blob: 480e1a13485966104362f265d842887640775781 [file] [log] [blame]
// 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 <>
* Copyright (C) 2008 Fabio Checconi <>
* Paolo Valente <>
* Copyright (C) 2010 Paolo Valente <>
* Arianna Avanzini <>
* Copyright (C) 2017 Paolo Valente <>
* 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.
* [2] Jon C.R. Bennett and H. Zhang, "Hierarchical Packet Fair Queueing
* Algorithms", IEEE/ACM Transactions on Networking, 5(5):675-689,
* Oct 1997.
* [3] I. Stoica and H. Abdel-Wahab, "Earliest Eligible Virtual Deadline
* First: A Flexible and Accurate Mechanism for Proportional Share
* Resource Allocation", technical report.
#include <linux/module.h>
#include <linux/slab.h>
#include <linux/blkdev.h>
#include <linux/cgroup.h>
#include <linux/elevator.h>
#include <linux/ktime.h>
#include <linux/rbtree.h>
#include <linux/ioprio.h>
#include <linux/sbitmap.h>
#include <linux/delay.h>
#include <linux/backing-dev.h>
#include <trace/events/block.h>
#include "blk.h"
#include "blk-mq.h"
#include "blk-mq-tag.h"
#include "blk-mq-sched.h"
#include "bfq-iosched.h"
#include "blk-wbt.h"
#define BFQ_BFQQ_FNS(name) \
void bfq_mark_bfqq_##name(struct bfq_queue *bfqq) \
{ \
__set_bit(BFQQF_##name, &(bfqq)->flags); \
} \
void bfq_clear_bfqq_##name(struct bfq_queue *bfqq) \
{ \
__clear_bit(BFQQF_##name, &(bfqq)->flags); \
} \
int bfq_bfqq_##name(const struct bfq_queue *bfqq) \
{ \
return test_bit(BFQQF_##name, &(bfqq)->flags); \
#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 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) > \
(!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 */
/* Min observation time interval required to perform a peak-rate update (ns) */
/* Target observation time interval for a peak-rate update (ns) */
* 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
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) icq_to_bic((rq)->elv.priv[0])
#define RQ_BFQQ(rq) ((rq)->elv.priv[1])
struct bfq_queue *bic_to_bfqq(struct bfq_io_cq *bic, bool is_sync)
return bic->bfqq[is_sync];
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)
* 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.
bic->bfqq[is_sync] = bfqq;
if (bfqq && bic->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.
bic->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.
* @bfqd: the lookup key.
* @ioc: the io_context of the process doing I/O.
* @q: the request queue.
static struct bfq_io_cq *bfq_bic_lookup(struct bfq_data *bfqd,
struct io_context *ioc,
struct request_queue *q)
if (ioc) {
unsigned long flags;
struct bfq_io_cq *icq;
spin_lock_irqsave(&q->queue_lock, flags);
icq = icq_to_bic(ioc_lookup_icq(ioc, q));
spin_unlock_irqrestore(&q->queue_lock, flags);
return icq;
return NULL;
* Scheduler run of queue, if there are requests pending and no one in the
* driver that will restart queueing.
void bfq_schedule_dispatch(struct bfq_data *bfqd)
if (bfqd->queued != 0) {
bfq_log(bfqd, "schedule dispatch");
blk_mq_run_hw_queues(bfqd->queue, true);
#define bfq_class_idle(bfqq) ((bfqq)->ioprio_class == IOPRIO_CLASS_IDLE)
#define bfq_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;
wrap |= BFQ_RQ1_WRAP;
if (s2 >= last)
d2 = s2 - last;
else if (s2 + back_max >= last)
d2 = (last - s2) * bfqd->bfq_back_penalty;
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;
return rq2;
case BFQ_RQ2_WRAP:
return rq1;
case BFQ_RQ1_WRAP:
return rq2;
case BFQ_RQ1_WRAP|BFQ_RQ2_WRAP: /* both rqs wrapped */
* 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;
return rq2;
* Async I/O can easily starve sync I/O (both sync reads and sync
* writes), by consuming all tags. Similarly, storms of sync writes,
* such as those that sync(2) may trigger, can starve sync reads.
* Limit depths of async I/O and sync writes so as to counter both
* problems.
static void bfq_limit_depth(unsigned int op, struct blk_mq_alloc_data *data)
struct bfq_data *bfqd = data->q->elevator->elevator_data;
if (op_is_sync(op) && !op_is_write(op))
data->shallow_depth =
bfq_log(bfqd, "[%s] wr_busy %d sync %d depth %u",
__func__, bfqd->wr_busy_queues, op_is_sync(op),
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;
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 +
* 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)
* 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))
if (bfq_class_idle(bfqq))
if (!bfqq->next_rq)
bfqq->pos_root = &bfq_bfqq_to_bfqg(bfqq)->rq_pos_tree;
__bfqq = bfq_rq_pos_tree_lookup(bfqd, bfqq->pos_root,
blk_rq_pos(bfqq->next_rq), &parent, &p);
if (!__bfqq) {
rb_link_node(&bfqq->pos_node, parent, p);
rb_insert_color(&bfqq->pos_node, bfqq->pos_root);
} else
bfqq->pos_root = NULL;
* The following function returns false either if every active queue
* must receive the same share of the throughput (symmetric scenario),
* or, as a special case, if bfqq must receive a share of the
* throughput lower than or equal to the share that every other active
* queue must receive. If bfqq does sync I/O, then these are the only
* two cases where bfqq happens to be guaranteed its share of the
* throughput even if I/O dispatching is not plugged when bfqq remains
* temporarily empty (for more details, see the comments in the
* function bfq_better_to_idle()). For this reason, the return value
* of this function is used to check whether I/O-dispatch plugging can
* be avoided.
* The above first case (symmetric scenario) occurs when:
* 1) all active queues have the same weight,
* 2) all active queues belong to the same I/O-priority class,
* 3) all active groups at the same level in the groups tree have the same
* weight,
* 4) all active groups at the same level in the groups tree have the same
* number of children.
* Unfortunately, keeping the necessary state for evaluating exactly
* the last two symmetry sub-conditions above would be quite complex
* and time consuming. Therefore this function evaluates, instead,
* only the following stronger three sub-conditions, for which it is
* much easier to maintain the needed state:
* 1) all active queues have the same weight,
* 2) all active queues belong to the same I/O-priority class,
* 3) there are no active groups.
* In particular, the last condition is always true if hierarchical
* support or the cgroups interface are not enabled, thus no state
* needs to be maintained in this case.
static bool bfq_asymmetric_scenario(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
bool smallest_weight = bfqq &&
bfqq->weight_counter &&
bfqq->weight_counter ==
struct bfq_weight_counter,
* 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 ||
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
|| bfqd->num_groups_with_pending_reqs > 0
* If the weight-counter tree passed as input contains no counter for
* the weight of the input queue, then add that counter; otherwise just
* increment the existing counter.
* Note that weight-counter trees contain few nodes in mostly symmetric
* scenarios. For example, if all queues have the same weight, then the
* weight-counter tree for the queues may contain at most one node.
* This holds even if low_latency is on, because weight-raised queues
* are not inserted in the tree.
* In most scenarios, the rate at which nodes are created/destroyed
* should be low too.
void bfq_weights_tree_add(struct bfq_data *bfqd, struct bfq_queue *bfqq,
struct rb_root_cached *root)
struct bfq_entity *entity = &bfqq->entity;
struct rb_node **new = &(root->rb_root.rb_node), *parent = NULL;
bool leftmost = true;
* Do not insert if the queue is already associated with a
* counter, which happens if:
* 1) a request arrival has caused the queue to become both
* non-weight-raised, and hence change its weight, and
* backlogged; in this respect, each of the two events
* causes an invocation of this function,
* 2) this is the invocation of this function caused by the
* second event. This second invocation is actually useless,
* and we handle this fact by exiting immediately. More
* efficient or clearer solutions might possibly be adopted.
if (bfqq->weight_counter)
while (*new) {
struct bfq_weight_counter *__counter = container_of(*new,
struct bfq_weight_counter,
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),
* 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))
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,
* Decrement the weight counter associated with the queue, and, if the
* counter reaches 0, remove the counter from the tree.
* See the comments to the function bfq_weights_tree_add() for considerations
* about overhead.
void __bfq_weights_tree_remove(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
struct rb_root_cached *root)
if (!bfqq->weight_counter)
if (bfqq->weight_counter->num_active > 0)
goto reset_entity_pointer;
rb_erase_cached(&bfqq->weight_counter->weights_node, root);
bfqq->weight_counter = NULL;
* Invoke __bfq_weights_tree_remove on bfqq and decrement the number
* of active groups for each queue's inactive parent entity.
void bfq_weights_tree_remove(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
struct bfq_entity *entity = bfqq->entity.parent;
for_each_entity(entity) {
struct bfq_sched_data *sd = entity->my_sched_data;
if (sd->next_in_service || sd->in_service_entity) {
* entity is still active, because either
* next_in_service or in_service_entity is not
* NULL (see the comments on the definition of
* next_in_service for details on why
* in_service_entity must be checked too).
* As a consequence, its parent entities are
* active as well, and thus this loop must
* stop here.
* The decrement of num_groups_with_pending_reqs is
* not performed immediately upon the deactivation of
* entity, but it is delayed to when it also happens
* that the first leaf descendant bfqq of entity gets
* all its pending requests completed. The following
* instructions perform this delayed decrement, if
* needed. See the comments on
* num_groups_with_pending_reqs for details.
if (entity->in_groups_with_pending_reqs) {
entity->in_groups_with_pending_reqs = false;
* Next function is invoked last, because it causes bfqq to be
* freed if the following holds: bfqq is not in service and
* has no dispatched request. DO NOT use bfqq after the next
* function invocation.
__bfq_weights_tree_remove(bfqd, bfqq,
* 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;
rq = rq_entry_fifo(bfqq->;
if (rq == last || ktime_get_ns() < rq->fifo_time)
return NULL;
bfq_log_bfqq(bfqq->bfqd, bfqq, "check_fifo: returned %p", rq);
return rq;
static struct request *bfq_find_next_rq(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
struct request *last)
struct rb_node *rbnext = rb_next(&last->rb_node);
struct rb_node *rbprev = rb_prev(&last->rb_node);
struct request *next, *prev = NULL;
/* Follow expired path, else get first next available. */
next = bfq_check_fifo(bfqq, last);
if (next)
return next;
if (rbprev)
prev = rb_entry_rq(rbprev);
if (rbnext)
next = rb_entry_rq(rbnext);
else {
rbnext = rb_first(&bfqq->sort_list);
if (rbnext && rbnext != &last->rb_node)
next = rb_entry_rq(rbnext);
return bfq_choose_req(bfqd, next, prev, blk_rq_pos(last));
/* see the definition of bfq_async_charge_factor for details */
static unsigned long bfq_serv_to_charge(struct request *rq,
struct bfq_queue *bfqq)
if (bfq_bfqq_sync(bfqq) || bfqq->wr_coeff > 1 ||
bfq_asymmetric_scenario(bfqq->bfqd, bfqq))
return blk_rq_sectors(rq);
return blk_rq_sectors(rq) * bfq_async_charge_factor;
* bfq_updated_next_req - update the queue after a new next_rq selection.
* @bfqd: the device data the queue belongs to.
* @bfqq: the queue to update.
* If the first request of a queue changes we make sure that the queue
* has enough budget to serve at least its first request (if the
* request has grown). We do this because if the queue has not enough
* budget for its first request, it has to go through two dispatch
* rounds to actually get it dispatched.
static void bfq_updated_next_req(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
struct bfq_entity *entity = &bfqq->entity;
struct request *next_rq = bfqq->next_rq;
unsigned long new_budget;
if (!next_rq)
if (bfqq == bfqd->in_service_queue)
* In order not to break guarantees, budgets cannot be
* changed after an entity has been selected.
new_budget = max_t(unsigned long,
max_t(unsigned long, bfqq->max_budget,
bfq_serv_to_charge(next_rq, bfqq)),
if (entity->budget != new_budget) {
entity->budget = new_budget;
bfq_log_bfqq(bfqd, bfqq, "updated next rq: new budget %lu",
bfq_requeue_bfqq(bfqd, bfqq, false);
static unsigned int bfq_wr_duration(struct bfq_data *bfqd)
u64 dur;
if (bfqd->bfq_wr_max_time > 0)
return bfqd->bfq_wr_max_time;
dur = bfqd->rate_dur_prod;
do_div(dur, bfqd->peak_rate);
* Limit duration between 3 and 25 seconds. The upper limit
* has been conservatively set after the following worst case:
* on a QEMU/KVM virtual machine
* - running in a slow PC
* - with a virtual disk stacked on a slow low-end 5400rpm HDD
* - serving a heavy I/O workload, such as the sequential reading
* of several files
* mplayer took 23 seconds to start, if constantly weight-raised.
* As for higher values than that accommodating the above bad
* scenario, tests show that higher values would often yield
* the opposite of the desired result, i.e., would worsen
* responsiveness by allowing non-interactive applications to
* preserve weight raising for too long.
* On the other end, lower values than 3 seconds make it
* difficult for most interactive tasks to complete their jobs
* before weight-raising finishes.
return clamp_val(dur, msecs_to_jiffies(3000), msecs_to_jiffies(25000));
/* switch back from soft real-time to interactive weight raising */
static void switch_back_to_interactive_wr(struct bfq_queue *bfqq,
struct bfq_data *bfqd)
bfqq->wr_coeff = bfqd->bfq_wr_coeff;
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd);
bfqq->last_wr_start_finish = bfqq->wr_start_at_switch_to_srt;
static void
bfq_bfqq_resume_state(struct bfq_queue *bfqq, struct bfq_data *bfqd,
struct bfq_io_cq *bic, bool bfq_already_existing)
unsigned int old_wr_coeff = 1;
bool busy = bfq_already_existing && bfq_bfqq_busy(bfqq);
if (bic->saved_has_short_ttime)
if (bic->saved_IO_bound)
bfqq->last_serv_time_ns = bic->saved_last_serv_time_ns;
bfqq->inject_limit = bic->saved_inject_limit;
bfqq->decrease_time_jif = bic->saved_decrease_time_jif;
bfqq->entity.new_weight = bic->saved_weight;
bfqq->ttime = bic->saved_ttime;
bfqq->io_start_time = bic->saved_io_start_time;
bfqq->tot_idle_time = bic->saved_tot_idle_time;
* Restore weight coefficient only if low_latency is on
if (bfqd->low_latency) {
old_wr_coeff = bfqq->wr_coeff;
bfqq->wr_coeff = bic->saved_wr_coeff;
bfqq->service_from_wr = bic->saved_service_from_wr;
bfqq->wr_start_at_switch_to_srt = bic->saved_wr_start_at_switch_to_srt;
bfqq->last_wr_start_finish = bic->saved_last_wr_start_finish;
bfqq->wr_cur_max_time = bic->saved_wr_cur_max_time;
if (bfqq->wr_coeff > 1 && (bfq_bfqq_in_large_burst(bfqq) ||
time_is_before_jiffies(bfqq->last_wr_start_finish +
bfqq->wr_cur_max_time))) {
if (bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time &&
!bfq_bfqq_in_large_burst(bfqq) &&
time_is_after_eq_jiffies(bfqq->wr_start_at_switch_to_srt +
bfq_wr_duration(bfqd))) {
switch_back_to_interactive_wr(bfqq, bfqd);
} else {
bfqq->wr_coeff = 1;
bfq_log_bfqq(bfqq->bfqd, bfqq,
"resume state: switching off wr");
/* make sure weight will be updated, however we got here */
bfqq->entity.prio_changed = 1;
if (likely(!busy))
if (old_wr_coeff == 1 && bfqq->wr_coeff > 1)
else if (old_wr_coeff > 1 && bfqq->wr_coeff == 1)
static int bfqq_process_refs(struct bfq_queue *bfqq)
return bfqq->ref - bfqq->allocated - bfqq->entity.on_st_or_in_serv -
(bfqq->weight_counter != NULL) - bfqq->stable_ref;
/* Empty burst list and add just bfqq (see comments on bfq_handle_burst) */
static void bfq_reset_burst_list(struct bfq_data *bfqd, struct bfq_queue *bfqq)
struct bfq_queue *item;
struct hlist_node *n;
hlist_for_each_entry_safe(item, n, &bfqd->burst_list, burst_list_node)
* 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 */
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,
* 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,
} 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 +
* 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) {
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);
* 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;
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;
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,
* 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));
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 =
bfqq->wr_coeff = bfqd->bfq_wr_coeff *
bfqq->wr_cur_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,
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->wr_cur_max_time =
bfqq->wr_coeff = bfqd->bfq_wr_coeff *
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 &&
bfqq->budget_timeout +
* 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;
bfqq_weight = bfqq->entity.weight;
if (in_serv_bfqq->entity.parent)
in_serv_weight = in_serv_bfqq->entity.parent->weight;
in_serv_weight = in_serv_bfqq->entity.weight;
return bfqq_weight > in_serv_weight;
static bool bfq_better_to_idle(struct bfq_queue *bfqq);
static void bfq_bfqq_handle_idle_busy_switch(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
int old_wr_coeff,
struct request *rq,
bool *interactive)
bool soft_rt, in_burst, wr_or_deserves_wr,
idle_for_long_time = bfq_bfqq_idle_for_long_time(bfqd, bfqq),
* See the comments on
* bfq_bfqq_update_budg_for_activation for
* details on the usage of the next variable.
arrived_in_time = ktime_get_ns() <=
bfqq->ttime.last_end_request +
bfqd->bfq_slice_idle * 3;
* bfqq deserves to be weight-raised if:
* - it is sync,
* - it does not belong to a large burst,
* - it has been idle for enough time or is soft real-time,
* - is linked to a bfq_io_cq (it is not shared in any sense),
* - has a default weight (otherwise we assume the user wanted
* to control its weight explicitly)
in_burst = bfq_bfqq_in_large_burst(bfqq);
soft_rt = bfqd->bfq_wr_max_softrt_rate > 0 &&
!in_burst &&
time_is_before_jiffies(bfqq->soft_rt_next_start) &&
bfqq->dispatched == 0 &&
bfqq->entity.new_weight == 40;
*interactive = !in_burst && idle_for_long_time &&
bfqq->entity.new_weight == 40;
* Merged bfq_queues are kept out of weight-raising
* (low-latency) mechanisms. The reason is that these queues
* are usually created for non-interactive and
* non-soft-real-time tasks. Yet this is not the case for
* stably-merged queues. These queues are merged just because
* they are created shortly after each other. So they may
* easily serve the I/O of an interactive or soft-real time
* application, if the application happens to spawn multiple
* processes. So let also stably-merged queued enjoy weight
* raising.
wr_or_deserves_wr = bfqd->low_latency &&
(bfqq->wr_coeff > 1 ||
(bfq_bfqq_sync(bfqq) &&
(bfqq->bic || RQ_BIC(rq)->stably_merged) &&
(*interactive || soft_rt)));
* Using the last flag, update budget and check whether bfqq
* may want to preempt the in-service queue.
bfqq_wants_to_preempt =
bfq_bfqq_update_budg_for_activation(bfqd, bfqq,
* 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 &&
bfqq->budget_timeout +
msecs_to_jiffies(10000))) {
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,
if (old_wr_coeff != bfqq->wr_coeff)
bfqq->entity.prio_changed = 1;
bfqq->last_idle_bklogged = jiffies;
bfqq->service_from_backlogged = 0;
bfq_add_bfqq_busy(bfqd, bfqq);
* Expire in-service queue if preemption may be needed for
* guarantees or throughput. As for guarantees, we care
* explicitly about two cases. The first is that bfqq has to
* recover a service hole, as explained in the comments on
* bfq_bfqq_update_budg_for_activation(), i.e., that
* bfqq_wants_to_preempt is true. However, if bfqq does not
* carry time-critical I/O, then bfqq's bandwidth is less
* important than that of queues that carry time-critical I/O.
* So, as a further constraint, we consider this case only if
* bfqq is at least as weight-raised, i.e., at least as time
* critical, as the in-service queue.
* The second case is that bfqq is in a higher priority class,
* or has a higher weight than the in-service queue. If this
* condition does not hold, we don't care because, even if
* bfqq does not start to be served immediately, the resulting
* delay for bfqq's I/O is however lower or much lower than
* the ideal completion time to be guaranteed to bfqq's I/O.
* In both cases, preemption is needed only if, according to
* the timestamps of both bfqq and of the in-service queue,
* bfqq actually is the next queue to serve. So, to reduce
* useless preemptions, the return value of
* next_queue_may_preempt() is considered in the next compound
* condition too. Yet next_queue_may_preempt() just checks a
* simple, necessary condition for bfqq to be the next queue
* to serve. In fact, to evaluate a sufficient condition, the
* timestamps of the in-service queue would need to be
* updated, and this operation is quite costly (see the
* comments on bfq_bfqq_update_budg_for_activation()).
* As for throughput, we ask bfq_better_to_idle() whether we
* still need to plug I/O dispatching. If bfq_better_to_idle()
* says no, then plugging is not needed any longer, either to
* boost throughput or to perserve service guarantees. Then
* the best option is to stop plugging I/O, as not doing so
* would certainly lower throughput. We may end up in this
* case if: (1) upon a dispatch attempt, we detected that it
* was better to plug I/O dispatch, and to wait for a new
* request to arrive for the currently in-service queue, but
* (2) this switch of bfqq to busy changes the scenario.
if (bfqd->in_service_queue &&
((bfqq_wants_to_preempt &&
bfqq->wr_coeff >= bfqd->in_service_queue->wr_coeff) ||
bfq_bfqq_higher_class_or_weight(bfqq, bfqd->in_service_queue) ||
!bfq_better_to_idle(bfqd->in_service_queue)) &&
bfq_bfqq_expire(bfqd, bfqd->in_service_queue,
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;
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)))
* 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)
* Keep an observation window of at most 200 ms in the past
* from now.
if (tot_io_time > 200 * NSEC_PER_MSEC) {
bfqq->io_start_time = now_ns - (tot_io_time>>1);
bfqq->tot_idle_time >>= 1;
* Detect whether bfqq's I/O seems synchronized with that of some
* other queue, i.e., whether bfqq, after remaining empty, happens to
* receive new I/O only right after some I/O request of the other
* queue has been completed. We call waker queue the other queue, and
* we assume, for simplicity, that bfqq may have at most one waker
* queue.
* A remarkable throughput boost can be reached by unconditionally
* injecting the I/O of the waker queue, every time a new
* bfq_dispatch_request happens to be invoked while I/O is being
* plugged for bfqq. In addition to boosting throughput, this
* unblocks bfqq's I/O, thereby improving bandwidth and latency for
* bfqq. Note that these same results may be achieved with the general
* injection mechanism, but less effectively. For details on this
* aspect, see the comments on the choice of the queue for injection
* in bfq_select_queue().
* Turning back to the detection of a waker queue, a queue Q is deemed
* as a waker queue for bfqq if, for three consecutive times, bfqq
* happens to become non empty right after a request of Q has been
* completed. In this respect, even if bfqq is empty, we do not check
* for a waker if it still has some in-flight I/O. In fact, in this
* case bfqq is actually still being served by the drive, and may
* receive new I/O on the completion of some of the in-flight
* requests. In particular, on the first time, Q is tentatively set as
* a candidate waker queue, while on the third consecutive time that Q
* is detected, the field waker_bfqq is set to Q, to confirm that Q is
* a waker queue for bfqq. These detection steps are performed only if
* bfqq has a long think time, so as to make it more likely that
* bfqq's I/O is actually being blocked by a synchronization. This
* last filter, plus the above three-times requirement, make false
* positives less likely.
* 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.
* On queue merging all waker information is lost.
static void bfq_check_waker(struct bfq_data *bfqd, struct bfq_queue *bfqq,
u64 now_ns)
if (!bfqd->last_completed_rq_bfqq ||
bfqd->last_completed_rq_bfqq == bfqq ||
bfq_bfqq_has_short_ttime(bfqq) ||
bfqq->dispatched > 0 ||
now_ns - bfqd->last_completion >= 4 * NSEC_PER_MSEC ||
bfqd->last_completed_rq_bfqq == bfqq->waker_bfqq)
if (bfqd->last_completed_rq_bfqq !=
bfqq->tentative_waker_bfqq) {
* First synchronization detected with a
* candidate waker queue, or with a different
* candidate waker queue from the current one.
bfqq->tentative_waker_bfqq =
bfqq->num_waker_detections = 1;
} else /* Same tentative waker queue detected again */
if (bfqq->num_waker_detections == 3) {
bfqq->waker_bfqq = bfqd->last_completed_rq_bfqq;
bfqq->tentative_waker_bfqq = NULL;
* If the waker queue disappears, then
* bfqq->waker_bfqq must be reset. To
* this goal, we maintain in each
* waker queue a list, woken_list, of
* all the queues that reference the
* waker queue through their
* waker_bfqq pointer. When the waker
* queue exits, the waker_bfqq pointer
* of all the queues in the woken_list
* is reset.
* In addition, if bfqq is already in
* the woken_list of a waker queue,
* then, before being inserted into
* the woken_list of a new waker
* queue, bfqq must be removed from
* the woken_list of the old waker
* queue.
if (!hlist_unhashed(&bfqq->woken_list_node))
static void bfq_add_request(struct request *rq)
struct bfq_queue *bfqq = RQ_BFQQ(rq);
struct bfq_data *bfqd = bfqq->bfqd;
struct request *next_rq, *prev;
unsigned int old_wr_coeff = bfqq->wr_coeff;
bool interactive = false;
u64 now_ns = ktime_get_ns();
bfq_log_bfqq(bfqd, bfqq, "add_request %d", rq_is_sync(rq));
if (RB_EMPTY_ROOT(&bfqq->sort_list) && bfq_bfqq_sync(bfqq)) {
bfq_check_waker(bfqd, bfqq, now_ns);
* Periodically reset inject limit, to make sure that
* the latter eventually drops in case workload
* changes, see step (3) in the comments on
* bfq_update_inject_limit().
if (time_is_before_eq_jiffies(bfqq->decrease_time_jif +
bfq_reset_inject_limit(bfqd, bfqq);
* The following conditions must hold to setup a new
* sampling of total service time, and then a new
* update of the inject limit:
* - bfqq is in service, because the total service
* time is evaluated only for the I/O requests of
* the queues in service;
* - this is the right occasion to compute or to
* lower the baseline total service time, because
* there are actually no requests in the drive,
* or
* the baseline total service time is available, and
* this is the right occasion to compute the other
* quantity needed to update the inject limit, i.e.,
* the total service time caused by the amount of
* injection allowed by the current value of the
* limit. It is the right occasion because injection
* has actually been performed during the service
* hole, and there are still in-flight requests,
* which are very likely to be exactly the injected
* requests, or part of them;
* - the minimum interval for sampling the total
* service time and updating the inject limit has
* elapsed.
if (bfqq == bfqd->in_service_queue &&
(bfqd->rq_in_driver == 0 ||
(bfqq->last_serv_time_ns > 0 &&
bfqd->rqs_injected && bfqd->rq_in_driver > 0)) &&
time_is_before_eq_jiffies(bfqq->decrease_time_jif +
msecs_to_jiffies(10))) {
bfqd->last_empty_occupied_ns = ktime_get_ns();
* Start the state machine for measuring the
* total service time of rq: setting
* wait_dispatch will cause bfqd->waited_rq to
* be set when rq will be dispatched.
bfqd->wait_dispatch = true;
* If there is no I/O in service in the drive,
* then possible injection occurred before the
* arrival of rq will not affect the total
* service time of rq. So the injection limit
* must not be updated as a function of such
* total service time, unless new injection
* occurs before rq is completed. To have the
* injection limit updated only in the latter
* case, reset rqs_injected here (rqs_injected
* will be set in case injection is performed
* on bfqq before rq is completed).
if (bfqd->rq_in_driver == 0)
bfqd->rqs_injected = false;
if (bfq_bfqq_sync(bfqq))
bfq_update_io_intensity(bfqq, now_ns);
elv_rb_add(&bfqq->sort_list, rq);
* Check if this request is a better next-serve candidate.
prev = bfqq->next_rq;
next_rq = bfq_choose_req(bfqd, bfqq->next_rq, rq, bfqd->last_position);
bfqq->next_rq = next_rq;
* Adjust priority tree position, if next_rq changes.
* See comments on bfq_pos_tree_add_move() for the unlikely().
if (unlikely(!bfqd->nonrot_with_queueing && prev != bfqq->next_rq))
bfq_pos_tree_add_move(bfqd, bfqq);
if (!bfq_bfqq_busy(bfqq)) /* switching to busy ... */
bfq_bfqq_handle_idle_busy_switch(bfqd, bfqq, old_wr_coeff,
rq, &interactive);
else {
if (bfqd->low_latency && old_wr_coeff == 1 && !rq_is_sync(rq) &&
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);
bfqq->entity.prio_changed = 1;
if (prev != bfqq->next_rq)
bfq_updated_next_req(bfqd, bfqq);
* Assign jiffies to last_wr_start_finish in the following
* cases:
* . if bfqq is not going to be weight-raised, because, for
* non weight-raised queues, last_wr_start_finish stores the
* arrival time of the last request; as of now, this piece
* of information is used only for deciding whether to
* weight-raise async queues
* . if bfqq is not weight-raised, because, if bfqq is now
* switching to weight-raised, then last_wr_start_finish
* stores the time when weight-raising starts
* . if bfqq is interactive, because, regardless of whether
* bfqq is currently weight-raised, the weight-raising
* period must start or restart (this case is considered
* separately because it is not detected by the above
* conditions, if bfqq is already weight-raised)
* last_wr_start_finish has to be updated also if bfqq is soft
* real-time, because the weight-raising period is constantly
* restarted on idle-to-busy transitions for these queues, but
* this is already done in bfq_bfqq_handle_idle_busy_switch if
* needed.
if (bfqd->low_latency &&
(old_wr_coeff == 1 || bfqq->wr_coeff == 1 || interactive))
bfqq->last_wr_start_finish = jiffies;
static struct request *bfq_find_rq_fmerge(struct bfq_data *bfqd,
struct bio *bio,
struct request_queue *q)
struct bfq_queue *bfqq = bfqd->bio_bfqq;
if (bfqq)
return elv_rb_find(&bfqq->sort_list, bio_end_sector(bio));
return NULL;
static sector_t get_sdist(sector_t last_pos, struct request *rq)
if (last_pos)
return abs(blk_rq_pos(rq) - last_pos);
return 0;
#if 0 /* Still not clear if we can do without next two functions */
static void bfq_activate_request(struct request_queue *q, struct request *rq)
struct bfq_data *bfqd = q->elevator->elevator_data;
static void bfq_deactivate_request(struct request_queue *q, struct request *rq)
struct bfq_data *bfqd = q->elevator->elevator_data;
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)
elv_rb_del(&bfqq->sort_list, rq);
elv_rqhash_del(q, rq);
if (q->last_merge == rq)
q->last_merge = NULL;
if (RB_EMPTY_ROOT(&bfqq->sort_list)) {
bfqq->next_rq = NULL;
if (bfq_bfqq_busy(bfqq) && bfqq != bfqd->in_service_queue) {
bfq_del_bfqq_busy(bfqd, bfqq, false);
* bfqq emptied. In normal operation, when
* bfqq is empty, bfqq->entity.service and
* bfqq->entity.budget must contain,
* respectively, the service received and the
* budget used last time bfqq emptied. These
* facts do not hold in this case, as at least
* this last removal occurred while bfqq is
* not in service. To avoid inconsistencies,
* reset both bfqq->entity.service and
* bfqq->entity.budget, if bfqq has still a
* process that may issue I/O requests to it.
bfqq->entity.budget = bfqq->entity.service = 0;
* Remove queue from request-position tree as it is empty.
if (bfqq->pos_root) {
rb_erase(&bfqq->pos_node, bfqq->pos_root);
bfqq->pos_root = NULL;
} else {
/* see comments on bfq_pos_tree_add_move() for the unlikely() */
if (unlikely(!bfqd->nonrot_with_queueing))
bfq_pos_tree_add_move(bfqd, bfqq);
if (rq->cmd_flags & REQ_META)
static bool bfq_bio_merge(struct request_queue *q, struct bio *bio,
unsigned int nr_segs)
struct bfq_data *bfqd = q->elevator->elevator_data;
struct request *free = NULL;
* bfq_bic_lookup grabs the queue_lock: invoke it now and
* store its return value for later use, to avoid nesting
* queue_lock inside the bfqd->lock. We assume that the bic
* returned by bfq_bic_lookup does not go away before
* bfqd->lock is taken.
struct bfq_io_cq *bic = bfq_bic_lookup(bfqd, current->io_context, q);
bool ret;
if (bic)
bfqd->bio_bfqq = bic_to_bfqq(bic, op_is_sync(bio->bi_opf));
bfqd->bio_bfqq = NULL;
bfqd->bio_bic = bic;
ret = blk_mq_sched_try_merge(q, bio, nr_segs, &free);
if (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))
static struct bfq_queue *bfq_init_rq(struct request *rq);
static void bfq_request_merged(struct request_queue *q, struct request *req,
enum elv_merge type)
rb_prev(&req->rb_node) &&
blk_rq_pos(req) <
struct request, rb_node))) {
struct bfq_queue *bfqq = bfq_init_rq(req);
struct bfq_data *bfqd;
struct request *prev, *next_rq;
if (!bfqq)
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,
bfqq->next_rq = next_rq;
* If next_rq changes, update both the queue's budget to
* fit the new request and the queue's position in its
* rq_pos_tree.
if (prev != bfqq->next_rq) {
bfq_updated_next_req(bfqd, bfqq);
* See comments on bfq_pos_tree_add_move() for
* the unlikely().
if (unlikely(!bfqd->nonrot_with_queueing))
bfq_pos_tree_add_move(bfqd, bfqq);
* This function is called to notify the scheduler that the requests
* rq and 'next' have been merged, with 'next' going away. BFQ
* exploits this hook to address the following issue: if 'next' has a
* fifo_time lower that rq, then the fifo_time of rq must be set to
* the value of 'next', to not forget the greater age of 'next'.
* NOTE: in this function we assume that rq is in a bfq_queue, basing
* on that rq is picked from the hash table q->elevator->hash, which,
* in its turn, is filled only with I/O requests present in
* bfq_queues, while BFQ is in use for the request queue q. In fact,
* the function that fills this hash table (elv_rqhash_add) is called
* only by bfq_insert_request.
static void bfq_requests_merged(struct request_queue *q, struct request *rq,
struct request *next)
struct bfq_queue *bfqq = bfq_init_rq(rq),
*next_bfqq = bfq_init_rq(next);
if (!bfqq)
goto remove;
* If next and rq belong to the same bfq_queue and next is older
* than rq, then reposition rq in the fifo (by substituting next
* with rq). Otherwise, if next and rq belong to different
* bfq_queues, never reposition rq: in fact, we would have to
* reposition it with respect to next's position in its own fifo,
* which would most certainly be too expensive with respect to
* the benefits.
if (bfqq == next_bfqq &&
!list_empty(&rq->queuelist) && !list_empty(&next->queuelist) &&
next->fifo_time < rq->fifo_time) {
list_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);
/* 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)
/* 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->soft_rt_next_start = jiffies;
if (bfq_bfqq_busy(bfqq))
bfqq->wr_coeff = 1;
bfqq->wr_cur_max_time = 0;
bfqq->last_wr_start_finish = jiffies;
* Trigger a weight change on the next invocation of
* __bfq_entity_update_weight_prio.
bfqq->entity.prio_changed = 1;
void bfq_end_wr_async_queues(struct bfq_data *bfqd,
struct bfq_group *bfqg)
int i, j;
for (i = 0; i < 2; i++)
for (j = 0; j < IOPRIO_NR_LEVELS; j++)
if (bfqg->async_bfqq[i][j])
if (bfqg->async_idle_bfqq)
static void bfq_end_wr(struct bfq_data *bfqd)
struct bfq_queue *bfqq;
list_for_each_entry(bfqq, &bfqd->active_list, bfqq_list)
list_for_each_entry(bfqq, &bfqd->idle_list, bfqq_list)
static sector_t bfq_io_struct_pos(void *io_struct, bool request)
if (request)
return blk_rq_pos(io_struct);
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) <=
static struct bfq_queue *bfqq_find_close(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
sector_t sector)
struct rb_root *root = &bfq_bfqq_to_bfqg(bfqq)->rq_pos_tree;
struct rb_node *parent, *node;
struct bfq_queue *__bfqq;
if (RB_EMPTY_ROOT(root))
return NULL;
* First, if we find a request starting at the end of the last
* request, choose it.
__bfqq = bfq_rq_pos_tree_lookup(bfqd, root, sector, &parent, NULL);
if (__bfqq)
return __bfqq;
* If the exact sector wasn't found, the parent of the NULL leaf
* will contain the closest sector (rq_pos_tree sorted by
* next_request position).
__bfqq = rb_entry(parent, struct bfq_queue, pos_node);
if (bfq_rq_close_to_sector(__bfqq->next_rq, true, sector))
return __bfqq;
if (blk_rq_pos(__bfqq->next_rq) < sector)
node = rb_next(&__bfqq->pos_node);
node = rb_prev(&__bfqq->pos_node);
if (!node)
return NULL;
__bfqq = rb_entry(node, struct bfq_queue, pos_node);
if (bfq_rq_close_to_sector(__bfqq->next_rq, true, sector))
return __bfqq;
return NULL;
static struct bfq_queue *bfq_find_close_cooperator(struct bfq_data *bfqd,
struct bfq_queue *cur_bfqq,
sector_t sector)
struct bfq_queue *bfqq;
* We shall notice if some of the queues are cooperating,
* e.g., working closely on the same area of the device. In
* that case, we can group them together and: 1) don't waste
* time idling, and 2) serve the union of their requests in
* the best possible order for throughput.
bfqq = bfqq_find_close(bfqd, cur_bfqq, sector);
if (!bfqq || bfqq == cur_bfqq)
return NULL;
return bfqq;
static struct bfq_queue *
bfq_setup_merge(struct bfq_queue *bfqq, struct bfq_queue *new_bfqq)
int process_refs, new_process_refs;
struct bfq_queue *__bfqq;
* If there are no process references on the new_bfqq, then it is
* unsafe to follow the ->new_bfqq chain as other bfqq's in the chain
* may have dropped their last reference (not just their last process
* reference).
if (!bfqq_process_refs(new_bfqq))
return NULL;
/* Avoid a circular list and skip interim queue merges. */
while ((__bfqq = new_bfqq->new_bfqq)) {
if (__bfqq == bfqq)
return NULL;
new_bfqq = __bfqq;
process_refs = bfqq_process_refs(bfqq);
new_process_refs = bfqq_process_refs(new_bfqq);
* If the process for the bfqq has gone away, there is no
* sense in merging the queues.
if (process_refs == 0 || new_process_refs == 0)
return NULL;
bfq_log_bfqq(bfqq->bfqd, bfqq, "scheduling merge with queue %d",
* Merging is just a redirection: the requests of the process
* owning one of the two queues are redirected to the other queue.
* The latter queue, in its turn, is set as shared if this is the
* first time that the requests of some process are redirected to
* it.
* We redirect bfqq to new_bfqq and not the opposite, because
* we are in the context of the process owning bfqq, thus we
* have the io_cq of this process. So we can immediately
* configure this io_cq to redirect the requests of the
* process to new_bfqq. In contrast, the io_cq of new_bfqq is
* not available any more (new_bfqq->bic == NULL).
* Anyway, even in case new_bfqq coincides with the in-service
* queue, redirecting requests the in-service queue is the
* best option, as we feed the in-service queue with new
* requests close to the last request served and, by doing so,
* are likely to increase the throughput.
bfqq->new_bfqq = new_bfqq;
new_bfqq->ref += process_refs;
return new_bfqq;
static bool bfq_may_be_close_cooperator(struct bfq_queue *bfqq,
struct bfq_queue *new_bfqq)
if (bfq_too_late_for_merging(new_bfqq))
return false;
if (bfq_class_idle(bfqq) || bfq_class_idle(new_bfqq) ||
(bfqq->ioprio_class != new_bfqq->ioprio_class))
return false;
* If either of the queues has already been detected as seeky,
* then merging it with the other queue is unlikely to lead to
* sequential I/O.
if (BFQQ_SEEKY(bfqq) || BFQQ_SEEKY(new_bfqq))
return false;
* Interleaved I/O is known to be done by (some) applications
* only for reads, so it does not make sense to merge async
* queues.
if (!bfq_bfqq_sync(bfqq) || !bfq_bfqq_sync(new_bfqq))
return false;
return true;
static bool idling_boosts_thr_without_issues(struct bfq_data *bfqd,
struct bfq_queue *bfqq);
* Attempt to schedule a merge of bfqq with the currently in-service
* queue or with a close queue among the scheduled queues. Return
* NULL if no merge was scheduled, a pointer to the shared bfq_queue
* structure otherwise.
* The OOM queue is not allowed to participate to cooperation: in fact, since
* the requests temporarily redirected to the OOM queue could be redirected
* again to dedicated queues at any time, the state needed to correctly
* handle merging with the OOM queue would be quite complex and expensive
* to maintain. Besides, in such a critical condition as an out of memory,
* the benefits of queue merging may be little relevant, or even negligible.
* WARNING: queue merging may impair fairness among non-weight raised
* queues, for at least two reasons: 1) the original weight of a
* merged queue may change during the merged state, 2) even being the
* weight the same, a merged queue may be bloated with many more
* requests than the ones produced by its originally-associated
* process.
static struct bfq_queue *
bfq_setup_cooperator(struct bfq_data *bfqd, struct bfq_queue *bfqq,
void *io_struct, bool request, struct bfq_io_cq *bic)
struct bfq_queue *in_service_bfqq, *new_bfqq;
* Check delayed stable merge for rotational or non-queueing
* devs. For this branch to be executed, bfqq must not be
* currently merged with some other queue (i.e., bfqq->bic
* must be non null). If we considered also merged queues,
* then we should also check whether bfqq has already been
* merged with bic->stable_merge_bfqq. But this would be
* costly and complicated.
if (unlikely(!bfqd->nonrot_with_queueing)) {
* Make sure also that bfqq is sync, because
* bic->stable_merge_bfqq may point to some queue (for
* stable merging) also if bic is associated with a
* sync queue, but this bfqq is async
if (bfq_bfqq_sync(bfqq) && bic->stable_merge_bfqq &&
!bfq_bfqq_just_created(bfqq) &&
time_is_before_jiffies(bfqq->split_time +
msecs_to_jiffies(bfq_late_stable_merging)) &&
time_is_before_jiffies(bfqq->creation_time +
msecs_to_jiffies(bfq_late_stable_merging))) {
struct bfq_queue *stable_merge_bfqq =
int proc_ref = min(bfqq_process_refs(bfqq),
/* deschedule stable merge, because done or aborted here */
bic->stable_merge_bfqq = NULL;
if (!idling_boosts_thr_without_issues(bfqd, bfqq) &&
proc_ref > 0) {
/* next function will take at least one ref */
struct bfq_queue *new_bfqq =
bfq_setup_merge(bfqq, stable_merge_bfqq);
bic->stably_merged = true;
if (new_bfqq && new_bfqq->bic)
new_bfqq->bic->stably_merged = true;
return new_bfqq;
} else
return NULL;
* Do not perform queue merging if the device is non
* rotational and performs internal queueing. In fact, such a
* device reaches a high speed through internal parallelism
* and pipelining. This means that, to reach a high
* throughput, it must have many requests enqueued at the same
* time. But, in this configuration, the internal scheduling
* algorithm of the device does exactly the job of queue
* merging: it reorders requests so as to obtain as much as
* possible a sequential I/O pattern. As a consequence, with
* the workload generated by processes doing interleaved I/O,
* the throughput reached by the device is likely to be the
* same, with and without queue merging.
* Disabling merging also provides a remarkable benefit in
* terms of throughput. Merging tends to make many workloads
* artificially more uneven, because of shared queues
* remaining non empty for incomparably more time than
* non-merged queues. This may accentuate workload
* asymmetries. For example, if one of the queues in a set of
* merged queues has a higher weight than a normal queue, then
* the shared queue may inherit such a high weight and, by
* staying almost always active, may force BFQ to perform I/O
* plugging most of the time. This evidently makes it harder
* for BFQ to let the device reach a high throughput.
* Finally, the likely() macro below is not used because one
* of the two branches is more likely than the other, but to
* have the code path after the following if() executed as
* fast as possible for the case of a non rotational device
* with queueing. We want it because this is the fastest kind
* of device. On the opposite end, the likely() may lengthen
* the execution time of BFQ for the case of slower devices
* (rotational or at least without queueing). But in this case
* the execution time of BFQ matters very little, if not at
* all.
if (likely(bfqd->nonrot_with_queueing))
return NULL;
* Prevent bfqq from being merged if it has been created too
* long ago. The idea is that true cooperating processes, and
* thus their associated bfq_queues, are supposed to be
* created shortly after each other. This is the case, e.g.,
* for KVM/QEMU and dump I/O threads. Basing on this
* assumption, the following filtering greatly reduces the
* probability that two non-cooperating processes, which just
* happen to do close I/O for some short time interval, have
* their queues merged by mistake.
if (bfq_too_late_for_merging(bfqq))
return NULL;
if (bfqq->new_bfqq)
return bfqq->new_bfqq;
if (!io_struct || unlikely(bfqq == &bfqd->oom_bfqq))
return NULL;
/* If there is only one backlogged queue, don't search. */
if (bfq_tot_busy_queues(bfqd) == 1)
return NULL;
in_service_bfqq = bfqd->in_service_queue;
if (in_service_bfqq && in_service_bfqq != bfqq &&
likely(in_service_bfqq != &bfqd->oom_bfqq) &&
bfq_rq_close_to_sector(io_struct, request,
bfqd->in_serv_last_pos) &&
bfqq->entity.parent == in_service_bfqq->entity.parent &&
bfq_may_be_close_cooperator(bfqq, in_service_bfqq)) {
new_bfqq = bfq_setup_merge(bfqq, in_service_bfqq);
if (new_bfqq)
return new_bfqq;
* Check whether there is a cooperator among currently scheduled
* queues. The only thing we need is that the bio/request is not
* NULL, as we need it to establish whether a cooperator exists.
new_bfqq = bfq_find_close_cooperator(bfqd, bfqq,
bfq_io_struct_pos(io_struct, request));
if (new_bfqq && likely(new_bfqq != &bfqd->oom_bfqq) &&
bfq_may_be_close_cooperator(bfqq, new_bfqq))
return bfq_setup_merge(bfqq, new_bfqq);
return NULL;
static void bfq_bfqq_save_state(struct bfq_queue *bfqq)
struct bfq_io_cq *bic = bfqq->bic;
* If !bfqq->bic, the queue is already shared or its requests
* have already been redirected to a shared queue; both idle window
* and weight raising state have already been saved. Do nothing.
if (!bic)
bic->saved_last_serv_time_ns = bfqq->last_serv_time_ns;
bic->saved_inject_limit = bfqq->inject_limit;
bic->saved_decrease_time_jif = bfqq->decrease_time_jif;
bic->saved_weight = bfqq->entity.orig_weight;
bic->saved_ttime = bfqq->ttime;
bic->saved_has_short_ttime = bfq_bfqq_has_short_ttime(bfqq);
bic->saved_IO_bound = bfq_bfqq_IO_bound(bfqq);
bic->saved_io_start_time = bfqq->io_start_time;
bic->saved_tot_idle_time = bfqq->tot_idle_time;
bic->saved_in_large_burst = bfq_bfqq_in_large_burst(bfqq);
bic->was_in_burst_list = !hlist_unhashed(&bfqq->burst_list_node);
if (unlikely(bfq_bfqq_just_created(bfqq) &&
!bfq_bfqq_in_large_burst(bfqq) &&
bfqq->bfqd->low_latency)) {
* bfqq being merged right after being created: bfqq
* would have deserved interactive weight raising, but
* did not make it to be set in a weight-raised state,
* because of this early merge. Store directly the
* weight-raising state that would have been assigned
* to bfqq, so that to avoid that bfqq unjustly fails
* to enjoy weight raising if split soon.
bic->saved_wr_coeff = bfqq->bfqd->bfq_wr_coeff;
bic->saved_wr_start_at_switch_to_srt = bfq_smallest_from_now();
bic->saved_wr_cur_max_time = bfq_wr_duration(bfqq->bfqd);
bic->saved_last_wr_start_finish = jiffies;
} else {
bic->saved_wr_coeff = bfqq->wr_coeff;
bic->saved_wr_start_at_switch_to_srt =
bic->saved_service_from_wr = bfqq->service_from_wr;
bic->saved_last_wr_start_finish = bfqq->last_wr_start_finish;
bic->saved_wr_cur_max_time = bfqq->wr_cur_max_time;
static void
bfq_reassign_last_bfqq(struct bfq_queue *cur_bfqq, struct bfq_queue *new_bfqq)
if (cur_bfqq->entity.parent &&
cur_bfqq->entity.parent->last_bfqq_created == cur_bfqq)
cur_bfqq->entity.parent->last_bfqq_created = new_bfqq;
else if (cur_bfqq->bfqd && cur_bfqq->bfqd->last_bfqq_created == cur_bfqq)
cur_bfqq->bfqd->last_bfqq_created = new_bfqq;
void bfq_release_process_ref(struct bfq_data *bfqd, struct bfq_queue *bfqq)
* To prevent bfqq's service guarantees from being violated,
* bfqq may be left busy, i.e., queued for service, even if
* empty (see comments in __bfq_bfqq_expire() for
* details). But, if no process will send requests to bfqq any
* longer, then there is no point in keeping bfqq queued for
* service. In addition, keeping bfqq queued for service, but
* with no process ref any longer, may have caused bfqq to be
* freed when dequeued from service. But this is assumed to
* never happen.
if (bfq_bfqq_busy(bfqq) && RB_EMPTY_ROOT(&bfqq->sort_list) &&
bfqq != bfqd->in_service_queue)
bfq_del_bfqq_busy(bfqd, bfqq, false);
bfq_reassign_last_bfqq(bfqq, NULL);
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 */
if (bfq_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.
* 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 =
if (bfq_bfqq_busy(new_bfqq))
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))
bfq_log_bfqq(bfqd, new_bfqq, "merge_bfqqs: wr_busy %d",
* Merge queues (that is, let bic redirect its requests to new_bfqq)
bic_set_bfqq(bic, new_bfqq, 1);
* 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,
* 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;
timeout_coeff = bfqq->entity.weight / bfqq->entity.orig_weight;
bfqd->last_budget_start = ktime_get();
bfqq->budget_timeout = jiffies +
bfqd->bfq_timeout * timeout_coeff;
static void __bfq_set_in_service_queue(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
if (bfqq) {
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 +=
jiffies - bfqq->budget_timeout;
bfqq->last_wr_start_finish = jiffies;
bfq_set_budget_timeout(bfqd, bfqq);
bfq_log_bfqq(bfqd, bfqq,
"set_in_service_queue, cur-budget = %d",
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;
* We don't want to idle for seeks, but we do want to allow
* fair distribution of slice time for a process doing back-to-back
* seeks. So allow a little bit of time for him to submit a new rq.
sl = bfqd->bfq_slice_idle;
* Unless the queue is being weight-raised or the scenario is
* asymmetric, grant only minimum idle time if the queue
* is seeky. A long idling is preserved for a weight-raised
* queue, or, more in general, in an asymmetric scenario,
* because a long idling is needed for guaranteeing to a queue
* its reserved share of the throughput (in particular, it is
* needed if the queue has a higher weight than some other
* queue).
if (BFQQ_SEEKY(bfqq) && bfqq->wr_coeff == 1 &&
!bfq_asymmetric_scenario(bfqd, bfqq))
sl = min_t(u64, sl, BFQ_MIN_TT);
else if (bfqq->wr_coeff > 1)
sl = max_t(u32, sl, 20ULL * NSEC_PER_MSEC);
bfqd->last_idling_start = ktime_get();
bfqd->last_idling_start_jiffies = jiffies;
hrtimer_start(&bfqd->idle_slice_timer, ns_to_ktime(sl),
* 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 *
* 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_log(bfqd, "new max_budget = %d", bfqd->bfq_max_budget);
static void bfq_reset_rate_computation(struct bfq_data *bfqd,
struct request *rq)
if (rq != NULL) { /* new rq dispatch now, reset accordingly */
bfqd->last_dispatch = bfqd->first_dispatch = ktime_get_ns();
bfqd->peak_rate_samples = 1;
bfqd->sequential_samples = 0;
bfqd->tot_sectors_dispatched = bfqd->last_rq_max_size =
} else /* no new rq dispatched, just reset the number of samples */
bfqd->peak_rate_samples = 0; /* full re-init on next disp. */
"reset_rate_computation at end, sample %u/%u tot_sects %llu",
bfqd->peak_rate_samples, bfqd->sequential_samples,
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,
* 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);
bfq_reset_rate_computation(bfqd, rq);
* Update the read/write peak rate (the main quantity used for
* auto-tuning, see update_thr_responsiveness_params()).
* It is not trivial to estimate the peak rate (correctly): because of
* the presence of sw and hw queues between the scheduler and the
* device components that finally serve I/O requests, it is hard to
* say exactly when a given dispatched request is served inside the
* device, and for how long. As a consequence, it is hard to know
* precisely at what rate a given set of requests is actually served
* by the device.
* On the opposite end, the dispatch time of any request is trivially
* available, and, from this piece of information, the "dispatch rate"
* of requests can be immediately computed. So, the idea in the next
* function is to use what is known, namely request dispatch times
* (plus, when useful, request completion times), to estimate what is
* unknown, namely in-device request service rate.
* The main issue is that, because of the above facts, the rate at
* which a certain set of requests is dispatched over a certain time
* interval can vary greatly with respect to the rate at which the
* same requests are then served. But, since the size of any
* intermediate queue is limited, and the service scheme is lossless
* (no request is silently dropped), the following obvious convergence
* property holds: the number of requests dispatched MUST become
* closer and closer to the number of requests completed as the
* observation interval grows. This is the key property used in
* the next function to estimate the peak service rate as a function
* of the observed dispatch rate. The function assumes to be invoked
* on every request dispatch.
static void bfq_update_peak_rate(struct bfq_data *bfqd, struct request *rq)
u64 now_ns = ktime_get_ns();
if (bfqd->peak_rate_samples == 0) { /* first dispatch