bio_end_sector(bio) gives the sector after the last one */ end = bio_end_sector(bio) - 1; for (i = 0; i < bfqd->num_actuators; i++) { if (end >= bfqd->sector[i] && end < bfqd->sector[i] + bfqd->nr_sectors[i]) return i; } WARN_ONCE(true, "bfq_actuator_index: bio sector out of ranges: end=%llu\n", end); return 0; } static bool bfq_better_to_idle(struct bfq_queue *bfqq); static void bfq_bfqq_handle_idle_busy_switch(struct bfq_data *bfqd, struct bfq_queue *bfqq, int old_wr_coeff, struct request *rq, bool *interactive) { bool soft_rt, in_burst, wr_or_deserves_wr, bfqq_wants_to_preempt, idle_for_long_time = bfq_bfqq_idle_for_long_time(bfqd, bfqq), /* * See the comments on * bfq_bfqq_update_budg_for_activation for * details on the usage of the next variable. */ arrived_in_time = blk_time_get_ns() <= bfqq->ttime.last_end_request + bfqd->bfq_slice_idle * 3; unsigned int act_idx = bfq_actuator_index(bfqd, rq->bio); bool bfqq_non_merged_or_stably_merged = bfqq->bic || RQ_BIC(rq)->bfqq_data[act_idx].stably_merged; /* * bfqq deserves to be weight-raised if: * - it is sync, * - it does not belong to a large burst, * - it has been idle for enough time or is soft real-time, * - is linked to a bfq_io_cq (it is not shared in any sense), * - has a default weight (otherwise we assume the user wanted * to control its weight explicitly) */ in_burst = bfq_bfqq_in_large_burst(bfqq); soft_rt = bfqd->bfq_wr_max_softrt_rate > 0 && !BFQQ_TOTALLY_SEEKY(bfqq) && !in_burst && time_is_before_jiffies(bfqq->soft_rt_next_start) && bfqq->dispatched == 0 && bfqq->entity.new_weight == 40; *interactive = !in_burst && idle_for_long_time && bfqq->entity.new_weight == 40; /* * Merged bfq_queues are kept out of weight-raising * (low-latency) mechanisms. The reason is that these queues * are usually created for non-interactive and * non-soft-real-time tasks. Yet this is not the case for * stably-merged queues. These queues are merged just because * they are created shortly after each other. So they may * easily serve the I/O of an interactive or soft-real time * application, if the application happens to spawn multiple * processes. So let also stably-merged queued enjoy weight * raising. */ wr_or_deserves_wr = bfqd->low_latency && (bfqq->wr_coeff > 1 || (bfq_bfqq_sync(bfqq) && bfqq_non_merged_or_stably_merged && (*interactive || soft_rt))); /* * Using the last flag, update budget and check whether bfqq * may want to preempt the in-service queue. */ bfqq_wants_to_preempt = bfq_bfqq_update_budg_for_activation(bfqd, bfqq, arrived_in_time); /* * If bfqq happened to be activated in a burst, but has been * idle for much more than an interactive queue, then we * assume that, in the overall I/O initiated in the burst, the * I/O associated with bfqq is finished. So bfqq does not need * to be treated as a queue belonging to a burst * anymore. Accordingly, we reset bfqq's in_large_burst flag * if set, and remove bfqq from the burst list if it's * there. We do not decrement burst_size, because the fact * that bfqq does not need to belong to the burst list any * more does not invalidate the fact that bfqq was created in * a burst. */ if (likely(!bfq_bfqq_just_created(bfqq)) && idle_for_long_time && time_is_before_jiffies( bfqq->budget_timeout + msecs_to_jiffies(10000))) { hlist_del_init(&bfqq->burst_list_node); bfq_clear_bfqq_in_large_burst(bfqq); } bfq_clear_bfqq_just_created(bfqq); if (bfqd->low_latency) { if (unlikely(time_is_after_jiffies(bfqq->split_time))) /* wraparound */ bfqq->split_time = jiffies - bfqd->bfq_wr_min_idle_time - 1; if (time_is_before_jiffies(bfqq->split_time + bfqd->bfq_wr_min_idle_time)) { bfq_update_bfqq_wr_on_rq_arrival(bfqd, bfqq, old_wr_coeff, wr_or_deserves_wr, *interactive, in_burst, soft_rt); if (old_wr_coeff != bfqq->wr_coeff) bfqq->entity.prio_changed = 1; } } bfqq->last_idle_bklogged = jiffies; bfqq->service_from_backlogged = 0; bfq_clear_bfqq_softrt_update(bfqq); bfq_add_bfqq_busy(bfqq); /* * Expire in-service queue if preemption may be needed for * guarantees or throughput. As for guarantees, we care * explicitly about two cases. The first is that bfqq has to * recover a service hole, as explained in the comments on * bfq_bfqq_update_budg_for_activation(), i.e., that * bfqq_wants_to_preempt is true. However, if bfqq does not * carry time-critical I/O, then bfqq's bandwidth is less * important than that of queues that carry time-critical I/O. * So, as a further constraint, we consider this case only if * bfqq is at least as weight-raised, i.e., at least as time * critical, as the in-service queue. * * The second case is that bfqq is in a higher priority class, * or has a higher weight than the in-service queue. If this * condition does not hold, we don't care because, even if * bfqq does not start to be served immediately, the resulting * delay for bfqq's I/O is however lower or much lower than * the ideal completion time to be guaranteed to bfqq's I/O. * * In both cases, preemption is needed only if, according to * the timestamps of both bfqq and of the in-service queue, * bfqq actually is the next queue to serve. So, to reduce * useless preemptions, the return value of * next_queue_may_preempt() is considered in the next compound * condition too. Yet next_queue_may_preempt() just checks a * simple, necessary condition for bfqq to be the next queue * to serve. In fact, to evaluate a sufficient condition, the * timestamps of the in-service queue would need to be * updated, and this operation is quite costly (see the * comments on bfq_bfqq_update_budg_for_activation()). * * As for throughput, we ask bfq_better_to_idle() whether we * still need to plug I/O dispatching. If bfq_better_to_idle() * says no, then plugging is not needed any longer, either to * boost throughput or to perserve service guarantees. Then * the best option is to stop plugging I/O, as not doing so * would certainly lower throughput. We may end up in this * case if: (1) upon a dispatch attempt, we detected that it * was better to plug I/O dispatch, and to wait for a new * request to arrive for the currently in-service queue, but * (2) this switch of bfqq to busy changes the scenario. */ if (bfqd->in_service_queue && ((bfqq_wants_to_preempt && bfqq->wr_coeff >= bfqd->in_service_queue->wr_coeff) || bfq_bfqq_higher_class_or_weight(bfqq, bfqd->in_service_queue) || !bfq_better_to_idle(bfqd->in_service_queue)) && next_queue_may_preempt(bfqd)) bfq_bfqq_expire(bfqd, bfqd->in_service_queue, false, BFQQE_PREEMPTED); } static void bfq_reset_inject_limit(struct bfq_data *bfqd, struct bfq_queue *bfqq) { /* invalidate baseline total service time */ bfqq->last_serv_time_ns = 0; /* * Reset pointer in case we are waiting for * some request completion. */ bfqd->waited_rq = NULL; /* * If bfqq has a short think time, then start by setting the * inject limit to 0 prudentially, because the service time of * an injected I/O request may be higher than the think time * of bfqq, and therefore, if one request was injected when * bfqq remains empty, this injected request might delay the * service of the next I/O request for bfqq significantly. In * case bfqq can actually tolerate some injection, then the * adaptive update will however raise the limit soon. This * lucky circumstance holds exactly because bfqq has a short * think time, and thus, after remaining empty, is likely to * get new I/O enqueued---and then completed---before being * expired. This is the very pattern that gives the * limit-update algorithm the chance to measure the effect of * injection on request service times, and then to update the * limit accordingly. * * However, in the following special case, the inject limit is * left to 1 even if the think time is short: bfqq's I/O is * synchronized with that of some other queue, i.e., bfqq may * receive new I/O only after the I/O of the other queue is * completed. Keeping the inject limit to 1 allows the * blocking I/O to be served while bfqq is in service. And * this is very convenient both for bfqq and for overall * throughput, as explained in detail in the comments in * bfq_update_has_short_ttime(). * * On the opposite end, if bfqq has a long think time, then * start directly by 1, because: * a) on the bright side, keeping at most one request in * service in the drive is unlikely to cause any harm to the * latency of bfqq's requests, as the service time of a single * request is likely to be lower than the think time of bfqq; * b) on the downside, after becoming empty, bfqq is likely to * expire before getting its next request. With this request * arrival pattern, it is very hard to sample total service * times and update the inject limit accordingly (see comments * on bfq_update_inject_limit()). So the limit is likely to be * never, or at least seldom, updated. As a consequence, by * setting the limit to 1, we avoid that no injection ever * occurs with bfqq. On the downside, this proactive step * further reduces chances to actually compute the baseline * total service time. Thus it reduces chances to execute the * limit-update algorithm and possibly raise the limit to more * than 1. */ if (bfq_bfqq_has_short_ttime(bfqq)) bfqq->inject_limit = 0; else bfqq->inject_limit = 1; bfqq->decrease_time_jif = jiffies; } static void bfq_update_io_intensity(struct bfq_queue *bfqq, u64 now_ns) { u64 tot_io_time = now_ns - bfqq->io_start_time; if (RB_EMPTY_ROOT(&bfqq->sort_list) && bfqq->dispatched == 0) bfqq->tot_idle_time += now_ns - bfqq->ttime.last_end_request; if (unlikely(bfq_bfqq_just_created(bfqq))) return; /* * Must be busy for at least about 80% of the time to be * considered I/O bound. */ if (bfqq->tot_idle_time * 5 > tot_io_time) bfq_clear_bfqq_IO_bound(bfqq); else bfq_mark_bfqq_IO_bound(bfqq); /* * Keep an observation window of at most 200 ms in the past * from now. */ if (tot_io_time > 200 * NSEC_PER_MSEC) { bfqq->io_start_time = now_ns - (tot_io_time>>1); bfqq->tot_idle_time >>= 1; } } /* * Detect whether bfqq's I/O seems synchronized with that of some * other queue, i.e., whether bfqq, after remaining empty, happens to * receive new I/O only right after some I/O request of the other * queue has been completed. We call waker queue the other queue, and * we assume, for simplicity, that bfqq may have at most one waker * queue. * * A remarkable throughput boost can be reached by unconditionally * injecting the I/O of the waker queue, every time a new * bfq_dispatch_request happens to be invoked while I/O is being * plugged for bfqq. In addition to boosting throughput, this * unblocks bfqq's I/O, thereby improving bandwidth and latency for * bfqq. Note that these same results may be achieved with the general * injection mechanism, but less effectively. For details on this * aspect, see the comments on the choice of the queue for injection * in bfq_select_queue(). * * Turning back to the detection of a waker queue, a queue Q is deemed as a * waker queue for bfqq if, for three consecutive times, bfqq happens to become * non empty right after a request of Q has been completed within given * timeout. In this respect, even if bfqq is empty, we do not check for a waker * if it still has some in-flight I/O. In fact, in this case bfqq is actually * still being served by the drive, and may receive new I/O on the completion * of some of the in-flight requests. In particular, on the first time, Q is * tentatively set as a candidate waker queue, while on the third consecutive * time that Q is detected, the field waker_bfqq is set to Q, to confirm that Q * is a waker queue for bfqq. These detection steps are performed only if bfqq * has a long think time, so as to make it more likely that bfqq's I/O is * actually being blocked by a synchronization. This last filter, plus the * above three-times requirement and time limit for detection, make false * positives less likely. * * NOTE * * The sooner a waker queue is detected, the sooner throughput can be * boosted by injecting I/O from the waker queue. Fortunately, * detection is likely to be actually fast, for the following * reasons. While blocked by synchronization, bfqq has a long think * time. This implies that bfqq's inject limit is at least equal to 1 * (see the comments in bfq_update_inject_limit()). So, thanks to * injection, the waker queue is likely to be served during the very * first I/O-plugging time interval for bfqq. This triggers the first * step of the detection mechanism. Thanks again to injection, the * candidate waker queue is then likely to be confirmed no later than * during the next I/O-plugging interval for bfqq. * * ISSUE * * On queue merging all waker information is lost. */ static void bfq_check_waker(struct bfq_data *bfqd, struct bfq_queue *bfqq, u64 now_ns) { char waker_name[MAX_BFQQ_NAME_LENGTH]; if (!bfqd->last_completed_rq_bfqq || bfqd->last_completed_rq_bfqq == bfqq || bfq_bfqq_has_short_ttime(bfqq) || now_ns - bfqd->last_completion >= 4 * NSEC_PER_MSEC || bfqd->last_completed_rq_bfqq == &bfqd->oom_bfqq || bfqq == &bfqd->oom_bfqq) return; /* * We reset waker detection logic also if too much time has passed * since the first detection. If wakeups are rare, pointless idling * doesn't hurt throughput that much. The condition below makes sure * we do not uselessly idle blocking waker in more than 1/64 cases. */ if (bfqd->last_completed_rq_bfqq != bfqq->tentative_waker_bfqq || now_ns > bfqq->waker_detection_started + 128 * (u64)bfqd->bfq_slice_idle) { /* * First synchronization detected with a * candidate waker queue, or with a different * candidate waker queue from the current one. */ bfqq->tentative_waker_bfqq = bfqd->last_completed_rq_bfqq; bfqq->num_waker_detections = 1; bfqq->waker_detection_started = now_ns; bfq_bfqq_name(bfqq->tentative_waker_bfqq, waker_name, MAX_BFQQ_NAME_LENGTH); bfq_log_bfqq(bfqd, bfqq, "set tentative waker %s", waker_name); } else /* Same tentative waker queue detected again */ bfqq->num_waker_detections++; if (bfqq->num_waker_detections == 3) { bfqq->waker_bfqq = bfqd->last_completed_rq_bfqq; bfqq->tentative_waker_bfqq = NULL; bfq_bfqq_name(bfqq->waker_bfqq, waker_name, MAX_BFQQ_NAME_LENGTH); bfq_log_bfqq(bfqd, bfqq, "set waker %s", waker_name); /* * If the waker queue disappears, then * bfqq->waker_bfqq must be reset. To * this goal, we maintain in each * waker queue a list, woken_list, of * all the queues that reference the * waker queue through their * waker_bfqq pointer. When the waker * queue exits, the waker_bfqq pointer * of all the queues in the woken_list * is reset. * * In addition, if bfqq is already in * the woken_list of a waker queue, * then, before being inserted into * the woken_list of a new waker * queue, bfqq must be removed from * the woken_list of the old waker * queue. */ if (!hlist_unhashed(&bfqq->woken_list_node)) hlist_del_init(&bfqq->woken_list_node); hlist_add_head(&bfqq->woken_list_node, &bfqd->last_completed_rq_bfqq->woken_list); } } static void bfq_add_request(struct request *rq) { struct bfq_queue *bfqq = RQ_BFQQ(rq); struct bfq_data *bfqd = bfqq->bfqd; struct request *next_rq, *prev; unsigned int old_wr_coeff = bfqq->wr_coeff; bool interactive = false; u64 now_ns = blk_time_get_ns(); bfq_log_bfqq(bfqd, bfqq, "add_request %d", rq_is_sync(rq)); bfqq->queued[rq_is_sync(rq)]++; /* * Updating of 'bfqd->queued' is protected by 'bfqd->lock', however, it * may be read without holding the lock in bfq_has_work(). */ WRITE_ONCE(bfqd->queued, bfqd->queued + 1); if (bfq_bfqq_sync(bfqq) && RQ_BIC(rq)->requests <= 1) { bfq_check_waker(bfqd, bfqq, now_ns); /* * Periodically reset inject limit, to make sure that * the latter eventually drops in case workload * changes, see step (3) in the comments on * bfq_update_inject_limit(). */ if (time_is_before_eq_jiffies(bfqq->decrease_time_jif + msecs_to_jiffies(1000))) bfq_reset_inject_limit(bfqd, bfqq); /* * The following conditions must hold to setup a new * sampling of total service time, and then a new * update of the inject limit: * - bfqq is in service, because the total service * time is evaluated only for the I/O requests of * the queues in service; * - this is the right occasion to compute or to * lower the baseline total service time, because * there are actually no requests in the drive, * or * the baseline total service time is available, and * this is the right occasion to compute the other * quantity needed to update the inject limit, i.e., * the total service time caused by the amount of * injection allowed by the current value of the * limit. It is the right occasion because injection * has actually been performed during the service * hole, and there are still in-flight requests, * which are very likely to be exactly the injected * requests, or part of them; * - the minimum interval for sampling the total * service time and updating the inject limit has * elapsed. */ if (bfqq == bfqd->in_service_queue && (bfqd->tot_rq_in_driver == 0 || (bfqq->last_serv_time_ns > 0 && bfqd->rqs_injected && bfqd->tot_rq_in_driver > 0)) && time_is_before_eq_jiffies(bfqq->decrease_time_jif + msecs_to_jiffies(10))) { bfqd->last_empty_occupied_ns = blk_time_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->tot_rq_in_driver == 0) bfqd->rqs_injected = false; } } if (bfq_bfqq_sync(bfqq)) bfq_update_io_intensity(bfqq, now_ns); elv_rb_add(&bfqq->sort_list, rq); /* * Check if this request is a better next-serve candidate. */ prev = bfqq->next_rq; next_rq = bfq_choose_req(bfqd, bfqq->next_rq, rq, bfqd->last_position); bfqq->next_rq = next_rq; /* * Adjust priority tree position, if next_rq changes. * See comments on bfq_pos_tree_add_move() for the unlikely(). */ if (unlikely(!bfqd->nonrot_with_queueing && prev != bfqq->next_rq)) bfq_pos_tree_add_move(bfqd, bfqq); if (!bfq_bfqq_busy(bfqq)) /* switching to busy ... */ bfq_bfqq_handle_idle_busy_switch(bfqd, bfqq, old_wr_coeff, rq, &interactive); else { if (bfqd->low_latency && old_wr_coeff == 1 && !rq_is_sync(rq) && time_is_before_jiffies( bfqq->last_wr_start_finish + bfqd->bfq_wr_min_inter_arr_async)) { bfqq->wr_coeff = bfqd->bfq_wr_coeff; bfqq->wr_cur_max_time = bfq_wr_duration(bfqd); bfqd->wr_busy_queues++; bfqq->entity.prio_changed = 1; } if (prev != bfqq->next_rq) bfq_updated_next_req(bfqd, bfqq); } /* * Assign jiffies to last_wr_start_finish in the following * cases: * * . if bfqq is not going to be weight-raised, because, for * non weight-raised queues, last_wr_start_finish stores the * arrival time of the last request; as of now, this piece * of information is used only for deciding whether to * weight-raise async queues * * . if bfqq is not weight-raised, because, if bfqq is now * switching to weight-raised, then last_wr_start_finish * stores the time when weight-raising starts * * . if bfqq is interactive, because, regardless of whether * bfqq is currently weight-raised, the weight-raising * period must start or restart (this case is considered * separately because it is not detected by the above * conditions, if bfqq is already weight-raised) * * last_wr_start_finish has to be updated also if bfqq is soft * real-time, because the weight-raising period is constantly * restarted on idle-to-busy transitions for these queues, but * this is already done in bfq_bfqq_handle_idle_busy_switch if * needed. */ if (bfqd->low_latency && (old_wr_coeff == 1 || bfqq->wr_coeff == 1 || interactive)) bfqq->last_wr_start_finish = jiffies; } static struct request *bfq_find_rq_fmerge(struct bfq_data *bfqd, struct bio *bio, struct request_queue *q) { struct bfq_queue *bfqq = bfqd->bio_bfqq; if (bfqq) return elv_rb_find(&bfqq->sort_list, bio_end_sector(bio)); return NULL; } static sector_t get_sdist(sector_t last_pos, struct request *rq) { if (last_pos) return abs(blk_rq_pos(rq) - last_pos); return 0; } 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]--; /* * Updating of 'bfqd->queued' is protected by 'bfqd->lock', however, it * may be read without holding the lock in bfq_has_work(). */ WRITE_ONCE(bfqd->queued, bfqd->queued - 1); 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(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(q); bool ret; spin_lock_irq(&bfqd->lock); if (bic) { /* * Make sure cgroup info is uptodate for current process before * considering the merge. */ bfq_bic_update_cgroup(bic, bio); bfqd->bio_bfqq = bic_to_bfqq(bic, op_is_sync(bio->bi_opf), bfq_actuator_index(bfqd, bio)); } else { bfqd->bio_bfqq = NULL; } bfqd->bio_bic = bic; ret = blk_mq_sched_try_merge(q, bio, nr_segs, &free); spin_unlock_irq(&bfqd->lock); if (free) blk_mq_free_request(free); return ret; } static int bfq_request_merge(struct request_queue *q, struct request **req, struct bio *bio) { struct bfq_data *bfqd = q->elevator->elevator_data; struct request *__rq; __rq = bfq_find_rq_fmerge(bfqd, bio, q); if (__rq && elv_bio_merge_ok(__rq, bio)) { *req = __rq; if (blk_discard_mergable(__rq)) return ELEVATOR_DISCARD_MERGE; return ELEVATOR_FRONT_MERGE; } return ELEVATOR_NO_MERGE; } static 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 = RQ_BFQQ(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 = RQ_BFQQ(rq), *next_bfqq = RQ_BFQQ(next); if (!bfqq) goto remove; /* * If next and rq belong to the same bfq_queue and next is older * than rq, then reposition rq in the fifo (by substituting next * with rq). Otherwise, if next and rq belong to different * bfq_queues, never reposition rq: in fact, we would have to * reposition it with respect to next's position in its own fifo, * which would most certainly be too expensive with respect to * the benefits. */ if (bfqq == next_bfqq && !list_empty(&rq->queuelist) && !list_empty(&next->queuelist) && next->fifo_time < rq->fifo_time) { list_del_init(&rq->queuelist); list_replace_init(&next->queuelist, &rq->queuelist); rq->fifo_time = next->fifo_time; } if (bfqq->next_rq == next) bfqq->next_rq = rq; bfqg_stats_update_io_merged(bfqq_group(bfqq), next->cmd_flags); remove: /* Merged request may be in the IO scheduler. Remove it. */ if (!RB_EMPTY_NODE(&next->rb_node)) { bfq_remove_request(next->q, next); if (next_bfqq) bfqg_stats_update_io_remove(bfqq_group(next_bfqq), next->cmd_flags); } } /* Must be called with bfqq != NULL */ static void bfq_bfqq_end_wr(struct bfq_queue *bfqq) { /* * If bfqq has been enjoying interactive weight-raising, then * reset soft_rt_next_start. We do it for the following * reason. bfqq may have been conveying the I/O needed to load * a soft real-time application. Such an application actually * exhibits a soft real-time I/O pattern after it finishes * loading, and finally starts doing its job. But, if bfqq has * been receiving a lot of bandwidth so far (likely to happen * on a fast device), then soft_rt_next_start now contains a * high value that. So, without this reset, bfqq would be * prevented from being possibly considered as soft_rt for a * very long time. */ if (bfqq->wr_cur_max_time != bfqq->bfqd->bfq_wr_rt_max_time) bfqq->soft_rt_next_start = jiffies; if (bfq_bfqq_busy(bfqq)) bfqq->bfqd->wr_busy_queues--; bfqq->wr_coeff = 1; bfqq->wr_cur_max_time = 0; bfqq->last_wr_start_finish = jiffies; /* * Trigger a weight change on the next invocation of * __bfq_entity_update_weight_prio. */ bfqq->entity.prio_changed = 1; } void bfq_end_wr_async_queues(struct bfq_data *bfqd, struct bfq_group *bfqg) { int i, j, k; for (k = 0; k < bfqd->num_actuators; k++) { for (i = 0; i < 2; i++) for (j = 0; j < IOPRIO_NR_LEVELS; j++) if (bfqg->async_bfqq[i][j][k]) bfq_bfqq_end_wr(bfqg->async_bfqq[i][j][k]); if (bfqg->async_idle_bfqq[k]) bfq_bfqq_end_wr(bfqg->async_idle_bfqq[k]); } } static void bfq_end_wr(struct bfq_data *bfqd) { struct bfq_queue *bfqq; int i; spin_lock_irq(&bfqd->lock); for (i = 0; i < bfqd->num_actuators; i++) { list_for_each_entry(bfqq, &bfqd->active_list[i], 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 = &bfqq_group(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; /* * Make sure merged queues belong to the same parent. Parents could * have changed since the time we decided the two queues are suitable * for merging. */ if (new_bfqq->entity.parent != bfqq->entity.parent) return NULL; bfq_log_bfqq(bfqq->bfqd, bfqq, "scheduling merge with queue %d", new_bfqq->pid); /* * Merging is just a redirection: the requests of the process * owning one of the two queues are redirected to the other queue. * The latter queue, in its turn, is set as shared if this is the * first time that the requests of some process are redirected to * it. * * We redirect bfqq to new_bfqq and not the opposite, because * we are in the context of the process owning bfqq, thus we * have the io_cq of this process. So we can immediately * configure this io_cq to redirect the requests of the * process to new_bfqq. In contrast, the io_cq of new_bfqq is * not available any more (new_bfqq->bic == NULL). * * Anyway, even in case new_bfqq coincides with the in-service * queue, redirecting requests the in-service queue is the * best option, as we feed the in-service queue with new * requests close to the last request served and, by doing so, * are likely to increase the throughput. */ bfqq->new_bfqq = new_bfqq; /* * The above assignment schedules the following redirections: * each time some I/O for bfqq arrives, the process that * generated that I/O is disassociated from bfqq and * associated with new_bfqq. Here we increases new_bfqq->ref * in advance, adding the number of processes that are * expected to be associated with new_bfqq as they happen to * issue I/O. */ new_bfqq->ref += process_refs; return new_bfqq; } static bool bfq_may_be_close_cooperator(struct bfq_queue *bfqq, struct bfq_queue *new_bfqq) { if (bfq_too_late_for_merging(new_bfqq)) return false; if (bfq_class_idle(bfqq) || bfq_class_idle(new_bfqq) || (bfqq->ioprio_class != new_bfqq->ioprio_class)) return false; /* * If either of the queues has already been detected as seeky, * then merging it with the other queue is unlikely to lead to * sequential I/O. */ if (BFQQ_SEEKY(bfqq) || BFQQ_SEEKY(new_bfqq)) return false; /* * Interleaved I/O is known to be done by (some) applications * only for reads, so it does not make sense to merge async * queues. */ if (!bfq_bfqq_sync(bfqq) || !bfq_bfqq_sync(new_bfqq)) return false; return true; } static bool idling_boosts_thr_without_issues(struct bfq_data *bfqd, struct bfq_queue *bfqq); static struct bfq_queue * bfq_setup_stable_merge(struct bfq_data *bfqd, struct bfq_queue *bfqq, struct bfq_queue *stable_merge_bfqq, struct bfq_iocq_bfqq_data *bfqq_data) { int proc_ref = min(bfqq_process_refs(bfqq), bfqq_process_refs(stable_merge_bfqq)); struct bfq_queue *new_bfqq = NULL; bfqq_data->stable_merge_bfqq = NULL; if (idling_boosts_thr_without_issues(bfqd, bfqq) || proc_ref == 0) goto out; /* next function will take at least one ref */ new_bfqq = bfq_setup_merge(bfqq, stable_merge_bfqq); if (new_bfqq) { bfqq_data->stably_merged = true; if (new_bfqq->bic) { unsigned int new_a_idx = new_bfqq->actuator_idx; struct bfq_iocq_bfqq_data *new_bfqq_data = &new_bfqq->bic->bfqq_data[new_a_idx]; new_bfqq_data->stably_merged = true; } } out: /* deschedule stable merge, because done or aborted here */ bfq_put_stable_ref(stable_merge_bfqq); return new_bfqq; } /* * Attempt to schedule a merge of bfqq with the currently in-service * queue or with a close queue among the scheduled queues. Return * NULL if no merge was scheduled, a pointer to the shared bfq_queue * structure otherwise. * * The OOM queue is not allowed to participate to cooperation: in fact, since * the requests temporarily redirected to the OOM queue could be redirected * again to dedicated queues at any time, the state needed to correctly * handle merging with the OOM queue would be quite complex and expensive * to maintain. Besides, in such a critical condition as an out of memory, * the benefits of queue merging may be little relevant, or even negligible. * * WARNING: queue merging may impair fairness among non-weight raised * queues, for at least two reasons: 1) the original weight of a * merged queue may change during the merged state, 2) even being the * weight the same, a merged queue may be bloated with many more * requests than the ones produced by its originally-associated * process. */ static struct bfq_queue * bfq_setup_cooperator(struct bfq_data *bfqd, struct bfq_queue *bfqq, void *io_struct, bool request, struct bfq_io_cq *bic) { struct bfq_queue *in_service_bfqq, *new_bfqq; unsigned int a_idx = bfqq->actuator_idx; struct bfq_iocq_bfqq_data *bfqq_data = &bic->bfqq_data[a_idx]; /* if a merge has already been setup, then proceed with that first */ if (bfqq->new_bfqq) return bfqq->new_bfqq; /* * Check delayed stable merge for rotational or non-queueing * devs. For this branch to be executed, bfqq must not be * currently merged with some other queue (i.e., bfqq->bic * must be non null). If we considered also merged queues, * then we should also check whether bfqq has already been * merged with bic->stable_merge_bfqq. But this would be * costly and complicated. */ if (unlikely(!bfqd->nonrot_with_queueing)) { /* * Make sure also that bfqq is sync, because * bic->stable_merge_bfqq may point to some queue (for * stable merging) also if bic is associated with a * sync queue, but this bfqq is async */ if (bfq_bfqq_sync(bfqq) && bfqq_data->stable_merge_bfqq && !bfq_bfqq_just_created(bfqq) && time_is_before_jiffies(bfqq->split_time + msecs_to_jiffies(bfq_late_stable_merging)) && time_is_before_jiffies(bfqq->creation_time + msecs_to_jiffies(bfq_late_stable_merging))) { struct bfq_queue *stable_merge_bfqq = bfqq_data->stable_merge_bfqq; return bfq_setup_stable_merge(bfqd, bfqq, stable_merge_bfqq, bfqq_data); } } /* * 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; unsigned int a_idx = bfqq->actuator_idx; struct bfq_iocq_bfqq_data *bfqq_data = &bic->bfqq_data[a_idx]; /* * 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; bfqq_data->saved_last_serv_time_ns = bfqq->last_serv_time_ns; bfqq_data->saved_inject_limit = bfqq->inject_limit; bfqq_data->saved_decrease_time_jif = bfqq->decrease_time_jif; bfqq_data->saved_weight = bfqq->entity.orig_weight; bfqq_data->saved_ttime = bfqq->ttime; bfqq_data->saved_has_short_ttime = bfq_bfqq_has_short_ttime(bfqq); bfqq_data->saved_IO_bound = bfq_bfqq_IO_bound(bfqq); bfqq_data->saved_io_start_time = bfqq->io_start_time; bfqq_data->saved_tot_idle_time = bfqq->tot_idle_time; bfqq_data->saved_in_large_burst = bfq_bfqq_in_large_burst(bfqq); bfqq_data->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. */ bfqq_data->saved_wr_coeff = bfqq->bfqd->bfq_wr_coeff; bfqq_data->saved_wr_start_at_switch_to_srt = bfq_smallest_from_now(); bfqq_data->saved_wr_cur_max_time = bfq_wr_duration(bfqq->bfqd); bfqq_data->saved_last_wr_start_finish = jiffies; } else { bfqq_data->saved_wr_coeff = bfqq->wr_coeff; bfqq_data->saved_wr_start_at_switch_to_srt = bfqq->wr_start_at_switch_to_srt; bfqq_data->saved_service_from_wr = bfqq->service_from_wr; bfqq_data->saved_last_wr_start_finish = bfqq->last_wr_start_finish; bfqq_data->saved_wr_cur_max_time = bfqq->wr_cur_max_time; } } static void bfq_reassign_last_bfqq(struct bfq_queue *cur_bfqq, struct bfq_queue *new_bfqq) { if (cur_bfqq->entity.parent && cur_bfqq->entity.parent->last_bfqq_created == cur_bfqq) cur_bfqq->entity.parent->last_bfqq_created = new_bfqq; else if (cur_bfqq->bfqd && cur_bfqq->bfqd->last_bfqq_created == cur_bfqq) cur_bfqq->bfqd->last_bfqq_created = new_bfqq; } void bfq_release_process_ref(struct bfq_data *bfqd, struct bfq_queue *bfqq) { /* * To prevent bfqq's service guarantees from being violated, * bfqq may be left busy, i.e., queued for service, even if * empty (see comments in __bfq_bfqq_expire() for * details). But, if no process will send requests to bfqq any * longer, then there is no point in keeping bfqq queued for * service. In addition, keeping bfqq queued for service, but * with no process ref any longer, may have caused bfqq to be * freed when dequeued from service. But this is assumed to * never happen. */ if (bfq_bfqq_busy(bfqq) && RB_EMPTY_ROOT(&bfqq->sort_list) && bfqq != bfqd->in_service_queue) bfq_del_bfqq_busy(bfqq, false); bfq_reassign_last_bfqq(bfqq, NULL); bfq_put_queue(bfqq); } static void bfq_merge_bfqqs(struct bfq_data *bfqd, struct bfq_io_cq *bic, struct bfq_queue *bfqq, struct bfq_queue *new_bfqq) { bfq_log_bfqq(bfqd, bfqq, "merging with queue %lu", (unsigned long)new_bfqq->pid); /* Save weight raising and idle window of the merged queues */ bfq_bfqq_save_state(bfqq); bfq_bfqq_save_state(new_bfqq); if (bfq_bfqq_IO_bound(bfqq)) bfq_mark_bfqq_IO_bound(new_bfqq); bfq_clear_bfqq_IO_bound(bfqq); /* * The processes associated with bfqq are cooperators of the * processes associated with new_bfqq. So, if bfqq has a * waker, then assume that all these processes will be happy * to let bfqq's waker freely inject I/O when they have no * I/O. */ if (bfqq->waker_bfqq && !new_bfqq->waker_bfqq && bfqq->waker_bfqq != new_bfqq) { new_bfqq->waker_bfqq = bfqq->waker_bfqq; new_bfqq->tentative_waker_bfqq = NULL; /* * If the waker queue disappears, then * new_bfqq->waker_bfqq must be reset. So insert * new_bfqq into the woken_list of the waker. See * bfq_check_waker for details. */ hlist_add_head(&new_bfqq->woken_list_node, &new_bfqq->waker_bfqq->woken_list); } /* * If bfqq is weight-raised, then let new_bfqq inherit * weight-raising. To reduce false positives, neglect the case * where bfqq has just been created, but has not yet made it * to be weight-raised (which may happen because EQM may merge * bfqq even before bfq_add_request is executed for the first * time for bfqq). Handling this case would however be very * easy, thanks to the flag just_created. */ if (new_bfqq->wr_coeff == 1 && bfqq->wr_coeff > 1) { new_bfqq->wr_coeff = bfqq->wr_coeff; new_bfqq->wr_cur_max_time = bfqq->wr_cur_max_time; new_bfqq->last_wr_start_finish = bfqq->last_wr_start_finish; new_bfqq->wr_start_at_switch_to_srt = bfqq->wr_start_at_switch_to_srt; if (bfq_bfqq_busy(new_bfqq)) bfqd->wr_busy_queues++; new_bfqq->entity.prio_changed = 1; } if (bfqq->wr_coeff > 1) { /* bfqq has given its wr to new_bfqq */ bfqq->wr_coeff = 1; bfqq->entity.prio_changed = 1; if (bfq_bfqq_busy(bfqq)) bfqd->wr_busy_queues--; } bfq_log_bfqq(bfqd, new_bfqq, "merge_bfqqs: wr_busy %d", bfqd->wr_busy_queues); /* * Merge queues (that is, let bic redirect its requests to new_bfqq) */ bic_set_bfqq(bic, new_bfqq, true, bfqq->actuator_idx); bfq_mark_bfqq_coop(new_bfqq); /* * new_bfqq now belongs to at least two bics (it is a shared queue): * set new_bfqq->bic to NULL. bfqq either: * - does not belong to any bic any more, and hence bfqq->bic must * be set to NULL, or * - is a queue whose owning bics have already been redirected to a * different queue, hence the queue is destined to not belong to * any bic soon and bfqq->bic is already NULL (therefore the next * assignment causes no harm). */ new_bfqq->bic = NULL; /* * If the queue is shared, the pid is the pid of one of the associated * processes. Which pid depends on the exact sequence of merge events * the queue underwent. So printing such a pid is useless and confusing * because it reports a random pid between those of the associated * processes. * We mark such a queue with a pid -1, and then print SHARED instead of * a pid in logging messages. */ new_bfqq->pid = -1; bfqq->bic = NULL; bfq_reassign_last_bfqq(bfqq, new_bfqq); bfq_release_process_ref(bfqd, bfqq); } static bool bfq_allow_bio_merge(struct request_queue *q, struct request *rq, struct bio *bio) { struct bfq_data *bfqd = q->elevator->elevator_data; bool is_sync = op_is_sync(bio->bi_opf); struct bfq_queue *bfqq = bfqd->bio_bfqq, *new_bfqq; /* * Disallow merge of a sync bio into an async request. */ if (is_sync && !rq_is_sync(rq)) return false; /* * Lookup the bfqq that this bio will be queued with. Allow * merge only if rq is queued there. */ if (!bfqq) return false; /* * We take advantage of this function to perform an early merge * of the queues of possible cooperating processes. */ new_bfqq = bfq_setup_cooperator(bfqd, bfqq, bio, false, bfqd->bio_bic); if (new_bfqq) { /* * bic still points to bfqq, then it has not yet been * redirected to some other bfq_queue, and a queue * merge between bfqq and new_bfqq can be safely * fulfilled, i.e., bic can be redirected to new_bfqq * and bfqq can be put. */ bfq_merge_bfqqs(bfqd, bfqd->bio_bic, bfqq, new_bfqq); /* * If we get here, bio will be queued into new_queue, * so use new_bfqq to decide whether bio and rq can be * merged. */ bfqq = new_bfqq; /* * Change also bqfd->bio_bfqq, as * bfqd->bio_bic now points to new_bfqq, and * this function may be invoked again (and then may * use again bqfd->bio_bfqq). */ bfqd->bio_bfqq = bfqq; } return bfqq == RQ_BFQQ(rq); } /* * Set the maximum time for the in-service queue to consume its * budget. This prevents seeky processes from lowering the throughput. * In practice, a time-slice service scheme is used with seeky * processes. */ static void bfq_set_budget_timeout(struct bfq_data *bfqd, struct bfq_queue *bfqq) { unsigned int timeout_coeff; if (bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time) timeout_coeff = 1; else timeout_coeff = bfqq->entity.weight / bfqq->entity.orig_weight; bfqd->last_budget_start = blk_time_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 = blk_time_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 = blk_time_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 = blk_time_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->tot_rq_in_driver == 0) goto update_rate_and_reset; /* Update sampling information */ bfqd->peak_rate_samples++; if ((bfqd->tot_rq_in_driver > 0 || now_ns - bfqd->last_completion < BFQ_MIN_TT) && !BFQ_RQ_SEEKY(bfqd, bfqd->last_position, rq)) bfqd->sequential_samples++; bfqd->tot_sectors_dispatched += blk_rq_sectors(rq); /* Reset max observed rq size every 32 dispatches */ if (likely(bfqd->peak_rate_samples % 32)) bfqd->last_rq_max_size = max_t(u32, blk_rq_sectors(rq), bfqd->last_rq_max_size); else bfqd->last_rq_max_size = blk_rq_sectors(rq); bfqd->delta_from_first = now_ns - bfqd->first_dispatch; /* Target observation interval not yet reached, go on sampling */ if (bfqd->delta_from_first < BFQ_RATE_REF_INTERVAL) goto update_last_values; update_rate_and_reset: bfq_update_rate_reset(bfqd, rq); update_last_values: bfqd->last_position = blk_rq_pos(rq) + blk_rq_sectors(rq); if (RQ_BFQQ(rq) == bfqd->in_service_queue) bfqd->in_serv_last_pos = bfqd->last_position; bfqd->last_dispatch = now_ns; } /* * Remove request from internal lists. */ static void bfq_dispatch_remove(struct request_queue *q, struct request *rq) { struct bfq_queue *bfqq = RQ_BFQQ(rq); /* * For consistency, the next instruction should have been * executed after removing the request from the queue and * dispatching it. We execute instead this instruction before * bfq_remove_request() (and hence introduce a temporary * inconsistency), for efficiency. In fact, should this * dispatch occur for a non in-service bfqq, this anticipated * increment prevents two counters related to bfqq->dispatched * from risking to be, first, uselessly decremented, and then * incremented again when the (new) value of bfqq->dispatched * happens to be taken into account. */ bfqq->dispatched++; bfq_update_peak_rate(q->elevator->elevator_data, rq); bfq_remove_request(q, rq); } /* * There is a case where idling does not have to be performed for * throughput concerns, but to preserve the throughput share of * the process associated with bfqq. * * To introduce this case, we can note that allowing the drive * to enqueue more than one request at a time, and hence * delegating de facto final scheduling decisions to the * drive's internal scheduler, entails loss of control on the * actual request service order. In particular, the critical * situation is when requests from different processes happen * to be present, at the same time, in the internal queue(s) * of the drive. In such a situation, the drive, by deciding * the service order of the internally-queued requests, does * determine also the actual throughput distribution among * these processes. But the drive typically has no notion or * concern about per-process throughput distribution, and * makes its decisions only on a per-request basis. Therefore, * the service distribution enforced by the drive's internal * scheduler is likely to coincide with the desired throughput * distribution only in a completely symmetric, or favorably * skewed scenario where: * (i-a) each of these processes must get the same throughput as * the others, * (i-b) in case (i-a) does not hold, it holds that the process * associated with bfqq must receive a lower or equal * throughput than any of the other processes; * (ii) the I/O of each process has the same properties, in * terms of locality (sequential or random), direction * (reads or writes), request sizes, greediness * (from I/O-bound to sporadic), and so on; * In fact, in such a scenario, the drive tends to treat the requests * of each process in about the same way as the requests of the * others, and thus to provide each of these processes with about the * same throughput. This is exactly the desired throughput * distribution if (i-a) holds, or, if (i-b) holds instead, this is an * even more convenient distribution for (the process associated with) * bfqq. * * In contrast, in any asymmetric or unfavorable scenario, device * idling (I/O-dispatch plugging) is certainly needed to guarantee * that bfqq receives its assigned fraction of the device throughput * (see [1] for details). * * The problem is that idling may significantly reduce throughput with * certain combinations of types of I/O and devices. An important * example is sync random I/O on flash storage with command * queueing. So, unless bfqq falls in cases where idling also boosts * throughput, it is important to check conditions (i-a), i(-b) and * (ii) accurately, so as to avoid idling when not strictly needed for * service guarantees. * * Unfortunately, it is extremely difficult to thoroughly check * condition (ii). And, in case there are active groups, it becomes * very difficult to check conditions (i-a) and (i-b) too. In fact, * if there are active groups, then, for conditions (i-a) or (i-b) to * become false 'indirectly', it is enough that an active group * contains more active processes or sub-groups than some other active * group. More precisely, for conditions (i-a) or (i-b) to become * false because of such a group, it is not even necessary that the * group is (still) active: it is sufficient that, even if the group * has become inactive, some of its descendant processes still have * some request already dispatched but still waiting for * completion. In fact, requests have still to be guaranteed their * share of the throughput even after being dispatched. In this * respect, it is easy to show that, if a group frequently becomes * inactive while still having in-flight requests, and if, when this * happens, the group is not considered in the calculation of whether * the scenario is asymmetric, then the group may fail to be * guaranteed its fair share of the throughput (basically because * idling may not be performed for the descendant processes of the * group, but it had to be). We address this issue with the following * bi-modal behavior, implemented in the function * bfq_asymmetric_scenario(). * * If there are groups with requests waiting for completion * (as commented above, some of these groups may even be * already inactive), then the scenario is tagged as * asymmetric, conservatively, without checking any of the * conditions (i-a), (i-b) or (ii). So the device is idled for bfqq. * This behavior matches also the fact that groups are created * exactly if controlling I/O is a primary concern (to * preserve bandwidth and latency guarantees). * * On the opposite end, if there are no groups with requests waiting * for completion, then only conditions (i-a) and (i-b) are actually * controlled, i.e., provided that conditions (i-a) or (i-b) holds, * idling is not performed, regardless of whether condition (ii) * holds. In other words, only if conditions (i-a) and (i-b) do not * hold, then idling is allowed, and the device tends to be prevented * from queueing many requests, possibly of several processes. Since * there are no groups with requests waiting for completion, then, to * control conditions (i-a) and (i-b) it is enough to check just * whether all the queues with requests waiting for completion also * have the same weight. * * Not checking condition (ii) evidently exposes bfqq to the * risk of getting less throughput than its fair share. * However, for queues with the same weight, a further * mechanism, preemption, mitigates or even eliminates this * problem. And it does so without consequences on overall * throughput. This mechanism and its benefits are explained * in the next three paragraphs. * * Even if a queue, say Q, is expired when it remains idle, Q * can still preempt the new in-service queue if the next * request of Q arrives soon (see the comments on * bfq_bfqq_update_budg_for_activation). If all queues and * groups have the same weight, this form of preemption, * combined with the hole-recovery heuristic described in the * comments on function bfq_bfqq_update_budg_for_activation, * are enough to preserve a correct bandwidth distribution in * the mid term, even without idling. In fact, even if not * idling allows the internal queues of the device to contain * many requests, and thus to reorder requests, we can rather * safely assume that the internal scheduler still preserves a * minimum of mid-term fairness. * * More precisely, this preemption-based, idleless approach * provides fairness in terms of IOPS, and not sectors per * second. This can be seen with a simple example. Suppose * that there are two queues with the same weight, but that * the first queue receives requests of 8 sectors, while the * second queue receives requests of 1024 sectors. In * addition, suppose that each of the two queues contains at * most one request at a time, which implies that each queue * always remains idle after it is served. Finally, after * remaining idle, each queue receives very quickly a new * request. It follows that the two queues are served * alternatively, preempting each other if needed. This * implies that, although both queues have the same weight, * the queue with large requests receives a service that is * 1024/8 times as high as the service received by the other * queue. * * The motivation for using preemption instead of idling (for * queues with the same weight) is that, by not idling, * service guarantees are preserved (completely or at least in * part) without minimally sacrificing throughput. And, if * there is no active group, then the primary expectation for * this device is probably a high throughput. * * We are now left only with explaining the two sub-conditions in the * additional compound condition that is checked below for deciding * whether the scenario is asymmetric. To explain the first * sub-condition, we need to add that the function * bfq_asymmetric_scenario checks the weights of only * non-weight-raised queues, for efficiency reasons (see comments on * bfq_weights_tree_add()). Then the fact that bfqq is weight-raised * is checked explicitly here. More precisely, the compound condition * below takes into account also the fact that, even if bfqq is being * weight-raised, the scenario is still symmetric if all queues with * requests waiting for completion happen to be * weight-raised. Actually, we should be even more precise here, and * differentiate between interactive weight raising and soft real-time * weight raising. * * The second sub-condition checked in the compound condition is * whether there is a fair amount of already in-flight I/O not * belonging to bfqq. If so, I/O dispatching is to be plugged, for the * following reason. The drive may decide to serve in-flight * non-bfqq's I/O requests before bfqq's ones, thereby delaying the * arrival of new I/O requests for bfqq (recall that bfqq is sync). If * I/O-dispatching is not plugged, then, while bfqq remains empty, a * basically uncontrolled amount of I/O from other queues may be * dispatched too, possibly causing the service of bfqq's I/O to be * delayed even longer in the drive. This problem gets more and more * serious as the speed and the queue depth of the drive grow, * because, as these two quantities grow, the probability to find no * queue busy but many requests in flight grows too. By contrast, * plugging I/O dispatching minimizes the delay induced by already * in-flight I/O, and enables bfqq to recover the bandwidth it may * lose because of this delay. * * As a side note, it is worth considering that the above * device-idling countermeasures may however fail in the following * unlucky scenario: if I/O-dispatch plugging is (correctly) disabled * in a time period during which all symmetry sub-conditions hold, and * therefore the device is allowed to enqueue many requests, but at * some later point in time some sub-condition stops to hold, then it * may become impossible to make requests be served in the desired * order until all the requests already queued in the device have been * served. The last sub-condition commented above somewhat mitigates * this problem for weight-raised queues. * * However, as an additional mitigation for this problem, we preserve * plugging for a special symmetric case that may suddenly turn into * asymmetric: the case where only bfqq is busy. In this case, not * expiring bfqq does not cause any harm to any other queues in terms * of service guarantees. In contrast, it avoids the following unlucky * sequence of events: (1) bfqq is expired, (2) a new queue with a * lower weight than bfqq becomes busy (or more queues), (3) the new * queue is served until a new request arrives for bfqq, (4) when bfqq * is finally served, there are so many requests of the new queue in * the drive that the pending requests for bfqq take a lot of time to * be served. In particular, event (2) may case even already * dispatched requests of bfqq to be delayed, inside the drive. So, to * avoid this series of events, the scenario is preventively declared * as asymmetric also if bfqq is the only busy queues */ static bool idling_needed_for_service_guarantees(struct bfq_data *bfqd, struct bfq_queue *bfqq) { int tot_busy_queues = bfq_tot_busy_queues(bfqd); /* No point in idling for bfqq if it won't get requests any longer */ if (unlikely(!bfqq_process_refs(bfqq))) return false; return (bfqq->wr_coeff > 1 && (bfqd->wr_busy_queues < tot_busy_queues || bfqd->tot_rq_in_driver >= bfqq->dispatched + 4)) || bfq_asymmetric_scenario(bfqd, bfqq) || tot_busy_queues == 1; } static bool __bfq_bfqq_expire(struct bfq_data *bfqd, struct bfq_queue *bfqq, enum bfqq_expiration reason) { /* * If this bfqq is shared between multiple processes, check * to make sure that those processes are still issuing I/Os * within the mean seek distance. If not, it may be time to * break the queues apart again. */ if (bfq_bfqq_coop(bfqq) && BFQQ_SEEKY(bfqq)) bfq_mark_bfqq_split_coop(bfqq); /* * Consider queues with a higher finish virtual time than * bfqq. If idling_needed_for_service_guarantees(bfqq) returns * true, then bfqq's bandwidth would be violated if an * uncontrolled amount of I/O from these queues were * dispatched while bfqq is waiting for its new I/O to * arrive. This is exactly what may happen if this is a forced * expiration caused by a preemption attempt, and if bfqq is * not re-scheduled. To prevent this from happening, re-queue * bfqq if it needs I/O-dispatch plugging, even if it is * empty. By doing so, bfqq is granted to be served before the * above queues (provided that bfqq is of course eligible). */ if (RB_EMPTY_ROOT(&bfqq->sort_list) && !(reason == BFQQE_PREEMPTED && idling_needed_for_service_guarantees(bfqd, bfqq))) { if (bfqq->dispatched == 0) /* * Overloading budget_timeout field to store * the time at which the queue remains with no * backlog and no outstanding request; used by * the weight-raising mechanism. */ bfqq->budget_timeout = jiffies; bfq_del_bfqq_busy(bfqq, true); } else { bfq_requeue_bfqq(bfqd, bfqq, true); /* * Resort priority tree of potential close cooperators. * See comments on bfq_pos_tree_add_move() for the unlikely(). */ if (unlikely(!bfqd->nonrot_with_queueing && !RB_EMPTY_ROOT(&bfqq->sort_list))) bfq_pos_tree_add_move(bfqd, bfqq); } /* * All in-service entities must have been properly deactivated * or requeued before executing the next function, which * resets all in-service entities as no more in service. This * may cause bfqq to be freed. If this happens, the next * function returns true. */ return __bfq_bfqd_reset_in_service(bfqd); } /** * __bfq_bfqq_recalc_budget - try to adapt the budget to the @bfqq behavior. * @bfqd: device data. * @bfqq: queue to update. * @reason: reason for expiration. * * Handle the feedback on @bfqq budget at queue expiration. * See the body for detailed comments. */ static void __bfq_bfqq_recalc_budget(struct bfq_data *bfqd, struct bfq_queue *bfqq, enum bfqq_expiration reason) { struct request *next_rq; int budget, min_budget; min_budget = bfq_min_budget(bfqd); if (bfqq->wr_coeff == 1) budget = bfqq->max_budget; else /* * Use a constant, low budget for weight-raised queues, * to help achieve a low latency. Keep it slightly higher * than the minimum possible budget, to cause a little * bit fewer expirations. */ budget = 2 * min_budget; bfq_log_bfqq(bfqd, bfqq, "recalc_budg: last budg %d, budg left %d", bfqq->entity.budget, bfq_bfqq_budget_left(bfqq)); bfq_log_bfqq(bfqd, bfqq, "recalc_budg: last max_budg %d, min budg %d", budget, bfq_min_budget(bfqd)); bfq_log_bfqq(bfqd, bfqq, "recalc_budg: sync %d, seeky %d", bfq_bfqq_sync(bfqq), BFQQ_SEEKY(bfqd->in_service_queue)); if (bfq_bfqq_sync(bfqq) && bfqq->wr_coeff == 1) { switch (reason) { /* * Caveat: in all the following cases we trade latency * for throughput. */ case BFQQE_TOO_IDLE: /* * This is the only case where we may reduce * the budget: if there is no request of the * process still waiting for completion, then * we assume (tentatively) that the timer has * expired because the batch of requests of * the process could have been served with a * smaller budget. Hence, betting that * process will behave in the same way when it * becomes backlogged again, we reduce its * next budget. As long as we guess right, * this budget cut reduces the latency * experienced by the process. * * However, if there are still outstanding * requests, then the process may have not yet * issued its next request just because it is * still waiting for the completion of some of * the still outstanding ones. So in this * subcase we do not reduce its budget, on the * contrary we increase it to possibly boost * the throughput, as discussed in the * comments to the BUDGET_TIMEOUT case. */ if (bfqq->dispatched > 0) /* still outstanding reqs */ budget = min(budget * 2, bfqd->bfq_max_budget); else { if (budget > 5 * min_budget) budget -= 4 * min_budget; else budget = min_budget; } break; case BFQQE_BUDGET_TIMEOUT: /* * We double the budget here because it gives * the chance to boost the throughput if this * is not a seeky process (and has bumped into * this timeout because of, e.g., ZBR). */ budget = min(budget * 2, bfqd->bfq_max_budget); break; case BFQQE_BUDGET_EXHAUSTED: /* * The process still has backlog, and did not * let either the budget timeout or the disk * idling timeout expire. Hence it is not * seeky, has a short thinktime and may be * happy with a higher budget too. So * definitely increase the budget of this good * candidate to boost the disk throughput. */ budget = min(budget * 4, bfqd->bfq_max_budget); break; case BFQQE_NO_MORE_REQUESTS: /* * For queues that expire for this reason, it * is particularly important to keep the * budget close to the actual service they * need. Doing so reduces the timestamp * misalignment problem described in the * comments in the body of * __bfq_activate_entity. In fact, suppose * that a queue systematically expires for * BFQQE_NO_MORE_REQUESTS and presents a * new request in time to enjoy timestamp * back-shifting. The larger the budget of the * queue is with respect to the service the * queue actually requests in each service * slot, the more times the queue can be * reactivated with the same virtual finish * time. It follows that, even if this finish * time is pushed to the system virtual time * to reduce the consequent timestamp * misalignment, the queue unjustly enjoys for * many re-activations a lower finish time * than all newly activated queues. * * The service needed by bfqq is measured * quite precisely by bfqq->entity.service. * Since bfqq does not enjoy device idling, * bfqq->entity.service is equal to the number * of sectors that the process associated with * bfqq requested to read/write before waiting * for request completions, or blocking for * other reasons. */ budget = max_t(int, bfqq->entity.service, min_budget); break; default: return; } } else if (!bfq_bfqq_sync(bfqq)) { /* * Async queues get always the maximum possible * budget, as for them we do not care about latency * (in addition, their ability to dispatch is limited * by the charging factor). */ budget = bfqd->bfq_max_budget; } bfqq->max_budget = budget; if (bfqd->budgets_assigned >= bfq_stats_min_budgets && !bfqd->bfq_user_max_budget) bfqq->max_budget = min(bfqq->max_budget, bfqd->bfq_max_budget); /* * If there is still backlog, then assign a new budget, making * sure that it is large enough for the next request. Since * the finish time of bfqq must be kept in sync with the * budget, be sure to call __bfq_bfqq_expire() *after* this * update. * * If there is no backlog, then no need to update the budget; * it will be updated on the arrival of a new request. */ next_rq = bfqq->next_rq; if (next_rq) bfqq->entity.budget = max_t(unsigned long, bfqq->max_budget, bfq_serv_to_charge(next_rq, bfqq)); bfq_log_bfqq(bfqd, bfqq, "head sect: %u, new budget %d", next_rq ? blk_rq_sectors(next_rq) : 0, bfqq->entity.budget); } /* * Return true if the process associated with bfqq is "slow". The slow * flag is used, in addition to the budget timeout, to reduce the * amount of service provided to seeky processes, and thus reduce * their chances to lower the throughput. More details in the comments * on the function bfq_bfqq_expire(). * * An important observation is in order: as discussed in the comments * on the function bfq_update_peak_rate(), with devices with internal * queues, it is hard if ever possible to know when and for how long * an I/O request is processed by the device (apart from the trivial * I/O pattern where a new request is dispatched only after the * previous one has been completed). This makes it hard to evaluate * the real rate at which the I/O requests of each bfq_queue are * served. In fact, for an I/O scheduler like BFQ, serving a * bfq_queue means just dispatching its requests during its service * slot (i.e., until the budget of the queue is exhausted, or the * queue remains idle, or, finally, a timeout fires). But, during the * service slot of a bfq_queue, around 100 ms at most, the device may * be even still processing requests of bfq_queues served in previous * service slots. On the opposite end, the requests of the in-service * bfq_queue may be completed after the service slot of the queue * finishes. * * Anyway, unless more sophisticated solutions are used * (where possible), the sum of the sizes of the requests dispatched * during the service slot of a bfq_queue is probably the only * approximation available for the service received by the bfq_queue * during its service slot. And this sum is the quantity used in this * function to evaluate the I/O speed of a process. */ static bool bfq_bfqq_is_slow(struct bfq_data *bfqd, struct bfq_queue *bfqq, bool compensate, unsigned long *delta_ms) { ktime_t delta_ktime; u32 delta_usecs; bool slow = BFQQ_SEEKY(bfqq); /* if delta too short, use seekyness */ if (!bfq_bfqq_sync(bfqq)) return false; if (compensate) delta_ktime = bfqd->last_idling_start; else delta_ktime = blk_time_get(); delta_ktime = ktime_sub(delta_ktime, bfqd->last_budget_start); delta_usecs = ktime_to_us(delta_ktime); /* don't use too short time intervals */ if (delta_usecs < 1000) { if (blk_queue_nonrot(bfqd->queue)) /* * give same worst-case guarantees as idling * for seeky */ *delta_ms = BFQ_MIN_TT / NSEC_PER_MSEC; else /* charge at least one seek */ *delta_ms = bfq_slice_idle / NSEC_PER_MSEC; return slow; } *delta_ms = delta_usecs / USEC_PER_MSEC; /* * Use only long (> 20ms) intervals to filter out excessive * spikes in service rate estimation. */ if (delta_usecs > 20000) { /* * Caveat for rotational devices: processes doing I/O * in the slower disk zones tend to be slow(er) even * if not seeky. In this respect, the estimated peak * rate is likely to be an average over the disk * surface. Accordingly, to not be too harsh with * unlucky processes, a process is deemed slow only if * its rate has been lower than half of the estimated * peak rate. */ slow = bfqq->entity.service < bfqd->bfq_max_budget / 2; } bfq_log_bfqq(bfqd, bfqq, "bfq_bfqq_is_slow: slow %d", slow); return slow; } /* * To be deemed as soft real-time, an application must meet two * requirements. First, the application must not require an average * bandwidth higher than the approximate bandwidth required to playback or * record a compressed high-definition video. * The next function is invoked on the completion of the last request of a * batch, to compute the next-start time instant, soft_rt_next_start, such * that, if the next request of the application does not arrive before * soft_rt_next_start, then the above requirement on the bandwidth is met. * * The second requirement is that the request pattern of the application is * isochronous, i.e., that, after issuing a request or a batch of requests, * the application stops issuing new requests until all its pending requests * have been completed. After that, the application may issue a new batch, * and so on. * For this reason the next function is invoked to compute * soft_rt_next_start only for applications that meet this requirement, * whereas soft_rt_next_start is set to infinity for applications that do * not. * * Unfortunately, even a greedy (i.e., I/O-bound) application may * happen to meet, occasionally or systematically, both the above * bandwidth and isochrony requirements. This may happen at least in * the following circumstances. First, if the CPU load is high. The * application may stop issuing requests while the CPUs are busy * serving other processes, then restart, then stop again for a while, * and so on. The other circumstances are related to the storage * device: the storage device is highly loaded or reaches a low-enough * throughput with the I/O of the application (e.g., because the I/O * is random and/or the device is slow). In all these cases, the * I/O of the application may be simply slowed down enough to meet * the bandwidth and isochrony requirements. To reduce the probability * that greedy applications are deemed as soft real-time in these * corner cases, a further rule is used in the computation of * soft_rt_next_start: the return value of this function is forced to * be higher than the maximum between the following two quantities. * * (a) Current time plus: (1) the maximum time for which the arrival * of a request is waited for when a sync queue becomes idle, * namely bfqd->bfq_slice_idle, and (2) a few extra jiffies. We * postpone for a moment the reason for adding a few extra * jiffies; we get back to it after next item (b). Lower-bounding * the return value of this function with the current time plus * bfqd->bfq_slice_idle tends to filter out greedy applications, * because the latter issue their next request as soon as possible * after the last one has been completed. In contrast, a soft * real-time application spends some time processing data, after a * batch of its requests has been completed. * * (b) Current value of bfqq->soft_rt_next_start. As pointed out * above, greedy applications may happen to meet both the * bandwidth and isochrony requirements under heavy CPU or * storage-device load. In more detail, in these scenarios, these * applications happen, only for limited time periods, to do I/O * slowly enough to meet all the requirements described so far, * including the filtering in above item (a). These slow-speed * time intervals are usually interspersed between other time * intervals during which these applications do I/O at a very high * speed. Fortunately, exactly because of the high speed of the * I/O in the high-speed intervals, the values returned by this * function happen to be so high, near the end of any such * high-speed interval, to be likely to fall *after* the end of * the low-speed time interval that follows. These high values are * stored in bfqq->soft_rt_next_start after each invocation of * this function. As a consequence, if the last value of * bfqq->soft_rt_next_start is constantly used to lower-bound the * next value that this function may return, then, from the very * beginning of a low-speed interval, bfqq->soft_rt_next_start is * likely to be constantly kept so high that any I/O request * issued during the low-speed interval is considered as arriving * to soon for the application to be deemed as soft * real-time. Then, in the high-speed interval that follows, the * application will not be deemed as soft real-time, just because * it will do I/O at a high speed. And so on. * * Getting back to the filtering in item (a), in the following two * cases this filtering might be easily passed by a greedy * application, if the reference quantity was just * bfqd->bfq_slice_idle: * 1) HZ is so low that the duration of a jiffy is comparable to or * higher than bfqd->bfq_slice_idle. This happens, e.g., on slow * devices with HZ=100. The time granularity may be so coarse * that the approximation, in jiffies, of bfqd->bfq_slice_idle * is rather lower than the exact value. * 2) jiffies, instead of increasing at a constant rate, may stop increasing * for a while, then suddenly 'jump' by several units to recover the lost * increments. This seems to happen, e.g., inside virtual machines. * To address this issue, in the filtering in (a) we do not use as a * reference time interval just bfqd->bfq_slice_idle, but * bfqd->bfq_slice_idle plus a few jiffies. In particular, we add the * minimum number of jiffies for which the filter seems to be quite * precise also in embedded systems and KVM/QEMU virtual machines. */ static unsigned long bfq_bfqq_softrt_next_start(struct bfq_data *bfqd, struct bfq_queue *bfqq) { return max3(bfqq->soft_rt_next_start, bfqq->last_idle_bklogged + HZ * bfqq->service_from_backlogged / bfqd->bfq_wr_max_softrt_rate, jiffies + nsecs_to_jiffies(bfqq->bfqd->bfq_slice_idle) + 4); } /** * bfq_bfqq_expire - expire a queue. * @bfqd: device owning the queue. * @bfqq: the queue to expire. * @compensate: if true, compensate for the time spent idling. * @reason: the reason causing the expiration. * * If the process associated with bfqq does slow I/O (e.g., because it * issues random requests), we charge bfqq with the time it has been * in service instead of the service it has received (see * bfq_bfqq_charge_time for details on how this goal is achieved). As * a consequence, bfqq will typically get higher timestamps upon * reactivation, and hence it will be rescheduled as if it had * received more service than what it has actually received. In the * end, bfqq receives less service in proportion to how slowly its * associated process consumes its budgets (and hence how seriously it * tends to lower the throughput). In addition, this time-charging * strategy guarantees time fairness among slow processes. In * contrast, if the process associated with bfqq is not slow, we * charge bfqq exactly with the service it has received. * * Charging time to the first type of queues and the exact service to * the other has the effect of using the WF2Q+ policy to schedule the * former on a timeslice basis, without violating service domain * guarantees among the latter. */ void bfq_bfqq_expire(struct bfq_data *bfqd, struct bfq_queue *bfqq, bool compensate, enum bfqq_expiration reason) { bool slow; unsigned long delta = 0; struct bfq_entity *entity = &bfqq->entity; /* * Check whether the process is slow (see bfq_bfqq_is_slow). */ slow = bfq_bfqq_is_slow(bfqd, bfqq, compensate, &delta); /* * As above explained, charge slow (typically seeky) and * timed-out queues with the time and not the service * received, to favor sequential workloads. * * Processes doing I/O in the slower disk zones will tend to * be slow(er) even if not seeky. Therefore, since the * estimated peak rate is actually an average over the disk * surface, these processes may timeout just for bad luck. To * avoid punishing them, do not charge time to processes that * succeeded in consuming at least 2/3 of their budget. This * allows BFQ to preserve enough elasticity to still perform * bandwidth, and not time, distribution with little unlucky * or quasi-sequential processes. */ if (bfqq->wr_coeff == 1 && (slow || (reason == BFQQE_BUDGET_TIMEOUT && bfq_bfqq_budget_left(bfqq) >= entity->budget / 3))) bfq_bfqq_charge_time(bfqd, bfqq, delta); if (bfqd->low_latency && bfqq->wr_coeff == 1) bfqq->last_wr_start_finish = jiffies; if (bfqd->low_latency && bfqd->bfq_wr_max_softrt_rate > 0 && RB_EMPTY_ROOT(&bfqq->sort_list)) { /* * If we get here, and there are no outstanding * requests, then the request pattern is isochronous * (see the comments on the function * bfq_bfqq_softrt_next_start()). Therefore we can * compute soft_rt_next_start. * * If, instead, the queue still has outstanding * requests, then we have to wait for the completion * of all the outstanding requests to discover whether * the request pattern is actually isochronous. */ if (bfqq->dispatched == 0) bfqq->soft_rt_next_start = bfq_bfqq_softrt_next_start(bfqd, bfqq); else if (bfqq->dispatched > 0) { /* * Schedule an update of soft_rt_next_start to when * the task may be discovered to be isochronous. */ bfq_mark_bfqq_softrt_update(bfqq); } } bfq_log_bfqq(bfqd, bfqq, "expire (%d, slow %d, num_disp %d, short_ttime %d)", reason, slow, bfqq->dispatched, bfq_bfqq_has_short_ttime(bfqq)); /* * bfqq expired, so no total service time needs to be computed * any longer: reset state machine for measuring total service * times. */ bfqd->rqs_injected = bfqd->wait_dispatch = false; bfqd->waited_rq = NULL; /* * Increase, decrease or leave budget unchanged according to * reason. */ __bfq_bfqq_recalc_budget(bfqd, bfqq, reason); if (__bfq_bfqq_expire(bfqd, bfqq, reason)) /* bfqq is gone, no more actions on it */ return; /* mark bfqq as waiting a request only if a bic still points to it */ if (!bfq_bfqq_busy(bfqq) && reason != BFQQE_BUDGET_TIMEOUT && reason != BFQQE_BUDGET_EXHAUSTED) { bfq_mark_bfqq_non_blocking_wait_rq(bfqq); /* * Not setting service to 0, because, if the next rq * arrives in time, the queue will go on receiving * service with this same budget (as if it never expired) */ } else entity->service = 0; /* * Reset the received-service counter for every parent entity. * Differently from what happens with bfqq->entity.service, * the resetting of this counter never needs to be postponed * for parent entities. In fact, in case bfqq may have a * chance to go on being served using the last, partially * consumed budget, bfqq->entity.service needs to be kept, * because if bfqq then actually goes on being served using * the same budget, the last value of bfqq->entity.service is * needed to properly decrement bfqq->entity.budget by the * portion already consumed. In contrast, it is not necessary * to keep entity->service for parent entities too, because * the bubble up of the new value of bfqq->entity.budget will * make sure that the budgets of parent entities are correct, * even in case bfqq and thus parent entities go on receiving * service with the same budget. */ entity = entity->parent; for_each_entity(entity) entity->service = 0; } /* * Budget timeout is not implemented through a dedicated timer, but * just checked on request arrivals and completions, as well as on * idle timer expirations. */ static bool bfq_bfqq_budget_timeout(struct bfq_queue *bfqq) { return time_is_before_eq_jiffies(bfqq->budget_timeout); } /* * If we expire a queue that is actively waiting (i.e., with the * device idled) for the arrival of a new request, then we may incur * the timestamp misalignment problem described in the body of the * function __bfq_activate_entity. Hence we return true only if this * condition does not hold, or if the queue is slow enough to deserve * only to be kicked off for preserving a high throughput. */ static bool bfq_may_expire_for_budg_timeout(struct bfq_queue *bfqq) { bfq_log_bfqq(bfqq->bfqd, bfqq, "may_budget_timeout: wait_request %d left %d timeout %d", bfq_bfqq_wait_request(bfqq), bfq_bfqq_budget_left(bfqq) >= bfqq->entity.budget / 3, bfq_bfqq_budget_timeout(bfqq)); return (!bfq_bfqq_wait_request(bfqq) || bfq_bfqq_budget_left(bfqq) >= bfqq->entity.budget / 3) && bfq_bfqq_budget_timeout(bfqq); } static bool idling_boosts_thr_without_issues(struct bfq_data *bfqd, struct bfq_queue *bfqq) { bool rot_without_queueing = !blk_queue_nonrot(bfqd->queue) && !bfqd->hw_tag, bfqq_sequential_and_IO_bound, idling_boosts_thr; /* No point in idling for bfqq if it won't get requests any longer */ if (unlikely(!bfqq_process_refs(bfqq))) return false; bfqq_sequential_and_IO_bound = !BFQQ_SEEKY(bfqq) && bfq_bfqq_IO_bound(bfqq) && bfq_bfqq_has_short_ttime(bfqq); /* * The next variable takes into account the cases where idling * boosts the throughput. * * The value of the variable is computed considering, first, that * idling is virtually always beneficial for the throughput if: * (a) the device is not NCQ-capable and rotational, or * (b) regardless of the presence of NCQ, the device is rotational and * the request pattern for bfqq is I/O-bound and sequential, or * (c) regardless of whether it is rotational, the device is * not NCQ-capable and the request pattern for bfqq is * I/O-bound and sequential. * * Secondly, and in contrast to the above item (b), idling an * NCQ-capable flash-based device would not boost the * throughput even with sequential I/O; rather it would lower * the throughput in proportion to how fast the device * is. Accordingly, the next variable is true if any of the * above conditions (a), (b) or (c) is true, and, in * particular, happens to be false if bfqd is an NCQ-capable * flash-based device. */ idling_boosts_thr = rot_without_queueing || ((!blk_queue_nonrot(bfqd->queue) || !bfqd->hw_tag) && bfqq_sequential_and_IO_bound); /* * The return value of this function is equal to that of * idling_boosts_thr, unless a special case holds. In this * special case, described below, idling may cause problems to * weight-raised queues. * * When the request pool is saturated (e.g., in the presence * of write hogs), if the processes associated with * non-weight-raised queues ask for requests at a lower rate, * then processes associated with weight-raised queues have a * higher probability to get a request from the pool * immediately (or at least soon) when they need one. Thus * they have a higher probability to actually get a fraction * of the device throughput proportional to their high * weight. This is especially true with NCQ-capable drives, * which enqueue several requests in advance, and further * reorder internally-queued requests. * * For this reason, we force to false the return value if * there are weight-raised busy queues. In this case, and if * bfqq is not weight-raised, this guarantees that the device * is not idled for bfqq (if, instead, bfqq is weight-raised, * then idling will be guaranteed by another variable, see * below). Combined with the timestamping rules of BFQ (see * [1] for details), this behavior causes bfqq, and hence any * sync non-weight-raised queue, to get a lower number of * requests served, and thus to ask for a lower number of * requests from the request pool, before the busy * weight-raised queues get served again. This often mitigates * starvation problems in the presence of heavy write * workloads and NCQ, thereby guaranteeing a higher * application and system responsiveness in these hostile * scenarios. */ return idling_boosts_thr && bfqd->wr_busy_queues == 0; } /* * For a queue that becomes empty, device idling is allowed only if * this function returns true for that queue. As a consequence, since * device idling plays a critical role for both throughput boosting * and service guarantees, the return value of this function plays a * critical role as well. * * In a nutshell, this function returns true only if idling is * beneficial for throughput or, even if detrimental for throughput, * idling is however necessary to preserve service guarantees (low * latency, desired throughput distribution, ...). In particular, on * NCQ-capable devices, this function tries to return false, so as to * help keep the drives' internal queues full, whenever this helps the * device boost the throughput without causing any service-guarantee * issue. * * Most of the issues taken into account to get the return value of * this function are not trivial. We discuss these issues in the two * functions providing the main pieces of information needed by this * function. */ static bool bfq_better_to_idle(struct bfq_queue *bfqq) { struct bfq_data *bfqd = bfqq->bfqd; bool idling_boosts_thr_with_no_issue, idling_needed_for_service_guar; /* No point in idling for bfqq if it won't get requests any longer */ if (unlikely(!bfqq_process_refs(bfqq))) return false; if (unlikely(bfqd->strict_guarantees)) return true; /* * Idling is performed only if slice_idle > 0. In addition, we * do not idle if * (a) bfqq is async * (b) bfqq is in the idle io prio class: in this case we do * not idle because we want to minimize the bandwidth that * queues in this class can steal to higher-priority queues */ if (bfqd->bfq_slice_idle == 0 || !bfq_bfqq_sync(bfqq) || bfq_class_idle(bfqq)) return false; idling_boosts_thr_with_no_issue = idling_boosts_thr_without_issues(bfqd, bfqq); idling_needed_for_service_guar = idling_needed_for_service_guarantees(bfqd, bfqq); /* * We have now the two components we need to compute the * return value of the function, which is true only if idling * either boosts the throughput (without issues), or is * necessary to preserve service guarantees. */ return idling_boosts_thr_with_no_issue || idling_needed_for_service_guar; } /* * If the in-service queue is empty but the function bfq_better_to_idle * returns true, then: * 1) the queue must remain in service and cannot be expired, and * 2) the device must be idled to wait for the possible arrival of a new * request for the queue. * See the comments on the function bfq_better_to_idle for the reasons * why performing device idling is the best choice to boost the throughput * and preserve service guarantees when bfq_better_to_idle itself * returns true. */ static bool bfq_bfqq_must_idle(struct bfq_queue *bfqq) { return RB_EMPTY_ROOT(&bfqq->sort_list) && bfq_better_to_idle(bfqq); } /* * This function chooses the queue from which to pick the next extra * I/O request to inject, if it finds a compatible queue. See the * comments on bfq_update_inject_limit() for details on the injection * mechanism, and for the definitions of the quantities mentioned * below. */ static struct bfq_queue * bfq_choose_bfqq_for_injection(struct bfq_data *bfqd) { struct bfq_queue *bfqq, *in_serv_bfqq = bfqd->in_service_queue; unsigned int limit = in_serv_bfqq->inject_limit; int i; /* * If * - bfqq is not weight-raised and therefore does not carry * time-critical I/O, * or * - regardless of whether bfqq is weight-raised, bfqq has * however a long think time, during which it can absorb the * effect of an appropriate number of extra I/O requests * from other queues (see bfq_update_inject_limit for * details on the computation of this number); * then injection can be performed without restrictions. */ bool in_serv_always_inject = in_serv_bfqq->wr_coeff == 1 || !bfq_bfqq_has_short_ttime(in_serv_bfqq); /* * If * - the baseline total service time could not be sampled yet, * so the inject limit happens to be still 0, and * - a lot of time has elapsed since the plugging of I/O * dispatching started, so drive speed is being wasted * significantly; * then temporarily raise inject limit to one request. */ if (limit == 0 && in_serv_bfqq->last_serv_time_ns == 0 && bfq_bfqq_wait_request(in_serv_bfqq) && time_is_before_eq_jiffies(bfqd->last_idling_start_jiffies + bfqd->bfq_slice_idle) ) limit = 1; if (bfqd->tot_rq_in_driver >= limit) return NULL; /* * Linear search of the source queue for injection; but, with * a high probability, very few steps are needed to find a * candidate queue, i.e., a queue with enough budget left for * its next request. In fact: * - BFQ dynamically updates the budget of every queue so as * to accommodate the expected backlog of the queue; * - if a queue gets all its requests dispatched as injected * service, then the queue is removed from the active list * (and re-added only if it gets new requests, but then it * is assigned again enough budget for its new backlog). */ for (i = 0; i < bfqd->num_actuators; i++) { list_for_each_entry(bfqq, &bfqd->active_list[i], bfqq_list) if (!RB_EMPTY_ROOT(&bfqq->sort_list) && (in_serv_always_inject || bfqq->wr_coeff > 1) && bfq_serv_to_charge(bfqq->next_rq, bfqq) <= bfq_bfqq_budget_left(bfqq)) { /* * Allow for only one large in-flight request * on non-rotational devices, for the * following reason. On non-rotationl drives, * large requests take much longer than * smaller requests to be served. In addition, * the drive prefers to serve large requests * w.r.t. to small ones, if it can choose. So, * having more than one large requests queued * in the drive may easily make the next first * request of the in-service queue wait for so * long to break bfqq's service guarantees. On * the bright side, large requests let the * drive reach a very high throughput, even if * there is only one in-flight large request * at a time. */ if (blk_queue_nonrot(bfqd->queue) && blk_rq_sectors(bfqq->next_rq) >= BFQQ_SECT_THR_NONROT && bfqd->tot_rq_in_driver >= 1) continue; else { bfqd->rqs_injected = true; return bfqq; } } } return NULL; } static struct bfq_queue * bfq_find_active_bfqq_for_actuator(struct bfq_data *bfqd, int idx) { struct bfq_queue *bfqq; if (bfqd->in_service_queue && bfqd->in_service_queue->actuator_idx == idx) return bfqd->in_service_queue; list_for_each_entry(bfqq, &bfqd->active_list[idx], bfqq_list) { if (!RB_EMPTY_ROOT(&bfqq->sort_list) && bfq_serv_to_charge(bfqq->next_rq, bfqq) <= bfq_bfqq_budget_left(bfqq)) { return bfqq; } } return NULL; } /* * Perform a linear scan of each actuator, until an actuator is found * for which the following three conditions hold: the load of the * actuator is below the threshold (see comments on * actuator_load_threshold for details) and lower than that of the * next actuator (comments on this extra condition below), and there * is a queue that contains I/O for that actuator. On success, return * that queue. * * Performing a plain linear scan entails a prioritization among * actuators. The extra condition above breaks this prioritization and * tends to distribute injection uniformly across actuators. */ static struct bfq_queue * bfq_find_bfqq_for_underused_actuator(struct bfq_data *bfqd) { int i; for (i = 0 ; i < bfqd->num_actuators; i++) { if (bfqd->rq_in_driver[i] < bfqd->actuator_load_threshold && (i == bfqd->num_actuators - 1 || bfqd->rq_in_driver[i] < bfqd->rq_in_driver[i+1])) { struct bfq_queue *bfqq = bfq_find_active_bfqq_for_actuator(bfqd, i); if (bfqq) return bfqq; } } return NULL; } /* * Select a queue for service. If we have a current queue in service, * check whether to continue servicing it, or retrieve and set a new one. */ static struct bfq_queue *bfq_select_queue(struct bfq_data *bfqd) { struct bfq_queue *bfqq, *inject_bfqq; struct request *next_rq; enum bfqq_expiration reason = BFQQE_BUDGET_TIMEOUT; bfqq = bfqd->in_service_queue; if (!bfqq) goto new_queue; bfq_log_bfqq(bfqd, bfqq, "select_queue: already in-service queue"); /* * Do not expire bfqq for budget timeout if bfqq may be about * to enjoy device idling. The reason why, in this case, we * prevent bfqq from expiring is the same as in the comments * on the case where bfq_bfqq_must_idle() returns true, in * bfq_completed_request(). */ if (bfq_may_expire_for_budg_timeout(bfqq) && !bfq_bfqq_must_idle(bfqq)) goto expire; check_queue: /* * If some actuator is underutilized, but the in-service * queue does not contain I/O for that actuator, then try to * inject I/O for that actuator. */ inject_bfqq = bfq_find_bfqq_for_underused_actuator(bfqd); if (inject_bfqq && inject_bfqq != bfqq) return inject_bfqq; /* * This loop is rarely executed more than once. Even when it * happens, it is much more convenient to re-execute this loop * than to return NULL and trigger a new dispatch to get a * request served. */ next_rq = bfqq->next_rq; /* * If bfqq has requests queued and it has enough budget left to * serve them, keep the queue, otherwise expire it. */ if (next_rq) { if (bfq_serv_to_charge(next_rq, bfqq) > bfq_bfqq_budget_left(bfqq)) { /* * Expire the queue for budget exhaustion, * which makes sure that the next budget is * enough to serve the next request, even if * it comes from the fifo expired path. */ reason = BFQQE_BUDGET_EXHAUSTED; goto expire; } else { /* * The idle timer may be pending because we may * not disable disk idling even when a new request * arrives. */ if (bfq_bfqq_wait_request(bfqq)) { /* * If we get here: 1) at least a new request * has arrived but we have not disabled the * timer because the request was too small, * 2) then the block layer has unplugged * the device, causing the dispatch to be * invoked. * * Since the device is unplugged, now the * requests are probably large enough to * provide a reasonable throughput. * So we disable idling. */ bfq_clear_bfqq_wait_request(bfqq); hrtimer_try_to_cancel(&bfqd->idle_slice_timer); } goto keep_queue; } } /* * No requests pending. However, if the in-service queue is idling * for a new request, or has requests waiting for a completion and * may idle after their completion, then keep it anyway. * * Yet, inject service from other queues if it boosts * throughput and is possible. */ if (bfq_bfqq_wait_request(bfqq) || (bfqq->dispatched != 0 && bfq_better_to_idle(bfqq))) { unsigned int act_idx = bfqq->actuator_idx; struct bfq_queue *async_bfqq = NULL; struct bfq_queue *blocked_bfqq = !hlist_empty(&bfqq->woken_list) ? container_of(bfqq->woken_list.first, struct bfq_queue, woken_list_node) : NULL; if (bfqq->bic && bfqq->bic->bfqq[0][act_idx] && bfq_bfqq_busy(bfqq->bic->bfqq[0][act_idx]) && bfqq->bic->bfqq[0][act_idx]->next_rq) async_bfqq = bfqq->bic->bfqq[0][act_idx]; /* * The next four mutually-exclusive ifs decide * whether to try injection, and choose the queue to * pick an I/O request from. * * The first if checks whether the process associated * with bfqq has also async I/O pending. If so, it * injects such I/O unconditionally. Injecting async * I/O from the same process can cause no harm to the * process. On the contrary, it can only increase * bandwidth and reduce latency for the process. * * The second if checks whether there happens to be a * non-empty waker queue for bfqq, i.e., a queue whose * I/O needs to be completed for bfqq to receive new * I/O. This happens, e.g., if bfqq is associated with * a process that does some sync. A sync generates * extra blocking I/O, which must be completed before * the process associated with bfqq can go on with its * I/O. If the I/O of the waker queue is not served, * then bfqq remains empty, and no I/O is dispatched, * until the idle timeout fires for bfqq. This is * likely to result in lower bandwidth and higher * latencies for bfqq, and in a severe loss of total * throughput. The best action to take is therefore to * serve the waker queue as soon as possible. So do it * (without relying on the third alternative below for * eventually serving waker_bfqq's I/O; see the last * paragraph for further details). This systematic * injection of I/O from the waker queue does not * cause any delay to bfqq's I/O. On the contrary, * next bfqq's I/O is brought forward dramatically, * for it is not blocked for milliseconds. * * The third if checks whether there is a queue woken * by bfqq, and currently with pending I/O. Such a * woken queue does not steal bandwidth from bfqq, * because it remains soon without I/O if bfqq is not * served. So there is virtually no risk of loss of * bandwidth for bfqq if this woken queue has I/O * dispatched while bfqq is waiting for new I/O. * * The fourth if checks whether bfqq is a queue for * which it is better to avoid injection. It is so if * bfqq delivers more throughput when served without * any further I/O from other queues in the middle, or * if the service times of bfqq's I/O requests both * count more than overall throughput, and may be * easily increased by injection (this happens if bfqq * has a short think time). If none of these * conditions holds, then a candidate queue for * injection is looked for through * bfq_choose_bfqq_for_injection(). Note that the * latter may return NULL (for example if the inject * limit for bfqq is currently 0). * * NOTE: motivation for the second alternative * * Thanks to the way the inject limit is updated in * bfq_update_has_short_ttime(), it is rather likely * that, if I/O is being plugged for bfqq and the * waker queue has pending I/O requests that are * blocking bfqq's I/O, then the fourth alternative * above lets the waker queue get served before the * I/O-plugging timeout fires. So one may deem the * second alternative superfluous. It is not, because * the fourth alternative may be way less effective in * case of a synchronization. For two main * reasons. First, throughput may be low because the * inject limit may be too low to guarantee the same * amount of injected I/O, from the waker queue or * other queues, that the second alternative * guarantees (the second alternative unconditionally * injects a pending I/O request of the waker queue * for each bfq_dispatch_request()). Second, with the * fourth alternative, the duration of the plugging, * i.e., the time before bfqq finally receives new I/O, * may not be minimized, because the waker queue may * happen to be served only after other queues. */ if (async_bfqq && icq_to_bic(async_bfqq->next_rq->elv.icq) == bfqq->bic && bfq_serv_to_charge(async_bfqq->next_rq, async_bfqq) <= bfq_bfqq_budget_left(async_bfqq)) bfqq = async_bfqq; else if (bfqq->waker_bfqq && bfq_bfqq_busy(bfqq->waker_bfqq) && bfqq->waker_bfqq->next_rq && bfq_serv_to_charge(bfqq->waker_bfqq->next_rq, bfqq->waker_bfqq) <= bfq_bfqq_budget_left(bfqq->waker_bfqq) ) bfqq = bfqq->waker_bfqq; else if (blocked_bfqq && bfq_bfqq_busy(blocked_bfqq) && blocked_bfqq->next_rq && bfq_serv_to_charge(blocked_bfqq->next_rq, blocked_bfqq) <= bfq_bfqq_budget_left(blocked_bfqq) ) bfqq = blocked_bfqq; else if (!idling_boosts_thr_without_issues(bfqd, bfqq) && (bfqq->wr_coeff == 1 || bfqd->wr_busy_queues > 1 || !bfq_bfqq_has_short_ttime(bfqq))) bfqq = bfq_choose_bfqq_for_injection(bfqd); else bfqq = NULL; goto keep_queue; } reason = BFQQE_NO_MORE_REQUESTS; expire: bfq_bfqq_expire(bfqd, bfqq, false, reason); new_queue: bfqq = bfq_set_in_service_queue(bfqd); if (bfqq) { bfq_log_bfqq(bfqd, bfqq, "select_queue: checking new queue"); goto check_queue; } keep_queue: if (bfqq) bfq_log_bfqq(bfqd, bfqq, "select_queue: returned this queue"); else bfq_log(bfqd, "select_queue: no queue returned"); return bfqq; } static void bfq_update_wr_data(struct bfq_data *bfqd, struct bfq_queue *bfqq) { struct bfq_entity *entity = &bfqq->entity; if (bfqq->wr_coeff > 1) { /* queue is being weight-raised */ bfq_log_bfqq(bfqd, bfqq, "raising period dur %u/%u msec, old coeff %u, w %d(%d)", jiffies_to_msecs(jiffies - bfqq->last_wr_start_finish), jiffies_to_msecs(bfqq->wr_cur_max_time), bfqq->wr_coeff, bfqq->entity.weight, bfqq->entity.orig_weight); if (entity->prio_changed) bfq_log_bfqq(bfqd, bfqq, "WARN: pending prio change"); /* * If the queue was activated in a burst, or too much * time has elapsed from the beginning of this * weight-raising period, then end weight raising. */ if (bfq_bfqq_in_large_burst(bfqq)) bfq_bfqq_end_wr(bfqq); else if (time_is_before_jiffies(bfqq->last_wr_start_finish + bfqq->wr_cur_max_time)) { if (bfqq->wr_cur_max_time != bfqd->bfq_wr_rt_max_time || time_is_before_jiffies(bfqq->wr_start_at_switch_to_srt + bfq_wr_duration(bfqd))) { /* * Either in interactive weight * raising, or in soft_rt weight * raising with the * interactive-weight-raising period * elapsed (so no switch back to * interactive weight raising). */ bfq_bfqq_end_wr(bfqq); } else { /* * soft_rt finishing while still in * interactive period, switch back to * interactive weight raising */ switch_back_to_interactive_wr(bfqq, bfqd); bfqq->entity.prio_changed = 1; } } if (bfqq->wr_coeff > 1 && bfqq->wr_cur_max_time != bfqd->bfq_wr_rt_max_time && bfqq->service_from_wr > max_service_from_wr) { /* see comments on max_service_from_wr */ bfq_bfqq_end_wr(bfqq); } } /* * To improve latency (for this or other queues), immediately * update weight both if it must be raised and if it must be * lowered. Since, entity may be on some active tree here, and * might have a pending change of its ioprio class, invoke * next function with the last parameter unset (see the * comments on the function). */ if ((entity->weight > entity->orig_weight) != (bfqq->wr_coeff > 1)) __bfq_entity_update_weight_prio(bfq_entity_service_tree(entity), entity, false); } /* * Dispatch next request from bfqq. */ static struct request *bfq_dispatch_rq_from_bfqq(struct bfq_data *bfqd, struct bfq_queue *bfqq) { struct request *rq = bfqq->next_rq; unsigned long service_to_charge; service_to_charge = bfq_serv_to_charge(rq, bfqq); bfq_bfqq_served(bfqq, service_to_charge); if (bfqq == bfqd->in_service_queue && bfqd->wait_dispatch) { bfqd->wait_dispatch = false; bfqd->waited_rq = rq; } bfq_dispatch_remove(bfqd->queue, rq); if (bfqq != bfqd->in_service_queue) return rq; /* * If weight raising has to terminate for bfqq, then next * function causes an immediate update of bfqq's weight, * without waiting for next activation. As a consequence, on * expiration, bfqq will be timestamped as if has never been * weight-raised during this service slot, even if it has * received part or even most of the service as a * weight-raised queue. This inflates bfqq's timestamps, which * is beneficial, as bfqq is then more willing to leave the * device immediately to possible other weight-raised queues. */ bfq_update_wr_data(bfqd, bfqq); /* * Expire bfqq, pretending that its budget expired, if bfqq * belongs to CLASS_IDLE and other queues are waiting for * service. */ if (bfq_tot_busy_queues(bfqd) > 1 && bfq_class_idle(bfqq)) bfq_bfqq_expire(bfqd, bfqq, false, BFQQE_BUDGET_EXHAUSTED); return rq; } static bool bfq_has_work(struct blk_mq_hw_ctx *hctx) { struct bfq_data *bfqd = hctx->queue->elevator->elevator_data; /* * Avoiding lock: a race on bfqd->queued should cause at * most a call to dispatch for nothing */ return !list_empty_careful(&bfqd->dispatch) || READ_ONCE(bfqd->queued); } static struct request *__bfq_dispatch_request(struct blk_mq_hw_ctx *hctx) { struct bfq_data *bfqd = hctx->queue->elevator->elevator_data; struct request *rq = NULL; struct bfq_queue *bfqq = NULL; if (!list_empty(&bfqd->dispatch)) { rq = list_first_entry(&bfqd->dispatch, struct request, queuelist); list_del_init(&rq->queuelist); bfqq = RQ_BFQQ(rq); if (bfqq) { /* * Increment counters here, because this * dispatch does not follow the standard * dispatch flow (where counters are * incremented) */ bfqq->dispatched++; goto inc_in_driver_start_rq; } /* * We exploit the bfq_finish_requeue_request hook to * decrement tot_rq_in_driver, but * bfq_finish_requeue_request will not be invoked on * this request. So, to avoid unbalance, just start * this request, without incrementing tot_rq_in_driver. As * a negative consequence, tot_rq_in_driver is deceptively * lower than it should be while this request is in * service. This may cause bfq_schedule_dispatch to be * invoked uselessly. * * As for implementing an exact solution, the * bfq_finish_requeue_request hook, if defined, is * probably invoked also on this request. So, by * exploiting this hook, we could 1) increment * tot_rq_in_driver here, and 2) decrement it in * bfq_finish_requeue_request. Such a solution would * let the value of the counter be always accurate, * but it would entail using an extra interface * function. This cost seems higher than the benefit, * being the frequency of non-elevator-private * requests very low. */ goto start_rq; } bfq_log(bfqd, "dispatch requests: %d busy queues", bfq_tot_busy_queues(bfqd)); if (bfq_tot_busy_queues(bfqd) == 0) goto exit; /* * Force device to serve one request at a time if * strict_guarantees is true. Forcing this service scheme is * currently the ONLY way to guarantee that the request * service order enforced by the scheduler is respected by a * queueing device. Otherwise the device is free even to make * some unlucky request wait for as long as the device * wishes. * * Of course, serving one request at a time may cause loss of * throughput. */ if (bfqd->strict_guarantees && bfqd->tot_rq_in_driver > 0) goto exit; bfqq = bfq_select_queue(bfqd); if (!bfqq) goto exit; rq = bfq_dispatch_rq_from_bfqq(bfqd, bfqq); if (rq) { inc_in_driver_start_rq: bfqd->rq_in_driver[bfqq->actuator_idx]++; bfqd->tot_rq_in_driver++; start_rq: rq->rq_flags |= RQF_STARTED; } exit: return rq; } #ifdef CONFIG_BFQ_CGROUP_DEBUG static void bfq_update_dispatch_stats(struct request_queue *q, struct request *rq, struct bfq_queue *in_serv_queue, bool idle_timer_disabled) { struct bfq_queue *bfqq = rq ? RQ_BFQQ(rq) : NULL; if (!idle_timer_disabled && !bfqq) return; /* * rq and bfqq are guaranteed to exist until this function * ends, for the following reasons. First, rq can be * dispatched to the device, and then can be completed and * freed, only after this function ends. Second, rq cannot be * merged (and thus freed because of a merge) any longer, * because it has already started. Thus rq cannot be freed * before this function ends, and, since rq has a reference to * bfqq, the same guarantee holds for bfqq too. * * In addition, the following queue lock guarantees that * bfqq_group(bfqq) exists as well. */ spin_lock_irq(&q->queue_lock); if (idle_timer_disabled) /* * Since the idle timer has been disabled, * in_serv_queue contained some request when * __bfq_dispatch_request was invoked above, which * implies that rq was picked exactly from * in_serv_queue. Thus in_serv_queue == bfqq, and is * therefore guaranteed to exist because of the above * arguments. */ bfqg_stats_update_idle_time(bfqq_group(in_serv_queue)); if (bfqq) { struct bfq_group *bfqg = bfqq_group(bfqq); bfqg_stats_update_avg_queue_size(bfqg); bfqg_stats_set_start_empty_time(bfqg); bfqg_stats_update_io_remove(bfqg, rq->cmd_flags); } spin_unlock_irq(&q->queue_lock); } #else static inline void bfq_update_dispatch_stats(struct request_queue *q, struct request *rq, struct bfq_queue *in_serv_queue, bool idle_timer_disabled) {} #endif /* CONFIG_BFQ_CGROUP_DEBUG */ static struct request *bfq_dispatch_request(struct blk_mq_hw_ctx *hctx) { struct bfq_data *bfqd = hctx->queue->elevator->elevator_data; struct request *rq; struct bfq_queue *in_serv_queue; bool waiting_rq, idle_timer_disabled = false; spin_lock_irq(&bfqd->lock); in_serv_queue = bfqd->in_service_queue; waiting_rq = in_serv_queue && bfq_bfqq_wait_request(in_serv_queue); rq = __bfq_dispatch_request(hctx); if (in_serv_queue == bfqd->in_service_queue) { idle_timer_disabled = waiting_rq && !bfq_bfqq_wait_request(in_serv_queue); } spin_unlock_irq(&bfqd->lock); bfq_update_dispatch_stats(hctx->queue, rq, idle_timer_disabled ? in_serv_queue : NULL, idle_timer_disabled); return rq; } /* * Task holds one reference to the queue, dropped when task exits. Each rq * in-flight on this queue also holds a reference, dropped when rq is freed. * * Scheduler lock must be held here. Recall not to use bfqq after calling * this function on it. */ void bfq_put_queue(struct bfq_queue *bfqq) { struct bfq_queue *item; struct hlist_node *n; struct bfq_group *bfqg = bfqq_group(bfqq); bfq_log_bfqq(bfqq->bfqd, bfqq, "put_queue: %p %d", bfqq, bfqq->ref); bfqq->ref--; if (bfqq->ref) return; if (!hlist_unhashed(&bfqq->burst_list_node)) { hlist_del_init(&bfqq->burst_list_node); /* * Decrement also burst size after the removal, if the * process associated with bfqq is exiting, and thus * does not contribute to the burst any longer. This * decrement helps filter out false positives of large * bursts, when some short-lived process (often due to * the execution of commands by some service) happens * to start and exit while a complex application is * starting, and thus spawning several processes that * do I/O (and that *must not* be treated as a large * burst, see comments on bfq_handle_burst). * * In particular, the decrement is performed only if: * 1) bfqq is not a merged queue, because, if it is, * then this free of bfqq is not triggered by the exit * of the process bfqq is associated with, but exactly * by the fact that bfqq has just been merged. * 2) burst_size is greater than 0, to handle * unbalanced decrements. Unbalanced decrements may * happen in te following case: bfqq is inserted into * the current burst list--without incrementing * bust_size--because of a split, but the current * burst list is not the burst list bfqq belonged to * (see comments on the case of a split in * bfq_set_request). */ if (bfqq->bic && bfqq->bfqd->burst_size > 0) bfqq->bfqd->burst_size--; } /* * bfqq does not exist any longer, so it cannot be woken by * any other queue, and cannot wake any other queue. Then bfqq * must be removed from the woken list of its possible waker * queue, and all queues in the woken list of bfqq must stop * having a waker queue. Strictly speaking, these updates * should be performed when bfqq remains with no I/O source * attached to it, which happens before bfqq gets freed. In * particular, this happens when the last process associated * with bfqq exits or gets associated with a different * queue. However, both events lead to bfqq being freed soon, * and dangling references would come out only after bfqq gets * freed. So these updates are done here, as a simple and safe * way to handle all cases. */ /* remove bfqq from woken list */ if (!hlist_unhashed(&bfqq->woken_list_node)) hlist_del_init(&bfqq->woken_list_node); /* reset waker for all queues in woken list */ hlist_for_each_entry_safe(item, n, &bfqq->woken_list, woken_list_node) { item->waker_bfqq = NULL; hlist_del_init(&item->woken_list_node); } if (bfqq->bfqd->last_completed_rq_bfqq == bfqq) bfqq->bfqd->last_completed_rq_bfqq = NULL; WARN_ON_ONCE(!list_empty(&bfqq->fifo)); WARN_ON_ONCE(!RB_EMPTY_ROOT(&bfqq->sort_list)); WARN_ON_ONCE(bfqq->dispatched); kmem_cache_free(bfq_pool, bfqq); bfqg_and_blkg_put(bfqg); } static void bfq_put_stable_ref(struct bfq_queue *bfqq) { bfqq->stable_ref--; bfq_put_queue(bfqq); } void bfq_put_cooperator(struct bfq_queue *bfqq) { struct bfq_queue *__bfqq, *next; /* * If this queue was scheduled to merge with another queue, be * sure to drop the reference taken on that queue (and others in * the merge chain). See bfq_setup_merge and bfq_merge_bfqqs. */ __bfqq = bfqq->new_bfqq; while (__bfqq) { next = __bfqq->new_bfqq; bfq_put_queue(__bfqq); __bfqq = next; } } static void bfq_exit_bfqq(struct bfq_data *bfqd, struct bfq_queue *bfqq) { if (bfqq == bfqd->in_service_queue) { __bfq_bfqq_expire(bfqd, bfqq, BFQQE_BUDGET_TIMEOUT); bfq_schedule_dispatch(bfqd); } bfq_log_bfqq(bfqd, bfqq, "exit_bfqq: %p, %d", bfqq, bfqq->ref); bfq_put_cooperator(bfqq); bfq_release_process_ref(bfqd, bfqq); } static void bfq_exit_icq_bfqq(struct bfq_io_cq *bic, bool is_sync, unsigned int actuator_idx) { struct bfq_queue *bfqq = bic_to_bfqq(bic, is_sync, actuator_idx); struct bfq_data *bfqd; if (bfqq) bfqd = bfqq->bfqd; /* NULL if scheduler already exited */ if (bfqq && bfqd) { bic_set_bfqq(bic, NULL, is_sync, actuator_idx); bfq_exit_bfqq(bfqd, bfqq); } } static void bfq_exit_icq(struct io_cq *icq) { struct bfq_io_cq *bic = icq_to_bic(icq); struct bfq_data *bfqd = bic_to_bfqd(bic); unsigned long flags; unsigned int act_idx; /* * If bfqd and thus bfqd->num_actuators is not available any * longer, then cycle over all possible per-actuator bfqqs in * next loop. We rely on bic being zeroed on creation, and * therefore on its unused per-actuator fields being NULL. */ unsigned int num_actuators = BFQ_MAX_ACTUATORS; struct bfq_iocq_bfqq_data *bfqq_data = bic->bfqq_data; /* * bfqd is NULL if scheduler already exited, and in that case * this is the last time these queues are accessed. */ if (bfqd) { spin_lock_irqsave(&bfqd->lock, flags); num_actuators = bfqd->num_actuators; } for (act_idx = 0; act_idx < num_actuators; act_idx++) { if (bfqq_data[act_idx].stable_merge_bfqq) bfq_put_stable_ref(bfqq_data[act_idx].stable_merge_bfqq); bfq_exit_icq_bfqq(bic, true, act_idx); bfq_exit_icq_bfqq(bic, false, act_idx); } if (bfqd) spin_unlock_irqrestore(&bfqd->lock, flags); } /* * Update the entity prio values; note that the new values will not * be used until the next (re)activation. */ static void bfq_set_next_ioprio_data(struct bfq_queue *bfqq, struct bfq_io_cq *bic) { struct task_struct *tsk = current; int ioprio_class; struct bfq_data *bfqd = bfqq->bfqd; if (!bfqd) return; ioprio_class = IOPRIO_PRIO_CLASS(bic->ioprio); switch (ioprio_class) { default: pr_err("bdi %s: bfq: bad prio class %d\n", bdi_dev_name(bfqq->bfqd->queue->disk->bdi), ioprio_class); fallthrough; case IOPRIO_CLASS_NONE: /* * No prio set, inherit CPU scheduling settings. */ bfqq->new_ioprio = task_nice_ioprio(tsk); bfqq->new_ioprio_class = task_nice_ioclass(tsk); break; case IOPRIO_CLASS_RT: bfqq->new_ioprio = IOPRIO_PRIO_LEVEL(bic->ioprio); bfqq->new_ioprio_class = IOPRIO_CLASS_RT; break; case IOPRIO_CLASS_BE: bfqq->new_ioprio = IOPRIO_PRIO_LEVEL(bic->ioprio); bfqq->new_ioprio_class = IOPRIO_CLASS_BE; break; case IOPRIO_CLASS_IDLE: bfqq->new_ioprio_class = IOPRIO_CLASS_IDLE; bfqq->new_ioprio = IOPRIO_NR_LEVELS - 1; break; } if (bfqq->new_ioprio >= IOPRIO_NR_LEVELS) { pr_crit("bfq_set_next_ioprio_data: new_ioprio %d\n", bfqq->new_ioprio); bfqq->new_ioprio = IOPRIO_NR_LEVELS - 1; } bfqq->entity.new_weight = bfq_ioprio_to_weight(bfqq->new_ioprio); bfq_log_bfqq(bfqd, bfqq, "new_ioprio %d new_weight %d", bfqq->new_ioprio, bfqq->entity.new_weight); bfqq->entity.prio_changed = 1; } static struct bfq_queue *bfq_get_queue(struct bfq_data *bfqd, struct bio *bio, bool is_sync, struct bfq_io_cq *bic, bool respawn); static void bfq_check_ioprio_change(struct bfq_io_cq *bic, struct bio *bio) { struct bfq_data *bfqd = bic_to_bfqd(bic); struct bfq_queue *bfqq; int ioprio = bic->icq.ioc->ioprio; /* * This condition may trigger on a newly created bic, be sure to * drop the lock before returning. */ if (unlikely(!bfqd) || likely(bic->ioprio == ioprio)) return; bic->ioprio = ioprio; bfqq = bic_to_bfqq(bic, false, bfq_actuator_index(bfqd, bio)); if (bfqq) { struct bfq_queue *old_bfqq = bfqq; bfqq = bfq_get_queue(bfqd, bio, false, bic, true); bic_set_bfqq(bic, bfqq, false, bfq_actuator_index(bfqd, bio)); bfq_release_process_ref(bfqd, old_bfqq); } bfqq = bic_to_bfqq(bic, true, bfq_actuator_index(bfqd, bio)); if (bfqq) bfq_set_next_ioprio_data(bfqq, bic); } static void bfq_init_bfqq(struct bfq_data *bfqd, struct bfq_queue *bfqq, struct bfq_io_cq *bic, pid_t pid, int is_sync, unsigned int act_idx) { u64 now_ns = blk_time_get_ns(); bfqq->actuator_idx = act_idx; RB_CLEAR_NODE(&bfqq->entity.rb_node); INIT_LIST_HEAD(&bfqq->fifo); INIT_HLIST_NODE(&bfqq->burst_list_node); INIT_HLIST_NODE(&bfqq->woken_list_node); INIT_HLIST_HEAD(&bfqq->woken_list); bfqq->ref = 0; bfqq->bfqd = bfqd; if (bic) bfq_set_next_ioprio_data(bfqq, bic); if (is_sync) { /* * No need to mark as has_short_ttime if in * idle_class, because no device idling is performed * for queues in idle class */ if (!bfq_class_idle(bfqq)) /* tentatively mark as has_short_ttime */ bfq_mark_bfqq_has_short_ttime(bfqq); bfq_mark_bfqq_sync(bfqq); bfq_mark_bfqq_just_created(bfqq); } else bfq_clear_bfqq_sync(bfqq); /* set end request to minus infinity from now */ bfqq->ttime.last_end_request = now_ns + 1; bfqq->creation_time = jiffies; bfqq->io_start_time = now_ns; bfq_mark_bfqq_IO_bound(bfqq); bfqq->pid = pid; /* Tentative initial value to trade off between thr and lat */ bfqq->max_budget = (2 * bfq_max_budget(bfqd)) / 3; bfqq->budget_timeout = bfq_smallest_from_now(); bfqq->wr_coeff = 1; bfqq->last_wr_start_finish = jiffies; bfqq->wr_start_at_switch_to_srt = bfq_smallest_from_now(); bfqq->split_time = bfq_smallest_from_now(); /* * To not forget the possibly high bandwidth consumed by a * process/queue in the recent past, * bfq_bfqq_softrt_next_start() returns a value at least equal * to the current value of bfqq->soft_rt_next_start (see * comments on bfq_bfqq_softrt_next_start). Set * soft_rt_next_start to now, to mean that bfqq has consumed * no bandwidth so far. */ bfqq->soft_rt_next_start = jiffies; /* first request is almost certainly seeky */ bfqq->seek_history = 1; bfqq->decrease_time_jif = jiffies; } static struct bfq_queue **bfq_async_queue_prio(struct bfq_data *bfqd, struct bfq_group *bfqg, int ioprio_class, int ioprio, int act_idx) { switch (ioprio_class) { case IOPRIO_CLASS_RT: return &bfqg->async_bfqq[0][ioprio][act_idx]; case IOPRIO_CLASS_NONE: ioprio = IOPRIO_BE_NORM; fallthrough; case IOPRIO_CLASS_BE: return &bfqg->async_bfqq[1][ioprio][act_idx]; case IOPRIO_CLASS_IDLE: return &bfqg->async_idle_bfqq[act_idx]; default: return NULL; } } static struct bfq_queue * bfq_do_early_stable_merge(struct bfq_data *bfqd, struct bfq_queue *bfqq, struct bfq_io_cq *bic, struct bfq_queue *last_bfqq_created) { unsigned int a_idx = last_bfqq_created->actuator_idx; struct bfq_queue *new_bfqq = bfq_setup_merge(bfqq, last_bfqq_created); if (!new_bfqq) return bfqq; if (new_bfqq->bic) new_bfqq->bic->bfqq_data[a_idx].stably_merged = true; bic->bfqq_data[a_idx].stably_merged = true; /* * Reusing merge functions. This implies that * bfqq->bic must be set too, for * bfq_merge_bfqqs to correctly save bfqq's * state before killing it. */ bfqq->bic = bic; bfq_merge_bfqqs(bfqd, bic, bfqq, new_bfqq); return new_bfqq; } /* * Many throughput-sensitive workloads are made of several parallel * I/O flows, with all flows generated by the same application, or * more generically by the same task (e.g., system boot). The most * counterproductive action with these workloads is plugging I/O * dispatch when one of the bfq_queues associated with these flows * remains temporarily empty. * * To avoid this plugging, BFQ has been using a burst-handling * mechanism for years now. This mechanism has proven effective for * throughput, and not detrimental for service guarantees. The * following function pushes this mechanism a little bit further, * basing on the following two facts. * * First, all the I/O flows of a the same application or task * contribute to the execution/completion of that common application * or task. So the performance figures that matter are total * throughput of the flows and task-wide I/O latency. In particular, * these flows do not need to be protected from each other, in terms * of individual bandwidth or latency. * * Second, the above fact holds regardless of the number of flows. * * Putting these two facts together, this commits merges stably the * bfq_queues associated with these I/O flows, i.e., with the * processes that generate these IO/ flows, regardless of how many the * involved processes are. * * To decide whether a set of bfq_queues is actually associated with * the I/O flows of a common application or task, and to merge these * queues stably, this function operates as follows: given a bfq_queue, * say Q2, currently being created, and the last bfq_queue, say Q1, * created before Q2, Q2 is merged stably with Q1 if * - very little time has elapsed since when Q1 was created * - Q2 has the same ioprio as Q1 * - Q2 belongs to the same group as Q1 * * Merging bfq_queues also reduces scheduling overhead. A fio test * with ten random readers on /dev/nullb shows a throughput boost of * 40%, with a quadcore. Since BFQ's execution time amounts to ~50% of * the total per-request processing time, the above throughput boost * implies that BFQ's overhead is reduced by more than 50%. * * This new mechanism most certainly obsoletes the current * burst-handling heuristics. We keep those heuristics for the moment. */ static struct bfq_queue *bfq_do_or_sched_stable_merge(struct bfq_data *bfqd, struct bfq_queue *bfqq, struct bfq_io_cq *bic) { struct bfq_queue **source_bfqq = bfqq->entity.parent ? &bfqq->entity.parent->last_bfqq_created : &bfqd->last_bfqq_created; struct bfq_queue *last_bfqq_created = *source_bfqq; /* * If last_bfqq_created has not been set yet, then init it. If * it has been set already, but too long ago, then move it * forward to bfqq. Finally, move also if bfqq belongs to a * different group than last_bfqq_created, or if bfqq has a * different ioprio, ioprio_class or actuator_idx. If none of * these conditions holds true, then try an early stable merge * or schedule a delayed stable merge. As for the condition on * actuator_idx, the reason is that, if queues associated with * different actuators are merged, then control is lost on * each actuator. Therefore some actuator may be * underutilized, and throughput may decrease. * * A delayed merge is scheduled (instead of performing an * early merge), in case bfqq might soon prove to be more * throughput-beneficial if not merged. Currently this is * possible only if bfqd is rotational with no queueing. For * such a drive, not merging bfqq is better for throughput if * bfqq happens to contain sequential I/O. So, we wait a * little bit for enough I/O to flow through bfqq. After that, * if such an I/O is sequential, then the merge is * canceled. Otherwise the merge is finally performed. */ if (!last_bfqq_created || time_before(last_bfqq_created->creation_time + msecs_to_jiffies(bfq_activation_stable_merging), bfqq->creation_time) || bfqq->entity.parent != last_bfqq_created->entity.parent || bfqq->ioprio != last_bfqq_created->ioprio || bfqq->ioprio_class != last_bfqq_created->ioprio_class || bfqq->actuator_idx != last_bfqq_created->actuator_idx) *source_bfqq = bfqq; else if (time_after_eq(last_bfqq_created->creation_time + bfqd->bfq_burst_interval, bfqq->creation_time)) { if (likely(bfqd->nonrot_with_queueing)) /* * With this type of drive, leaving * bfqq alone may provide no * throughput benefits compared with * merging bfqq. So merge bfqq now. */ bfqq = bfq_do_early_stable_merge(bfqd, bfqq, bic, last_bfqq_created); else { /* schedule tentative stable merge */ /* * get reference on last_bfqq_created, * to prevent it from being freed, * until we decide whether to merge */ last_bfqq_created->ref++; /* * need to keep track of stable refs, to * compute process refs correctly */ last_bfqq_created->stable_ref++; /* * Record the bfqq to merge to. */ bic->bfqq_data[last_bfqq_created->actuator_idx].stable_merge_bfqq = last_bfqq_created; } } return bfqq; } static struct bfq_queue *bfq_get_queue(struct bfq_data *bfqd, struct bio *bio, bool is_sync, struct bfq_io_cq *bic, bool respawn) { const int ioprio = IOPRIO_PRIO_LEVEL(bic->ioprio); const int ioprio_class = IOPRIO_PRIO_CLASS(bic->ioprio); struct bfq_queue **async_bfqq = NULL; struct bfq_queue *bfqq; struct bfq_group *bfqg; bfqg = bfq_bio_bfqg(bfqd, bio); if (!is_sync) { async_bfqq = bfq_async_queue_prio(bfqd, bfqg, ioprio_class, ioprio, bfq_actuator_index(bfqd, bio)); bfqq = *async_bfqq; if (bfqq) goto out; } bfqq = kmem_cache_alloc_node(bfq_pool, GFP_NOWAIT | __GFP_ZERO | __GFP_NOWARN, bfqd->queue->node); if (bfqq) { bfq_init_bfqq(bfqd, bfqq, bic, current->pid, is_sync, bfq_actuator_index(bfqd, bio)); bfq_init_entity(&bfqq->entity, bfqg); bfq_log_bfqq(bfqd, bfqq, "allocated"); } else { bfqq = &bfqd->oom_bfqq; bfq_log_bfqq(bfqd, bfqq, "using oom bfqq"); goto out; } /* * Pin the queue now that it's allocated, scheduler exit will * prune it. */ if (async_bfqq) { bfqq->ref++; /* * Extra group reference, w.r.t. sync * queue. This extra reference is removed * only if bfqq->bfqg disappears, to * guarantee that this queue is not freed * until its group goes away. */ bfq_log_bfqq(bfqd, bfqq, "get_queue, bfqq not in async: %p, %d", bfqq, bfqq->ref); *async_bfqq = bfqq; } out: bfqq->ref++; /* get a process reference to this queue */ if (bfqq != &bfqd->oom_bfqq && is_sync && !respawn) bfqq = bfq_do_or_sched_stable_merge(bfqd, bfqq, bic); return bfqq; } static void bfq_update_io_thinktime(struct bfq_data *bfqd, struct bfq_queue *bfqq) { struct bfq_ttime *ttime = &bfqq->ttime; u64 elapsed; /* * We are really interested in how long it takes for the queue to * become busy when there is no outstanding IO for this queue. So * ignore cases when the bfq queue has already IO queued. */ if (bfqq->dispatched || bfq_bfqq_busy(bfqq)) return; elapsed = blk_time_get_ns() - bfqq->ttime.last_end_request; elapsed = min_t(u64, elapsed, 2ULL * bfqd->bfq_slice_idle); ttime->ttime_samples = (7*ttime->ttime_samples + 256) / 8; ttime->ttime_total = div_u64(7*ttime->ttime_total + 256*elapsed, 8); ttime->ttime_mean = div64_ul(ttime->ttime_total + 128, ttime->ttime_samples); } static void bfq_update_io_seektime(struct bfq_data *bfqd, struct bfq_queue *bfqq, struct request *rq) { bfqq->seek_history <<= 1; bfqq->seek_history |= BFQ_RQ_SEEKY(bfqd, bfqq->last_request_pos, rq); if (bfqq->wr_coeff > 1 && bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time && BFQQ_TOTALLY_SEEKY(bfqq)) { if (time_is_before_jiffies(bfqq->wr_start_at_switch_to_srt + bfq_wr_duration(bfqd))) { /* * In soft_rt weight raising with the * interactive-weight-raising period * elapsed (so no switch back to * interactive weight raising). */ bfq_bfqq_end_wr(bfqq); } else { /* * stopping soft_rt weight raising * while still in interactive period, * switch back to interactive weight * raising */ switch_back_to_interactive_wr(bfqq, bfqd); bfqq->entity.prio_changed = 1; } } } static void bfq_update_has_short_ttime(struct bfq_data *bfqd, struct bfq_queue *bfqq, struct bfq_io_cq *bic) { bool has_short_ttime = true, state_changed; /* * No need to update has_short_ttime if bfqq is async or in * idle io prio class, or if bfq_slice_idle is zero, because * no device idling is performed for bfqq in this case. */ if (!bfq_bfqq_sync(bfqq) || bfq_class_idle(bfqq) || bfqd->bfq_slice_idle == 0) return; /* Idle window just restored, statistics are meaningless. */ if (time_is_after_eq_jiffies(bfqq->split_time + bfqd->bfq_wr_min_idle_time)) return; /* Think time is infinite if no process is linked to * bfqq. Otherwise check average think time to decide whether * to mark as has_short_ttime. To this goal, compare average * think time with half the I/O-plugging timeout. */ if (atomic_read(&bic->icq.ioc->active_ref) == 0 || (bfq_sample_valid(bfqq->ttime.ttime_samples) && bfqq->ttime.ttime_mean > bfqd->bfq_slice_idle>>1)) has_short_ttime = false; state_changed = has_short_ttime != bfq_bfqq_has_short_ttime(bfqq); if (has_short_ttime) bfq_mark_bfqq_has_short_ttime(bfqq); else bfq_clear_bfqq_has_short_ttime(bfqq); /* * Until the base value for the total service time gets * finally computed for bfqq, the inject limit does depend on * the think-time state (short|long). In particular, the limit * is 0 or 1 if the think time is deemed, respectively, as * short or long (details in the comments in * bfq_update_inject_limit()). Accordingly, the next * instructions reset the inject limit if the think-time state * has changed and the above base value is still to be * computed. * * However, the reset is performed only if more than 100 ms * have elapsed since the last update of the inject limit, or * (inclusive) if the change is from short to long think * time. The reason for this waiting is as follows. * * bfqq may have a long think time because of a * synchronization with some other queue, i.e., because the * I/O of some other queue may need to be completed for bfqq * to receive new I/O. Details in the comments on the choice * of the queue for injection in bfq_select_queue(). * * As stressed in those comments, if such a synchronization is * actually in place, then, without injection on bfqq, the * blocking I/O cannot happen to served while bfqq is in * service. As a consequence, if bfqq is granted * I/O-dispatch-plugging, then bfqq remains empty, and no I/O * is dispatched, until the idle timeout fires. This is likely * to result in lower bandwidth and higher latencies for bfqq, * and in a severe loss of total throughput. * * On the opposite end, a non-zero inject limit may allow the * I/O that blocks bfqq to be executed soon, and therefore * bfqq to receive new I/O soon. * * But, if the blocking gets actually eliminated, then the * next think-time sample for bfqq may be very low. This in * turn may cause bfqq's think time to be deemed * short. Without the 100 ms barrier, this new state change * would cause the body of the next if to be executed * immediately. But this would set to 0 the inject * limit. Without injection, the blocking I/O would cause the * think time of bfqq to become long again, and therefore the * inject limit to be raised again, and so on. The only effect * of such a steady oscillation between the two think-time * states would be to prevent effective injection on bfqq. * * In contrast, if the inject limit is not reset during such a * long time interval as 100 ms, then the number of short * think time samples can grow significantly before the reset * is performed. As a consequence, the think time state can * become stable before the reset. Therefore there will be no * state change when the 100 ms elapse, and no reset of the * inject limit. The inject limit remains steadily equal to 1 * both during and after the 100 ms. So injection can be * performed at all times, and throughput gets boosted. * * An inject limit equal to 1 is however in conflict, in * general, with the fact that the think time of bfqq is * short, because injection may be likely to delay bfqq's I/O * (as explained in the comments in * bfq_update_inject_limit()). But this does not happen in * this special case, because bfqq's low think time is due to * an effective handling of a synchronization, through * injection. In this special case, bfqq's I/O does not get * delayed by injection; on the contrary, bfqq's I/O is * brought forward, because it is not blocked for * milliseconds. * * In addition, serving the blocking I/O much sooner, and much * more frequently than once per I/O-plugging timeout, makes * it much quicker to detect a waker queue (the concept of * waker queue is defined in the comments in * bfq_add_request()). This makes it possible to start sooner * to boost throughput more effectively, by injecting the I/O * of the waker queue unconditionally on every * bfq_dispatch_request(). * * One last, important benefit of not resetting the inject * limit before 100 ms is that, during this time interval, the * base value for the total service time is likely to get * finally computed for bfqq, freeing the inject limit from * its relation with the think time. */ if (state_changed && bfqq->last_serv_time_ns == 0 && (time_is_before_eq_jiffies(bfqq->decrease_time_jif + msecs_to_jiffies(100)) || !has_short_ttime)) bfq_reset_inject_limit(bfqd, bfqq); } /* * Called when a new fs request (rq) is added to bfqq. Check if there's * something we should do about it. */ static void bfq_rq_enqueued(struct bfq_data *bfqd, struct bfq_queue *bfqq, struct request *rq) { if (rq->cmd_flags & REQ_META) bfqq->meta_pending++; bfqq->last_request_pos = blk_rq_pos(rq) + blk_rq_sectors(rq); if (bfqq == bfqd->in_service_queue && bfq_bfqq_wait_request(bfqq)) { bool small_req = bfqq->queued[rq_is_sync(rq)] == 1 && blk_rq_sectors(rq) < 32; bool budget_timeout = bfq_bfqq_budget_timeout(bfqq); /* * There is just this request queued: if * - the request is small, and * - we are idling to boost throughput, and * - the queue is not to be expired, * then just exit. * * In this way, if the device is being idled to wait * for a new request from the in-service queue, we * avoid unplugging the device and committing the * device to serve just a small request. In contrast * we wait for the block layer to decide when to * unplug the device: hopefully, new requests will be * merged to this one quickly, then the device will be * unplugged and larger requests will be dispatched. */ if (small_req && idling_boosts_thr_without_issues(bfqd, bfqq) && !budget_timeout) return; /* * A large enough request arrived, or idling is being * performed to preserve service guarantees, or * finally the queue is to be expired: in all these * cases disk idling is to be stopped, so clear * wait_request flag and reset timer. */ bfq_clear_bfqq_wait_request(bfqq); hrtimer_try_to_cancel(&bfqd->idle_slice_timer); /* * The queue is not empty, because a new request just * arrived. Hence we can safely expire the queue, in * case of budget timeout, without risking that the * timestamps of the queue are not updated correctly. * See [1] for more details. */ if (budget_timeout) bfq_bfqq_expire(bfqd, bfqq, false, BFQQE_BUDGET_TIMEOUT); } } static void bfqq_request_allocated(struct bfq_queue *bfqq) { struct bfq_entity *entity = &bfqq->entity; for_each_entity(entity) entity->allocated++; } static void bfqq_request_freed(struct bfq_queue *bfqq) { struct bfq_entity *entity = &bfqq->entity; for_each_entity(entity) entity->allocated--; } /* returns true if it causes the idle timer to be disabled */ static bool __bfq_insert_request(struct bfq_data *bfqd, struct request *rq) { struct bfq_queue *bfqq = RQ_BFQQ(rq), *new_bfqq = bfq_setup_cooperator(bfqd, bfqq, rq, true, RQ_BIC(rq)); bool waiting, idle_timer_disabled = false; if (new_bfqq) { /* * Release the request's reference to the old bfqq * and make sure one is taken to the shared queue. */ bfqq_request_allocated(new_bfqq); bfqq_request_freed(bfqq); new_bfqq->ref++; /* * If the bic associated with the process * issuing this request still points to bfqq * (and thus has not been already redirected * to new_bfqq or even some other bfq_queue), * then complete the merge and redirect it to * new_bfqq. */ if (bic_to_bfqq(RQ_BIC(rq), true, bfq_actuator_index(bfqd, rq->bio)) == bfqq) bfq_merge_bfqqs(bfqd, RQ_BIC(rq), bfqq, new_bfqq); bfq_clear_bfqq_just_created(bfqq); /* * rq is about to be enqueued into new_bfqq, * release rq reference on bfqq */ bfq_put_queue(bfqq); rq->elv.priv[1] = new_bfqq; bfqq = new_bfqq; } bfq_update_io_thinktime(bfqd, bfqq); bfq_update_has_short_ttime(bfqd, bfqq, RQ_BIC(rq)); bfq_update_io_seektime(bfqd, bfqq, rq); waiting = bfqq && bfq_bfqq_wait_request(bfqq); bfq_add_request(rq); idle_timer_disabled = waiting && !bfq_bfqq_wait_request(bfqq); rq->fifo_time = blk_time_get_ns() + bfqd->bfq_fifo_expire[rq_is_sync(rq)]; list_add_tail(&rq->queuelist, &bfqq->fifo); bfq_rq_enqueued(bfqd, bfqq, rq); return idle_timer_disabled; } #ifdef CONFIG_BFQ_CGROUP_DEBUG static void bfq_update_insert_stats(struct request_queue *q, struct bfq_queue *bfqq, bool idle_timer_disabled, blk_opf_t cmd_flags) { if (!bfqq) return; /* * bfqq still exists, because it can disappear only after * either it is merged with another queue, or the process it * is associated with exits. But both actions must be taken by * the same process currently executing this flow of * instructions. * * In addition, the following queue lock guarantees that * bfqq_group(bfqq) exists as well. */ spin_lock_irq(&q->queue_lock); bfqg_stats_update_io_add(bfqq_group(bfqq), bfqq, cmd_flags); if (idle_timer_disabled) bfqg_stats_update_idle_time(bfqq_group(bfqq)); spin_unlock_irq(&q->queue_lock); } #else static inline void bfq_update_insert_stats(struct request_queue *q, struct bfq_queue *bfqq, bool idle_timer_disabled, blk_opf_t cmd_flags) {} #endif /* CONFIG_BFQ_CGROUP_DEBUG */ static struct bfq_queue *bfq_init_rq(struct request *rq); static void bfq_insert_request(struct blk_mq_hw_ctx *hctx, struct request *rq, blk_insert_t flags) { struct request_queue *q = hctx->queue; struct bfq_data *bfqd = q->elevator->elevator_data; struct bfq_queue *bfqq; bool idle_timer_disabled = false; blk_opf_t cmd_flags; LIST_HEAD(free); #ifdef CONFIG_BFQ_GROUP_IOSCHED if (!cgroup_subsys_on_dfl(io_cgrp_subsys) && rq->bio) bfqg_stats_update_legacy_io(q, rq); #endif spin_lock_irq(&bfqd->lock); bfqq = bfq_init_rq(rq); if (blk_mq_sched_try_insert_merge(q, rq, &free)) { spin_unlock_irq(&bfqd->lock); blk_mq_free_requests(&free); return; } trace_block_rq_insert(rq); if (flags & BLK_MQ_INSERT_AT_HEAD) { list_add(&rq->queuelist, &bfqd->dispatch); } else if (!bfqq) { list_add_tail(&rq->queuelist, &bfqd->dispatch); } else { idle_timer_disabled = __bfq_insert_request(bfqd, rq); /* * Update bfqq, because, if a queue merge has occurred * in __bfq_insert_request, then rq has been * redirected into a new queue. */ bfqq = RQ_BFQQ(rq); if (rq_mergeable(rq)) { elv_rqhash_add(q, rq); if (!q->last_merge) q->last_merge = rq; } } /* * Cache cmd_flags before releasing scheduler lock, because rq * may disappear afterwards (for example, because of a request * merge). */ cmd_flags = rq->cmd_flags; spin_unlock_irq(&bfqd->lock); bfq_update_insert_stats(q, bfqq, idle_timer_disabled, cmd_flags); } static void bfq_insert_requests(struct blk_mq_hw_ctx *hctx, struct list_head *list, blk_insert_t flags) { while (!list_empty(list)) { struct request *rq; rq = list_first_entry(list, struct request, queuelist); list_del_init(&rq->queuelist); bfq_insert_request(hctx, rq, flags); } } static void bfq_update_hw_tag(struct bfq_data *bfqd) { struct bfq_queue *bfqq = bfqd->in_service_queue; bfqd->max_rq_in_driver = max_t(int, bfqd->max_rq_in_driver, bfqd->tot_rq_in_driver); if (bfqd->hw_tag == 1) return; /* * This sample is valid if the number of outstanding requests * is large enough to allow a queueing behavior. Note that the * sum is not exact, as it's not taking into account deactivated * requests. */ if (bfqd->tot_rq_in_driver + bfqd->queued <= BFQ_HW_QUEUE_THRESHOLD) return; /* * If active queue hasn't enough requests and can idle, bfq might not * dispatch sufficient requests to hardware. Don't zero hw_tag in this * case */ if (bfqq && bfq_bfqq_has_short_ttime(bfqq) && bfqq->dispatched + bfqq->queued[0] + bfqq->queued[1] < BFQ_HW_QUEUE_THRESHOLD && bfqd->tot_rq_in_driver < BFQ_HW_QUEUE_THRESHOLD) return; if (bfqd->hw_tag_samples++ < BFQ_HW_QUEUE_SAMPLES) return; bfqd->hw_tag = bfqd->max_rq_in_driver > BFQ_HW_QUEUE_THRESHOLD; bfqd->max_rq_in_driver = 0; bfqd->hw_tag_samples = 0; bfqd->nonrot_with_queueing = blk_queue_nonrot(bfqd->queue) && bfqd->hw_tag; } static void bfq_completed_request(struct bfq_queue *bfqq, struct bfq_data *bfqd) { u64 now_ns; u32 delta_us; bfq_update_hw_tag(bfqd); bfqd->rq_in_driver[bfqq->actuator_idx]--; bfqd->tot_rq_in_driver--; bfqq->dispatched--; if (!bfqq->dispatched && !bfq_bfqq_busy(bfqq)) { /* * Set budget_timeout (which we overload to store the * time at which the queue remains with no backlog and * no outstanding request; used by the weight-raising * mechanism). */ bfqq->budget_timeout = jiffies; bfq_del_bfqq_in_groups_with_pending_reqs(bfqq); bfq_weights_tree_remove(bfqq); } now_ns = blk_time_get_ns(); bfqq->ttime.last_end_request = now_ns; /* * Using us instead of ns, to get a reasonable precision in * computing rate in next check. */ delta_us = div_u64(now_ns - bfqd->last_completion, NSEC_PER_USEC); /* * If the request took rather long to complete, and, according * to the maximum request size recorded, this completion latency * implies that the request was certainly served at a very low * rate (less than 1M sectors/sec), then the whole observation * interval that lasts up to this time instant cannot be a * valid time interval for computing a new peak rate. Invoke * bfq_update_rate_reset to have the following three steps * taken: * - close the observation interval at the last (previous) * request dispatch or completion * - compute rate, if possible, for that observation interval * - reset to zero samples, which will trigger a proper * re-initialization of the observation interval on next * dispatch */ if (delta_us > BFQ_MIN_TT/NSEC_PER_USEC && (bfqd->last_rq_max_size<<BFQ_RATE_SHIFT)/delta_us < 1UL<<(BFQ_RATE_SHIFT - 10)) bfq_update_rate_reset(bfqd, NULL); bfqd->last_completion = now_ns; /* * Shared queues are likely to receive I/O at a high * rate. This may deceptively let them be considered as wakers * of other queues. But a false waker will unjustly steal * bandwidth to its supposedly woken queue. So considering * also shared queues in the waking mechanism may cause more * control troubles than throughput benefits. Then reset * last_completed_rq_bfqq if bfqq is a shared queue. */ if (!bfq_bfqq_coop(bfqq)) bfqd->last_completed_rq_bfqq = bfqq; else bfqd->last_completed_rq_bfqq = NULL; /* * If we are waiting to discover whether the request pattern * of the task associated with the queue is actually * isochronous, and both requisites for this condition to hold * are now satisfied, then compute soft_rt_next_start (see the * comments on the function bfq_bfqq_softrt_next_start()). We * do not compute soft_rt_next_start if bfqq is in interactive * weight raising (see the comments in bfq_bfqq_expire() for * an explanation). We schedule this delayed update when bfqq * expires, if it still has in-flight requests. */ if (bfq_bfqq_softrt_update(bfqq) && bfqq->dispatched == 0 && RB_EMPTY_ROOT(&bfqq->sort_list) && bfqq->wr_coeff != bfqd->bfq_wr_coeff) bfqq->soft_rt_next_start = bfq_bfqq_softrt_next_start(bfqd, bfqq); /* * If this is the in-service queue, check if it needs to be expired, * or if we want to idle in case it has no pending requests. */ if (bfqd->in_service_queue == bfqq) { if (bfq_bfqq_must_idle(bfqq)) { if (bfqq->dispatched == 0) bfq_arm_slice_timer(bfqd); /* * If we get here, we do not expire bfqq, even * if bfqq was in budget timeout or had no * more requests (as controlled in the next * conditional instructions). The reason for * not expiring bfqq is as follows. * * Here bfqq->dispatched > 0 holds, but * bfq_bfqq_must_idle() returned true. This * implies that, even if no request arrives * for bfqq before bfqq->dispatched reaches 0, * bfqq will, however, not be expired on the * completion event that causes bfqq->dispatch * to reach zero. In contrast, on this event, * bfqq will start enjoying device idling * (I/O-dispatch plugging). * * But, if we expired bfqq here, bfqq would * not have the chance to enjoy device idling * when bfqq->dispatched finally reaches * zero. This would expose bfqq to violation * of its reserved service guarantees. */ return; } else if (bfq_may_expire_for_budg_timeout(bfqq)) bfq_bfqq_expire(bfqd, bfqq, false, BFQQE_BUDGET_TIMEOUT); else if (RB_EMPTY_ROOT(&bfqq->sort_list) && (bfqq->dispatched == 0 || !bfq_better_to_idle(bfqq))) bfq_bfqq_expire(bfqd, bfqq, false, BFQQE_NO_MORE_REQUESTS); } if (!bfqd->tot_rq_in_driver) bfq_schedule_dispatch(bfqd); } /* * The processes associated with bfqq may happen to generate their * cumulative I/O at a lower rate than the rate at which the device * could serve the same I/O. This is rather probable, e.g., if only * one process is associated with bfqq and the device is an SSD. It * results in bfqq becoming often empty while in service. In this * respect, if BFQ is allowed to switch to another queue when bfqq * remains empty, then the device goes on being fed with I/O requests, * and the throughput is not affected. In contrast, if BFQ is not * allowed to switch to another queue---because bfqq is sync and * I/O-dispatch needs to be plugged while bfqq is temporarily * empty---then, during the service of bfqq, there will be frequent * "service holes", i.e., time intervals during which bfqq gets empty * and the device can only consume the I/O already queued in its * hardware queues. During service holes, the device may even get to * remaining idle. In the end, during the service of bfqq, the device * is driven at a lower speed than the one it can reach with the kind * of I/O flowing through bfqq. * * To counter this loss of throughput, BFQ implements a "request * injection mechanism", which tries to fill the above service holes * with I/O requests taken from other queues. The hard part in this * mechanism is finding the right amount of I/O to inject, so as to * both boost throughput and not break bfqq's bandwidth and latency * guarantees. In this respect, the mechanism maintains a per-queue * inject limit, computed as below. While bfqq is empty, the injection * mechanism dispatches extra I/O requests only until the total number * of I/O requests in flight---i.e., already dispatched but not yet * completed---remains lower than this limit. * * A first definition comes in handy to introduce the algorithm by * which the inject limit is computed. We define as first request for * bfqq, an I/O request for bfqq that arrives while bfqq is in * service, and causes bfqq to switch from empty to non-empty. The * algorithm updates the limit as a function of the effect of * injection on the service times of only the first requests of * bfqq. The reason for this restriction is that these are the * requests whose service time is affected most, because they are the * first to arrive after injection possibly occurred. * * To evaluate the effect of injection, the algorithm measures the * "total service time" of first requests. We define as total service * time of an I/O request, the time that elapses since when the * request is enqueued into bfqq, to when it is completed. This * quantity allows the whole effect of injection to be measured. It is * easy to see why. Suppose that some requests of other queues are * actually injected while bfqq is empty, and that a new request R * then arrives for bfqq. If the device does start to serve all or * part of the injected requests during the service hole, then, * because of this extra service, it may delay the next invocation of * the dispatch hook of BFQ. Then, even after R gets eventually * dispatched, the device may delay the actual service of R if it is * still busy serving the extra requests, or if it decides to serve, * before R, some extra request still present in its queues. As a * conclusion, the cumulative extra delay caused by injection can be * easily evaluated by just comparing the total service time of first * requests with and without injection. * * The limit-update algorithm works as follows. On the arrival of a * first request of bfqq, the algorithm measures the total time of the * request only if one of the three cases below holds, and, for each * case, it updates the limit as described below: * * (1) If there is no in-flight request. This gives a baseline for the * total service time of the requests of bfqq. If the baseline has * not been computed yet, then, after computing it, the limit is * set to 1, to start boosting throughput, and to prepare the * ground for the next case. If the baseline has already been * computed, then it is updated, in case it results to be lower * than the previous value. * * (2) If the limit is higher than 0 and there are in-flight * requests. By comparing the total service time in this case with * the above baseline, it is possible to know at which extent the * current value of the limit is inflating the total service * time. If the inflation is below a certain threshold, then bfqq * is assumed to be suffering from no perceivable loss of its * service guarantees, and the limit is even tentatively * increased. If the inflation is above the threshold, then the * limit is decreased. Due to the lack of any hysteresis, this * logic makes the limit oscillate even in steady workload * conditions. Yet we opted for it, because it is fast in reaching * the best value for the limit, as a function of the current I/O * workload. To reduce oscillations, this step is disabled for a * short time interval after the limit happens to be decreased. * * (3) Periodically, after resetting the limit, to make sure that the * limit eventually drops in case the workload changes. This is * needed because, after the limit has gone safely up for a * certain workload, it is impossible to guess whether the * baseline total service time may have changed, without measuring * it again without injection. A more effective version of this * step might be to just sample the baseline, by interrupting * injection only once, and then to reset/lower the limit only if * the total service time with the current limit does happen to be * too large. * * More details on each step are provided in the comments on the * pieces of code that implement these steps: the branch handling the * transition from empty to non empty in bfq_add_request(), the branch * handling injection in bfq_select_queue(), and the function * bfq_choose_bfqq_for_injection(). These comments also explain some * exceptions, made by the injection mechanism in some special cases. */ static void bfq_update_inject_limit(struct bfq_data *bfqd, struct bfq_queue *bfqq) { u64 tot_time_ns = blk_time_get_ns() - bfqd->last_empty_occupied_ns; unsigned int old_limit = bfqq->inject_limit; if (bfqq->last_serv_time_ns > 0 && bfqd->rqs_injected) { u64 threshold = (bfqq->last_serv_time_ns * 3)>>1; if (tot_time_ns >= threshold && old_limit > 0) { bfqq->inject_limit--; bfqq->decrease_time_jif = jiffies; } else if (tot_time_ns < threshold && old_limit <= bfqd->max_rq_in_driver) bfqq->inject_limit++; } /* * Either we still have to compute the base value for the * total service time, and there seem to be the right * conditions to do it, or we can lower the last base value * computed. * * NOTE: (bfqd->tot_rq_in_driver == 1) means that there is no I/O * request in flight, because this function is in the code * path that handles the completion of a request of bfqq, and, * in particular, this function is executed before * bfqd->tot_rq_in_driver is decremented in such a code path. */ if ((bfqq->last_serv_time_ns == 0 && bfqd->tot_rq_in_driver == 1) || tot_time_ns < bfqq->last_serv_time_ns) { if (bfqq->last_serv_time_ns == 0) { /* * Now we certainly have a base value: make sure we * start trying injection. */ bfqq->inject_limit = max_t(unsigned int, 1, old_limit); } bfqq->last_serv_time_ns = tot_time_ns; } else if (!bfqd->rqs_injected && bfqd->tot_rq_in_driver == 1) /* * No I/O injected and no request still in service in * the drive: these are the exact conditions for * computing the base value of the total service time * for bfqq. So let's update this value, because it is * rather variable. For example, it varies if the size * or the spatial locality of the I/O requests in bfqq * change. */ bfqq->last_serv_time_ns = tot_time_ns; /* update complete, not waiting for any request completion any longer */ bfqd->waited_rq = NULL; bfqd->rqs_injected = false; } /* * Handle either a requeue or a finish for rq. The things to do are * the same in both cases: all references to rq are to be dropped. In * particular, rq is considered completed from the point of view of * the scheduler. */ static void bfq_finish_requeue_request(struct request *rq) { struct bfq_queue *bfqq = RQ_BFQQ(rq); struct bfq_data *bfqd; unsigned long flags; /* * rq either is not associated with any icq, or is an already * requeued request that has not (yet) been re-inserted into * a bfq_queue. */ if (!rq->elv.icq || !bfqq) return; bfqd = bfqq->bfqd; if (rq->rq_flags & RQF_STARTED) bfqg_stats_update_completion(bfqq_group(bfqq), rq->start_time_ns, rq->io_start_time_ns, rq->cmd_flags); spin_lock_irqsave(&bfqd->lock, flags); if (likely(rq->rq_flags & RQF_STARTED)) { if (rq == bfqd->waited_rq) bfq_update_inject_limit(bfqd, bfqq); bfq_completed_request(bfqq, bfqd); } bfqq_request_freed(bfqq); bfq_put_queue(bfqq); RQ_BIC(rq)->requests--; spin_unlock_irqrestore(&bfqd->lock, flags); /* * Reset private fields. In case of a requeue, this allows * this function to correctly do nothing if it is spuriously * invoked again on this same request (see the check at the * beginning of the function). Probably, a better general * design would be to prevent blk-mq from invoking the requeue * or finish hooks of an elevator, for a request that is not * referred by that elevator. * * Resetting the following fields would break the * request-insertion logic if rq is re-inserted into a bfq * internal queue, without a re-preparation. Here we assume * that re-insertions of requeued requests, without * re-preparation, can happen only for pass_through or at_head * requests (which are not re-inserted into bfq internal * queues). */ rq->elv.priv[0] = NULL; rq->elv.priv[1] = NULL; } static void bfq_finish_request(struct request *rq) { bfq_finish_requeue_request(rq); if (rq->elv.icq) { put_io_context(rq->elv.icq->ioc); rq->elv.icq = NULL; } } /* * Removes the association between the current task and bfqq, assuming * that bic points to the bfq iocontext of the task. * Returns NULL if a new bfqq should be allocated, or the old bfqq if this * was the last process referring to that bfqq. */ static struct bfq_queue * bfq_split_bfqq(struct bfq_io_cq *bic, struct bfq_queue *bfqq) { bfq_log_bfqq(bfqq->bfqd, bfqq, "splitting queue"); if (bfqq_process_refs(bfqq) == 1) { bfqq->pid = current->pid; bfq_clear_bfqq_coop(bfqq); bfq_clear_bfqq_split_coop(bfqq); return bfqq; } bic_set_bfqq(bic, NULL, true, bfqq->actuator_idx); bfq_put_cooperator(bfqq); bfq_release_process_ref(bfqq->bfqd, bfqq); return NULL; } static struct bfq_queue *bfq_get_bfqq_handle_split(struct bfq_data *bfqd, struct bfq_io_cq *bic, struct bio *bio, bool split, bool is_sync, bool *new_queue) { unsigned int act_idx = bfq_actuator_index(bfqd, bio); struct bfq_queue *bfqq = bic_to_bfqq(bic, is_sync, act_idx); struct bfq_iocq_bfqq_data *bfqq_data = &bic->bfqq_data[act_idx]; if (likely(bfqq && bfqq != &bfqd->oom_bfqq)) return bfqq; if (new_queue) *new_queue = true; if (bfqq) bfq_put_queue(bfqq); bfqq = bfq_get_queue(bfqd, bio, is_sync, bic, split); bic_set_bfqq(bic, bfqq, is_sync, act_idx); if (split && is_sync) { if ((bfqq_data->was_in_burst_list && bfqd->large_burst) || bfqq_data->saved_in_large_burst) bfq_mark_bfqq_in_large_burst(bfqq); else { bfq_clear_bfqq_in_large_burst(bfqq); if (bfqq_data->was_in_burst_list) /* * If bfqq was in the current * burst list before being * merged, then we have to add * it back. And we do not need * to increase burst_size, as * we did not decrement * burst_size when we removed * bfqq from the burst list as * a consequence of a merge * (see comments in * bfq_put_queue). In this * respect, it would be rather * costly to know whether the * current burst list is still * the same burst list from * which bfqq was removed on * the merge. To avoid this * cost, if bfqq was in a * burst list, then we add * bfqq to the current burst * list without any further * check. This can cause * inappropriate insertions, * but rarely enough to not * harm the detection of large * bursts significantly. */ hlist_add_head(&bfqq->burst_list_node, &bfqd->burst_list); } bfqq->split_time = jiffies; } return bfqq; } /* * Only reset private fields. The actual request preparation will be * performed by bfq_init_rq, when rq is either inserted or merged. See * comments on bfq_init_rq for the reason behind this delayed * preparation. */ static void bfq_prepare_request(struct request *rq) { rq->elv.icq = ioc_find_get_icq(rq->q); /* * Regardless of whether we have an icq attached, we have to * clear the scheduler pointers, as they might point to * previously allocated bic/bfqq structs. */ rq->elv.priv[0] = rq->elv.priv[1] = NULL; } /* * If needed, init rq, allocate bfq data structures associated with * rq, and increment reference counters in the destination bfq_queue * for rq. Return the destination bfq_queue for rq, or NULL is rq is * not associated with any bfq_queue. * * This function is invoked by the functions that perform rq insertion * or merging. One may have expected the above preparation operations * to be performed in bfq_prepare_request, and not delayed to when rq * is inserted or merged. The rationale behind this delayed * preparation is that, after the prepare_request hook is invoked for * rq, rq may still be transformed into a request with no icq, i.e., a * request not associated with any queue. No bfq hook is invoked to * signal this transformation. As a consequence, should these * preparation operations be performed when the prepare_request hook * is invoked, and should rq be transformed one moment later, bfq * would end up in an inconsistent state, because it would have * incremented some queue counters for an rq destined to * transformation, without any chance to correctly lower these * counters back. In contrast, no transformation can still happen for * rq after rq has been inserted or merged. So, it is safe to execute * these preparation operations when rq is finally inserted or merged. */ static struct bfq_queue *bfq_init_rq(struct request *rq) { struct request_queue *q = rq->q; struct bio *bio = rq->bio; struct bfq_data *bfqd = q->elevator->elevator_data; struct bfq_io_cq *bic; const int is_sync = rq_is_sync(rq); struct bfq_queue *bfqq; bool new_queue = false; bool bfqq_already_existing = false, split = false; unsigned int a_idx = bfq_actuator_index(bfqd, bio); if (unlikely(!rq->elv.icq)) return NULL; /* * Assuming that RQ_BFQQ(rq) is set only if everything is set * for this rq. This holds true, because this function is * invoked only for insertion or merging, and, after such * events, a request cannot be manipulated any longer before * being removed from bfq. */ if (RQ_BFQQ(rq)) return RQ_BFQQ(rq); bic = icq_to_bic(rq->elv.icq); bfq_check_ioprio_change(bic, bio); bfq_bic_update_cgroup(bic, bio); bfqq = bfq_get_bfqq_handle_split(bfqd, bic, bio, false, is_sync, &new_queue); if (likely(!new_queue)) { /* If the queue was seeky for too long, break it apart. */ if (bfq_bfqq_coop(bfqq) && bfq_bfqq_split_coop(bfqq) && !bic->bfqq_data[a_idx].stably_merged) { struct bfq_queue *old_bfqq = bfqq; /* Update bic before losing reference to bfqq */ if (bfq_bfqq_in_large_burst(bfqq)) bic->bfqq_data[a_idx].saved_in_large_burst = true; bfqq = bfq_split_bfqq(bic, bfqq); split = true; if (!bfqq) { bfqq = bfq_get_bfqq_handle_split(bfqd, bic, bio, true, is_sync, NULL); if (unlikely(bfqq == &bfqd->oom_bfqq)) bfqq_already_existing = true; } else bfqq_already_existing = true; if (!bfqq_already_existing) { bfqq->waker_bfqq = old_bfqq->waker_bfqq; bfqq->tentative_waker_bfqq = NULL; /* * If the waker queue disappears, then * new_bfqq->waker_bfqq must be * reset. So insert new_bfqq into the * woken_list of the waker. See * bfq_check_waker for details. */ if (bfqq->waker_bfqq) hlist_add_head(&bfqq->woken_list_node, &bfqq->waker_bfqq->woken_list); } } } bfqq_request_allocated(bfqq); bfqq->ref++; bic->requests++; bfq_log_bfqq(bfqd, bfqq, "get_request %p: bfqq %p, %d", rq, bfqq, bfqq->ref); rq->elv.priv[0] = bic; rq->elv.priv[1] = bfqq; /* * If a bfq_queue has only one process reference, it is owned * by only this bic: we can then set bfqq->bic = bic. in * addition, if the queue has also just been split, we have to * resume its state. */ if (likely(bfqq != &bfqd->oom_bfqq) && bfqq_process_refs(bfqq) == 1) { bfqq->bic = bic; if (split) { /* * The queue has just been split from a shared * queue: restore the idle window and the * possible weight raising period. */ bfq_bfqq_resume_state(bfqq, bfqd, bic, bfqq_already_existing); } } /* * Consider bfqq as possibly belonging to a burst of newly * created queues only if: * 1) A burst is actually happening (bfqd->burst_size > 0) * or * 2) There is no other active queue. In fact, if, in * contrast, there are active queues not belonging to the * possible burst bfqq may belong to, then there is no gain * in considering bfqq as belonging to a burst, and * therefore in not weight-raising bfqq. See comments on * bfq_handle_burst(). * * This filtering also helps eliminating false positives, * occurring when bfqq does not belong to an actual large * burst, but some background task (e.g., a service) happens * to trigger the creation of new queues very close to when * bfqq and its possible companion queues are created. See * comments on bfq_handle_burst() for further details also on * this issue. */ if (unlikely(bfq_bfqq_just_created(bfqq) && (bfqd->burst_size > 0 || bfq_tot_busy_queues(bfqd) == 0))) bfq_handle_burst(bfqd, bfqq); return bfqq; } static void bfq_idle_slice_timer_body(struct bfq_data *bfqd, struct bfq_queue *bfqq) { enum bfqq_expiration reason; unsigned long flags; spin_lock_irqsave(&bfqd->lock, flags); /* * Considering that bfqq may be in race, we should firstly check * whether bfqq is in service before doing something on it. If * the bfqq in race is not in service, it has already been expired * through __bfq_bfqq_expire func and its wait_request flags has * been cleared in __bfq_bfqd_reset_in_service func. */ if (bfqq != bfqd->in_service_queue) { spin_unlock_irqrestore(&bfqd->lock, flags); return; } bfq_clear_bfqq_wait_request(bfqq); if (bfq_bfqq_budget_timeout(bfqq)) /* * Also here the queue can be safely expired * for budget timeout without wasting * guarantees */ reason = BFQQE_BUDGET_TIMEOUT; else if (bfqq->queued[0] == 0 && bfqq->queued[1] == 0) /* * The queue may not be empty upon timer expiration, * because we may not disable the timer when the * first request of the in-service queue arrives * during disk idling. */ reason = BFQQE_TOO_IDLE; else goto schedule_dispatch; bfq_bfqq_expire(bfqd, bfqq, true, reason); schedule_dispatch: bfq_schedule_dispatch(bfqd); spin_unlock_irqrestore(&bfqd->lock, flags); } /* * Handler of the expiration of the timer running if the in-service queue * is idling inside its time slice. */ static enum hrtimer_restart bfq_idle_slice_timer(struct hrtimer *timer) { struct bfq_data *bfqd = container_of(timer, struct bfq_data, idle_slice_timer); struct bfq_queue *bfqq = bfqd->in_service_queue; /* * Theoretical race here: the in-service queue can be NULL or * different from the queue that was idling if a new request * arrives for the current queue and there is a full dispatch * cycle that changes the in-service queue. This can hardly * happen, but in the worst case we just expire a queue too * early. */ if (bfqq) bfq_idle_slice_timer_body(bfqd, bfqq); return HRTIMER_NORESTART; } static void __bfq_put_async_bfqq(struct bfq_data *bfqd, struct bfq_queue **bfqq_ptr) { struct bfq_queue *bfqq = *bfqq_ptr; bfq_log(bfqd, "put_async_bfqq: %p", bfqq); if (bfqq) { bfq_bfqq_move(bfqd, bfqq, bfqd->root_group); bfq_log_bfqq(bfqd, bfqq, "put_async_bfqq: putting %p, %d", bfqq, bfqq->ref); bfq_put_queue(bfqq); *bfqq_ptr = NULL; } } /* * Release all the bfqg references to its async queues. If we are * deallocating the group these queues may still contain requests, so * we reparent them to the root cgroup (i.e., the only one that will * exist for sure until all the requests on a device are gone). */ void bfq_put_async_queues(struct bfq_data *bfqd, struct bfq_group *bfqg) { int i, j, k; for (k = 0; k < bfqd->num_actuators; k++) { for (i = 0; i < 2; i++) for (j = 0; j < IOPRIO_NR_LEVELS; j++) __bfq_put_async_bfqq(bfqd, &bfqg->async_bfqq[i][j][k]); __bfq_put_async_bfqq(bfqd, &bfqg->async_idle_bfqq[k]); } } /* * See the comments on bfq_limit_depth for the purpose of * the depths set in the function. Return minimum shallow depth we'll use. */ static void bfq_update_depths(struct bfq_data *bfqd, struct sbitmap_queue *bt) { unsigned int depth = 1U << bt->sb.shift; bfqd->full_depth_shift = bt->sb.shift; /* * In-word depths if no bfq_queue is being weight-raised: * leaving 25% of tags only for sync reads. * * In next formulas, right-shift the value * (1U<<bt->sb.shift), instead of computing directly * (1U<<(bt->sb.shift - something)), to be robust against * any possible value of bt->sb.shift, without having to * limit 'something'. */ /* no more than 50% of tags for async I/O */ bfqd->word_depths[0][0] = max(depth >> 1, 1U); /* * no more than 75% of tags for sync writes (25% extra tags * w.r.t. async I/O, to prevent async I/O from starving sync * writes) */ bfqd->word_depths[0][1] = max((depth * 3) >> 2, 1U); /* * In-word depths in case some bfq_queue is being weight- * raised: leaving ~63% of tags for sync reads. This is the * highest percentage for which, in our tests, application * start-up times didn't suffer from any regression due to tag * shortage. */ /* no more than ~18% of tags for async I/O */ bfqd->word_depths[1][0] = max((depth * 3) >> 4, 1U); /* no more than ~37% of tags for sync writes (~20% extra tags) */ bfqd->word_depths[1][1] = max((depth * 6) >> 4, 1U); } static void bfq_depth_updated(struct blk_mq_hw_ctx *hctx) { struct bfq_data *bfqd = hctx->queue->elevator->elevator_data; struct blk_mq_tags *tags = hctx->sched_tags; bfq_update_depths(bfqd, &tags->bitmap_tags); sbitmap_queue_min_shallow_depth(&tags->bitmap_tags, 1); } static int bfq_init_hctx(struct blk_mq_hw_ctx *hctx, unsigned int index) { bfq_depth_updated(hctx); return 0; } static void bfq_exit_queue(struct elevator_queue *e) { struct bfq_data *bfqd = e->elevator_data; struct bfq_queue *bfqq, *n; unsigned int actuator; hrtimer_cancel(&bfqd->idle_slice_timer); spin_lock_irq(&bfqd->lock); list_for_each_entry_safe(bfqq, n, &bfqd->idle_list, bfqq_list) bfq_deactivate_bfqq(bfqd, bfqq, false, false); spin_unlock_irq(&bfqd->lock); for (actuator = 0; actuator < bfqd->num_actuators; actuator++) WARN_ON_ONCE(bfqd->rq_in_driver[actuator]); WARN_ON_ONCE(bfqd->tot_rq_in_driver); hrtimer_cancel(&bfqd->idle_slice_timer); /* release oom-queue reference to root group */ bfqg_and_blkg_put(bfqd->root_group); #ifdef CONFIG_BFQ_GROUP_IOSCHED blkcg_deactivate_policy(bfqd->queue->disk, &blkcg_policy_bfq); #else spin_lock_irq(&bfqd->lock); bfq_put_async_queues(bfqd, bfqd->root_group); kfree(bfqd->root_group); spin_unlock_irq(&bfqd->lock); #endif blk_stat_disable_accounting(bfqd->queue); clear_bit(ELEVATOR_FLAG_DISABLE_WBT, &e->flags); wbt_enable_default(bfqd->queue->disk); kfree(bfqd); } static void bfq_init_root_group(struct bfq_group *root_group, struct bfq_data *bfqd) { int i; #ifdef CONFIG_BFQ_GROUP_IOSCHED root_group->entity.parent = NULL; root_group->my_entity = NULL; root_group->bfqd = bfqd; #endif root_group->rq_pos_tree = RB_ROOT; for (i = 0; i < BFQ_IOPRIO_CLASSES; i++) root_group->sched_data.service_tree[i] = BFQ_SERVICE_TREE_INIT; root_group->sched_data.bfq_class_idle_last_service = jiffies; } static int bfq_init_queue(struct request_queue *q, struct elevator_type *e) { struct bfq_data *bfqd; struct elevator_queue *eq; unsigned int i; struct blk_independent_access_ranges *ia_ranges = q->disk->ia_ranges; eq = elevator_alloc(q, e); if (!eq) return -ENOMEM; bfqd = kzalloc_node(sizeof(*bfqd), GFP_KERNEL, q->node); if (!bfqd) { kobject_put(&eq->kobj); return -ENOMEM; } eq->elevator_data = bfqd; spin_lock_irq(&q->queue_lock); q->elevator = eq; spin_unlock_irq(&q->queue_lock); /* * Our fallback bfqq if bfq_find_alloc_queue() runs into OOM issues. * Grab a permanent reference to it, so that the normal code flow * will not attempt to free it. * Set zero as actuator index: we will pretend that * all I/O requests are for the same actuator. */ bfq_init_bfqq(bfqd, &bfqd->oom_bfqq, NULL, 1, 0, 0); bfqd->oom_bfqq.ref++; bfqd->oom_bfqq.new_ioprio = BFQ_DEFAULT_QUEUE_IOPRIO; bfqd->oom_bfqq.new_ioprio_class = IOPRIO_CLASS_BE; bfqd->oom_bfqq.entity.new_weight = bfq_ioprio_to_weight(bfqd->oom_bfqq.new_ioprio); /* oom_bfqq does not participate to bursts */ bfq_clear_bfqq_just_created(&bfqd->oom_bfqq); /* * Trigger weight initialization, according to ioprio, at the * oom_bfqq's first activation. The oom_bfqq's ioprio and ioprio * class won't be changed any more. */ bfqd->oom_bfqq.entity.prio_changed = 1; bfqd->queue = q; bfqd->num_actuators = 1; /* * If the disk supports multiple actuators, copy independent * access ranges from the request queue structure. */ spin_lock_irq(&q->queue_lock); if (ia_ranges) { /* * Check if the disk ia_ranges size exceeds the current bfq * actuator limit. */ if (ia_ranges->nr_ia_ranges > BFQ_MAX_ACTUATORS) { pr_crit("nr_ia_ranges higher than act limit: iars=%d, max=%d.\n", ia_ranges->nr_ia_ranges, BFQ_MAX_ACTUATORS); pr_crit("Falling back to single actuator mode.\n"); } else { bfqd->num_actuators = ia_ranges->nr_ia_ranges; for (i = 0; i < bfqd->num_actuators; i++) { bfqd->sector[i] = ia_ranges->ia_range[i].sector; bfqd->nr_sectors[i] = ia_ranges->ia_range[i].nr_sectors; } } } /* Otherwise use single-actuator dev info */ if (bfqd->num_actuators == 1) { bfqd->sector[0] = 0; bfqd->nr_sectors[0] = get_capacity(q->disk); } spin_unlock_irq(&q->queue_lock); INIT_LIST_HEAD(&bfqd->dispatch); hrtimer_init(&bfqd->idle_slice_timer, CLOCK_MONOTONIC, HRTIMER_MODE_REL); bfqd->idle_slice_timer.function = bfq_idle_slice_timer; bfqd->queue_weights_tree = RB_ROOT_CACHED; #ifdef CONFIG_BFQ_GROUP_IOSCHED bfqd->num_groups_with_pending_reqs = 0; #endif INIT_LIST_HEAD(&bfqd->active_list[0]); INIT_LIST_HEAD(&bfqd->active_list[1]); INIT_LIST_HEAD(&bfqd->idle_list); INIT_HLIST_HEAD(&bfqd->burst_list); bfqd->hw_tag = -1; bfqd->nonrot_with_queueing = blk_queue_nonrot(bfqd->queue); bfqd->bfq_max_budget = bfq_default_max_budget; bfqd->bfq_fifo_expire[0] = bfq_fifo_expire[0]; bfqd->bfq_fifo_expire[1] = bfq_fifo_expire[1]; bfqd->bfq_back_max = bfq_back_max; bfqd->bfq_back_penalty = bfq_back_penalty; bfqd->bfq_slice_idle = bfq_slice_idle; bfqd->bfq_timeout = bfq_timeout; bfqd->bfq_large_burst_thresh = 8; bfqd->bfq_burst_interval = msecs_to_jiffies(180); bfqd->low_latency = true; /* * Trade-off between responsiveness and fairness. */ bfqd->bfq_wr_coeff = 30; bfqd->bfq_wr_rt_max_time = msecs_to_jiffies(300); bfqd->bfq_wr_min_idle_time = msecs_to_jiffies(2000); bfqd->bfq_wr_min_inter_arr_async = msecs_to_jiffies(500); bfqd->bfq_wr_max_softrt_rate = 7000; /* * Approximate rate required * to playback or record a * high-definition compressed * video. */ bfqd->wr_busy_queues = 0; /* * Begin by assuming, optimistically, that the device peak * rate is equal to 2/3 of the highest reference rate. */ bfqd->rate_dur_prod = ref_rate[blk_queue_nonrot(bfqd->queue)] * ref_wr_duration[blk_queue_nonrot(bfqd->queue)]; bfqd->peak_rate = ref_rate[blk_queue_nonrot(bfqd->queue)] * 2 / 3; /* see comments on the definition of next field inside bfq_data */ bfqd->actuator_load_threshold = 4; spin_lock_init(&bfqd->lock); /* * The invocation of the next bfq_create_group_hierarchy * function is the head of a chain of function calls * (bfq_create_group_hierarchy->blkcg_activate_policy-> * blk_mq_freeze_queue) that may lead to the invocation of the * has_work hook function. For this reason, * bfq_create_group_hierarchy is invoked only after all * scheduler data has been initialized, apart from the fields * that can be initialized only after invoking * bfq_create_group_hierarchy. This, in particular, enables * has_work to correctly return false. Of course, to avoid * other inconsistencies, the blk-mq stack must then refrain * from invoking further scheduler hooks before this init * function is finished. */ bfqd->root_group = bfq_create_group_hierarchy(bfqd, q->node); if (!bfqd->root_group) goto out_free; bfq_init_root_group(bfqd->root_group, bfqd); bfq_init_entity(&bfqd->oom_bfqq.entity, bfqd->root_group); /* We dispatch from request queue wide instead of hw queue */ blk_queue_flag_set(QUEUE_FLAG_SQ_SCHED, q); set_bit(ELEVATOR_FLAG_DISABLE_WBT, &eq->flags); wbt_disable_default(q->disk); blk_stat_enable_accounting(q); return 0; out_free: kfree(bfqd); kobject_put(&eq->kobj); return -ENOMEM; } static void bfq_slab_kill(void) { kmem_cache_destroy(bfq_pool); } static int __init bfq_slab_setup(void) { bfq_pool = KMEM_CACHE(bfq_queue, 0); if (!bfq_pool) return -ENOMEM; return 0; } static ssize_t bfq_var_show(unsigned int var, char *page) { return sprintf(page, "%u\n", var); } static int bfq_var_store(unsigned long *var, const char *page) { unsigned long new_val; int ret = kstrtoul(page, 10, &new_val); if (ret) return ret; *var = new_val; return 0; } #define SHOW_FUNCTION(__FUNC, __VAR, __CONV) \ static ssize_t __FUNC(struct elevator_queue *e, char *page) \ { \ struct bfq_data *bfqd = e->elevator_data; \ u64 __data = __VAR; \ if (__CONV == 1) \ __data = jiffies_to_msecs(__data); \ else if (__CONV == 2) \ __data = div_u64(__data, NSEC_PER_MSEC); \ return bfq_var_show(__data, (page)); \ } SHOW_FUNCTION(bfq_fifo_expire_sync_show, bfqd->bfq_fifo_expire[1], 2); SHOW_FUNCTION(bfq_fifo_expire_async_show, bfqd->bfq_fifo_expire[0], 2); SHOW_FUNCTION(bfq_back_seek_max_show, bfqd->bfq_back_max, 0); SHOW_FUNCTION(bfq_back_seek_penalty_show, bfqd->bfq_back_penalty, 0); SHOW_FUNCTION(bfq_slice_idle_show, bfqd->bfq_slice_idle, 2); SHOW_FUNCTION(bfq_max_budget_show, bfqd->bfq_user_max_budget, 0); SHOW_FUNCTION(bfq_timeout_sync_show, bfqd->bfq_timeout, 1); SHOW_FUNCTION(bfq_strict_guarantees_show, bfqd->strict_guarantees, 0); SHOW_FUNCTION(bfq_low_latency_show, bfqd->low_latency, 0); #undef SHOW_FUNCTION #define USEC_SHOW_FUNCTION(__FUNC, __VAR) \ static ssize_t __FUNC(struct elevator_queue *e, char *page) \ { \ struct bfq_data *bfqd = e->elevator_data; \ u64 __data = __VAR; \ __data = div_u64(__data, NSEC_PER_USEC); \ return bfq_var_show(__data, (page)); \ } USEC_SHOW_FUNCTION(bfq_slice_idle_us_show, bfqd->bfq_slice_idle); #undef USEC_SHOW_FUNCTION #define STORE_FUNCTION(__FUNC, __PTR, MIN, MAX, __CONV) \ static ssize_t \ __FUNC(struct elevator_queue *e, const char *page, size_t count) \ { \ struct bfq_data *bfqd = e->elevator_data; \ unsigned long __data, __min = (MIN), __max = (MAX); \ int ret; \ \ ret = bfq_var_store(&__data, (page)); \ if (ret) \ return ret; \ if (__data < __min) \ __data = __min; \ else if (__data > __max) \ __data = __max; \ if (__CONV == 1) \ *(__PTR) = msecs_to_jiffies(__data); \ else if (__CONV == 2) \ *(__PTR) = (u64)__data * NSEC_PER_MSEC; \ else \ *(__PTR) = __data; \ return count; \ } STORE_FUNCTION(bfq_fifo_expire_sync_store, &bfqd->bfq_fifo_expire[1], 1, INT_MAX, 2); STORE_FUNCTION(bfq_fifo_expire_async_store, &bfqd->bfq_fifo_expire[0], 1, INT_MAX, 2); STORE_FUNCTION(bfq_back_seek_max_store, &bfqd->bfq_back_max, 0, INT_MAX, 0); STORE_FUNCTION(bfq_back_seek_penalty_store, &bfqd->bfq_back_penalty, 1, INT_MAX, 0); STORE_FUNCTION(bfq_slice_idle_store, &bfqd->bfq_slice_idle, 0, INT_MAX, 2); #undef STORE_FUNCTION #define USEC_STORE_FUNCTION(__FUNC, __PTR, MIN, MAX) \ static ssize_t __FUNC(struct elevator_queue *e, const char *page, size_t count)\ { \ struct bfq_data *bfqd = e->elevator_data; \ unsigned long __data, __min = (MIN), __max = (MAX); \ int ret; \ \ ret = bfq_var_store(&__data, (page)); \ if (ret) \ return ret; \ if (__data < __min) \ __data = __min; \ else if (__data > __max) \ __data = __max; \ *(__PTR) = (u64)__data * NSEC_PER_USEC; \ return count; \ } USEC_STORE_FUNCTION(bfq_slice_idle_us_store, &bfqd->bfq_slice_idle, 0, UINT_MAX); #undef USEC_STORE_FUNCTION static ssize_t bfq_max_budget_store(struct elevator_queue *e, const char *page, size_t count) { struct bfq_data *bfqd = e->elevator_data; unsigned long __data; int ret; ret = bfq_var_store(&__data, (page)); if (ret) return ret; if (__data == 0) bfqd->bfq_max_budget = bfq_calc_max_budget(bfqd); else { if (__data > INT_MAX) __data = INT_MAX; bfqd->bfq_max_budget = __data; } bfqd->bfq_user_max_budget = __data; return count; } /* * Leaving this name to preserve name compatibility with cfq * parameters, but this timeout is used for both sync and async. */ static ssize_t bfq_timeout_sync_store(struct elevator_queue *e, const char *page, size_t count) { struct bfq_data *bfqd = e->elevator_data; unsigned long __data; int ret; ret = bfq_var_store(&__data, (page)); if (ret) return ret; if (__data < 1) __data = 1; else if (__data > INT_MAX) __data = INT_MAX; bfqd->bfq_timeout = msecs_to_jiffies(__data); if (bfqd->bfq_user_max_budget == 0) bfqd->bfq_max_budget = bfq_calc_max_budget(bfqd); return count; } static ssize_t bfq_strict_guarantees_store(struct elevator_queue *e, const char *page, size_t count) { struct bfq_data *bfqd = e->elevator_data; unsigned long __data; int ret; ret = bfq_var_store(&__data, (page)); if (ret) return ret; if (__data > 1) __data = 1; if (!bfqd->strict_guarantees && __data == 1 && bfqd->bfq_slice_idle < 8 * NSEC_PER_MSEC) bfqd->bfq_slice_idle = 8 * NSEC_PER_MSEC; bfqd->strict_guarantees = __data; return count; } static ssize_t bfq_low_latency_store(struct elevator_queue *e, const char *page, size_t count) { struct bfq_data *bfqd = e->elevator_data; unsigned long __data; int ret; ret = bfq_var_store(&__data, (page)); if (ret) return ret; if (__data > 1) __data = 1; if (__data == 0 && bfqd->low_latency != 0) bfq_end_wr(bfqd); bfqd->low_latency = __data; return count; } #define BFQ_ATTR(name) \ __ATTR(name, 0644, bfq_##name##_show, bfq_##name##_store) static struct elv_fs_entry bfq_attrs[] = { BFQ_ATTR(fifo_expire_sync), BFQ_ATTR(fifo_expire_async), BFQ_ATTR(back_seek_max), BFQ_ATTR(back_seek_penalty), BFQ_ATTR(slice_idle), BFQ_ATTR(slice_idle_us), BFQ_ATTR(max_budget), BFQ_ATTR(timeout_sync), BFQ_ATTR(strict_guarantees), BFQ_ATTR(low_latency), __ATTR_NULL }; static struct elevator_type iosched_bfq_mq = { .ops = { .limit_depth = bfq_limit_depth, .prepare_request = bfq_prepare_request, .requeue_request = bfq_finish_requeue_request, .finish_request = bfq_finish_request, .exit_icq = bfq_exit_icq, .insert_requests = bfq_insert_requests, .dispatch_request = bfq_dispatch_request, .next_request = elv_rb_latter_request, .former_request = elv_rb_former_request, .allow_merge = bfq_allow_bio_merge, .bio_merge = bfq_bio_merge, .request_merge = bfq_request_merge, .requests_merged = bfq_requests_merged, .request_merged = bfq_request_merged, .has_work = bfq_has_work, .depth_updated = bfq_depth_updated, .init_hctx = bfq_init_hctx, .init_sched = bfq_init_queue, .exit_sched = bfq_exit_queue, }, .icq_size = sizeof(struct bfq_io_cq), .icq_align = __alignof__(struct bfq_io_cq), .elevator_attrs = bfq_attrs, .elevator_name = "bfq", .elevator_owner = THIS_MODULE, }; MODULE_ALIAS("bfq-iosched"); static int __init bfq_init(void) { int ret; #ifdef CONFIG_BFQ_GROUP_IOSCHED ret = blkcg_policy_register(&blkcg_policy_bfq); if (ret) return ret; #endif ret = -ENOMEM; if (bfq_slab_setup()) goto err_pol_unreg; /* * Times to load large popular applications for the typical * systems installed on the reference devices (see the * comments before the definition of the next * array). Actually, we use slightly lower values, as the * estimated peak rate tends to be smaller than the actual * peak rate. The reason for this last fact is that estimates * are computed over much shorter time intervals than the long * intervals typically used for benchmarking. Why? First, to * adapt more quickly to variations. Second, because an I/O * scheduler cannot rely on a peak-rate-evaluation workload to * be run for a long time. */ ref_wr_duration[0] = msecs_to_jiffies(7000); /* actually 8 sec */ ref_wr_duration[1] = msecs_to_jiffies(2500); /* actually 3 sec */ ret = elv_register(&iosched_bfq_mq); if (ret) goto slab_kill; return 0; slab_kill: bfq_slab_kill(); err_pol_unreg: #ifdef CONFIG_BFQ_GROUP_IOSCHED blkcg_policy_unregister(&blkcg_policy_bfq); #endif return ret; } static void __exit bfq_exit(void) { elv_unregister(&iosched_bfq_mq); #ifdef CONFIG_BFQ_GROUP_IOSCHED blkcg_policy_unregister(&blkcg_policy_bfq); #endif bfq_slab_kill(); } module_init(bfq_init); module_exit(bfq_exit); MODULE_AUTHOR("Paolo Valente"); MODULE_LICENSE("GPL"); MODULE_DESCRIPTION("MQ Budget Fair Queueing I/O Scheduler");
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