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linux/block/bfq-iosched.c

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// SPDX-License-Identifier: GPL-2.0-or-later
/*
* Budget Fair Queueing (BFQ) I/O scheduler.
*
* Based on ideas and code from CFQ:
* Copyright (C) 2003 Jens Axboe <axboe@kernel.dk>
*
* Copyright (C) 2008 Fabio Checconi <fabio@gandalf.sssup.it>
* Paolo Valente <paolo.valente@unimore.it>
*
* Copyright (C) 2010 Paolo Valente <paolo.valente@unimore.it>
* Arianna Avanzini <avanzini@google.com>
*
* Copyright (C) 2017 Paolo Valente <paolo.valente@linaro.org>
*
* BFQ is a proportional-share I/O scheduler, with some extra
* low-latency capabilities. BFQ also supports full hierarchical
* scheduling through cgroups. Next paragraphs provide an introduction
* on BFQ inner workings. Details on BFQ benefits, usage and
* limitations can be found in Documentation/block/bfq-iosched.rst.
*
* BFQ is a proportional-share storage-I/O scheduling algorithm based
* on the slice-by-slice service scheme of CFQ. But BFQ assigns
* budgets, measured in number of sectors, to processes instead of
* time slices. The device is not granted to the in-service process
* for a given time slice, but until it has exhausted its assigned
* budget. This change from the time to the service domain enables BFQ
* to distribute the device throughput among processes as desired,
* without any distortion due to throughput fluctuations, or to device
* internal queueing. BFQ uses an ad hoc internal scheduler, called
* B-WF2Q+, to schedule processes according to their budgets. More
* precisely, BFQ schedules queues associated with processes. Each
* process/queue is assigned a user-configurable weight, and B-WF2Q+
* guarantees that each queue receives a fraction of the throughput
* proportional to its weight. Thanks to the accurate policy of
* B-WF2Q+, BFQ can afford to assign high budgets to I/O-bound
* processes issuing sequential requests (to boost the throughput),
* and yet guarantee a low latency to interactive and soft real-time
* applications.
*
* In particular, to provide these low-latency guarantees, BFQ
* explicitly privileges the I/O of two classes of time-sensitive
* applications: interactive and soft real-time. In more detail, BFQ
* behaves this way if the low_latency parameter is set (default
* configuration). This feature enables BFQ to provide applications in
* these classes with a very low latency.
*
* To implement this feature, BFQ constantly tries to detect whether
* the I/O requests in a bfq_queue come from an interactive or a soft
* real-time application. For brevity, in these cases, the queue is
* said to be interactive or soft real-time. In both cases, BFQ
* privileges the service of the queue, over that of non-interactive
* and non-soft-real-time queues. This privileging is performed,
* mainly, by raising the weight of the queue. So, for brevity, we
* call just weight-raising periods the time periods during which a
* queue is privileged, because deemed interactive or soft real-time.
*
* The detection of soft real-time queues/applications is described in
* detail in the comments on the function
* bfq_bfqq_softrt_next_start. On the other hand, the detection of an
* interactive queue works as follows: a queue is deemed interactive
* if it is constantly non empty only for a limited time interval,
* after which it does become empty. The queue may be deemed
* interactive again (for a limited time), if it restarts being
* constantly non empty, provided that this happens only after the
* queue has remained empty for a given minimum idle time.
*
* By default, BFQ computes automatically the above maximum time
* interval, i.e., the time interval after which a constantly
* non-empty queue stops being deemed interactive. Since a queue is
* weight-raised while it is deemed interactive, this maximum time
* interval happens to coincide with the (maximum) duration of the
* weight-raising for interactive queues.
*
* Finally, BFQ also features additional heuristics for
* preserving both a low latency and a high throughput on NCQ-capable,
* rotational or flash-based devices, and to get the job done quickly
* for applications consisting in many I/O-bound processes.
*
* NOTE: if the main or only goal, with a given device, is to achieve
* the maximum-possible throughput at all times, then do switch off
* all low-latency heuristics for that device, by setting low_latency
* to 0.
*
* BFQ is described in [1], where also a reference to the initial,
* more theoretical paper on BFQ can be found. The interested reader
* can find in the latter paper full details on the main algorithm, as
* well as formulas of the guarantees and formal proofs of all the
* properties. With respect to the version of BFQ presented in these
* papers, this implementation adds a few more heuristics, such as the
* ones that guarantee a low latency to interactive and soft real-time
* applications, and a hierarchical extension based on H-WF2Q+.
*
* B-WF2Q+ is based on WF2Q+, which is described in [2], together with
* H-WF2Q+, while the augmented tree used here to implement B-WF2Q+
* with O(log N) complexity derives from the one introduced with EEVDF
* in [3].
*
* [1] P. Valente, A. Avanzini, "Evolution of the BFQ Storage I/O
* Scheduler", Proceedings of the First Workshop on Mobile System
* Technologies (MST-2015), May 2015.
* http://algogroup.unimore.it/people/paolo/disk_sched/mst-2015.pdf
*
* [2] Jon C.R. Bennett and H. Zhang, "Hierarchical Packet Fair Queueing
* Algorithms", IEEE/ACM Transactions on Networking, 5(5):675-689,
* Oct 1997.
*
* http://www.cs.cmu.edu/~hzhang/papers/TON-97-Oct.ps.gz
*
* [3] I. Stoica and H. Abdel-Wahab, "Earliest Eligible Virtual Deadline
* First: A Flexible and Accurate Mechanism for Proportional Share
* Resource Allocation", technical report.
*
* http://www.cs.berkeley.edu/~istoica/papers/eevdf-tr-95.pdf
*/
#include <linux/module.h>
#include <linux/slab.h>
#include <linux/blkdev.h>
#include <linux/cgroup.h>
#include <linux/elevator.h>
#include <linux/ktime.h>
#include <linux/rbtree.h>
#include <linux/ioprio.h>
#include <linux/sbitmap.h>
#include <linux/delay.h>
#include "blk.h"
#include "blk-mq.h"
#include "blk-mq-tag.h"
#include "blk-mq-sched.h"
#include "bfq-iosched.h"
#include "blk-wbt.h"
#define BFQ_BFQQ_FNS(name) \
void bfq_mark_bfqq_##name(struct bfq_queue *bfqq) \
{ \
__set_bit(BFQQF_##name, &(bfqq)->flags); \
} \
void bfq_clear_bfqq_##name(struct bfq_queue *bfqq) \
{ \
__clear_bit(BFQQF_##name, &(bfqq)->flags); \
} \
int bfq_bfqq_##name(const struct bfq_queue *bfqq) \
{ \
return test_bit(BFQQF_##name, &(bfqq)->flags); \
}
BFQ_BFQQ_FNS(just_created);
BFQ_BFQQ_FNS(busy);
BFQ_BFQQ_FNS(wait_request);
BFQ_BFQQ_FNS(non_blocking_wait_rq);
BFQ_BFQQ_FNS(fifo_expire);
BFQ_BFQQ_FNS(has_short_ttime);
BFQ_BFQQ_FNS(sync);
BFQ_BFQQ_FNS(IO_bound);
BFQ_BFQQ_FNS(in_large_burst);
BFQ_BFQQ_FNS(coop);
BFQ_BFQQ_FNS(split_coop);
BFQ_BFQQ_FNS(softrt_update);
BFQ_BFQQ_FNS(has_waker);
#undef BFQ_BFQQ_FNS \
/* Expiration time of sync (0) and async (1) requests, in ns. */
static const u64 bfq_fifo_expire[2] = { NSEC_PER_SEC / 4, NSEC_PER_SEC / 8 };
/* Maximum backwards seek (magic number lifted from CFQ), in KiB. */
static const int bfq_back_max = 16 * 1024;
/* Penalty of a backwards seek, in number of sectors. */
static const int bfq_back_penalty = 2;
/* Idling period duration, in ns. */
static u64 bfq_slice_idle = NSEC_PER_SEC / 125;
/* Minimum number of assigned budgets for which stats are safe to compute. */
static const int bfq_stats_min_budgets = 194;
/* Default maximum budget values, in sectors and number of requests. */
static const int bfq_default_max_budget = 16 * 1024;
/*
* When a sync request is dispatched, the queue that contains that
* request, and all the ancestor entities of that queue, are charged
* with the number of sectors of the request. In contrast, if the
* request is async, then the queue and its ancestor entities are
* charged with the number of sectors of the request, multiplied by
* the factor below. This throttles the bandwidth for async I/O,
* w.r.t. to sync I/O, and it is done to counter the tendency of async
* writes to steal I/O throughput to reads.
*
* The current value of this parameter is the result of a tuning with
* several hardware and software configurations. We tried to find the
* lowest value for which writes do not cause noticeable problems to
* reads. In fact, the lower this parameter, the stabler I/O control,
* in the following respect. The lower this parameter is, the less
* the bandwidth enjoyed by a group decreases
* - when the group does writes, w.r.t. to when it does reads;
* - when other groups do reads, w.r.t. to when they do writes.
*/
static const int bfq_async_charge_factor = 3;
/* Default timeout values, in jiffies, approximating CFQ defaults. */
const int bfq_timeout = HZ / 8;
/*
* Time limit for merging (see comments in bfq_setup_cooperator). Set
* to the slowest value that, in our tests, proved to be effective in
* removing false positives, while not causing true positives to miss
* queue merging.
*
* As can be deduced from the low time limit below, queue merging, if
* successful, happens at the very beginning of the I/O of the involved
* cooperating processes, as a consequence of the arrival of the very
* first requests from each cooperator. After that, there is very
* little chance to find cooperators.
*/
static const unsigned long bfq_merge_time_limit = HZ/10;
static struct kmem_cache *bfq_pool;
/* Below this threshold (in ns), we consider thinktime immediate. */
#define BFQ_MIN_TT (2 * NSEC_PER_MSEC)
/* hw_tag detection: parallel requests threshold and min samples needed. */
#define BFQ_HW_QUEUE_THRESHOLD 3
#define BFQ_HW_QUEUE_SAMPLES 32
#define BFQQ_SEEK_THR (sector_t)(8 * 100)
#define BFQQ_SECT_THR_NONROT (sector_t)(2 * 32)
#define BFQ_RQ_SEEKY(bfqd, last_pos, rq) \
(get_sdist(last_pos, rq) > \
BFQQ_SEEK_THR && \
(!blk_queue_nonrot(bfqd->queue) || \
blk_rq_sectors(rq) < BFQQ_SECT_THR_NONROT))
#define BFQQ_CLOSE_THR (sector_t)(8 * 1024)
#define BFQQ_SEEKY(bfqq) (hweight32(bfqq->seek_history) > 19)
/*
* Sync random I/O is likely to be confused with soft real-time I/O,
* because it is characterized by limited throughput and apparently
* isochronous arrival pattern. To avoid false positives, queues
* containing only random (seeky) I/O are prevented from being tagged
* as soft real-time.
*/
#define BFQQ_TOTALLY_SEEKY(bfqq) (bfqq->seek_history == -1)
/* Min number of samples required to perform peak-rate update */
#define BFQ_RATE_MIN_SAMPLES 32
/* Min observation time interval required to perform a peak-rate update (ns) */
#define BFQ_RATE_MIN_INTERVAL (300*NSEC_PER_MSEC)
/* Target observation time interval for a peak-rate update (ns) */
#define BFQ_RATE_REF_INTERVAL NSEC_PER_SEC
/*
* Shift used for peak-rate fixed precision calculations.
* With
* - the current shift: 16 positions
* - the current type used to store rate: u32
* - the current unit of measure for rate: [sectors/usec], or, more precisely,
* [(sectors/usec) / 2^BFQ_RATE_SHIFT] to take into account the shift,
* the range of rates that can be stored is
* [1 / 2^BFQ_RATE_SHIFT, 2^(32 - BFQ_RATE_SHIFT)] sectors/usec =
* [1 / 2^16, 2^16] sectors/usec = [15e-6, 65536] sectors/usec =
* [15, 65G] sectors/sec
* Which, assuming a sector size of 512B, corresponds to a range of
* [7.5K, 33T] B/sec
*/
#define BFQ_RATE_SHIFT 16
/*
* When configured for computing the duration of the weight-raising
* for interactive queues automatically (see the comments at the
* beginning of this file), BFQ does it using the following formula:
* duration = (ref_rate / r) * ref_wr_duration,
* where r is the peak rate of the device, and ref_rate and
* ref_wr_duration are two reference parameters. In particular,
* ref_rate is the peak rate of the reference storage device (see
* below), and ref_wr_duration is about the maximum time needed, with
* BFQ and while reading two files in parallel, to load typical large
* applications on the reference device (see the comments on
* max_service_from_wr below, for more details on how ref_wr_duration
* is obtained). In practice, the slower/faster the device at hand
* is, the more/less it takes to load applications with respect to the
* reference device. Accordingly, the longer/shorter BFQ grants
* weight raising to interactive applications.
*
* BFQ uses two different reference pairs (ref_rate, ref_wr_duration),
* depending on whether the device is rotational or non-rotational.
*
* In the following definitions, ref_rate[0] and ref_wr_duration[0]
* are the reference values for a rotational device, whereas
* ref_rate[1] and ref_wr_duration[1] are the reference values for a
* non-rotational device. The reference rates are not the actual peak
* rates of the devices used as a reference, but slightly lower
* values. The reason for using slightly lower values is that the
* peak-rate estimator tends to yield slightly lower values than the
* actual peak rate (it can yield the actual peak rate only if there
* is only one process doing I/O, and the process does sequential
* I/O).
*
* The reference peak rates are measured in sectors/usec, left-shifted
* by BFQ_RATE_SHIFT.
*/
static int ref_rate[2] = {14000, 33000};
/*
* To improve readability, a conversion function is used to initialize
* the following array, which entails that the array can be
* initialized only in a function.
*/
static int ref_wr_duration[2];
/*
* BFQ uses the above-detailed, time-based weight-raising mechanism to
* privilege interactive tasks. This mechanism is vulnerable to the
* following false positives: I/O-bound applications that will go on
* doing I/O for much longer than the duration of weight
* raising. These applications have basically no benefit from being
* weight-raised at the beginning of their I/O. On the opposite end,
* while being weight-raised, these applications
* a) unjustly steal throughput to applications that may actually need
* low latency;
* b) make BFQ uselessly perform device idling; device idling results
* in loss of device throughput with most flash-based storage, and may
* increase latencies when used purposelessly.
*
* BFQ tries to reduce these problems, by adopting the following
* countermeasure. To introduce this countermeasure, we need first to
* finish explaining how the duration of weight-raising for
* interactive tasks is computed.
*
* For a bfq_queue deemed as interactive, the duration of weight
* raising is dynamically adjusted, as a function of the estimated
* peak rate of the device, so as to be equal to the time needed to
* execute the 'largest' interactive task we benchmarked so far. By
* largest task, we mean the task for which each involved process has
* to do more I/O than for any of the other tasks we benchmarked. This
* reference interactive task is the start-up of LibreOffice Writer,
* and in this task each process/bfq_queue needs to have at most ~110K
* sectors transferred.
*
* This last piece of information enables BFQ to reduce the actual
* duration of weight-raising for at least one class of I/O-bound
* applications: those doing sequential or quasi-sequential I/O. An
* example is file copy. In fact, once started, the main I/O-bound
* processes of these applications usually consume the above 110K
* sectors in much less time than the processes of an application that
* is starting, because these I/O-bound processes will greedily devote
* almost all their CPU cycles only to their target,
* throughput-friendly I/O operations. This is even more true if BFQ
* happens to be underestimating the device peak rate, and thus
* overestimating the duration of weight raising. But, according to
* our measurements, once transferred 110K sectors, these processes
* have no right to be weight-raised any longer.
*
* Basing on the last consideration, BFQ ends weight-raising for a
* bfq_queue if the latter happens to have received an amount of
* service at least equal to the following constant. The constant is
* set to slightly more than 110K, to have a minimum safety margin.
*
* This early ending of weight-raising reduces the amount of time
* during which interactive false positives cause the two problems
* described at the beginning of these comments.
*/
static const unsigned long max_service_from_wr = 120000;
#define RQ_BIC(rq) icq_to_bic((rq)->elv.priv[0])
#define RQ_BFQQ(rq) ((rq)->elv.priv[1])
struct bfq_queue *bic_to_bfqq(struct bfq_io_cq *bic, bool is_sync)
{
return bic->bfqq[is_sync];
}
void bic_set_bfqq(struct bfq_io_cq *bic, struct bfq_queue *bfqq, bool is_sync)
{
bic->bfqq[is_sync] = bfqq;
}
struct bfq_data *bic_to_bfqd(struct bfq_io_cq *bic)
{
return bic->icq.q->elevator->elevator_data;
}
/**
* icq_to_bic - convert iocontext queue structure to bfq_io_cq.
* @icq: the iocontext queue.
*/
static struct bfq_io_cq *icq_to_bic(struct io_cq *icq)
{
/* bic->icq is the first member, %NULL will convert to %NULL */
return container_of(icq, struct bfq_io_cq, icq);
}
/**
* bfq_bic_lookup - search into @ioc a bic associated to @bfqd.
* @bfqd: the lookup key.
* @ioc: the io_context of the process doing I/O.
* @q: the request queue.
*/
static struct bfq_io_cq *bfq_bic_lookup(struct bfq_data *bfqd,
struct io_context *ioc,
struct request_queue *q)
{
if (ioc) {
unsigned long flags;
struct bfq_io_cq *icq;
spin_lock_irqsave(&q->queue_lock, flags);
icq = icq_to_bic(ioc_lookup_icq(ioc, q));
spin_unlock_irqrestore(&q->queue_lock, flags);
return icq;
}
return NULL;
}
/*
* Scheduler run of queue, if there are requests pending and no one in the
* driver that will restart queueing.
*/
void bfq_schedule_dispatch(struct bfq_data *bfqd)
{
if (bfqd->queued != 0) {
bfq_log(bfqd, "schedule dispatch");
blk_mq_run_hw_queues(bfqd->queue, true);
}
}
#define bfq_class_idle(bfqq) ((bfqq)->ioprio_class == IOPRIO_CLASS_IDLE)
#define bfq_class_rt(bfqq) ((bfqq)->ioprio_class == IOPRIO_CLASS_RT)
#define bfq_sample_valid(samples) ((samples) > 80)
/*
* Lifted from AS - choose which of rq1 and rq2 that is best served now.
* We choose the request that is closer to the head right now. Distance
* behind the head is penalized and only allowed to a certain extent.
*/
static struct request *bfq_choose_req(struct bfq_data *bfqd,
struct request *rq1,
struct request *rq2,
sector_t last)
{
sector_t s1, s2, d1 = 0, d2 = 0;
unsigned long back_max;
#define BFQ_RQ1_WRAP 0x01 /* request 1 wraps */
#define BFQ_RQ2_WRAP 0x02 /* request 2 wraps */
unsigned int wrap = 0; /* bit mask: requests behind the disk head? */
if (!rq1 || rq1 == rq2)
return rq2;
if (!rq2)
return rq1;
if (rq_is_sync(rq1) && !rq_is_sync(rq2))
return rq1;
else if (rq_is_sync(rq2) && !rq_is_sync(rq1))
return rq2;
if ((rq1->cmd_flags & REQ_META) && !(rq2->cmd_flags & REQ_META))
return rq1;
else if ((rq2->cmd_flags & REQ_META) && !(rq1->cmd_flags & REQ_META))
return rq2;
s1 = blk_rq_pos(rq1);
s2 = blk_rq_pos(rq2);
/*
* By definition, 1KiB is 2 sectors.
*/
back_max = bfqd->bfq_back_max * 2;
/*
* Strict one way elevator _except_ in the case where we allow
* short backward seeks which are biased as twice the cost of a
* similar forward seek.
*/
if (s1 >= last)
d1 = s1 - last;
else if (s1 + back_max >= last)
d1 = (last - s1) * bfqd->bfq_back_penalty;
else
wrap |= BFQ_RQ1_WRAP;
if (s2 >= last)
d2 = s2 - last;
else if (s2 + back_max >= last)
d2 = (last - s2) * bfqd->bfq_back_penalty;
else
wrap |= BFQ_RQ2_WRAP;
/* Found required data */
/*
* By doing switch() on the bit mask "wrap" we avoid having to
* check two variables for all permutations: --> faster!
*/
switch (wrap) {
case 0: /* common case for CFQ: rq1 and rq2 not wrapped */
if (d1 < d2)
return rq1;
else if (d2 < d1)
return rq2;
if (s1 >= s2)
return rq1;
else
return rq2;
case BFQ_RQ2_WRAP:
return rq1;
case BFQ_RQ1_WRAP:
return rq2;
case BFQ_RQ1_WRAP|BFQ_RQ2_WRAP: /* both rqs wrapped */
default:
/*
* Since both rqs are wrapped,
* start with the one that's further behind head
* (--> only *one* back seek required),
* since back seek takes more time than forward.
*/
if (s1 <= s2)
return rq1;
else
return rq2;
}
}
/*
* Async I/O can easily starve sync I/O (both sync reads and sync
* writes), by consuming all tags. Similarly, storms of sync writes,
* such as those that sync(2) may trigger, can starve sync reads.
* Limit depths of async I/O and sync writes so as to counter both
* problems.
*/
static void bfq_limit_depth(unsigned int op, struct blk_mq_alloc_data *data)
{
struct bfq_data *bfqd = data->q->elevator->elevator_data;
if (op_is_sync(op) && !op_is_write(op))
return;
data->shallow_depth =
bfqd->word_depths[!!bfqd->wr_busy_queues][op_is_sync(op)];
bfq_log(bfqd, "[%s] wr_busy %d sync %d depth %u",
__func__, bfqd->wr_busy_queues, op_is_sync(op),
data->shallow_depth);
}
static struct bfq_queue *
bfq_rq_pos_tree_lookup(struct bfq_data *bfqd, struct rb_root *root,
sector_t sector, struct rb_node **ret_parent,
struct rb_node ***rb_link)
{
struct rb_node **p, *parent;
struct bfq_queue *bfqq = NULL;
parent = NULL;
p = &root->rb_node;
while (*p) {
struct rb_node **n;
parent = *p;
bfqq = rb_entry(parent, struct bfq_queue, pos_node);
/*
* Sort strictly based on sector. Smallest to the left,
* largest to the right.
*/
if (sector > blk_rq_pos(bfqq->next_rq))
n = &(*p)->rb_right;
else if (sector < blk_rq_pos(bfqq->next_rq))
n = &(*p)->rb_left;
else
break;
p = n;
bfqq = NULL;
}
*ret_parent = parent;
if (rb_link)
*rb_link = p;
bfq_log(bfqd, "rq_pos_tree_lookup %llu: returning %d",
(unsigned long long)sector,
bfqq ? bfqq->pid : 0);
return bfqq;
}
static bool bfq_too_late_for_merging(struct bfq_queue *bfqq)
{
return bfqq->service_from_backlogged > 0 &&
time_is_before_jiffies(bfqq->first_IO_time +
bfq_merge_time_limit);
}
/*
* The following function is not marked as __cold because it is
* actually cold, but for the same performance goal described in the
* comments on the likely() at the beginning of
* bfq_setup_cooperator(). Unexpectedly, to reach an even lower
* execution time for the case where this function is not invoked, we
* had to add an unlikely() in each involved if().
*/
void __cold
bfq_pos_tree_add_move(struct bfq_data *bfqd, struct bfq_queue *bfqq)
{
struct rb_node **p, *parent;
struct bfq_queue *__bfqq;
if (bfqq->pos_root) {
rb_erase(&bfqq->pos_node, bfqq->pos_root);
bfqq->pos_root = NULL;
}
/*
* bfqq cannot be merged any longer (see comments in
* bfq_setup_cooperator): no point in adding bfqq into the
* position tree.
*/
if (bfq_too_late_for_merging(bfqq))
return;
if (bfq_class_idle(bfqq))
return;
if (!bfqq->next_rq)
return;
bfqq->pos_root = &bfq_bfqq_to_bfqg(bfqq)->rq_pos_tree;
__bfqq = bfq_rq_pos_tree_lookup(bfqd, bfqq->pos_root,
blk_rq_pos(bfqq->next_rq), &parent, &p);
if (!__bfqq) {
rb_link_node(&bfqq->pos_node, parent, p);
rb_insert_color(&bfqq->pos_node, bfqq->pos_root);
} else
bfqq->pos_root = NULL;
}
/*
* The following function returns false either if every active queue
* must receive the same share of the throughput (symmetric scenario),
* or, as a special case, if bfqq must receive a share of the
* throughput lower than or equal to the share that every other active
* queue must receive. If bfqq does sync I/O, then these are the only
* two cases where bfqq happens to be guaranteed its share of the
* throughput even if I/O dispatching is not plugged when bfqq remains
* temporarily empty (for more details, see the comments in the
* function bfq_better_to_idle()). For this reason, the return value
* of this function is used to check whether I/O-dispatch plugging can
* be avoided.
*
* The above first case (symmetric scenario) occurs when:
* 1) all active queues have the same weight,
* 2) all active queues belong to the same I/O-priority class,
* 3) all active groups at the same level in the groups tree have the same
* weight,
* 4) all active groups at the same level in the groups tree have the same
* number of children.
*
* Unfortunately, keeping the necessary state for evaluating exactly
* the last two symmetry sub-conditions above would be quite complex
* and time consuming. Therefore this function evaluates, instead,
* only the following stronger three sub-conditions, for which it is
* much easier to maintain the needed state:
* 1) all active queues have the same weight,
* 2) all active queues belong to the same I/O-priority class,
* 3) there are no active groups.
* In particular, the last condition is always true if hierarchical
* support or the cgroups interface are not enabled, thus no state
* needs to be maintained in this case.
*/
static bool bfq_asymmetric_scenario(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
bool smallest_weight = bfqq &&
bfqq->weight_counter &&
bfqq->weight_counter ==
container_of(
rb_first_cached(&bfqd->queue_weights_tree),
struct bfq_weight_counter,
weights_node);
/*
* For queue weights to differ, queue_weights_tree must contain
* at least two nodes.
*/
bool varied_queue_weights = !smallest_weight &&
!RB_EMPTY_ROOT(&bfqd->queue_weights_tree.rb_root) &&
(bfqd->queue_weights_tree.rb_root.rb_node->rb_left ||
bfqd->queue_weights_tree.rb_root.rb_node->rb_right);
bool multiple_classes_busy =
(bfqd->busy_queues[0] && bfqd->busy_queues[1]) ||
(bfqd->busy_queues[0] && bfqd->busy_queues[2]) ||
(bfqd->busy_queues[1] && bfqd->busy_queues[2]);
return varied_queue_weights || multiple_classes_busy
#ifdef CONFIG_BFQ_GROUP_IOSCHED
|| bfqd->num_groups_with_pending_reqs > 0
#endif
;
}
/*
* If the weight-counter tree passed as input contains no counter for
* the weight of the input queue, then add that counter; otherwise just
* increment the existing counter.
*
* Note that weight-counter trees contain few nodes in mostly symmetric
* scenarios. For example, if all queues have the same weight, then the
* weight-counter tree for the queues may contain at most one node.
* This holds even if low_latency is on, because weight-raised queues
* are not inserted in the tree.
* In most scenarios, the rate at which nodes are created/destroyed
* should be low too.
*/
void bfq_weights_tree_add(struct bfq_data *bfqd, struct bfq_queue *bfqq,
struct rb_root_cached *root)
{
struct bfq_entity *entity = &bfqq->entity;
struct rb_node **new = &(root->rb_root.rb_node), *parent = NULL;
bool leftmost = true;
/*
* Do not insert if the queue is already associated with a
* counter, which happens if:
* 1) a request arrival has caused the queue to become both
* non-weight-raised, and hence change its weight, and
* backlogged; in this respect, each of the two events
* causes an invocation of this function,
* 2) this is the invocation of this function caused by the
* second event. This second invocation is actually useless,
* and we handle this fact by exiting immediately. More
* efficient or clearer solutions might possibly be adopted.
*/
if (bfqq->weight_counter)
return;
while (*new) {
struct bfq_weight_counter *__counter = container_of(*new,
struct bfq_weight_counter,
weights_node);
parent = *new;
if (entity->weight == __counter->weight) {
bfqq->weight_counter = __counter;
goto inc_counter;
}
if (entity->weight < __counter->weight)
new = &((*new)->rb_left);
else {
new = &((*new)->rb_right);
leftmost = false;
}
}
bfqq->weight_counter = kzalloc(sizeof(struct bfq_weight_counter),
GFP_ATOMIC);
/*
* In the unlucky event of an allocation failure, we just
* exit. This will cause the weight of queue to not be
* considered in bfq_asymmetric_scenario, which, in its turn,
* causes the scenario to be deemed wrongly symmetric in case
* bfqq's weight would have been the only weight making the
* scenario asymmetric. On the bright side, no unbalance will
* however occur when bfqq becomes inactive again (the
* invocation of this function is triggered by an activation
* of queue). In fact, bfq_weights_tree_remove does nothing
* if !bfqq->weight_counter.
*/
if (unlikely(!bfqq->weight_counter))
return;
bfqq->weight_counter->weight = entity->weight;
rb_link_node(&bfqq->weight_counter->weights_node, parent, new);
rb_insert_color_cached(&bfqq->weight_counter->weights_node, root,
leftmost);
inc_counter:
bfqq->weight_counter->num_active++;
bfqq->ref++;
}
/*
* Decrement the weight counter associated with the queue, and, if the
* counter reaches 0, remove the counter from the tree.
* See the comments to the function bfq_weights_tree_add() for considerations
* about overhead.
*/
void __bfq_weights_tree_remove(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
struct rb_root_cached *root)
{
if (!bfqq->weight_counter)
return;
bfqq->weight_counter->num_active--;
if (bfqq->weight_counter->num_active > 0)
goto reset_entity_pointer;
rb_erase_cached(&bfqq->weight_counter->weights_node, root);
kfree(bfqq->weight_counter);
reset_entity_pointer:
bfqq->weight_counter = NULL;
bfq_put_queue(bfqq);
}
/*
* Invoke __bfq_weights_tree_remove on bfqq and decrement the number
* of active groups for each queue's inactive parent entity.
*/
void bfq_weights_tree_remove(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
struct bfq_entity *entity = bfqq->entity.parent;
for_each_entity(entity) {
struct bfq_sched_data *sd = entity->my_sched_data;
if (sd->next_in_service || sd->in_service_entity) {
/*
* entity is still active, because either
* next_in_service or in_service_entity is not
* NULL (see the comments on the definition of
* next_in_service for details on why
* in_service_entity must be checked too).
*
* As a consequence, its parent entities are
* active as well, and thus this loop must
* stop here.
*/
break;
}
/*
* The decrement of num_groups_with_pending_reqs is
* not performed immediately upon the deactivation of
* entity, but it is delayed to when it also happens
* that the first leaf descendant bfqq of entity gets
* all its pending requests completed. The following
* instructions perform this delayed decrement, if
* needed. See the comments on
* num_groups_with_pending_reqs for details.
*/
if (entity->in_groups_with_pending_reqs) {
entity->in_groups_with_pending_reqs = false;
bfqd->num_groups_with_pending_reqs--;
}
}
/*
* Next function is invoked last, because it causes bfqq to be
* freed if the following holds: bfqq is not in service and
* has no dispatched request. DO NOT use bfqq after the next
* function invocation.
*/
__bfq_weights_tree_remove(bfqd, bfqq,
&bfqd->queue_weights_tree);
}
/*
* Return expired entry, or NULL to just start from scratch in rbtree.
*/
static struct request *bfq_check_fifo(struct bfq_queue *bfqq,
struct request *last)
{
struct request *rq;
if (bfq_bfqq_fifo_expire(bfqq))
return NULL;
bfq_mark_bfqq_fifo_expire(bfqq);
rq = rq_entry_fifo(bfqq->fifo.next);
if (rq == last || ktime_get_ns() < rq->fifo_time)
return NULL;
bfq_log_bfqq(bfqq->bfqd, bfqq, "check_fifo: returned %p", rq);
return rq;
}
static struct request *bfq_find_next_rq(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
struct request *last)
{
struct rb_node *rbnext = rb_next(&last->rb_node);
struct rb_node *rbprev = rb_prev(&last->rb_node);
struct request *next, *prev = NULL;
/* Follow expired path, else get first next available. */
next = bfq_check_fifo(bfqq, last);
if (next)
return next;
if (rbprev)
prev = rb_entry_rq(rbprev);
if (rbnext)
next = rb_entry_rq(rbnext);
else {
rbnext = rb_first(&bfqq->sort_list);
if (rbnext && rbnext != &last->rb_node)
next = rb_entry_rq(rbnext);
}
return bfq_choose_req(bfqd, next, prev, blk_rq_pos(last));
}
/* see the definition of bfq_async_charge_factor for details */
static unsigned long bfq_serv_to_charge(struct request *rq,
struct bfq_queue *bfqq)
{
if (bfq_bfqq_sync(bfqq) || bfqq->wr_coeff > 1 ||
bfq_asymmetric_scenario(bfqq->bfqd, bfqq))
return blk_rq_sectors(rq);
return blk_rq_sectors(rq) * bfq_async_charge_factor;
}
/**
* bfq_updated_next_req - update the queue after a new next_rq selection.
* @bfqd: the device data the queue belongs to.
* @bfqq: the queue to update.
*
* If the first request of a queue changes we make sure that the queue
* has enough budget to serve at least its first request (if the
* request has grown). We do this because if the queue has not enough
* budget for its first request, it has to go through two dispatch
* rounds to actually get it dispatched.
*/
static void bfq_updated_next_req(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
struct bfq_entity *entity = &bfqq->entity;
struct request *next_rq = bfqq->next_rq;
unsigned long new_budget;
if (!next_rq)
return;
if (bfqq == bfqd->in_service_queue)
/*
* In order not to break guarantees, budgets cannot be
* changed after an entity has been selected.
*/
return;
new_budget = max_t(unsigned long,
max_t(unsigned long, bfqq->max_budget,
bfq_serv_to_charge(next_rq, bfqq)),
entity->service);
if (entity->budget != new_budget) {
entity->budget = new_budget;
bfq_log_bfqq(bfqd, bfqq, "updated next rq: new budget %lu",
new_budget);
bfq_requeue_bfqq(bfqd, bfqq, false);
}
}
static unsigned int bfq_wr_duration(struct bfq_data *bfqd)
{
u64 dur;
if (bfqd->bfq_wr_max_time > 0)
return bfqd->bfq_wr_max_time;
dur = bfqd->rate_dur_prod;
do_div(dur, bfqd->peak_rate);
/*
* Limit duration between 3 and 25 seconds. The upper limit
* has been conservatively set after the following worst case:
* on a QEMU/KVM virtual machine
* - running in a slow PC
* - with a virtual disk stacked on a slow low-end 5400rpm HDD
* - serving a heavy I/O workload, such as the sequential reading
* of several files
* mplayer took 23 seconds to start, if constantly weight-raised.
*
* As for higher values than that accommodating the above bad
* scenario, tests show that higher values would often yield
* the opposite of the desired result, i.e., would worsen
* responsiveness by allowing non-interactive applications to
* preserve weight raising for too long.
*
* On the other end, lower values than 3 seconds make it
* difficult for most interactive tasks to complete their jobs
* before weight-raising finishes.
*/
return clamp_val(dur, msecs_to_jiffies(3000), msecs_to_jiffies(25000));
}
/* switch back from soft real-time to interactive weight raising */
static void switch_back_to_interactive_wr(struct bfq_queue *bfqq,
struct bfq_data *bfqd)
{
bfqq->wr_coeff = bfqd->bfq_wr_coeff;
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd);
bfqq->last_wr_start_finish = bfqq->wr_start_at_switch_to_srt;
}
static void
bfq_bfqq_resume_state(struct bfq_queue *bfqq, struct bfq_data *bfqd,
struct bfq_io_cq *bic, bool bfq_already_existing)
{
unsigned int old_wr_coeff = bfqq->wr_coeff;
bool busy = bfq_already_existing && bfq_bfqq_busy(bfqq);
if (bic->saved_has_short_ttime)
bfq_mark_bfqq_has_short_ttime(bfqq);
else
bfq_clear_bfqq_has_short_ttime(bfqq);
if (bic->saved_IO_bound)
bfq_mark_bfqq_IO_bound(bfqq);
else
bfq_clear_bfqq_IO_bound(bfqq);
bfqq->entity.new_weight = bic->saved_weight;
bfqq->ttime = bic->saved_ttime;
bfqq->wr_coeff = bic->saved_wr_coeff;
bfqq->wr_start_at_switch_to_srt = bic->saved_wr_start_at_switch_to_srt;
bfqq->last_wr_start_finish = bic->saved_last_wr_start_finish;
bfqq->wr_cur_max_time = bic->saved_wr_cur_max_time;
if (bfqq->wr_coeff > 1 && (bfq_bfqq_in_large_burst(bfqq) ||
time_is_before_jiffies(bfqq->last_wr_start_finish +
bfqq->wr_cur_max_time))) {
if (bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time &&
!bfq_bfqq_in_large_burst(bfqq) &&
time_is_after_eq_jiffies(bfqq->wr_start_at_switch_to_srt +
bfq_wr_duration(bfqd))) {
switch_back_to_interactive_wr(bfqq, bfqd);
} else {
bfqq->wr_coeff = 1;
bfq_log_bfqq(bfqq->bfqd, bfqq,
"resume state: switching off wr");
}
}
/* make sure weight will be updated, however we got here */
bfqq->entity.prio_changed = 1;
if (likely(!busy))
return;
if (old_wr_coeff == 1 && bfqq->wr_coeff > 1)
bfqd->wr_busy_queues++;
else if (old_wr_coeff > 1 && bfqq->wr_coeff == 1)
bfqd->wr_busy_queues--;
}
static int bfqq_process_refs(struct bfq_queue *bfqq)
{
return bfqq->ref - bfqq->allocated - bfqq->entity.on_st -
(bfqq->weight_counter != NULL);
}
/* Empty burst list and add just bfqq (see comments on bfq_handle_burst) */
static void bfq_reset_burst_list(struct bfq_data *bfqd, struct bfq_queue *bfqq)
{
struct bfq_queue *item;
struct hlist_node *n;
hlist_for_each_entry_safe(item, n, &bfqd->burst_list, burst_list_node)
hlist_del_init(&item->burst_list_node);
/*
* Start the creation of a new burst list only if there is no
* active queue. See comments on the conditional invocation of
* bfq_handle_burst().
*/
if (bfq_tot_busy_queues(bfqd) == 0) {
hlist_add_head(&bfqq->burst_list_node, &bfqd->burst_list);
bfqd->burst_size = 1;
} else
bfqd->burst_size = 0;
bfqd->burst_parent_entity = bfqq->entity.parent;
}
/* Add bfqq to the list of queues in current burst (see bfq_handle_burst) */
static void bfq_add_to_burst(struct bfq_data *bfqd, struct bfq_queue *bfqq)
{
/* Increment burst size to take into account also bfqq */
bfqd->burst_size++;
if (bfqd->burst_size == bfqd->bfq_large_burst_thresh) {
struct bfq_queue *pos, *bfqq_item;
struct hlist_node *n;
/*
* Enough queues have been activated shortly after each
* other to consider this burst as large.
*/
bfqd->large_burst = true;
/*
* We can now mark all queues in the burst list as
* belonging to a large burst.
*/
hlist_for_each_entry(bfqq_item, &bfqd->burst_list,
burst_list_node)
bfq_mark_bfqq_in_large_burst(bfqq_item);
bfq_mark_bfqq_in_large_burst(bfqq);
/*
* From now on, and until the current burst finishes, any
* new queue being activated shortly after the last queue
* was inserted in the burst can be immediately marked as
* belonging to a large burst. So the burst list is not
* needed any more. Remove it.
*/
hlist_for_each_entry_safe(pos, n, &bfqd->burst_list,
burst_list_node)
hlist_del_init(&pos->burst_list_node);
} else /*
* Burst not yet large: add bfqq to the burst list. Do
* not increment the ref counter for bfqq, because bfqq
* is removed from the burst list before freeing bfqq
* in put_queue.
*/
hlist_add_head(&bfqq->burst_list_node, &bfqd->burst_list);
}
/*
* If many queues belonging to the same group happen to be created
* shortly after each other, then the processes associated with these
* queues have typically a common goal. In particular, bursts of queue
* creations are usually caused by services or applications that spawn
* many parallel threads/processes. Examples are systemd during boot,
* or git grep. To help these processes get their job done as soon as
* possible, it is usually better to not grant either weight-raising
* or device idling to their queues, unless these queues must be
* protected from the I/O flowing through other active queues.
*
* In this comment we describe, firstly, the reasons why this fact
* holds, and, secondly, the next function, which implements the main
* steps needed to properly mark these queues so that they can then be
* treated in a different way.
*
* The above services or applications benefit mostly from a high
* throughput: the quicker the requests of the activated queues are
* cumulatively served, the sooner the target job of these queues gets
* completed. As a consequence, weight-raising any of these queues,
* which also implies idling the device for it, is almost always
* counterproductive, unless there are other active queues to isolate
* these new queues from. If there no other active queues, then
* weight-raising these new queues just lowers throughput in most
* cases.
*
* On the other hand, a burst of queue creations may be caused also by
* the start of an application that does not consist of a lot of
* parallel I/O-bound threads. In fact, with a complex application,
* several short processes may need to be executed to start-up the
* application. In this respect, to start an application as quickly as
* possible, the best thing to do is in any case to privilege the I/O
* related to the application with respect to all other
* I/O. Therefore, the best strategy to start as quickly as possible
* an application that causes a burst of queue creations is to
* weight-raise all the queues created during the burst. This is the
* exact opposite of the best strategy for the other type of bursts.
*
* In the end, to take the best action for each of the two cases, the
* two types of bursts need to be distinguished. Fortunately, this
* seems relatively easy, by looking at the sizes of the bursts. In
* particular, we found a threshold such that only bursts with a
* larger size than that threshold are apparently caused by
* services or commands such as systemd or git grep. For brevity,
* hereafter we call just 'large' these bursts. BFQ *does not*
* weight-raise queues whose creation occurs in a large burst. In
* addition, for each of these queues BFQ performs or does not perform
* idling depending on which choice boosts the throughput more. The
* exact choice depends on the device and request pattern at
* hand.
*
* Unfortunately, false positives may occur while an interactive task
* is starting (e.g., an application is being started). The
* consequence is that the queues associated with the task do not
* enjoy weight raising as expected. Fortunately these false positives
* are very rare. They typically occur if some service happens to
* start doing I/O exactly when the interactive task starts.
*
* Turning back to the next function, it is invoked only if there are
* no active queues (apart from active queues that would belong to the
* same, possible burst bfqq would belong to), and it implements all
* the steps needed to detect the occurrence of a large burst and to
* properly mark all the queues belonging to it (so that they can then
* be treated in a different way). This goal is achieved by
* maintaining a "burst list" that holds, temporarily, the queues that
* belong to the burst in progress. The list is then used to mark
* these queues as belonging to a large burst if the burst does become
* large. The main steps are the following.
*
* . when the very first queue is created, the queue is inserted into the
* list (as it could be the first queue in a possible burst)
*
* . if the current burst has not yet become large, and a queue Q that does
* not yet belong to the burst is activated shortly after the last time
* at which a new queue entered the burst list, then the function appends
* Q to the burst list
*
* . if, as a consequence of the previous step, the burst size reaches
* the large-burst threshold, then
*
* . all the queues in the burst list are marked as belonging to a
* large burst
*
* . the burst list is deleted; in fact, the burst list already served
* its purpose (keeping temporarily track of the queues in a burst,
* so as to be able to mark them as belonging to a large burst in the
* previous sub-step), and now is not needed any more
*
* . the device enters a large-burst mode
*
* . if a queue Q that does not belong to the burst is created while
* the device is in large-burst mode and shortly after the last time
* at which a queue either entered the burst list or was marked as
* belonging to the current large burst, then Q is immediately marked
* as belonging to a large burst.
*
* . if a queue Q that does not belong to the burst is created a while
* later, i.e., not shortly after, than the last time at which a queue
* either entered the burst list or was marked as belonging to the
* current large burst, then the current burst is deemed as finished and:
*
* . the large-burst mode is reset if set
*
* . the burst list is emptied
*
* . Q is inserted in the burst list, as Q may be the first queue
* in a possible new burst (then the burst list contains just Q
* after this step).
*/
static void bfq_handle_burst(struct bfq_data *bfqd, struct bfq_queue *bfqq)
{
/*
* If bfqq is already in the burst list or is part of a large
* burst, or finally has just been split, then there is
* nothing else to do.
*/
if (!hlist_unhashed(&bfqq->burst_list_node) ||
bfq_bfqq_in_large_burst(bfqq) ||
time_is_after_eq_jiffies(bfqq->split_time +
msecs_to_jiffies(10)))
return;
/*
* If bfqq's creation happens late enough, or bfqq belongs to
* a different group than the burst group, then the current
* burst is finished, and related data structures must be
* reset.
*
* In this respect, consider the special case where bfqq is
* the very first queue created after BFQ is selected for this
* device. In this case, last_ins_in_burst and
* burst_parent_entity are not yet significant when we get
* here. But it is easy to verify that, whether or not the
* following condition is true, bfqq will end up being
* inserted into the burst list. In particular the list will
* happen to contain only bfqq. And this is exactly what has
* to happen, as bfqq may be the first queue of the first
* burst.
*/
if (time_is_before_jiffies(bfqd->last_ins_in_burst +
bfqd->bfq_burst_interval) ||
bfqq->entity.parent != bfqd->burst_parent_entity) {
bfqd->large_burst = false;
bfq_reset_burst_list(bfqd, bfqq);
goto end;
}
/*
* If we get here, then bfqq is being activated shortly after the
* last queue. So, if the current burst is also large, we can mark
* bfqq as belonging to this large burst immediately.
*/
if (bfqd->large_burst) {
bfq_mark_bfqq_in_large_burst(bfqq);
goto end;
}
/*
* If we get here, then a large-burst state has not yet been
* reached, but bfqq is being activated shortly after the last
* queue. Then we add bfqq to the burst.
*/
bfq_add_to_burst(bfqd, bfqq);
end:
/*
* At this point, bfqq either has been added to the current
* burst or has caused the current burst to terminate and a
* possible new burst to start. In particular, in the second
* case, bfqq has become the first queue in the possible new
* burst. In both cases last_ins_in_burst needs to be moved
* forward.
*/
bfqd->last_ins_in_burst = jiffies;
}
static int bfq_bfqq_budget_left(struct bfq_queue *bfqq)
{
struct bfq_entity *entity = &bfqq->entity;
return entity->budget - entity->service;
}
/*
* If enough samples have been computed, return the current max budget
* stored in bfqd, which is dynamically updated according to the
* estimated disk peak rate; otherwise return the default max budget
*/
static int bfq_max_budget(struct bfq_data *bfqd)
{
if (bfqd->budgets_assigned < bfq_stats_min_budgets)
return bfq_default_max_budget;
else
return bfqd->bfq_max_budget;
}
/*
* Return min budget, which is a fraction of the current or default
* max budget (trying with 1/32)
*/
static int bfq_min_budget(struct bfq_data *bfqd)
{
if (bfqd->budgets_assigned < bfq_stats_min_budgets)
return bfq_default_max_budget / 32;
else
return bfqd->bfq_max_budget / 32;
}
/*
* The next function, invoked after the input queue bfqq switches from
* idle to busy, updates the budget of bfqq. The function also tells
* whether the in-service queue should be expired, by returning
* true. The purpose of expiring the in-service queue is to give bfqq
* the chance to possibly preempt the in-service queue, and the reason
* for preempting the in-service queue is to achieve one of the two
* goals below.
*
* 1. Guarantee to bfqq its reserved bandwidth even if bfqq has
* expired because it has remained idle. In particular, bfqq may have
* expired for one of the following two reasons:
*
* - BFQQE_NO_MORE_REQUESTS bfqq did not enjoy any device idling
* and did not make it to issue a new request before its last
* request was served;
*
* - BFQQE_TOO_IDLE bfqq did enjoy device idling, but did not issue
* a new request before the expiration of the idling-time.
*
* Even if bfqq has expired for one of the above reasons, the process
* associated with the queue may be however issuing requests greedily,
* and thus be sensitive to the bandwidth it receives (bfqq may have
* remained idle for other reasons: CPU high load, bfqq not enjoying
* idling, I/O throttling somewhere in the path from the process to
* the I/O scheduler, ...). But if, after every expiration for one of
* the above two reasons, bfqq has to wait for the service of at least
* one full budget of another queue before being served again, then
* bfqq is likely to get a much lower bandwidth or resource time than
* its reserved ones. To address this issue, two countermeasures need
* to be taken.
*
* First, the budget and the timestamps of bfqq need to be updated in
* a special way on bfqq reactivation: they need to be updated as if
* bfqq did not remain idle and did not expire. In fact, if they are
* computed as if bfqq expired and remained idle until reactivation,
* then the process associated with bfqq is treated as if, instead of
* being greedy, it stopped issuing requests when bfqq remained idle,
* and restarts issuing requests only on this reactivation. In other
* words, the scheduler does not help the process recover the "service
* hole" between bfqq expiration and reactivation. As a consequence,
* the process receives a lower bandwidth than its reserved one. In
* contrast, to recover this hole, the budget must be updated as if
* bfqq was not expired at all before this reactivation, i.e., it must
* be set to the value of the remaining budget when bfqq was
* expired. Along the same line, timestamps need to be assigned the
* value they had the last time bfqq was selected for service, i.e.,
* before last expiration. Thus timestamps need to be back-shifted
* with respect to their normal computation (see [1] for more details
* on this tricky aspect).
*
* Secondly, to allow the process to recover the hole, the in-service
* queue must be expired too, to give bfqq the chance to preempt it
* immediately. In fact, if bfqq has to wait for a full budget of the
* in-service queue to be completed, then it may become impossible to
* let the process recover the hole, even if the back-shifted
* timestamps of bfqq are lower than those of the in-service queue. If
* this happens for most or all of the holes, then the process may not
* receive its reserved bandwidth. In this respect, it is worth noting
* that, being the service of outstanding requests unpreemptible, a
* little fraction of the holes may however be unrecoverable, thereby
* causing a little loss of bandwidth.
*
* The last important point is detecting whether bfqq does need this
* bandwidth recovery. In this respect, the next function deems the
* process associated with bfqq greedy, and thus allows it to recover
* the hole, if: 1) the process is waiting for the arrival of a new
* request (which implies that bfqq expired for one of the above two
* reasons), and 2) such a request has arrived soon. The first
* condition is controlled through the flag non_blocking_wait_rq,
* while the second through the flag arrived_in_time. If both
* conditions hold, then the function computes the budget in the
* above-described special way, and signals that the in-service queue
* should be expired. Timestamp back-shifting is done later in
* __bfq_activate_entity.
*
* 2. Reduce latency. Even if timestamps are not backshifted to let
* the process associated with bfqq recover a service hole, bfqq may
* however happen to have, after being (re)activated, a lower finish
* timestamp than the in-service queue. That is, the next budget of
* bfqq may have to be completed before the one of the in-service
* queue. If this is the case, then preempting the in-service queue
* allows this goal to be achieved, apart from the unpreemptible,
* outstanding requests mentioned above.
*
* Unfortunately, regardless of which of the above two goals one wants
* to achieve, service trees need first to be updated to know whether
* the in-service queue must be preempted. To have service trees
* correctly updated, the in-service queue must be expired and
* rescheduled, and bfqq must be scheduled too. This is one of the
* most costly operations (in future versions, the scheduling
* mechanism may be re-designed in such a way to make it possible to
* know whether preemption is needed without needing to update service
* trees). In addition, queue preemptions almost always cause random
* I/O, which may in turn cause loss of throughput. Finally, there may
* even be no in-service queue when the next function is invoked (so,
* no queue to compare timestamps with). Because of these facts, the
* next function adopts the following simple scheme to avoid costly
* operations, too frequent preemptions and too many dependencies on
* the state of the scheduler: it requests the expiration of the
* in-service queue (unconditionally) only for queues that need to
* recover a hole. Then it delegates to other parts of the code the
* responsibility of handling the above case 2.
*/
static bool bfq_bfqq_update_budg_for_activation(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
bool arrived_in_time)
{
struct bfq_entity *entity = &bfqq->entity;
/*
* In the next compound condition, we check also whether there
* is some budget left, because otherwise there is no point in
* trying to go on serving bfqq with this same budget: bfqq
* would be expired immediately after being selected for
* service. This would only cause useless overhead.
*/
if (bfq_bfqq_non_blocking_wait_rq(bfqq) && arrived_in_time &&
bfq_bfqq_budget_left(bfqq) > 0) {
/*
* We do not clear the flag non_blocking_wait_rq here, as
* the latter is used in bfq_activate_bfqq to signal
* that timestamps need to be back-shifted (and is
* cleared right after).
*/
/*
* In next assignment we rely on that either
* entity->service or entity->budget are not updated
* on expiration if bfqq is empty (see
* __bfq_bfqq_recalc_budget). Thus both quantities
* remain unchanged after such an expiration, and the
* following statement therefore assigns to
* entity->budget the remaining budget on such an
* expiration.
*/
entity->budget = min_t(unsigned long,
bfq_bfqq_budget_left(bfqq),
bfqq->max_budget);
/*
* At this point, we have used entity->service to get
* the budget left (needed for updating
* entity->budget). Thus we finally can, and have to,
* reset entity->service. The latter must be reset
* because bfqq would otherwise be charged again for
* the service it has received during its previous
* service slot(s).
*/
entity->service = 0;
return true;
}
/*
* We can finally complete expiration, by setting service to 0.
*/
entity->service = 0;
entity->budget = max_t(unsigned long, bfqq->max_budget,
bfq_serv_to_charge(bfqq->next_rq, bfqq));
bfq_clear_bfqq_non_blocking_wait_rq(bfqq);
return false;
}
/*
* Return the farthest past time instant according to jiffies
* macros.
*/
static unsigned long bfq_smallest_from_now(void)
{
return jiffies - MAX_JIFFY_OFFSET;
}
static void bfq_update_bfqq_wr_on_rq_arrival(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
unsigned int old_wr_coeff,
bool wr_or_deserves_wr,
bool interactive,
bool in_burst,
bool soft_rt)
{
if (old_wr_coeff == 1 && wr_or_deserves_wr) {
/* start a weight-raising period */
if (interactive) {
bfqq->service_from_wr = 0;
bfqq->wr_coeff = bfqd->bfq_wr_coeff;
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd);
} else {
/*
* No interactive weight raising in progress
* here: assign minus infinity to
* wr_start_at_switch_to_srt, to make sure
* that, at the end of the soft-real-time
* weight raising periods that is starting
* now, no interactive weight-raising period
* may be wrongly considered as still in
* progress (and thus actually started by
* mistake).
*/
bfqq->wr_start_at_switch_to_srt =
bfq_smallest_from_now();
bfqq->wr_coeff = bfqd->bfq_wr_coeff *
BFQ_SOFTRT_WEIGHT_FACTOR;
bfqq->wr_cur_max_time =
bfqd->bfq_wr_rt_max_time;
}
/*
* If needed, further reduce budget to make sure it is
* close to bfqq's backlog, so as to reduce the
* scheduling-error component due to a too large
* budget. Do not care about throughput consequences,
* but only about latency. Finally, do not assign a
* too small budget either, to avoid increasing
* latency by causing too frequent expirations.
*/
bfqq->entity.budget = min_t(unsigned long,
bfqq->entity.budget,
2 * bfq_min_budget(bfqd));
} else if (old_wr_coeff > 1) {
if (interactive) { /* update wr coeff and duration */
bfqq->wr_coeff = bfqd->bfq_wr_coeff;
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd);
} else if (in_burst)
bfqq->wr_coeff = 1;
else if (soft_rt) {
/*
* The application is now or still meeting the
* requirements for being deemed soft rt. We
* can then correctly and safely (re)charge
* the weight-raising duration for the
* application with the weight-raising
* duration for soft rt applications.
*
* In particular, doing this recharge now, i.e.,
* before the weight-raising period for the
* application finishes, reduces the probability
* of the following negative scenario:
* 1) the weight of a soft rt application is
* raised at startup (as for any newly
* created application),
* 2) since the application is not interactive,
* at a certain time weight-raising is
* stopped for the application,
* 3) at that time the application happens to
* still have pending requests, and hence
* is destined to not have a chance to be
* deemed soft rt before these requests are
* completed (see the comments to the
* function bfq_bfqq_softrt_next_start()
* for details on soft rt detection),
* 4) these pending requests experience a high
* latency because the application is not
* weight-raised while they are pending.
*/
if (bfqq->wr_cur_max_time !=
bfqd->bfq_wr_rt_max_time) {
bfqq->wr_start_at_switch_to_srt =
bfqq->last_wr_start_finish;
bfqq->wr_cur_max_time =
bfqd->bfq_wr_rt_max_time;
bfqq->wr_coeff = bfqd->bfq_wr_coeff *
BFQ_SOFTRT_WEIGHT_FACTOR;
}
bfqq->last_wr_start_finish = jiffies;
}
}
}
static bool bfq_bfqq_idle_for_long_time(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
return bfqq->dispatched == 0 &&
time_is_before_jiffies(
bfqq->budget_timeout +
bfqd->bfq_wr_min_idle_time);
}
/*
* Return true if bfqq is in a higher priority class, or has a higher
* weight than the in-service queue.
*/
static bool bfq_bfqq_higher_class_or_weight(struct bfq_queue *bfqq,
struct bfq_queue *in_serv_bfqq)
{
int bfqq_weight, in_serv_weight;
if (bfqq->ioprio_class < in_serv_bfqq->ioprio_class)
return true;
if (in_serv_bfqq->entity.parent == bfqq->entity.parent) {
bfqq_weight = bfqq->entity.weight;
in_serv_weight = in_serv_bfqq->entity.weight;
} else {
if (bfqq->entity.parent)
bfqq_weight = bfqq->entity.parent->weight;
else
bfqq_weight = bfqq->entity.weight;
if (in_serv_bfqq->entity.parent)
in_serv_weight = in_serv_bfqq->entity.parent->weight;
else
in_serv_weight = in_serv_bfqq->entity.weight;
}
return bfqq_weight > in_serv_weight;
}
static void bfq_bfqq_handle_idle_busy_switch(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
int old_wr_coeff,
struct request *rq,
bool *interactive)
{
bool soft_rt, in_burst, wr_or_deserves_wr,
bfqq_wants_to_preempt,
idle_for_long_time = bfq_bfqq_idle_for_long_time(bfqd, bfqq),
/*
* See the comments on
* bfq_bfqq_update_budg_for_activation for
* details on the usage of the next variable.
*/
arrived_in_time = ktime_get_ns() <=
bfqq->ttime.last_end_request +
bfqd->bfq_slice_idle * 3;
/*
* bfqq deserves to be weight-raised if:
* - it is sync,
* - it does not belong to a large burst,
* - it has been idle for enough time or is soft real-time,
* - is linked to a bfq_io_cq (it is not shared in any sense).
*/
in_burst = bfq_bfqq_in_large_burst(bfqq);
soft_rt = bfqd->bfq_wr_max_softrt_rate > 0 &&
!BFQQ_TOTALLY_SEEKY(bfqq) &&
!in_burst &&
time_is_before_jiffies(bfqq->soft_rt_next_start) &&
bfqq->dispatched == 0;
*interactive = !in_burst && idle_for_long_time;
wr_or_deserves_wr = bfqd->low_latency &&
(bfqq->wr_coeff > 1 ||
(bfq_bfqq_sync(bfqq) &&
bfqq->bic && (*interactive || soft_rt)));
/*
* Using the last flag, update budget and check whether bfqq
* may want to preempt the in-service queue.
*/
bfqq_wants_to_preempt =
bfq_bfqq_update_budg_for_activation(bfqd, bfqq,
arrived_in_time);
/*
* If bfqq happened to be activated in a burst, but has been
* idle for much more than an interactive queue, then we
* assume that, in the overall I/O initiated in the burst, the
* I/O associated with bfqq is finished. So bfqq does not need
* to be treated as a queue belonging to a burst
* anymore. Accordingly, we reset bfqq's in_large_burst flag
* if set, and remove bfqq from the burst list if it's
* there. We do not decrement burst_size, because the fact
* that bfqq does not need to belong to the burst list any
* more does not invalidate the fact that bfqq was created in
* a burst.
*/
if (likely(!bfq_bfqq_just_created(bfqq)) &&
idle_for_long_time &&
time_is_before_jiffies(
bfqq->budget_timeout +
msecs_to_jiffies(10000))) {
hlist_del_init(&bfqq->burst_list_node);
bfq_clear_bfqq_in_large_burst(bfqq);
}
bfq_clear_bfqq_just_created(bfqq);
if (!bfq_bfqq_IO_bound(bfqq)) {
if (arrived_in_time) {
bfqq->requests_within_timer++;
if (bfqq->requests_within_timer >=
bfqd->bfq_requests_within_timer)
bfq_mark_bfqq_IO_bound(bfqq);
} else
bfqq->requests_within_timer = 0;
}
if (bfqd->low_latency) {
if (unlikely(time_is_after_jiffies(bfqq->split_time)))
/* wraparound */
bfqq->split_time =
jiffies - bfqd->bfq_wr_min_idle_time - 1;
if (time_is_before_jiffies(bfqq->split_time +
bfqd->bfq_wr_min_idle_time)) {
bfq_update_bfqq_wr_on_rq_arrival(bfqd, bfqq,
old_wr_coeff,
wr_or_deserves_wr,
*interactive,
in_burst,
soft_rt);
if (old_wr_coeff != bfqq->wr_coeff)
bfqq->entity.prio_changed = 1;
}
}
bfqq->last_idle_bklogged = jiffies;
bfqq->service_from_backlogged = 0;
bfq_clear_bfqq_softrt_update(bfqq);
bfq_add_bfqq_busy(bfqd, bfqq);
/*
* Expire in-service queue only if preemption may be needed
* for guarantees. In particular, we care only about two
* cases. The first is that bfqq has to recover a service
* hole, as explained in the comments on
* bfq_bfqq_update_budg_for_activation(), i.e., that
* bfqq_wants_to_preempt is true. However, if bfqq does not
* carry time-critical I/O, then bfqq's bandwidth is less
* important than that of queues that carry time-critical I/O.
* So, as a further constraint, we consider this case only if
* bfqq is at least as weight-raised, i.e., at least as time
* critical, as the in-service queue.
*
* The second case is that bfqq is in a higher priority class,
* or has a higher weight than the in-service queue. If this
* condition does not hold, we don't care because, even if
* bfqq does not start to be served immediately, the resulting
* delay for bfqq's I/O is however lower or much lower than
* the ideal completion time to be guaranteed to bfqq's I/O.
*
* In both cases, preemption is needed only if, according to
* the timestamps of both bfqq and of the in-service queue,
* bfqq actually is the next queue to serve. So, to reduce
* useless preemptions, the return value of
* next_queue_may_preempt() is considered in the next compound
* condition too. Yet next_queue_may_preempt() just checks a
* simple, necessary condition for bfqq to be the next queue
* to serve. In fact, to evaluate a sufficient condition, the
* timestamps of the in-service queue would need to be
* updated, and this operation is quite costly (see the
* comments on bfq_bfqq_update_budg_for_activation()).
*/
if (bfqd->in_service_queue &&
((bfqq_wants_to_preempt &&
bfqq->wr_coeff >= bfqd->in_service_queue->wr_coeff) ||
bfq_bfqq_higher_class_or_weight(bfqq, bfqd->in_service_queue)) &&
next_queue_may_preempt(bfqd))
bfq_bfqq_expire(bfqd, bfqd->in_service_queue,
false, BFQQE_PREEMPTED);
}
static void bfq_reset_inject_limit(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
/* invalidate baseline total service time */
bfqq->last_serv_time_ns = 0;
/*
* Reset pointer in case we are waiting for
* some request completion.
*/
bfqd->waited_rq = NULL;
/*
* If bfqq has a short think time, then start by setting the
* inject limit to 0 prudentially, because the service time of
* an injected I/O request may be higher than the think time
* of bfqq, and therefore, if one request was injected when
* bfqq remains empty, this injected request might delay the
* service of the next I/O request for bfqq significantly. In
* case bfqq can actually tolerate some injection, then the
* adaptive update will however raise the limit soon. This
* lucky circumstance holds exactly because bfqq has a short
* think time, and thus, after remaining empty, is likely to
* get new I/O enqueued---and then completed---before being
* expired. This is the very pattern that gives the
* limit-update algorithm the chance to measure the effect of
* injection on request service times, and then to update the
* limit accordingly.
*
* However, in the following special case, the inject limit is
* left to 1 even if the think time is short: bfqq's I/O is
* synchronized with that of some other queue, i.e., bfqq may
* receive new I/O only after the I/O of the other queue is
* completed. Keeping the inject limit to 1 allows the
* blocking I/O to be served while bfqq is in service. And
* this is very convenient both for bfqq and for overall
* throughput, as explained in detail in the comments in
* bfq_update_has_short_ttime().
*
* On the opposite end, if bfqq has a long think time, then
* start directly by 1, because:
* a) on the bright side, keeping at most one request in
* service in the drive is unlikely to cause any harm to the
* latency of bfqq's requests, as the service time of a single
* request is likely to be lower than the think time of bfqq;
* b) on the downside, after becoming empty, bfqq is likely to
* expire before getting its next request. With this request
* arrival pattern, it is very hard to sample total service
* times and update the inject limit accordingly (see comments
* on bfq_update_inject_limit()). So the limit is likely to be
* never, or at least seldom, updated. As a consequence, by
* setting the limit to 1, we avoid that no injection ever
* occurs with bfqq. On the downside, this proactive step
* further reduces chances to actually compute the baseline
* total service time. Thus it reduces chances to execute the
* limit-update algorithm and possibly raise the limit to more
* than 1.
*/
if (bfq_bfqq_has_short_ttime(bfqq))
bfqq->inject_limit = 0;
else
bfqq->inject_limit = 1;
bfqq->decrease_time_jif = jiffies;
}
static void bfq_add_request(struct request *rq)
{
struct bfq_queue *bfqq = RQ_BFQQ(rq);
struct bfq_data *bfqd = bfqq->bfqd;
struct request *next_rq, *prev;
unsigned int old_wr_coeff = bfqq->wr_coeff;
bool interactive = false;
bfq_log_bfqq(bfqd, bfqq, "add_request %d", rq_is_sync(rq));
bfqq->queued[rq_is_sync(rq)]++;
bfqd->queued++;
if (RB_EMPTY_ROOT(&bfqq->sort_list) && bfq_bfqq_sync(bfqq)) {
/*
* Detect whether bfqq's I/O seems synchronized with
* that of some other queue, i.e., whether bfqq, after
* remaining empty, happens to receive new I/O only
* right after some I/O request of the other queue has
* been completed. We call waker queue the other
* queue, and we assume, for simplicity, that bfqq may
* have at most one waker queue.
*
* A remarkable throughput boost can be reached by
* unconditionally injecting the I/O of the waker
* queue, every time a new bfq_dispatch_request
* happens to be invoked while I/O is being plugged
* for bfqq. In addition to boosting throughput, this
* unblocks bfqq's I/O, thereby improving bandwidth
* and latency for bfqq. Note that these same results
* may be achieved with the general injection
* mechanism, but less effectively. For details on
* this aspect, see the comments on the choice of the
* queue for injection in bfq_select_queue().
*
* Turning back to the detection of a waker queue, a
* queue Q is deemed as a waker queue for bfqq if, for
* two consecutive times, bfqq happens to become non
* empty right after a request of Q has been
* completed. In particular, on the first time, Q is
* tentatively set as a candidate waker queue, while
* on the second time, the flag
* bfq_bfqq_has_waker(bfqq) is set to confirm that Q
* is a waker queue for bfqq. These detection steps
* are performed only if bfqq has a long think time,
* so as to make it more likely that bfqq's I/O is
* actually being blocked by a synchronization. This
* last filter, plus the above two-times requirement,
* make false positives less likely.
*
* NOTE
*
* The sooner a waker queue is detected, the sooner
* throughput can be boosted by injecting I/O from the
* waker queue. Fortunately, detection is likely to be
* actually fast, for the following reasons. While
* blocked by synchronization, bfqq has a long think
* time. This implies that bfqq's inject limit is at
* least equal to 1 (see the comments in
* bfq_update_inject_limit()). So, thanks to
* injection, the waker queue is likely to be served
* during the very first I/O-plugging time interval
* for bfqq. This triggers the first step of the
* detection mechanism. Thanks again to injection, the
* candidate waker queue is then likely to be
* confirmed no later than during the next
* I/O-plugging interval for bfqq.
*/
if (bfqd->last_completed_rq_bfqq &&
!bfq_bfqq_has_short_ttime(bfqq) &&
ktime_get_ns() - bfqd->last_completion <
200 * NSEC_PER_USEC) {
if (bfqd->last_completed_rq_bfqq != bfqq &&
bfqd->last_completed_rq_bfqq !=
bfqq->waker_bfqq) {
/*
* First synchronization detected with
* a candidate waker queue, or with a
* different candidate waker queue
* from the current one.
*/
bfqq->waker_bfqq = bfqd->last_completed_rq_bfqq;
/*
* If the waker queue disappears, then
* bfqq->waker_bfqq must be reset. To
* this goal, we maintain in each
* waker queue a list, woken_list, of
* all the queues that reference the
* waker queue through their
* waker_bfqq pointer. When the waker
* queue exits, the waker_bfqq pointer
* of all the queues in the woken_list
* is reset.
*
* In addition, if bfqq is already in
* the woken_list of a waker queue,
* then, before being inserted into
* the woken_list of a new waker
* queue, bfqq must be removed from
* the woken_list of the old waker
* queue.
*/
if (!hlist_unhashed(&bfqq->woken_list_node))
hlist_del_init(&bfqq->woken_list_node);
hlist_add_head(&bfqq->woken_list_node,
&bfqd->last_completed_rq_bfqq->woken_list);
bfq_clear_bfqq_has_waker(bfqq);
} else if (bfqd->last_completed_rq_bfqq ==
bfqq->waker_bfqq &&
!bfq_bfqq_has_waker(bfqq)) {
/*
* synchronization with waker_bfqq
* seen for the second time
*/
bfq_mark_bfqq_has_waker(bfqq);
}
}
/*
* Periodically reset inject limit, to make sure that
* the latter eventually drops in case workload
* changes, see step (3) in the comments on
* bfq_update_inject_limit().
*/
if (time_is_before_eq_jiffies(bfqq->decrease_time_jif +
msecs_to_jiffies(1000)))
bfq_reset_inject_limit(bfqd, bfqq);
/*
* The following conditions must hold to setup a new
* sampling of total service time, and then a new
* update of the inject limit:
* - bfqq is in service, because the total service
* time is evaluated only for the I/O requests of
* the queues in service;
* - this is the right occasion to compute or to
* lower the baseline total service time, because
* there are actually no requests in the drive,
* or
* the baseline total service time is available, and
* this is the right occasion to compute the other
* quantity needed to update the inject limit, i.e.,
* the total service time caused by the amount of
* injection allowed by the current value of the
* limit. It is the right occasion because injection
* has actually been performed during the service
* hole, and there are still in-flight requests,
* which are very likely to be exactly the injected
* requests, or part of them;
* - the minimum interval for sampling the total
* service time and updating the inject limit has
* elapsed.
*/
if (bfqq == bfqd->in_service_queue &&
(bfqd->rq_in_driver == 0 ||
(bfqq->last_serv_time_ns > 0 &&
bfqd->rqs_injected && bfqd->rq_in_driver > 0)) &&
time_is_before_eq_jiffies(bfqq->decrease_time_jif +
msecs_to_jiffies(100))) {
bfqd->last_empty_occupied_ns = ktime_get_ns();
/*
* Start the state machine for measuring the
* total service time of rq: setting
* wait_dispatch will cause bfqd->waited_rq to
* be set when rq will be dispatched.
*/
bfqd->wait_dispatch = true;
bfqd->rqs_injected = false;
}
}
elv_rb_add(&bfqq->sort_list, rq);
/*
* Check if this request is a better next-serve candidate.
*/
prev = bfqq->next_rq;
next_rq = bfq_choose_req(bfqd, bfqq->next_rq, rq, bfqd->last_position);
bfqq->next_rq = next_rq;
/*
* Adjust priority tree position, if next_rq changes.
* See comments on bfq_pos_tree_add_move() for the unlikely().
*/
if (unlikely(!bfqd->nonrot_with_queueing && prev != bfqq->next_rq))
bfq_pos_tree_add_move(bfqd, bfqq);
if (!bfq_bfqq_busy(bfqq)) /* switching to busy ... */
bfq_bfqq_handle_idle_busy_switch(bfqd, bfqq, old_wr_coeff,
rq, &interactive);
else {
if (bfqd->low_latency && old_wr_coeff == 1 && !rq_is_sync(rq) &&
time_is_before_jiffies(
bfqq->last_wr_start_finish +
bfqd->bfq_wr_min_inter_arr_async)) {
bfqq->wr_coeff = bfqd->bfq_wr_coeff;
bfqq->wr_cur_max_time = bfq_wr_duration(bfqd);
bfqd->wr_busy_queues++;
bfqq->entity.prio_changed = 1;
}
if (prev != bfqq->next_rq)
bfq_updated_next_req(bfqd, bfqq);
}
/*
* Assign jiffies to last_wr_start_finish in the following
* cases:
*
* . if bfqq is not going to be weight-raised, because, for
* non weight-raised queues, last_wr_start_finish stores the
* arrival time of the last request; as of now, this piece
* of information is used only for deciding whether to
* weight-raise async queues
*
* . if bfqq is not weight-raised, because, if bfqq is now
* switching to weight-raised, then last_wr_start_finish
* stores the time when weight-raising starts
*
* . if bfqq is interactive, because, regardless of whether
* bfqq is currently weight-raised, the weight-raising
* period must start or restart (this case is considered
* separately because it is not detected by the above
* conditions, if bfqq is already weight-raised)
*
* last_wr_start_finish has to be updated also if bfqq is soft
* real-time, because the weight-raising period is constantly
* restarted on idle-to-busy transitions for these queues, but
* this is already done in bfq_bfqq_handle_idle_busy_switch if
* needed.
*/
if (bfqd->low_latency &&
(old_wr_coeff == 1 || bfqq->wr_coeff == 1 || interactive))
bfqq->last_wr_start_finish = jiffies;
}
static struct request *bfq_find_rq_fmerge(struct bfq_data *bfqd,
struct bio *bio,
struct request_queue *q)
{
struct bfq_queue *bfqq = bfqd->bio_bfqq;
if (bfqq)
return elv_rb_find(&bfqq->sort_list, bio_end_sector(bio));
return NULL;
}
static sector_t get_sdist(sector_t last_pos, struct request *rq)
{
if (last_pos)
return abs(blk_rq_pos(rq) - last_pos);
return 0;
}
#if 0 /* Still not clear if we can do without next two functions */
static void bfq_activate_request(struct request_queue *q, struct request *rq)
{
struct bfq_data *bfqd = q->elevator->elevator_data;
bfqd->rq_in_driver++;
}
static void bfq_deactivate_request(struct request_queue *q, struct request *rq)
{
struct bfq_data *bfqd = q->elevator->elevator_data;
bfqd->rq_in_driver--;
}
#endif
static void bfq_remove_request(struct request_queue *q,
struct request *rq)
{
struct bfq_queue *bfqq = RQ_BFQQ(rq);
struct bfq_data *bfqd = bfqq->bfqd;
const int sync = rq_is_sync(rq);
if (bfqq->next_rq == rq) {
bfqq->next_rq = bfq_find_next_rq(bfqd, bfqq, rq);
bfq_updated_next_req(bfqd, bfqq);
}
if (rq->queuelist.prev != &rq->queuelist)
list_del_init(&rq->queuelist);
bfqq->queued[sync]--;
bfqd->queued--;
elv_rb_del(&bfqq->sort_list, rq);
elv_rqhash_del(q, rq);
if (q->last_merge == rq)
q->last_merge = NULL;
if (RB_EMPTY_ROOT(&bfqq->sort_list)) {
bfqq->next_rq = NULL;
if (bfq_bfqq_busy(bfqq) && bfqq != bfqd->in_service_queue) {
bfq_del_bfqq_busy(bfqd, bfqq, false);
/*
* bfqq emptied. In normal operation, when
* bfqq is empty, bfqq->entity.service and
* bfqq->entity.budget must contain,
* respectively, the service received and the
* budget used last time bfqq emptied. These
* facts do not hold in this case, as at least
* this last removal occurred while bfqq is
* not in service. To avoid inconsistencies,
* reset both bfqq->entity.service and
* bfqq->entity.budget, if bfqq has still a
* process that may issue I/O requests to it.
*/
bfqq->entity.budget = bfqq->entity.service = 0;
}
/*
* Remove queue from request-position tree as it is empty.
*/
if (bfqq->pos_root) {
rb_erase(&bfqq->pos_node, bfqq->pos_root);
bfqq->pos_root = NULL;
}
} else {
/* see comments on bfq_pos_tree_add_move() for the unlikely() */
if (unlikely(!bfqd->nonrot_with_queueing))
bfq_pos_tree_add_move(bfqd, bfqq);
}
if (rq->cmd_flags & REQ_META)
bfqq->meta_pending--;
}
static bool bfq_bio_merge(struct blk_mq_hw_ctx *hctx, struct bio *bio,
unsigned int nr_segs)
{
struct request_queue *q = hctx->queue;
struct bfq_data *bfqd = q->elevator->elevator_data;
struct request *free = NULL;
/*
* bfq_bic_lookup grabs the queue_lock: invoke it now and
* store its return value for later use, to avoid nesting
* queue_lock inside the bfqd->lock. We assume that the bic
* returned by bfq_bic_lookup does not go away before
* bfqd->lock is taken.
*/
struct bfq_io_cq *bic = bfq_bic_lookup(bfqd, current->io_context, q);
bool ret;
spin_lock_irq(&bfqd->lock);
if (bic)
bfqd->bio_bfqq = bic_to_bfqq(bic, op_is_sync(bio->bi_opf));
else
bfqd->bio_bfqq = NULL;
bfqd->bio_bic = bic;
ret = blk_mq_sched_try_merge(q, bio, nr_segs, &free);
if (free)
blk_mq_free_request(free);
spin_unlock_irq(&bfqd->lock);
return ret;
}
static int bfq_request_merge(struct request_queue *q, struct request **req,
struct bio *bio)
{
struct bfq_data *bfqd = q->elevator->elevator_data;
struct request *__rq;
__rq = bfq_find_rq_fmerge(bfqd, bio, q);
if (__rq && elv_bio_merge_ok(__rq, bio)) {
*req = __rq;
return ELEVATOR_FRONT_MERGE;
}
return ELEVATOR_NO_MERGE;
}
static struct bfq_queue *bfq_init_rq(struct request *rq);
static void bfq_request_merged(struct request_queue *q, struct request *req,
enum elv_merge type)
{
if (type == ELEVATOR_FRONT_MERGE &&
rb_prev(&req->rb_node) &&
blk_rq_pos(req) <
blk_rq_pos(container_of(rb_prev(&req->rb_node),
struct request, rb_node))) {
struct bfq_queue *bfqq = bfq_init_rq(req);
struct bfq_data *bfqd;
struct request *prev, *next_rq;
if (!bfqq)
return;
bfqd = bfqq->bfqd;
/* Reposition request in its sort_list */
elv_rb_del(&bfqq->sort_list, req);
elv_rb_add(&bfqq->sort_list, req);
/* Choose next request to be served for bfqq */
prev = bfqq->next_rq;
next_rq = bfq_choose_req(bfqd, bfqq->next_rq, req,
bfqd->last_position);
bfqq->next_rq = next_rq;
/*
* If next_rq changes, update both the queue's budget to
* fit the new request and the queue's position in its
* rq_pos_tree.
*/
if (prev != bfqq->next_rq) {
bfq_updated_next_req(bfqd, bfqq);
/*
* See comments on bfq_pos_tree_add_move() for
* the unlikely().
*/
if (unlikely(!bfqd->nonrot_with_queueing))
bfq_pos_tree_add_move(bfqd, bfqq);
}
}
}
/*
* This function is called to notify the scheduler that the requests
* rq and 'next' have been merged, with 'next' going away. BFQ
* exploits this hook to address the following issue: if 'next' has a
* fifo_time lower that rq, then the fifo_time of rq must be set to
* the value of 'next', to not forget the greater age of 'next'.
*
* NOTE: in this function we assume that rq is in a bfq_queue, basing
* on that rq is picked from the hash table q->elevator->hash, which,
* in its turn, is filled only with I/O requests present in
* bfq_queues, while BFQ is in use for the request queue q. In fact,
* the function that fills this hash table (elv_rqhash_add) is called
* only by bfq_insert_request.
*/
static void bfq_requests_merged(struct request_queue *q, struct request *rq,
struct request *next)
{
struct bfq_queue *bfqq = bfq_init_rq(rq),
*next_bfqq = bfq_init_rq(next);
if (!bfqq)
return;
/*
* If next and rq belong to the same bfq_queue and next is older
* than rq, then reposition rq in the fifo (by substituting next
* with rq). Otherwise, if next and rq belong to different
* bfq_queues, never reposition rq: in fact, we would have to
* reposition it with respect to next's position in its own fifo,
* which would most certainly be too expensive with respect to
* the benefits.
*/
if (bfqq == next_bfqq &&
!list_empty(&rq->queuelist) && !list_empty(&next->queuelist) &&
next->fifo_time < rq->fifo_time) {
list_del_init(&rq->queuelist);
list_replace_init(&next->queuelist, &rq->queuelist);
rq->fifo_time = next->fifo_time;
}
if (bfqq->next_rq == next)
bfqq->next_rq = rq;
bfqg_stats_update_io_merged(bfqq_group(bfqq), next->cmd_flags);
}
/* Must be called with bfqq != NULL */
static void bfq_bfqq_end_wr(struct bfq_queue *bfqq)
{
if (bfq_bfqq_busy(bfqq))
bfqq->bfqd->wr_busy_queues--;
bfqq->wr_coeff = 1;
bfqq->wr_cur_max_time = 0;
bfqq->last_wr_start_finish = jiffies;
/*
* Trigger a weight change on the next invocation of
* __bfq_entity_update_weight_prio.
*/
bfqq->entity.prio_changed = 1;
}
void bfq_end_wr_async_queues(struct bfq_data *bfqd,
struct bfq_group *bfqg)
{
int i, j;
for (i = 0; i < 2; i++)
for (j = 0; j < IOPRIO_BE_NR; j++)
if (bfqg->async_bfqq[i][j])
bfq_bfqq_end_wr(bfqg->async_bfqq[i][j]);
if (bfqg->async_idle_bfqq)
bfq_bfqq_end_wr(bfqg->async_idle_bfqq);
}
static void bfq_end_wr(struct bfq_data *bfqd)
{
struct bfq_queue *bfqq;
spin_lock_irq(&bfqd->lock);
list_for_each_entry(bfqq, &bfqd->active_list, bfqq_list)
bfq_bfqq_end_wr(bfqq);
list_for_each_entry(bfqq, &bfqd->idle_list, bfqq_list)
bfq_bfqq_end_wr(bfqq);
bfq_end_wr_async(bfqd);
spin_unlock_irq(&bfqd->lock);
}
static sector_t bfq_io_struct_pos(void *io_struct, bool request)
{
if (request)
return blk_rq_pos(io_struct);
else
return ((struct bio *)io_struct)->bi_iter.bi_sector;
}
static int bfq_rq_close_to_sector(void *io_struct, bool request,
sector_t sector)
{
return abs(bfq_io_struct_pos(io_struct, request) - sector) <=
BFQQ_CLOSE_THR;
}
static struct bfq_queue *bfqq_find_close(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
sector_t sector)
{
struct rb_root *root = &bfq_bfqq_to_bfqg(bfqq)->rq_pos_tree;
struct rb_node *parent, *node;
struct bfq_queue *__bfqq;
if (RB_EMPTY_ROOT(root))
return NULL;
/*
* First, if we find a request starting at the end of the last
* request, choose it.
*/
__bfqq = bfq_rq_pos_tree_lookup(bfqd, root, sector, &parent, NULL);
if (__bfqq)
return __bfqq;
/*
* If the exact sector wasn't found, the parent of the NULL leaf
* will contain the closest sector (rq_pos_tree sorted by
* next_request position).
*/
__bfqq = rb_entry(parent, struct bfq_queue, pos_node);
if (bfq_rq_close_to_sector(__bfqq->next_rq, true, sector))
return __bfqq;
if (blk_rq_pos(__bfqq->next_rq) < sector)
node = rb_next(&__bfqq->pos_node);
else
node = rb_prev(&__bfqq->pos_node);
if (!node)
return NULL;
__bfqq = rb_entry(node, struct bfq_queue, pos_node);
if (bfq_rq_close_to_sector(__bfqq->next_rq, true, sector))
return __bfqq;
return NULL;
}
static struct bfq_queue *bfq_find_close_cooperator(struct bfq_data *bfqd,
struct bfq_queue *cur_bfqq,
sector_t sector)
{
struct bfq_queue *bfqq;
/*
* We shall notice if some of the queues are cooperating,
* e.g., working closely on the same area of the device. In
* that case, we can group them together and: 1) don't waste
* time idling, and 2) serve the union of their requests in
* the best possible order for throughput.
*/
bfqq = bfqq_find_close(bfqd, cur_bfqq, sector);
if (!bfqq || bfqq == cur_bfqq)
return NULL;
return bfqq;
}
static struct bfq_queue *
bfq_setup_merge(struct bfq_queue *bfqq, struct bfq_queue *new_bfqq)
{
int process_refs, new_process_refs;
struct bfq_queue *__bfqq;
/*
* If there are no process references on the new_bfqq, then it is
* unsafe to follow the ->new_bfqq chain as other bfqq's in the chain
* may have dropped their last reference (not just their last process
* reference).
*/
if (!bfqq_process_refs(new_bfqq))
return NULL;
/* Avoid a circular list and skip interim queue merges. */
while ((__bfqq = new_bfqq->new_bfqq)) {
if (__bfqq == bfqq)
return NULL;
new_bfqq = __bfqq;
}
process_refs = bfqq_process_refs(bfqq);
new_process_refs = bfqq_process_refs(new_bfqq);
/*
* If the process for the bfqq has gone away, there is no
* sense in merging the queues.
*/
if (process_refs == 0 || new_process_refs == 0)
return NULL;
bfq_log_bfqq(bfqq->bfqd, bfqq, "scheduling merge with queue %d",
new_bfqq->pid);
/*
* Merging is just a redirection: the requests of the process
* owning one of the two queues are redirected to the other queue.
* The latter queue, in its turn, is set as shared if this is the
* first time that the requests of some process are redirected to
* it.
*
* We redirect bfqq to new_bfqq and not the opposite, because
* we are in the context of the process owning bfqq, thus we
* have the io_cq of this process. So we can immediately
* configure this io_cq to redirect the requests of the
* process to new_bfqq. In contrast, the io_cq of new_bfqq is
* not available any more (new_bfqq->bic == NULL).
*
* Anyway, even in case new_bfqq coincides with the in-service
* queue, redirecting requests the in-service queue is the
* best option, as we feed the in-service queue with new
* requests close to the last request served and, by doing so,
* are likely to increase the throughput.
*/
bfqq->new_bfqq = new_bfqq;
new_bfqq->ref += process_refs;
return new_bfqq;
}
static bool bfq_may_be_close_cooperator(struct bfq_queue *bfqq,
struct bfq_queue *new_bfqq)
{
if (bfq_too_late_for_merging(new_bfqq))
return false;
if (bfq_class_idle(bfqq) || bfq_class_idle(new_bfqq) ||
(bfqq->ioprio_class != new_bfqq->ioprio_class))
return false;
/*
* If either of the queues has already been detected as seeky,
* then merging it with the other queue is unlikely to lead to
* sequential I/O.
*/
if (BFQQ_SEEKY(bfqq) || BFQQ_SEEKY(new_bfqq))
return false;
/*
* Interleaved I/O is known to be done by (some) applications
* only for reads, so it does not make sense to merge async
* queues.
*/
if (!bfq_bfqq_sync(bfqq) || !bfq_bfqq_sync(new_bfqq))
return false;
return true;
}
/*
* Attempt to schedule a merge of bfqq with the currently in-service
* queue or with a close queue among the scheduled queues. Return
* NULL if no merge was scheduled, a pointer to the shared bfq_queue
* structure otherwise.
*
* The OOM queue is not allowed to participate to cooperation: in fact, since
* the requests temporarily redirected to the OOM queue could be redirected
* again to dedicated queues at any time, the state needed to correctly
* handle merging with the OOM queue would be quite complex and expensive
* to maintain. Besides, in such a critical condition as an out of memory,
* the benefits of queue merging may be little relevant, or even negligible.
*
* WARNING: queue merging may impair fairness among non-weight raised
* queues, for at least two reasons: 1) the original weight of a
* merged queue may change during the merged state, 2) even being the
* weight the same, a merged queue may be bloated with many more
* requests than the ones produced by its originally-associated
* process.
*/
static struct bfq_queue *
bfq_setup_cooperator(struct bfq_data *bfqd, struct bfq_queue *bfqq,
void *io_struct, bool request)
{
struct bfq_queue *in_service_bfqq, *new_bfqq;
/*
* Do not perform queue merging if the device is non
* rotational and performs internal queueing. In fact, such a
* device reaches a high speed through internal parallelism
* and pipelining. This means that, to reach a high
* throughput, it must have many requests enqueued at the same
* time. But, in this configuration, the internal scheduling
* algorithm of the device does exactly the job of queue
* merging: it reorders requests so as to obtain as much as
* possible a sequential I/O pattern. As a consequence, with
* the workload generated by processes doing interleaved I/O,
* the throughput reached by the device is likely to be the
* same, with and without queue merging.
*
* Disabling merging also provides a remarkable benefit in
* terms of throughput. Merging tends to make many workloads
* artificially more uneven, because of shared queues
* remaining non empty for incomparably more time than
* non-merged queues. This may accentuate workload
* asymmetries. For example, if one of the queues in a set of
* merged queues has a higher weight than a normal queue, then
* the shared queue may inherit such a high weight and, by
* staying almost always active, may force BFQ to perform I/O
* plugging most of the time. This evidently makes it harder
* for BFQ to let the device reach a high throughput.
*
* Finally, the likely() macro below is not used because one
* of the two branches is more likely than the other, but to
* have the code path after the following if() executed as
* fast as possible for the case of a non rotational device
* with queueing. We want it because this is the fastest kind
* of device. On the opposite end, the likely() may lengthen
* the execution time of BFQ for the case of slower devices
* (rotational or at least without queueing). But in this case
* the execution time of BFQ matters very little, if not at
* all.
*/
if (likely(bfqd->nonrot_with_queueing))
return NULL;
/*
* Prevent bfqq from being merged if it has been created too
* long ago. The idea is that true cooperating processes, and
* thus their associated bfq_queues, are supposed to be
* created shortly after each other. This is the case, e.g.,
* for KVM/QEMU and dump I/O threads. Basing on this
* assumption, the following filtering greatly reduces the
* probability that two non-cooperating processes, which just
* happen to do close I/O for some short time interval, have
* their queues merged by mistake.
*/
if (bfq_too_late_for_merging(bfqq))
return NULL;
if (bfqq->new_bfqq)
return bfqq->new_bfqq;
if (!io_struct || unlikely(bfqq == &bfqd->oom_bfqq))
return NULL;
/* If there is only one backlogged queue, don't search. */
if (bfq_tot_busy_queues(bfqd) == 1)
return NULL;
in_service_bfqq = bfqd->in_service_queue;
if (in_service_bfqq && in_service_bfqq != bfqq &&
likely(in_service_bfqq != &bfqd->oom_bfqq) &&
bfq_rq_close_to_sector(io_struct, request,
bfqd->in_serv_last_pos) &&
bfqq->entity.parent == in_service_bfqq->entity.parent &&
bfq_may_be_close_cooperator(bfqq, in_service_bfqq)) {
new_bfqq = bfq_setup_merge(bfqq, in_service_bfqq);
if (new_bfqq)
return new_bfqq;
}
/*
* Check whether there is a cooperator among currently scheduled
* queues. The only thing we need is that the bio/request is not
* NULL, as we need it to establish whether a cooperator exists.
*/
new_bfqq = bfq_find_close_cooperator(bfqd, bfqq,
bfq_io_struct_pos(io_struct, request));
if (new_bfqq && likely(new_bfqq != &bfqd->oom_bfqq) &&
bfq_may_be_close_cooperator(bfqq, new_bfqq))
return bfq_setup_merge(bfqq, new_bfqq);
return NULL;
}
static void bfq_bfqq_save_state(struct bfq_queue *bfqq)
{
struct bfq_io_cq *bic = bfqq->bic;
/*
* If !bfqq->bic, the queue is already shared or its requests
* have already been redirected to a shared queue; both idle window
* and weight raising state have already been saved. Do nothing.
*/
if (!bic)
return;
bic->saved_weight = bfqq->entity.orig_weight;
bic->saved_ttime = bfqq->ttime;
bic->saved_has_short_ttime = bfq_bfqq_has_short_ttime(bfqq);
bic->saved_IO_bound = bfq_bfqq_IO_bound(bfqq);
bic->saved_in_large_burst = bfq_bfqq_in_large_burst(bfqq);
bic->was_in_burst_list = !hlist_unhashed(&bfqq->burst_list_node);
if (unlikely(bfq_bfqq_just_created(bfqq) &&
!bfq_bfqq_in_large_burst(bfqq) &&
bfqq->bfqd->low_latency)) {
/*
* bfqq being merged right after being created: bfqq
* would have deserved interactive weight raising, but
* did not make it to be set in a weight-raised state,
* because of this early merge. Store directly the
* weight-raising state that would have been assigned
* to bfqq, so that to avoid that bfqq unjustly fails
* to enjoy weight raising if split soon.
*/
bic->saved_wr_coeff = bfqq->bfqd->bfq_wr_coeff;
bic->saved_wr_start_at_switch_to_srt = bfq_smallest_from_now();
bic->saved_wr_cur_max_time = bfq_wr_duration(bfqq->bfqd);
bic->saved_last_wr_start_finish = jiffies;
} else {
bic->saved_wr_coeff = bfqq->wr_coeff;
bic->saved_wr_start_at_switch_to_srt =
bfqq->wr_start_at_switch_to_srt;
bic->saved_last_wr_start_finish = bfqq->last_wr_start_finish;
bic->saved_wr_cur_max_time = bfqq->wr_cur_max_time;
}
}
static void
bfq_merge_bfqqs(struct bfq_data *bfqd, struct bfq_io_cq *bic,
struct bfq_queue *bfqq, struct bfq_queue *new_bfqq)
{
bfq_log_bfqq(bfqd, bfqq, "merging with queue %lu",
(unsigned long)new_bfqq->pid);
/* Save weight raising and idle window of the merged queues */
bfq_bfqq_save_state(bfqq);
bfq_bfqq_save_state(new_bfqq);
if (bfq_bfqq_IO_bound(bfqq))
bfq_mark_bfqq_IO_bound(new_bfqq);
bfq_clear_bfqq_IO_bound(bfqq);
/*
* If bfqq is weight-raised, then let new_bfqq inherit
* weight-raising. To reduce false positives, neglect the case
* where bfqq has just been created, but has not yet made it
* to be weight-raised (which may happen because EQM may merge
* bfqq even before bfq_add_request is executed for the first
* time for bfqq). Handling this case would however be very
* easy, thanks to the flag just_created.
*/
if (new_bfqq->wr_coeff == 1 && bfqq->wr_coeff > 1) {
new_bfqq->wr_coeff = bfqq->wr_coeff;
new_bfqq->wr_cur_max_time = bfqq->wr_cur_max_time;
new_bfqq->last_wr_start_finish = bfqq->last_wr_start_finish;
new_bfqq->wr_start_at_switch_to_srt =
bfqq->wr_start_at_switch_to_srt;
if (bfq_bfqq_busy(new_bfqq))
bfqd->wr_busy_queues++;
new_bfqq->entity.prio_changed = 1;
}
if (bfqq->wr_coeff > 1) { /* bfqq has given its wr to new_bfqq */
bfqq->wr_coeff = 1;
bfqq->entity.prio_changed = 1;
if (bfq_bfqq_busy(bfqq))
bfqd->wr_busy_queues--;
}
bfq_log_bfqq(bfqd, new_bfqq, "merge_bfqqs: wr_busy %d",
bfqd->wr_busy_queues);
/*
* Merge queues (that is, let bic redirect its requests to new_bfqq)
*/
bic_set_bfqq(bic, new_bfqq, 1);
bfq_mark_bfqq_coop(new_bfqq);
/*
* new_bfqq now belongs to at least two bics (it is a shared queue):
* set new_bfqq->bic to NULL. bfqq either:
* - does not belong to any bic any more, and hence bfqq->bic must
* be set to NULL, or
* - is a queue whose owning bics have already been redirected to a
* different queue, hence the queue is destined to not belong to
* any bic soon and bfqq->bic is already NULL (therefore the next
* assignment causes no harm).
*/
new_bfqq->bic = NULL;
/*
* If the queue is shared, the pid is the pid of one of the associated
* processes. Which pid depends on the exact sequence of merge events
* the queue underwent. So printing such a pid is useless and confusing
* because it reports a random pid between those of the associated
* processes.
* We mark such a queue with a pid -1, and then print SHARED instead of
* a pid in logging messages.
*/
new_bfqq->pid = -1;
bfqq->bic = NULL;
/* release process reference to bfqq */
bfq_put_queue(bfqq);
}
static bool bfq_allow_bio_merge(struct request_queue *q, struct request *rq,
struct bio *bio)
{
struct bfq_data *bfqd = q->elevator->elevator_data;
bool is_sync = op_is_sync(bio->bi_opf);
struct bfq_queue *bfqq = bfqd->bio_bfqq, *new_bfqq;
/*
* Disallow merge of a sync bio into an async request.
*/
if (is_sync && !rq_is_sync(rq))
return false;
/*
* Lookup the bfqq that this bio will be queued with. Allow
* merge only if rq is queued there.
*/
if (!bfqq)
return false;
/*
* We take advantage of this function to perform an early merge
* of the queues of possible cooperating processes.
*/
new_bfqq = bfq_setup_cooperator(bfqd, bfqq, bio, false);
if (new_bfqq) {
/*
* bic still points to bfqq, then it has not yet been
* redirected to some other bfq_queue, and a queue
* merge between bfqq and new_bfqq can be safely
* fulfilled, i.e., bic can be redirected to new_bfqq
* and bfqq can be put.
*/
bfq_merge_bfqqs(bfqd, bfqd->bio_bic, bfqq,
new_bfqq);
/*
* If we get here, bio will be queued into new_queue,
* so use new_bfqq to decide whether bio and rq can be
* merged.
*/
bfqq = new_bfqq;
/*
* Change also bqfd->bio_bfqq, as
* bfqd->bio_bic now points to new_bfqq, and
* this function may be invoked again (and then may
* use again bqfd->bio_bfqq).
*/
bfqd->bio_bfqq = bfqq;
}
return bfqq == RQ_BFQQ(rq);
}
/*
* Set the maximum time for the in-service queue to consume its
* budget. This prevents seeky processes from lowering the throughput.
* In practice, a time-slice service scheme is used with seeky
* processes.
*/
static void bfq_set_budget_timeout(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
unsigned int timeout_coeff;
if (bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time)
timeout_coeff = 1;
else
timeout_coeff = bfqq->entity.weight / bfqq->entity.orig_weight;
bfqd->last_budget_start = ktime_get();
bfqq->budget_timeout = jiffies +
bfqd->bfq_timeout * timeout_coeff;
}
static void __bfq_set_in_service_queue(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
if (bfqq) {
bfq_clear_bfqq_fifo_expire(bfqq);
bfqd->budgets_assigned = (bfqd->budgets_assigned * 7 + 256) / 8;
if (time_is_before_jiffies(bfqq->last_wr_start_finish) &&
bfqq->wr_coeff > 1 &&
bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time &&
time_is_before_jiffies(bfqq->budget_timeout)) {
/*
* For soft real-time queues, move the start
* of the weight-raising period forward by the
* time the queue has not received any
* service. Otherwise, a relatively long
* service delay is likely to cause the
* weight-raising period of the queue to end,
* because of the short duration of the
* weight-raising period of a soft real-time
* queue. It is worth noting that this move
* is not so dangerous for the other queues,
* because soft real-time queues are not
* greedy.
*
* To not add a further variable, we use the
* overloaded field budget_timeout to
* determine for how long the queue has not
* received service, i.e., how much time has
* elapsed since the queue expired. However,
* this is a little imprecise, because
* budget_timeout is set to jiffies if bfqq
* not only expires, but also remains with no
* request.
*/
if (time_after(bfqq->budget_timeout,
bfqq->last_wr_start_finish))
bfqq->last_wr_start_finish +=
jiffies - bfqq->budget_timeout;
else
bfqq->last_wr_start_finish = jiffies;
}
bfq_set_budget_timeout(bfqd, bfqq);
bfq_log_bfqq(bfqd, bfqq,
"set_in_service_queue, cur-budget = %d",
bfqq->entity.budget);
}
bfqd->in_service_queue = bfqq;
}
/*
* Get and set a new queue for service.
*/
static struct bfq_queue *bfq_set_in_service_queue(struct bfq_data *bfqd)
{
struct bfq_queue *bfqq = bfq_get_next_queue(bfqd);
__bfq_set_in_service_queue(bfqd, bfqq);
return bfqq;
}
static void bfq_arm_slice_timer(struct bfq_data *bfqd)
{
struct bfq_queue *bfqq = bfqd->in_service_queue;
u32 sl;
bfq_mark_bfqq_wait_request(bfqq);
/*
* We don't want to idle for seeks, but we do want to allow
* fair distribution of slice time for a process doing back-to-back
* seeks. So allow a little bit of time for him to submit a new rq.
*/
sl = bfqd->bfq_slice_idle;
/*
* Unless the queue is being weight-raised or the scenario is
* asymmetric, grant only minimum idle time if the queue
* is seeky. A long idling is preserved for a weight-raised
* queue, or, more in general, in an asymmetric scenario,
* because a long idling is needed for guaranteeing to a queue
* its reserved share of the throughput (in particular, it is
* needed if the queue has a higher weight than some other
* queue).
*/
if (BFQQ_SEEKY(bfqq) && bfqq->wr_coeff == 1 &&
!bfq_asymmetric_scenario(bfqd, bfqq))
sl = min_t(u64, sl, BFQ_MIN_TT);
else if (bfqq->wr_coeff > 1)
sl = max_t(u32, sl, 20ULL * NSEC_PER_MSEC);
bfqd->last_idling_start = ktime_get();
bfqd->last_idling_start_jiffies = jiffies;
hrtimer_start(&bfqd->idle_slice_timer, ns_to_ktime(sl),
HRTIMER_MODE_REL);
bfqg_stats_set_start_idle_time(bfqq_group(bfqq));
}
/*
* In autotuning mode, max_budget is dynamically recomputed as the
* amount of sectors transferred in timeout at the estimated peak
* rate. This enables BFQ to utilize a full timeslice with a full
* budget, even if the in-service queue is served at peak rate. And
* this maximises throughput with sequential workloads.
*/
static unsigned long bfq_calc_max_budget(struct bfq_data *bfqd)
{
return (u64)bfqd->peak_rate * USEC_PER_MSEC *
jiffies_to_msecs(bfqd->bfq_timeout)>>BFQ_RATE_SHIFT;
}
/*
* Update parameters related to throughput and responsiveness, as a
* function of the estimated peak rate. See comments on
* bfq_calc_max_budget(), and on the ref_wr_duration array.
*/
static void update_thr_responsiveness_params(struct bfq_data *bfqd)
{
if (bfqd->bfq_user_max_budget == 0) {
bfqd->bfq_max_budget =
bfq_calc_max_budget(bfqd);
bfq_log(bfqd, "new max_budget = %d", bfqd->bfq_max_budget);
}
}
static void bfq_reset_rate_computation(struct bfq_data *bfqd,
struct request *rq)
{
if (rq != NULL) { /* new rq dispatch now, reset accordingly */
bfqd->last_dispatch = bfqd->first_dispatch = ktime_get_ns();
bfqd->peak_rate_samples = 1;
bfqd->sequential_samples = 0;
bfqd->tot_sectors_dispatched = bfqd->last_rq_max_size =
blk_rq_sectors(rq);
} else /* no new rq dispatched, just reset the number of samples */
bfqd->peak_rate_samples = 0; /* full re-init on next disp. */
bfq_log(bfqd,
"reset_rate_computation at end, sample %u/%u tot_sects %llu",
bfqd->peak_rate_samples, bfqd->sequential_samples,
bfqd->tot_sectors_dispatched);
}
static void bfq_update_rate_reset(struct bfq_data *bfqd, struct request *rq)
{
u32 rate, weight, divisor;
/*
* For the convergence property to hold (see comments on
* bfq_update_peak_rate()) and for the assessment to be
* reliable, a minimum number of samples must be present, and
* a minimum amount of time must have elapsed. If not so, do
* not compute new rate. Just reset parameters, to get ready
* for a new evaluation attempt.
*/
if (bfqd->peak_rate_samples < BFQ_RATE_MIN_SAMPLES ||
bfqd->delta_from_first < BFQ_RATE_MIN_INTERVAL)
goto reset_computation;
/*
* If a new request completion has occurred after last
* dispatch, then, to approximate the rate at which requests
* have been served by the device, it is more precise to
* extend the observation interval to the last completion.
*/
bfqd->delta_from_first =
max_t(u64, bfqd->delta_from_first,
bfqd->last_completion - bfqd->first_dispatch);
/*
* Rate computed in sects/usec, and not sects/nsec, for
* precision issues.
*/
rate = div64_ul(bfqd->tot_sectors_dispatched<<BFQ_RATE_SHIFT,
div_u64(bfqd->delta_from_first, NSEC_PER_USEC));
/*
* Peak rate not updated if:
* - the percentage of sequential dispatches is below 3/4 of the
* total, and rate is below the current estimated peak rate
* - rate is unreasonably high (> 20M sectors/sec)
*/
if ((bfqd->sequential_samples < (3 * bfqd->peak_rate_samples)>>2 &&
rate <= bfqd->peak_rate) ||
rate > 20<<BFQ_RATE_SHIFT)
goto reset_computation;
/*
* We have to update the peak rate, at last! To this purpose,
* we use a low-pass filter. We compute the smoothing constant
* of the filter as a function of the 'weight' of the new
* measured rate.
*
* As can be seen in next formulas, we define this weight as a
* quantity proportional to how sequential the workload is,
* and to how long the observation time interval is.
*
* The weight runs from 0 to 8. The maximum value of the
* weight, 8, yields the minimum value for the smoothing
* constant. At this minimum value for the smoothing constant,
* the measured rate contributes for half of the next value of
* the estimated peak rate.
*
* So, the first step is to compute the weight as a function
* of how sequential the workload is. Note that the weight
* cannot reach 9, because bfqd->sequential_samples cannot
* become equal to bfqd->peak_rate_samples, which, in its
* turn, holds true because bfqd->sequential_samples is not
* incremented for the first sample.
*/
weight = (9 * bfqd->sequential_samples) / bfqd->peak_rate_samples;
/*
* Second step: further refine the weight as a function of the
* duration of the observation interval.
*/
weight = min_t(u32, 8,
div_u64(weight * bfqd->delta_from_first,
BFQ_RATE_REF_INTERVAL));
/*
* Divisor ranging from 10, for minimum weight, to 2, for
* maximum weight.
*/
divisor = 10 - weight;
/*
* Finally, update peak rate:
*
* peak_rate = peak_rate * (divisor-1) / divisor + rate / divisor
*/
bfqd->peak_rate *= divisor-1;
bfqd->peak_rate /= divisor;
rate /= divisor; /* smoothing constant alpha = 1/divisor */
bfqd->peak_rate += rate;
/*
* For a very slow device, bfqd->peak_rate can reach 0 (see
* the minimum representable values reported in the comments
* on BFQ_RATE_SHIFT). Push to 1 if this happens, to avoid
* divisions by zero where bfqd->peak_rate is used as a
* divisor.
*/
bfqd->peak_rate = max_t(u32, 1, bfqd->peak_rate);
update_thr_responsiveness_params(bfqd);
reset_computation:
bfq_reset_rate_computation(bfqd, rq);
}
/*
* Update the read/write peak rate (the main quantity used for
* auto-tuning, see update_thr_responsiveness_params()).
*
* It is not trivial to estimate the peak rate (correctly): because of
* the presence of sw and hw queues between the scheduler and the
* device components that finally serve I/O requests, it is hard to
* say exactly when a given dispatched request is served inside the
* device, and for how long. As a consequence, it is hard to know
* precisely at what rate a given set of requests is actually served
* by the device.
*
* On the opposite end, the dispatch time of any request is trivially
* available, and, from this piece of information, the "dispatch rate"
* of requests can be immediately computed. So, the idea in the next
* function is to use what is known, namely request dispatch times
* (plus, when useful, request completion times), to estimate what is
* unknown, namely in-device request service rate.
*
* The main issue is that, because of the above facts, the rate at
* which a certain set of requests is dispatched over a certain time
* interval can vary greatly with respect to the rate at which the
* same requests are then served. But, since the size of any
* intermediate queue is limited, and the service scheme is lossless
* (no request is silently dropped), the following obvious convergence
* property holds: the number of requests dispatched MUST become
* closer and closer to the number of requests completed as the
* observation interval grows. This is the key property used in
* the next function to estimate the peak service rate as a function
* of the observed dispatch rate. The function assumes to be invoked
* on every request dispatch.
*/
static void bfq_update_peak_rate(struct bfq_data *bfqd, struct request *rq)
{
u64 now_ns = ktime_get_ns();
if (bfqd->peak_rate_samples == 0) { /* first dispatch */
bfq_log(bfqd, "update_peak_rate: goto reset, samples %d",
bfqd->peak_rate_samples);
bfq_reset_rate_computation(bfqd, rq);
goto update_last_values; /* will add one sample */
}
/*
* Device idle for very long: the observation interval lasting
* up to this dispatch cannot be a valid observation interval
* for computing a new peak rate (similarly to the late-
* completion event in bfq_completed_request()). Go to
* update_rate_and_reset to have the following three steps
* taken:
* - close the observation interval at the last (previous)
* request dispatch or completion
* - compute rate, if possible, for that observation interval
* - start a new observation interval with this dispatch
*/
if (now_ns - bfqd->last_dispatch > 100*NSEC_PER_MSEC &&
bfqd->rq_in_driver == 0)
goto update_rate_and_reset;
/* Update sampling information */
bfqd->peak_rate_samples++;
if ((bfqd->rq_in_driver > 0 ||
now_ns - bfqd->last_completion < BFQ_MIN_TT)
&& !BFQ_RQ_SEEKY(bfqd, bfqd->last_position, rq))
bfqd->sequential_samples++;
bfqd->tot_sectors_dispatched += blk_rq_sectors(rq);
/* Reset max observed rq size every 32 dispatches */
if (likely(bfqd->peak_rate_samples % 32))
bfqd->last_rq_max_size =
max_t(u32, blk_rq_sectors(rq), bfqd->last_rq_max_size);
else
bfqd->last_rq_max_size = blk_rq_sectors(rq);
bfqd->delta_from_first = now_ns - bfqd->first_dispatch;
/* Target observation interval not yet reached, go on sampling */
if (bfqd->delta_from_first < BFQ_RATE_REF_INTERVAL)
goto update_last_values;
update_rate_and_reset:
bfq_update_rate_reset(bfqd, rq);
update_last_values:
bfqd->last_position = blk_rq_pos(rq) + blk_rq_sectors(rq);
if (RQ_BFQQ(rq) == bfqd->in_service_queue)
bfqd->in_serv_last_pos = bfqd->last_position;
bfqd->last_dispatch = now_ns;
}
/*
* Remove request from internal lists.
*/
static void bfq_dispatch_remove(struct request_queue *q, struct request *rq)
{
struct bfq_queue *bfqq = RQ_BFQQ(rq);
/*
* For consistency, the next instruction should have been
* executed after removing the request from the queue and
* dispatching it. We execute instead this instruction before
* bfq_remove_request() (and hence introduce a temporary
* inconsistency), for efficiency. In fact, should this
* dispatch occur for a non in-service bfqq, this anticipated
* increment prevents two counters related to bfqq->dispatched
* from risking to be, first, uselessly decremented, and then
* incremented again when the (new) value of bfqq->dispatched
* happens to be taken into account.
*/
bfqq->dispatched++;
bfq_update_peak_rate(q->elevator->elevator_data, rq);
bfq_remove_request(q, rq);
}
/*
* There is a case where idling does not have to be performed for
* throughput concerns, but to preserve the throughput share of
* the process associated with bfqq.
*
* To introduce this case, we can note that allowing the drive
* to enqueue more than one request at a time, and hence
* delegating de facto final scheduling decisions to the
* drive's internal scheduler, entails loss of control on the
* actual request service order. In particular, the critical
* situation is when requests from different processes happen
* to be present, at the same time, in the internal queue(s)
* of the drive. In such a situation, the drive, by deciding
* the service order of the internally-queued requests, does
* determine also the actual throughput distribution among
* these processes. But the drive typically has no notion or
* concern about per-process throughput distribution, and
* makes its decisions only on a per-request basis. Therefore,
* the service distribution enforced by the drive's internal
* scheduler is likely to coincide with the desired throughput
* distribution only in a completely symmetric, or favorably
* skewed scenario where:
* (i-a) each of these processes must get the same throughput as
* the others,
* (i-b) in case (i-a) does not hold, it holds that the process
* associated with bfqq must receive a lower or equal
* throughput than any of the other processes;
* (ii) the I/O of each process has the same properties, in
* terms of locality (sequential or random), direction
* (reads or writes), request sizes, greediness
* (from I/O-bound to sporadic), and so on;
* In fact, in such a scenario, the drive tends to treat the requests
* of each process in about the same way as the requests of the
* others, and thus to provide each of these processes with about the
* same throughput. This is exactly the desired throughput
* distribution if (i-a) holds, or, if (i-b) holds instead, this is an
* even more convenient distribution for (the process associated with)
* bfqq.
*
* In contrast, in any asymmetric or unfavorable scenario, device
* idling (I/O-dispatch plugging) is certainly needed to guarantee
* that bfqq receives its assigned fraction of the device throughput
* (see [1] for details).
*
* The problem is that idling may significantly reduce throughput with
* certain combinations of types of I/O and devices. An important
* example is sync random I/O on flash storage with command
* queueing. So, unless bfqq falls in cases where idling also boosts
* throughput, it is important to check conditions (i-a), i(-b) and
* (ii) accurately, so as to avoid idling when not strictly needed for
* service guarantees.
*
* Unfortunately, it is extremely difficult to thoroughly check
* condition (ii). And, in case there are active groups, it becomes
* very difficult to check conditions (i-a) and (i-b) too. In fact,
* if there are active groups, then, for conditions (i-a) or (i-b) to
* become false 'indirectly', it is enough that an active group
* contains more active processes or sub-groups than some other active
* group. More precisely, for conditions (i-a) or (i-b) to become
* false because of such a group, it is not even necessary that the
* group is (still) active: it is sufficient that, even if the group
* has become inactive, some of its descendant processes still have
* some request already dispatched but still waiting for
* completion. In fact, requests have still to be guaranteed their
* share of the throughput even after being dispatched. In this
* respect, it is easy to show that, if a group frequently becomes
* inactive while still having in-flight requests, and if, when this
* happens, the group is not considered in the calculation of whether
* the scenario is asymmetric, then the group may fail to be
* guaranteed its fair share of the throughput (basically because
* idling may not be performed for the descendant processes of the
* group, but it had to be). We address this issue with the following
* bi-modal behavior, implemented in the function
* bfq_asymmetric_scenario().
*
* If there are groups with requests waiting for completion
* (as commented above, some of these groups may even be
* already inactive), then the scenario is tagged as
* asymmetric, conservatively, without checking any of the
* conditions (i-a), (i-b) or (ii). So the device is idled for bfqq.
* This behavior matches also the fact that groups are created
* exactly if controlling I/O is a primary concern (to
* preserve bandwidth and latency guarantees).
*
* On the opposite end, if there are no groups with requests waiting
* for completion, then only conditions (i-a) and (i-b) are actually
* controlled, i.e., provided that conditions (i-a) or (i-b) holds,
* idling is not performed, regardless of whether condition (ii)
* holds. In other words, only if conditions (i-a) and (i-b) do not
* hold, then idling is allowed, and the device tends to be prevented
* from queueing many requests, possibly of several processes. Since
* there are no groups with requests waiting for completion, then, to
* control conditions (i-a) and (i-b) it is enough to check just
* whether all the queues with requests waiting for completion also
* have the same weight.
*
* Not checking condition (ii) evidently exposes bfqq to the
* risk of getting less throughput than its fair share.
* However, for queues with the same weight, a further
* mechanism, preemption, mitigates or even eliminates this
* problem. And it does so without consequences on overall
* throughput. This mechanism and its benefits are explained
* in the next three paragraphs.
*
* Even if a queue, say Q, is expired when it remains idle, Q
* can still preempt the new in-service queue if the next
* request of Q arrives soon (see the comments on
* bfq_bfqq_update_budg_for_activation). If all queues and
* groups have the same weight, this form of preemption,
* combined with the hole-recovery heuristic described in the
* comments on function bfq_bfqq_update_budg_for_activation,
* are enough to preserve a correct bandwidth distribution in
* the mid term, even without idling. In fact, even if not
* idling allows the internal queues of the device to contain
* many requests, and thus to reorder requests, we can rather
* safely assume that the internal scheduler still preserves a
* minimum of mid-term fairness.
*
* More precisely, this preemption-based, idleless approach
* provides fairness in terms of IOPS, and not sectors per
* second. This can be seen with a simple example. Suppose
* that there are two queues with the same weight, but that
* the first queue receives requests of 8 sectors, while the
* second queue receives requests of 1024 sectors. In
* addition, suppose that each of the two queues contains at
* most one request at a time, which implies that each queue
* always remains idle after it is served. Finally, after
* remaining idle, each queue receives very quickly a new
* request. It follows that the two queues are served
* alternatively, preempting each other if needed. This
* implies that, although both queues have the same weight,
* the queue with large requests receives a service that is
* 1024/8 times as high as the service received by the other
* queue.
*
* The motivation for using preemption instead of idling (for
* queues with the same weight) is that, by not idling,
* service guarantees are preserved (completely or at least in
* part) without minimally sacrificing throughput. And, if
* there is no active group, then the primary expectation for
* this device is probably a high throughput.
*
* We are now left only with explaining the two sub-conditions in the
* additional compound condition that is checked below for deciding
* whether the scenario is asymmetric. To explain the first
* sub-condition, we need to add that the function
* bfq_asymmetric_scenario checks the weights of only
* non-weight-raised queues, for efficiency reasons (see comments on
* bfq_weights_tree_add()). Then the fact that bfqq is weight-raised
* is checked explicitly here. More precisely, the compound condition
* below takes into account also the fact that, even if bfqq is being
* weight-raised, the scenario is still symmetric if all queues with
* requests waiting for completion happen to be
* weight-raised. Actually, we should be even more precise here, and
* differentiate between interactive weight raising and soft real-time
* weight raising.
*
* The second sub-condition checked in the compound condition is
* whether there is a fair amount of already in-flight I/O not
* belonging to bfqq. If so, I/O dispatching is to be plugged, for the
* following reason. The drive may decide to serve in-flight
* non-bfqq's I/O requests before bfqq's ones, thereby delaying the
* arrival of new I/O requests for bfqq (recall that bfqq is sync). If
* I/O-dispatching is not plugged, then, while bfqq remains empty, a
* basically uncontrolled amount of I/O from other queues may be
* dispatched too, possibly causing the service of bfqq's I/O to be
* delayed even longer in the drive. This problem gets more and more
* serious as the speed and the queue depth of the drive grow,
* because, as these two quantities grow, the probability to find no
* queue busy but many requests in flight grows too. By contrast,
* plugging I/O dispatching minimizes the delay induced by already
* in-flight I/O, and enables bfqq to recover the bandwidth it may
* lose because of this delay.
*
* As a side note, it is worth considering that the above
* device-idling countermeasures may however fail in the following
* unlucky scenario: if I/O-dispatch plugging is (correctly) disabled
* in a time period during which all symmetry sub-conditions hold, and
* therefore the device is allowed to enqueue many requests, but at
* some later point in time some sub-condition stops to hold, then it
* may become impossible to make requests be served in the desired
* order until all the requests already queued in the device have been
* served. The last sub-condition commented above somewhat mitigates
* this problem for weight-raised queues.
*/
static bool idling_needed_for_service_guarantees(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
return (bfqq->wr_coeff > 1 &&
(bfqd->wr_busy_queues <
bfq_tot_busy_queues(bfqd) ||
bfqd->rq_in_driver >=
bfqq->dispatched + 4)) ||
bfq_asymmetric_scenario(bfqd, bfqq);
}
static bool __bfq_bfqq_expire(struct bfq_data *bfqd, struct bfq_queue *bfqq,
enum bfqq_expiration reason)
{
/*
* If this bfqq is shared between multiple processes, check
* to make sure that those processes are still issuing I/Os
* within the mean seek distance. If not, it may be time to
* break the queues apart again.
*/
if (bfq_bfqq_coop(bfqq) && BFQQ_SEEKY(bfqq))
bfq_mark_bfqq_split_coop(bfqq);
/*
* Consider queues with a higher finish virtual time than
* bfqq. If idling_needed_for_service_guarantees(bfqq) returns
* true, then bfqq's bandwidth would be violated if an
* uncontrolled amount of I/O from these queues were
* dispatched while bfqq is waiting for its new I/O to
* arrive. This is exactly what may happen if this is a forced
* expiration caused by a preemption attempt, and if bfqq is
* not re-scheduled. To prevent this from happening, re-queue
* bfqq if it needs I/O-dispatch plugging, even if it is
* empty. By doing so, bfqq is granted to be served before the
* above queues (provided that bfqq is of course eligible).
*/
if (RB_EMPTY_ROOT(&bfqq->sort_list) &&
!(reason == BFQQE_PREEMPTED &&
idling_needed_for_service_guarantees(bfqd, bfqq))) {
if (bfqq->dispatched == 0)
/*
* Overloading budget_timeout field to store
* the time at which the queue remains with no
* backlog and no outstanding request; used by
* the weight-raising mechanism.
*/
bfqq->budget_timeout = jiffies;
bfq_del_bfqq_busy(bfqd, bfqq, true);
} else {
bfq_requeue_bfqq(bfqd, bfqq, true);
/*
* Resort priority tree of potential close cooperators.
* See comments on bfq_pos_tree_add_move() for the unlikely().
*/
if (unlikely(!bfqd->nonrot_with_queueing &&
!RB_EMPTY_ROOT(&bfqq->sort_list)))
bfq_pos_tree_add_move(bfqd, bfqq);
}
/*
* All in-service entities must have been properly deactivated
* or requeued before executing the next function, which
* resets all in-service entities as no more in service. This
* may cause bfqq to be freed. If this happens, the next
* function returns true.
*/
return __bfq_bfqd_reset_in_service(bfqd);
}
/**
* __bfq_bfqq_recalc_budget - try to adapt the budget to the @bfqq behavior.
* @bfqd: device data.
* @bfqq: queue to update.
* @reason: reason for expiration.
*
* Handle the feedback on @bfqq budget at queue expiration.
* See the body for detailed comments.
*/
static void __bfq_bfqq_recalc_budget(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
enum bfqq_expiration reason)
{
struct request *next_rq;
int budget, min_budget;
min_budget = bfq_min_budget(bfqd);
if (bfqq->wr_coeff == 1)
budget = bfqq->max_budget;
else /*
* Use a constant, low budget for weight-raised queues,
* to help achieve a low latency. Keep it slightly higher
* than the minimum possible budget, to cause a little
* bit fewer expirations.
*/
budget = 2 * min_budget;
bfq_log_bfqq(bfqd, bfqq, "recalc_budg: last budg %d, budg left %d",
bfqq->entity.budget, bfq_bfqq_budget_left(bfqq));
bfq_log_bfqq(bfqd, bfqq, "recalc_budg: last max_budg %d, min budg %d",
budget, bfq_min_budget(bfqd));
bfq_log_bfqq(bfqd, bfqq, "recalc_budg: sync %d, seeky %d",
bfq_bfqq_sync(bfqq), BFQQ_SEEKY(bfqd->in_service_queue));
if (bfq_bfqq_sync(bfqq) && bfqq->wr_coeff == 1) {
switch (reason) {
/*
* Caveat: in all the following cases we trade latency
* for throughput.
*/
case BFQQE_TOO_IDLE:
/*
* This is the only case where we may reduce
* the budget: if there is no request of the
* process still waiting for completion, then
* we assume (tentatively) that the timer has
* expired because the batch of requests of
* the process could have been served with a
* smaller budget. Hence, betting that
* process will behave in the same way when it
* becomes backlogged again, we reduce its
* next budget. As long as we guess right,
* this budget cut reduces the latency
* experienced by the process.
*
* However, if there are still outstanding
* requests, then the process may have not yet
* issued its next request just because it is
* still waiting for the completion of some of
* the still outstanding ones. So in this
* subcase we do not reduce its budget, on the
* contrary we increase it to possibly boost
* the throughput, as discussed in the
* comments to the BUDGET_TIMEOUT case.
*/
if (bfqq->dispatched > 0) /* still outstanding reqs */
budget = min(budget * 2, bfqd->bfq_max_budget);
else {
if (budget > 5 * min_budget)
budget -= 4 * min_budget;
else
budget = min_budget;
}
break;
case BFQQE_BUDGET_TIMEOUT:
/*
* We double the budget here because it gives
* the chance to boost the throughput if this
* is not a seeky process (and has bumped into
* this timeout because of, e.g., ZBR).
*/
budget = min(budget * 2, bfqd->bfq_max_budget);
break;
case BFQQE_BUDGET_EXHAUSTED:
/*
* The process still has backlog, and did not
* let either the budget timeout or the disk
* idling timeout expire. Hence it is not
* seeky, has a short thinktime and may be
* happy with a higher budget too. So
* definitely increase the budget of this good
* candidate to boost the disk throughput.
*/
budget = min(budget * 4, bfqd->bfq_max_budget);
break;
case BFQQE_NO_MORE_REQUESTS:
/*
* For queues that expire for this reason, it
* is particularly important to keep the
* budget close to the actual service they
* need. Doing so reduces the timestamp
* misalignment problem described in the
* comments in the body of
* __bfq_activate_entity. In fact, suppose
* that a queue systematically expires for
* BFQQE_NO_MORE_REQUESTS and presents a
* new request in time to enjoy timestamp
* back-shifting. The larger the budget of the
* queue is with respect to the service the
* queue actually requests in each service
* slot, the more times the queue can be
* reactivated with the same virtual finish
* time. It follows that, even if this finish
* time is pushed to the system virtual time
* to reduce the consequent timestamp
* misalignment, the queue unjustly enjoys for
* many re-activations a lower finish time
* than all newly activated queues.
*
* The service needed by bfqq is measured
* quite precisely by bfqq->entity.service.
* Since bfqq does not enjoy device idling,
* bfqq->entity.service is equal to the number
* of sectors that the process associated with
* bfqq requested to read/write before waiting
* for request completions, or blocking for
* other reasons.
*/
budget = max_t(int, bfqq->entity.service, min_budget);
break;
default:
return;
}
} else if (!bfq_bfqq_sync(bfqq)) {
/*
* Async queues get always the maximum possible
* budget, as for them we do not care about latency
* (in addition, their ability to dispatch is limited
* by the charging factor).
*/
budget = bfqd->bfq_max_budget;
}
bfqq->max_budget = budget;
if (bfqd->budgets_assigned >= bfq_stats_min_budgets &&
!bfqd->bfq_user_max_budget)
bfqq->max_budget = min(bfqq->max_budget, bfqd->bfq_max_budget);
/*
* If there is still backlog, then assign a new budget, making
* sure that it is large enough for the next request. Since
* the finish time of bfqq must be kept in sync with the
* budget, be sure to call __bfq_bfqq_expire() *after* this
* update.
*
* If there is no backlog, then no need to update the budget;
* it will be updated on the arrival of a new request.
*/
next_rq = bfqq->next_rq;
if (next_rq)
bfqq->entity.budget = max_t(unsigned long, bfqq->max_budget,
bfq_serv_to_charge(next_rq, bfqq));
bfq_log_bfqq(bfqd, bfqq, "head sect: %u, new budget %d",
next_rq ? blk_rq_sectors(next_rq) : 0,
bfqq->entity.budget);
}
/*
* Return true if the process associated with bfqq is "slow". The slow
* flag is used, in addition to the budget timeout, to reduce the
* amount of service provided to seeky processes, and thus reduce
* their chances to lower the throughput. More details in the comments
* on the function bfq_bfqq_expire().
*
* An important observation is in order: as discussed in the comments
* on the function bfq_update_peak_rate(), with devices with internal
* queues, it is hard if ever possible to know when and for how long
* an I/O request is processed by the device (apart from the trivial
* I/O pattern where a new request is dispatched only after the
* previous one has been completed). This makes it hard to evaluate
* the real rate at which the I/O requests of each bfq_queue are
* served. In fact, for an I/O scheduler like BFQ, serving a
* bfq_queue means just dispatching its requests during its service
* slot (i.e., until the budget of the queue is exhausted, or the
* queue remains idle, or, finally, a timeout fires). But, during the
* service slot of a bfq_queue, around 100 ms at most, the device may
* be even still processing requests of bfq_queues served in previous
* service slots. On the opposite end, the requests of the in-service
* bfq_queue may be completed after the service slot of the queue
* finishes.
*
* Anyway, unless more sophisticated solutions are used
* (where possible), the sum of the sizes of the requests dispatched
* during the service slot of a bfq_queue is probably the only
* approximation available for the service received by the bfq_queue
* during its service slot. And this sum is the quantity used in this
* function to evaluate the I/O speed of a process.
*/
static bool bfq_bfqq_is_slow(struct bfq_data *bfqd, struct bfq_queue *bfqq,
bool compensate, enum bfqq_expiration reason,
unsigned long *delta_ms)
{
ktime_t delta_ktime;
u32 delta_usecs;
bool slow = BFQQ_SEEKY(bfqq); /* if delta too short, use seekyness */
if (!bfq_bfqq_sync(bfqq))
return false;
if (compensate)
delta_ktime = bfqd->last_idling_start;
else
delta_ktime = ktime_get();
delta_ktime = ktime_sub(delta_ktime, bfqd->last_budget_start);
delta_usecs = ktime_to_us(delta_ktime);
/* don't use too short time intervals */
if (delta_usecs < 1000) {
if (blk_queue_nonrot(bfqd->queue))
/*
* give same worst-case guarantees as idling
* for seeky
*/
*delta_ms = BFQ_MIN_TT / NSEC_PER_MSEC;
else /* charge at least one seek */
*delta_ms = bfq_slice_idle / NSEC_PER_MSEC;
return slow;
}
*delta_ms = delta_usecs / USEC_PER_MSEC;
/*
* Use only long (> 20ms) intervals to filter out excessive
* spikes in service rate estimation.
*/
if (delta_usecs > 20000) {
/*
* Caveat for rotational devices: processes doing I/O
* in the slower disk zones tend to be slow(er) even
* if not seeky. In this respect, the estimated peak
* rate is likely to be an average over the disk
* surface. Accordingly, to not be too harsh with
* unlucky processes, a process is deemed slow only if
* its rate has been lower than half of the estimated
* peak rate.
*/
slow = bfqq->entity.service < bfqd->bfq_max_budget / 2;
}
bfq_log_bfqq(bfqd, bfqq, "bfq_bfqq_is_slow: slow %d", slow);
return slow;
}
/*
* To be deemed as soft real-time, an application must meet two
* requirements. First, the application must not require an average
* bandwidth higher than the approximate bandwidth required to playback or
* record a compressed high-definition video.
* The next function is invoked on the completion of the last request of a
* batch, to compute the next-start time instant, soft_rt_next_start, such
* that, if the next request of the application does not arrive before
* soft_rt_next_start, then the above requirement on the bandwidth is met.
*
* The second requirement is that the request pattern of the application is
* isochronous, i.e., that, after issuing a request or a batch of requests,
* the application stops issuing new requests until all its pending requests
* have been completed. After that, the application may issue a new batch,
* and so on.
* For this reason the next function is invoked to compute
* soft_rt_next_start only for applications that meet this requirement,
* whereas soft_rt_next_start is set to infinity for applications that do
* not.
*
* Unfortunately, even a greedy (i.e., I/O-bound) application may
* happen to meet, occasionally or systematically, both the above
* bandwidth and isochrony requirements. This may happen at least in
* the following circumstances. First, if the CPU load is high. The
* application may stop issuing requests while the CPUs are busy
* serving other processes, then restart, then stop again for a while,
* and so on. The other circumstances are related to the storage
* device: the storage device is highly loaded or reaches a low-enough
* throughput with the I/O of the application (e.g., because the I/O
* is random and/or the device is slow). In all these cases, the
* I/O of the application may be simply slowed down enough to meet
* the bandwidth and isochrony requirements. To reduce the probability
* that greedy applications are deemed as soft real-time in these
* corner cases, a further rule is used in the computation of
* soft_rt_next_start: the return value of this function is forced to
* be higher than the maximum between the following two quantities.
*
* (a) Current time plus: (1) the maximum time for which the arrival
* of a request is waited for when a sync queue becomes idle,
* namely bfqd->bfq_slice_idle, and (2) a few extra jiffies. We
* postpone for a moment the reason for adding a few extra
* jiffies; we get back to it after next item (b). Lower-bounding
* the return value of this function with the current time plus
* bfqd->bfq_slice_idle tends to filter out greedy applications,
* because the latter issue their next request as soon as possible
* after the last one has been completed. In contrast, a soft
* real-time application spends some time processing data, after a
* batch of its requests has been completed.
*
* (b) Current value of bfqq->soft_rt_next_start. As pointed out
* above, greedy applications may happen to meet both the
* bandwidth and isochrony requirements under heavy CPU or
* storage-device load. In more detail, in these scenarios, these
* applications happen, only for limited time periods, to do I/O
* slowly enough to meet all the requirements described so far,
* including the filtering in above item (a). These slow-speed
* time intervals are usually interspersed between other time
* intervals during which these applications do I/O at a very high
* speed. Fortunately, exactly because of the high speed of the
* I/O in the high-speed intervals, the values returned by this
* function happen to be so high, near the end of any such
* high-speed interval, to be likely to fall *after* the end of
* the low-speed time interval that follows. These high values are
* stored in bfqq->soft_rt_next_start after each invocation of
* this function. As a consequence, if the last value of
* bfqq->soft_rt_next_start is constantly used to lower-bound the
* next value that this function may return, then, from the very
* beginning of a low-speed interval, bfqq->soft_rt_next_start is
* likely to be constantly kept so high that any I/O request
* issued during the low-speed interval is considered as arriving
* to soon for the application to be deemed as soft
* real-time. Then, in the high-speed interval that follows, the
* application will not be deemed as soft real-time, just because
* it will do I/O at a high speed. And so on.
*
* Getting back to the filtering in item (a), in the following two
* cases this filtering might be easily passed by a greedy
* application, if the reference quantity was just
* bfqd->bfq_slice_idle:
* 1) HZ is so low that the duration of a jiffy is comparable to or
* higher than bfqd->bfq_slice_idle. This happens, e.g., on slow
* devices with HZ=100. The time granularity may be so coarse
* that the approximation, in jiffies, of bfqd->bfq_slice_idle
* is rather lower than the exact value.
* 2) jiffies, instead of increasing at a constant rate, may stop increasing
* for a while, then suddenly 'jump' by several units to recover the lost
* increments. This seems to happen, e.g., inside virtual machines.
* To address this issue, in the filtering in (a) we do not use as a
* reference time interval just bfqd->bfq_slice_idle, but
* bfqd->bfq_slice_idle plus a few jiffies. In particular, we add the
* minimum number of jiffies for which the filter seems to be quite
* precise also in embedded systems and KVM/QEMU virtual machines.
*/
static unsigned long bfq_bfqq_softrt_next_start(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
return max3(bfqq->soft_rt_next_start,
bfqq->last_idle_bklogged +
HZ * bfqq->service_from_backlogged /
bfqd->bfq_wr_max_softrt_rate,
jiffies + nsecs_to_jiffies(bfqq->bfqd->bfq_slice_idle) + 4);
}
/**
* bfq_bfqq_expire - expire a queue.
* @bfqd: device owning the queue.
* @bfqq: the queue to expire.
* @compensate: if true, compensate for the time spent idling.
* @reason: the reason causing the expiration.
*
* If the process associated with bfqq does slow I/O (e.g., because it
* issues random requests), we charge bfqq with the time it has been
* in service instead of the service it has received (see
* bfq_bfqq_charge_time for details on how this goal is achieved). As
* a consequence, bfqq will typically get higher timestamps upon
* reactivation, and hence it will be rescheduled as if it had
* received more service than what it has actually received. In the
* end, bfqq receives less service in proportion to how slowly its
* associated process consumes its budgets (and hence how seriously it
* tends to lower the throughput). In addition, this time-charging
* strategy guarantees time fairness among slow processes. In
* contrast, if the process associated with bfqq is not slow, we
* charge bfqq exactly with the service it has received.
*
* Charging time to the first type of queues and the exact service to
* the other has the effect of using the WF2Q+ policy to schedule the
* former on a timeslice basis, without violating service domain
* guarantees among the latter.
*/
void bfq_bfqq_expire(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
bool compensate,
enum bfqq_expiration reason)
{
bool slow;
unsigned long delta = 0;
struct bfq_entity *entity = &bfqq->entity;
/*
* Check whether the process is slow (see bfq_bfqq_is_slow).
*/
slow = bfq_bfqq_is_slow(bfqd, bfqq, compensate, reason, &delta);
/*
* As above explained, charge slow (typically seeky) and
* timed-out queues with the time and not the service
* received, to favor sequential workloads.
*
* Processes doing I/O in the slower disk zones will tend to
* be slow(er) even if not seeky. Therefore, since the
* estimated peak rate is actually an average over the disk
* surface, these processes may timeout just for bad luck. To
* avoid punishing them, do not charge time to processes that
* succeeded in consuming at least 2/3 of their budget. This
* allows BFQ to preserve enough elasticity to still perform
* bandwidth, and not time, distribution with little unlucky
* or quasi-sequential processes.
*/
if (bfqq->wr_coeff == 1 &&
(slow ||
(reason == BFQQE_BUDGET_TIMEOUT &&
bfq_bfqq_budget_left(bfqq) >= entity->budget / 3)))
bfq_bfqq_charge_time(bfqd, bfqq, delta);
if (reason == BFQQE_TOO_IDLE &&
entity->service <= 2 * entity->budget / 10)
bfq_clear_bfqq_IO_bound(bfqq);
if (bfqd->low_latency && bfqq->wr_coeff == 1)
bfqq->last_wr_start_finish = jiffies;
if (bfqd->low_latency && bfqd->bfq_wr_max_softrt_rate > 0 &&
RB_EMPTY_ROOT(&bfqq->sort_list)) {
/*
* If we get here, and there are no outstanding
* requests, then the request pattern is isochronous
* (see the comments on the function
* bfq_bfqq_softrt_next_start()). Thus we can compute
* soft_rt_next_start. And we do it, unless bfqq is in
* interactive weight raising. We do not do it in the
* latter subcase, for the following reason. bfqq may
* be conveying the I/O needed to load a soft
* real-time application. Such an application will
* actually exhibit a soft real-time I/O pattern after
* it finally starts doing its job. But, if
* soft_rt_next_start is computed here for an
* interactive bfqq, and bfqq had received a lot of
* service before remaining with no outstanding
* request (likely to happen on a fast device), then
* soft_rt_next_start would be assigned such a high
* value that, for a very long time, bfqq would be
* prevented from being possibly considered as soft
* real time.
*
* If, instead, the queue still has outstanding
* requests, then we have to wait for the completion
* of all the outstanding requests to discover whether
* the request pattern is actually isochronous.
*/
if (bfqq->dispatched == 0 &&
bfqq->wr_coeff != bfqd->bfq_wr_coeff)
bfqq->soft_rt_next_start =
bfq_bfqq_softrt_next_start(bfqd, bfqq);
else if (bfqq->dispatched > 0) {
/*
* Schedule an update of soft_rt_next_start to when
* the task may be discovered to be isochronous.
*/
bfq_mark_bfqq_softrt_update(bfqq);
}
}
bfq_log_bfqq(bfqd, bfqq,
"expire (%d, slow %d, num_disp %d, short_ttime %d)", reason,
slow, bfqq->dispatched, bfq_bfqq_has_short_ttime(bfqq));
/*
* bfqq expired, so no total service time needs to be computed
* any longer: reset state machine for measuring total service
* times.
*/
bfqd->rqs_injected = bfqd->wait_dispatch = false;
bfqd->waited_rq = NULL;
/*
* Increase, decrease or leave budget unchanged according to
* reason.
*/
__bfq_bfqq_recalc_budget(bfqd, bfqq, reason);
if (__bfq_bfqq_expire(bfqd, bfqq, reason))
/* bfqq is gone, no more actions on it */
return;
/* mark bfqq as waiting a request only if a bic still points to it */
if (!bfq_bfqq_busy(bfqq) &&
reason != BFQQE_BUDGET_TIMEOUT &&
reason != BFQQE_BUDGET_EXHAUSTED) {
bfq_mark_bfqq_non_blocking_wait_rq(bfqq);
/*
* Not setting service to 0, because, if the next rq
* arrives in time, the queue will go on receiving
* service with this same budget (as if it never expired)
*/
} else
entity->service = 0;
/*
* Reset the received-service counter for every parent entity.
* Differently from what happens with bfqq->entity.service,
* the resetting of this counter never needs to be postponed
* for parent entities. In fact, in case bfqq may have a
* chance to go on being served using the last, partially
* consumed budget, bfqq->entity.service needs to be kept,
* because if bfqq then actually goes on being served using
* the same budget, the last value of bfqq->entity.service is
* needed to properly decrement bfqq->entity.budget by the
* portion already consumed. In contrast, it is not necessary
* to keep entity->service for parent entities too, because
* the bubble up of the new value of bfqq->entity.budget will
* make sure that the budgets of parent entities are correct,
* even in case bfqq and thus parent entities go on receiving
* service with the same budget.
*/
entity = entity->parent;
for_each_entity(entity)
entity->service = 0;
}
/*
* Budget timeout is not implemented through a dedicated timer, but
* just checked on request arrivals and completions, as well as on
* idle timer expirations.
*/
static bool bfq_bfqq_budget_timeout(struct bfq_queue *bfqq)
{
return time_is_before_eq_jiffies(bfqq->budget_timeout);
}
/*
* If we expire a queue that is actively waiting (i.e., with the
* device idled) for the arrival of a new request, then we may incur
* the timestamp misalignment problem described in the body of the
* function __bfq_activate_entity. Hence we return true only if this
* condition does not hold, or if the queue is slow enough to deserve
* only to be kicked off for preserving a high throughput.
*/
static bool bfq_may_expire_for_budg_timeout(struct bfq_queue *bfqq)
{
bfq_log_bfqq(bfqq->bfqd, bfqq,
"may_budget_timeout: wait_request %d left %d timeout %d",
bfq_bfqq_wait_request(bfqq),
bfq_bfqq_budget_left(bfqq) >= bfqq->entity.budget / 3,
bfq_bfqq_budget_timeout(bfqq));
return (!bfq_bfqq_wait_request(bfqq) ||
bfq_bfqq_budget_left(bfqq) >= bfqq->entity.budget / 3)
&&
bfq_bfqq_budget_timeout(bfqq);
}
static bool idling_boosts_thr_without_issues(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
bool rot_without_queueing =
!blk_queue_nonrot(bfqd->queue) && !bfqd->hw_tag,
bfqq_sequential_and_IO_bound,
idling_boosts_thr;
bfqq_sequential_and_IO_bound = !BFQQ_SEEKY(bfqq) &&
bfq_bfqq_IO_bound(bfqq) && bfq_bfqq_has_short_ttime(bfqq);
/*
* The next variable takes into account the cases where idling
* boosts the throughput.
*
* The value of the variable is computed considering, first, that
* idling is virtually always beneficial for the throughput if:
* (a) the device is not NCQ-capable and rotational, or
* (b) regardless of the presence of NCQ, the device is rotational and
* the request pattern for bfqq is I/O-bound and sequential, or
* (c) regardless of whether it is rotational, the device is
* not NCQ-capable and the request pattern for bfqq is
* I/O-bound and sequential.
*
* Secondly, and in contrast to the above item (b), idling an
* NCQ-capable flash-based device would not boost the
* throughput even with sequential I/O; rather it would lower
* the throughput in proportion to how fast the device
* is. Accordingly, the next variable is true if any of the
* above conditions (a), (b) or (c) is true, and, in
* particular, happens to be false if bfqd is an NCQ-capable
* flash-based device.
*/
idling_boosts_thr = rot_without_queueing ||
((!blk_queue_nonrot(bfqd->queue) || !bfqd->hw_tag) &&
bfqq_sequential_and_IO_bound);
/*
* The return value of this function is equal to that of
* idling_boosts_thr, unless a special case holds. In this
* special case, described below, idling may cause problems to
* weight-raised queues.
*
* When the request pool is saturated (e.g., in the presence
* of write hogs), if the processes associated with
* non-weight-raised queues ask for requests at a lower rate,
* then processes associated with weight-raised queues have a
* higher probability to get a request from the pool
* immediately (or at least soon) when they need one. Thus
* they have a higher probability to actually get a fraction
* of the device throughput proportional to their high
* weight. This is especially true with NCQ-capable drives,
* which enqueue several requests in advance, and further
* reorder internally-queued requests.
*
* For this reason, we force to false the return value if
* there are weight-raised busy queues. In this case, and if
* bfqq is not weight-raised, this guarantees that the device
* is not idled for bfqq (if, instead, bfqq is weight-raised,
* then idling will be guaranteed by another variable, see
* below). Combined with the timestamping rules of BFQ (see
* [1] for details), this behavior causes bfqq, and hence any
* sync non-weight-raised queue, to get a lower number of
* requests served, and thus to ask for a lower number of
* requests from the request pool, before the busy
* weight-raised queues get served again. This often mitigates
* starvation problems in the presence of heavy write
* workloads and NCQ, thereby guaranteeing a higher
* application and system responsiveness in these hostile
* scenarios.
*/
return idling_boosts_thr &&
bfqd->wr_busy_queues == 0;
}
/*
* For a queue that becomes empty, device idling is allowed only if
* this function returns true for that queue. As a consequence, since
* device idling plays a critical role for both throughput boosting
* and service guarantees, the return value of this function plays a
* critical role as well.
*
* In a nutshell, this function returns true only if idling is
* beneficial for throughput or, even if detrimental for throughput,
* idling is however necessary to preserve service guarantees (low
* latency, desired throughput distribution, ...). In particular, on
* NCQ-capable devices, this function tries to return false, so as to
* help keep the drives' internal queues full, whenever this helps the
* device boost the throughput without causing any service-guarantee
* issue.
*
* Most of the issues taken into account to get the return value of
* this function are not trivial. We discuss these issues in the two
* functions providing the main pieces of information needed by this
* function.
*/
static bool bfq_better_to_idle(struct bfq_queue *bfqq)
{
struct bfq_data *bfqd = bfqq->bfqd;
bool idling_boosts_thr_with_no_issue, idling_needed_for_service_guar;
if (unlikely(bfqd->strict_guarantees))
return true;
/*
* Idling is performed only if slice_idle > 0. In addition, we
* do not idle if
* (a) bfqq is async
* (b) bfqq is in the idle io prio class: in this case we do
* not idle because we want to minimize the bandwidth that
* queues in this class can steal to higher-priority queues
*/
if (bfqd->bfq_slice_idle == 0 || !bfq_bfqq_sync(bfqq) ||
bfq_class_idle(bfqq))
return false;
idling_boosts_thr_with_no_issue =
idling_boosts_thr_without_issues(bfqd, bfqq);
idling_needed_for_service_guar =
idling_needed_for_service_guarantees(bfqd, bfqq);
/*
* We have now the two components we need to compute the
* return value of the function, which is true only if idling
* either boosts the throughput (without issues), or is
* necessary to preserve service guarantees.
*/
return idling_boosts_thr_with_no_issue ||
idling_needed_for_service_guar;
}
/*
* If the in-service queue is empty but the function bfq_better_to_idle
* returns true, then:
* 1) the queue must remain in service and cannot be expired, and
* 2) the device must be idled to wait for the possible arrival of a new
* request for the queue.
* See the comments on the function bfq_better_to_idle for the reasons
* why performing device idling is the best choice to boost the throughput
* and preserve service guarantees when bfq_better_to_idle itself
* returns true.
*/
static bool bfq_bfqq_must_idle(struct bfq_queue *bfqq)
{
return RB_EMPTY_ROOT(&bfqq->sort_list) && bfq_better_to_idle(bfqq);
}
/*
* This function chooses the queue from which to pick the next extra
* I/O request to inject, if it finds a compatible queue. See the
* comments on bfq_update_inject_limit() for details on the injection
* mechanism, and for the definitions of the quantities mentioned
* below.
*/
static struct bfq_queue *
bfq_choose_bfqq_for_injection(struct bfq_data *bfqd)
{
struct bfq_queue *bfqq, *in_serv_bfqq = bfqd->in_service_queue;
unsigned int limit = in_serv_bfqq->inject_limit;
/*
* If
* - bfqq is not weight-raised and therefore does not carry
* time-critical I/O,
* or
* - regardless of whether bfqq is weight-raised, bfqq has
* however a long think time, during which it can absorb the
* effect of an appropriate number of extra I/O requests
* from other queues (see bfq_update_inject_limit for
* details on the computation of this number);
* then injection can be performed without restrictions.
*/
bool in_serv_always_inject = in_serv_bfqq->wr_coeff == 1 ||
!bfq_bfqq_has_short_ttime(in_serv_bfqq);
/*
* If
* - the baseline total service time could not be sampled yet,
* so the inject limit happens to be still 0, and
* - a lot of time has elapsed since the plugging of I/O
* dispatching started, so drive speed is being wasted
* significantly;
* then temporarily raise inject limit to one request.
*/
if (limit == 0 && in_serv_bfqq->last_serv_time_ns == 0 &&
bfq_bfqq_wait_request(in_serv_bfqq) &&
time_is_before_eq_jiffies(bfqd->last_idling_start_jiffies +
bfqd->bfq_slice_idle)
)
limit = 1;
if (bfqd->rq_in_driver >= limit)
return NULL;
/*
* Linear search of the source queue for injection; but, with
* a high probability, very few steps are needed to find a
* candidate queue, i.e., a queue with enough budget left for
* its next request. In fact:
* - BFQ dynamically updates the budget of every queue so as
* to accommodate the expected backlog of the queue;
* - if a queue gets all its requests dispatched as injected
* service, then the queue is removed from the active list
* (and re-added only if it gets new requests, but then it
* is assigned again enough budget for its new backlog).
*/
list_for_each_entry(bfqq, &bfqd->active_list, bfqq_list)
if (!RB_EMPTY_ROOT(&bfqq->sort_list) &&
(in_serv_always_inject || bfqq->wr_coeff > 1) &&
bfq_serv_to_charge(bfqq->next_rq, bfqq) <=
bfq_bfqq_budget_left(bfqq)) {
/*
* Allow for only one large in-flight request
* on non-rotational devices, for the
* following reason. On non-rotationl drives,
* large requests take much longer than
* smaller requests to be served. In addition,
* the drive prefers to serve large requests
* w.r.t. to small ones, if it can choose. So,
* having more than one large requests queued
* in the drive may easily make the next first
* request of the in-service queue wait for so
* long to break bfqq's service guarantees. On
* the bright side, large requests let the
* drive reach a very high throughput, even if
* there is only one in-flight large request
* at a time.
*/
if (blk_queue_nonrot(bfqd->queue) &&
blk_rq_sectors(bfqq->next_rq) >=
BFQQ_SECT_THR_NONROT)
limit = min_t(unsigned int, 1, limit);
else
limit = in_serv_bfqq->inject_limit;
if (bfqd->rq_in_driver < limit) {
bfqd->rqs_injected = true;
return bfqq;
}
}
return NULL;
}
/*
* Select a queue for service. If we have a current queue in service,
* check whether to continue servicing it, or retrieve and set a new one.
*/
static struct bfq_queue *bfq_select_queue(struct bfq_data *bfqd)
{
struct bfq_queue *bfqq;
struct request *next_rq;
enum bfqq_expiration reason = BFQQE_BUDGET_TIMEOUT;
bfqq = bfqd->in_service_queue;
if (!bfqq)
goto new_queue;
bfq_log_bfqq(bfqd, bfqq, "select_queue: already in-service queue");
/*
* Do not expire bfqq for budget timeout if bfqq may be about
* to enjoy device idling. The reason why, in this case, we
* prevent bfqq from expiring is the same as in the comments
* on the case where bfq_bfqq_must_idle() returns true, in
* bfq_completed_request().
*/
if (bfq_may_expire_for_budg_timeout(bfqq) &&
!bfq_bfqq_must_idle(bfqq))
goto expire;
check_queue:
/*
* This loop is rarely executed more than once. Even when it
* happens, it is much more convenient to re-execute this loop
* than to return NULL and trigger a new dispatch to get a
* request served.
*/
next_rq = bfqq->next_rq;
/*
* If bfqq has requests queued and it has enough budget left to
* serve them, keep the queue, otherwise expire it.
*/
if (next_rq) {
if (bfq_serv_to_charge(next_rq, bfqq) >
bfq_bfqq_budget_left(bfqq)) {
/*
* Expire the queue for budget exhaustion,
* which makes sure that the next budget is
* enough to serve the next request, even if
* it comes from the fifo expired path.
*/
reason = BFQQE_BUDGET_EXHAUSTED;
goto expire;
} else {
/*
* The idle timer may be pending because we may
* not disable disk idling even when a new request
* arrives.
*/
if (bfq_bfqq_wait_request(bfqq)) {
/*
* If we get here: 1) at least a new request
* has arrived but we have not disabled the
* timer because the request was too small,
* 2) then the block layer has unplugged
* the device, causing the dispatch to be
* invoked.
*
* Since the device is unplugged, now the
* requests are probably large enough to
* provide a reasonable throughput.
* So we disable idling.
*/
bfq_clear_bfqq_wait_request(bfqq);
hrtimer_try_to_cancel(&bfqd->idle_slice_timer);
}
goto keep_queue;
}
}
/*
* No requests pending. However, if the in-service queue is idling
* for a new request, or has requests waiting for a completion and
* may idle after their completion, then keep it anyway.
*
* Yet, inject service from other queues if it boosts
* throughput and is possible.
*/
if (bfq_bfqq_wait_request(bfqq) ||
(bfqq->dispatched != 0 && bfq_better_to_idle(bfqq))) {
struct bfq_queue *async_bfqq =
bfqq->bic && bfqq->bic->bfqq[0] &&
bfq_bfqq_busy(bfqq->bic->bfqq[0]) &&
bfqq->bic->bfqq[0]->next_rq ?
bfqq->bic->bfqq[0] : NULL;
/*
* The next three mutually-exclusive ifs decide
* whether to try injection, and choose the queue to
* pick an I/O request from.
*
* The first if checks whether the process associated
* with bfqq has also async I/O pending. If so, it
* injects such I/O unconditionally. Injecting async
* I/O from the same process can cause no harm to the
* process. On the contrary, it can only increase
* bandwidth and reduce latency for the process.
*
* The second if checks whether there happens to be a
* non-empty waker queue for bfqq, i.e., a queue whose
* I/O needs to be completed for bfqq to receive new
* I/O. This happens, e.g., if bfqq is associated with
* a process that does some sync. A sync generates
* extra blocking I/O, which must be completed before
* the process associated with bfqq can go on with its
* I/O. If the I/O of the waker queue is not served,
* then bfqq remains empty, and no I/O is dispatched,
* until the idle timeout fires for bfqq. This is
* likely to result in lower bandwidth and higher
* latencies for bfqq, and in a severe loss of total
* throughput. The best action to take is therefore to
* serve the waker queue as soon as possible. So do it
* (without relying on the third alternative below for
* eventually serving waker_bfqq's I/O; see the last
* paragraph for further details). This systematic
* injection of I/O from the waker queue does not
* cause any delay to bfqq's I/O. On the contrary,
* next bfqq's I/O is brought forward dramatically,
* for it is not blocked for milliseconds.
*
* The third if checks whether bfqq is a queue for
* which it is better to avoid injection. It is so if
* bfqq delivers more throughput when served without
* any further I/O from other queues in the middle, or
* if the service times of bfqq's I/O requests both
* count more than overall throughput, and may be
* easily increased by injection (this happens if bfqq
* has a short think time). If none of these
* conditions holds, then a candidate queue for
* injection is looked for through
* bfq_choose_bfqq_for_injection(). Note that the
* latter may return NULL (for example if the inject
* limit for bfqq is currently 0).
*
* NOTE: motivation for the second alternative
*
* Thanks to the way the inject limit is updated in
* bfq_update_has_short_ttime(), it is rather likely
* that, if I/O is being plugged for bfqq and the
* waker queue has pending I/O requests that are
* blocking bfqq's I/O, then the third alternative
* above lets the waker queue get served before the
* I/O-plugging timeout fires. So one may deem the
* second alternative superfluous. It is not, because
* the third alternative may be way less effective in
* case of a synchronization. For two main
* reasons. First, throughput may be low because the
* inject limit may be too low to guarantee the same
* amount of injected I/O, from the waker queue or
* other queues, that the second alternative
* guarantees (the second alternative unconditionally
* injects a pending I/O request of the waker queue
* for each bfq_dispatch_request()). Second, with the
* third alternative, the duration of the plugging,
* i.e., the time before bfqq finally receives new I/O,
* may not be minimized, because the waker queue may
* happen to be served only after other queues.
*/
if (async_bfqq &&
icq_to_bic(async_bfqq->next_rq->elv.icq) == bfqq->bic &&
bfq_serv_to_charge(async_bfqq->next_rq, async_bfqq) <=
bfq_bfqq_budget_left(async_bfqq))
bfqq = bfqq->bic->bfqq[0];
else if (bfq_bfqq_has_waker(bfqq) &&
bfq_bfqq_busy(bfqq->waker_bfqq) &&
bfqq->next_rq &&
bfq_serv_to_charge(bfqq->waker_bfqq->next_rq,
bfqq->waker_bfqq) <=
bfq_bfqq_budget_left(bfqq->waker_bfqq)
)
bfqq = bfqq->waker_bfqq;
else if (!idling_boosts_thr_without_issues(bfqd, bfqq) &&
(bfqq->wr_coeff == 1 || bfqd->wr_busy_queues > 1 ||
!bfq_bfqq_has_short_ttime(bfqq)))
bfqq = bfq_choose_bfqq_for_injection(bfqd);
else
bfqq = NULL;
goto keep_queue;
}
reason = BFQQE_NO_MORE_REQUESTS;
expire:
bfq_bfqq_expire(bfqd, bfqq, false, reason);
new_queue:
bfqq = bfq_set_in_service_queue(bfqd);
if (bfqq) {
bfq_log_bfqq(bfqd, bfqq, "select_queue: checking new queue");
goto check_queue;
}
keep_queue:
if (bfqq)
bfq_log_bfqq(bfqd, bfqq, "select_queue: returned this queue");
else
bfq_log(bfqd, "select_queue: no queue returned");
return bfqq;
}
static void bfq_update_wr_data(struct bfq_data *bfqd, struct bfq_queue *bfqq)
{
struct bfq_entity *entity = &bfqq->entity;
if (bfqq->wr_coeff > 1) { /* queue is being weight-raised */
bfq_log_bfqq(bfqd, bfqq,
"raising period dur %u/%u msec, old coeff %u, w %d(%d)",
jiffies_to_msecs(jiffies - bfqq->last_wr_start_finish),
jiffies_to_msecs(bfqq->wr_cur_max_time),
bfqq->wr_coeff,
bfqq->entity.weight, bfqq->entity.orig_weight);
if (entity->prio_changed)
bfq_log_bfqq(bfqd, bfqq, "WARN: pending prio change");
/*
* If the queue was activated in a burst, or too much
* time has elapsed from the beginning of this
* weight-raising period, then end weight raising.
*/
if (bfq_bfqq_in_large_burst(bfqq))
bfq_bfqq_end_wr(bfqq);
else if (time_is_before_jiffies(bfqq->last_wr_start_finish +
bfqq->wr_cur_max_time)) {
if (bfqq->wr_cur_max_time != bfqd->bfq_wr_rt_max_time ||
time_is_before_jiffies(bfqq->wr_start_at_switch_to_srt +
bfq_wr_duration(bfqd)))
bfq_bfqq_end_wr(bfqq);
else {
switch_back_to_interactive_wr(bfqq, bfqd);
bfqq->entity.prio_changed = 1;
}
}
if (bfqq->wr_coeff > 1 &&
bfqq->wr_cur_max_time != bfqd->bfq_wr_rt_max_time &&
bfqq->service_from_wr > max_service_from_wr) {
/* see comments on max_service_from_wr */
bfq_bfqq_end_wr(bfqq);
}
}
/*
* To improve latency (for this or other queues), immediately
* update weight both if it must be raised and if it must be
* lowered. Since, entity may be on some active tree here, and
* might have a pending change of its ioprio class, invoke
* next function with the last parameter unset (see the
* comments on the function).
*/
if ((entity->weight > entity->orig_weight) != (bfqq->wr_coeff > 1))
__bfq_entity_update_weight_prio(bfq_entity_service_tree(entity),
entity, false);
}
/*
* Dispatch next request from bfqq.
*/
static struct request *bfq_dispatch_rq_from_bfqq(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
struct request *rq = bfqq->next_rq;
unsigned long service_to_charge;
service_to_charge = bfq_serv_to_charge(rq, bfqq);
bfq_bfqq_served(bfqq, service_to_charge);
if (bfqq == bfqd->in_service_queue && bfqd->wait_dispatch) {
bfqd->wait_dispatch = false;
bfqd->waited_rq = rq;
}
bfq_dispatch_remove(bfqd->queue, rq);
if (bfqq != bfqd->in_service_queue)
goto return_rq;
/*
* If weight raising has to terminate for bfqq, then next
* function causes an immediate update of bfqq's weight,
* without waiting for next activation. As a consequence, on
* expiration, bfqq will be timestamped as if has never been
* weight-raised during this service slot, even if it has
* received part or even most of the service as a
* weight-raised queue. This inflates bfqq's timestamps, which
* is beneficial, as bfqq is then more willing to leave the
* device immediately to possible other weight-raised queues.
*/
bfq_update_wr_data(bfqd, bfqq);
/*
* Expire bfqq, pretending that its budget expired, if bfqq
* belongs to CLASS_IDLE and other queues are waiting for
* service.
*/
if (!(bfq_tot_busy_queues(bfqd) > 1 && bfq_class_idle(bfqq)))
goto return_rq;
bfq_bfqq_expire(bfqd, bfqq, false, BFQQE_BUDGET_EXHAUSTED);
return_rq:
return rq;
}
static bool bfq_has_work(struct blk_mq_hw_ctx *hctx)
{
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data;
/*
* Avoiding lock: a race on bfqd->busy_queues should cause at
* most a call to dispatch for nothing
*/
return !list_empty_careful(&bfqd->dispatch) ||
bfq_tot_busy_queues(bfqd) > 0;
}
static struct request *__bfq_dispatch_request(struct blk_mq_hw_ctx *hctx)
{
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data;
struct request *rq = NULL;
struct bfq_queue *bfqq = NULL;
if (!list_empty(&bfqd->dispatch)) {
rq = list_first_entry(&bfqd->dispatch, struct request,
queuelist);
list_del_init(&rq->queuelist);
bfqq = RQ_BFQQ(rq);
if (bfqq) {
/*
* Increment counters here, because this
* dispatch does not follow the standard
* dispatch flow (where counters are
* incremented)
*/
bfqq->dispatched++;
goto inc_in_driver_start_rq;
}
/*
* We exploit the bfq_finish_requeue_request hook to
* decrement rq_in_driver, but
* bfq_finish_requeue_request will not be invoked on
* this request. So, to avoid unbalance, just start
* this request, without incrementing rq_in_driver. As
* a negative consequence, rq_in_driver is deceptively
* lower than it should be while this request is in
* service. This may cause bfq_schedule_dispatch to be
* invoked uselessly.
*
* As for implementing an exact solution, the
* bfq_finish_requeue_request hook, if defined, is
* probably invoked also on this request. So, by
* exploiting this hook, we could 1) increment
* rq_in_driver here, and 2) decrement it in
* bfq_finish_requeue_request. Such a solution would
* let the value of the counter be always accurate,
* but it would entail using an extra interface
* function. This cost seems higher than the benefit,
* being the frequency of non-elevator-private
* requests very low.
*/
goto start_rq;
}
bfq_log(bfqd, "dispatch requests: %d busy queues",
bfq_tot_busy_queues(bfqd));
if (bfq_tot_busy_queues(bfqd) == 0)
goto exit;
/*
* Force device to serve one request at a time if
* strict_guarantees is true. Forcing this service scheme is
* currently the ONLY way to guarantee that the request
* service order enforced by the scheduler is respected by a
* queueing device. Otherwise the device is free even to make
* some unlucky request wait for as long as the device
* wishes.
*
* Of course, serving one request at at time may cause loss of
* throughput.
*/
if (bfqd->strict_guarantees && bfqd->rq_in_driver > 0)
goto exit;
bfqq = bfq_select_queue(bfqd);
if (!bfqq)
goto exit;
rq = bfq_dispatch_rq_from_bfqq(bfqd, bfqq);
if (rq) {
inc_in_driver_start_rq:
bfqd->rq_in_driver++;
start_rq:
rq->rq_flags |= RQF_STARTED;
}
exit:
return rq;
}
#ifdef CONFIG_BFQ_CGROUP_DEBUG
static void bfq_update_dispatch_stats(struct request_queue *q,
struct request *rq,
struct bfq_queue *in_serv_queue,
bool idle_timer_disabled)
{
struct bfq_queue *bfqq = rq ? RQ_BFQQ(rq) : NULL;
if (!idle_timer_disabled && !bfqq)
return;
/*
* rq and bfqq are guaranteed to exist until this function
* ends, for the following reasons. First, rq can be
* dispatched to the device, and then can be completed and
* freed, only after this function ends. Second, rq cannot be
* merged (and thus freed because of a merge) any longer,
* because it has already started. Thus rq cannot be freed
* before this function ends, and, since rq has a reference to
* bfqq, the same guarantee holds for bfqq too.
*
* In addition, the following queue lock guarantees that
* bfqq_group(bfqq) exists as well.
*/
spin_lock_irq(&q->queue_lock);
if (idle_timer_disabled)
/*
* Since the idle timer has been disabled,
* in_serv_queue contained some request when
* __bfq_dispatch_request was invoked above, which
* implies that rq was picked exactly from
* in_serv_queue. Thus in_serv_queue == bfqq, and is
* therefore guaranteed to exist because of the above
* arguments.
*/
bfqg_stats_update_idle_time(bfqq_group(in_serv_queue));
if (bfqq) {
struct bfq_group *bfqg = bfqq_group(bfqq);
bfqg_stats_update_avg_queue_size(bfqg);
bfqg_stats_set_start_empty_time(bfqg);
bfqg_stats_update_io_remove(bfqg, rq->cmd_flags);
}
spin_unlock_irq(&q->queue_lock);
}
#else
static inline void bfq_update_dispatch_stats(struct request_queue *q,
struct request *rq,
struct bfq_queue *in_serv_queue,
bool idle_timer_disabled) {}
#endif /* CONFIG_BFQ_CGROUP_DEBUG */
static struct request *bfq_dispatch_request(struct blk_mq_hw_ctx *hctx)
{
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data;
struct request *rq;
struct bfq_queue *in_serv_queue;
bool waiting_rq, idle_timer_disabled;
spin_lock_irq(&bfqd->lock);
in_serv_queue = bfqd->in_service_queue;
waiting_rq = in_serv_queue && bfq_bfqq_wait_request(in_serv_queue);
rq = __bfq_dispatch_request(hctx);
idle_timer_disabled =
waiting_rq && !bfq_bfqq_wait_request(in_serv_queue);
spin_unlock_irq(&bfqd->lock);
bfq_update_dispatch_stats(hctx->queue, rq, in_serv_queue,
idle_timer_disabled);
return rq;
}
/*
* Task holds one reference to the queue, dropped when task exits. Each rq
* in-flight on this queue also holds a reference, dropped when rq is freed.
*
* Scheduler lock must be held here. Recall not to use bfqq after calling
* this function on it.
*/
void bfq_put_queue(struct bfq_queue *bfqq)
{
struct bfq_queue *item;
struct hlist_node *n;
#ifdef CONFIG_BFQ_GROUP_IOSCHED
struct bfq_group *bfqg = bfqq_group(bfqq);
#endif
if (bfqq->bfqd)
bfq_log_bfqq(bfqq->bfqd, bfqq, "put_queue: %p %d",
bfqq, bfqq->ref);
bfqq->ref--;
if (bfqq->ref)
return;
if (!hlist_unhashed(&bfqq->burst_list_node)) {
hlist_del_init(&bfqq->burst_list_node);
/*
* Decrement also burst size after the removal, if the
* process associated with bfqq is exiting, and thus
* does not contribute to the burst any longer. This
* decrement helps filter out false positives of large
* bursts, when some short-lived process (often due to
* the execution of commands by some service) happens
* to start and exit while a complex application is
* starting, and thus spawning several processes that
* do I/O (and that *must not* be treated as a large
* burst, see comments on bfq_handle_burst).
*
* In particular, the decrement is performed only if:
* 1) bfqq is not a merged queue, because, if it is,
* then this free of bfqq is not triggered by the exit
* of the process bfqq is associated with, but exactly
* by the fact that bfqq has just been merged.
* 2) burst_size is greater than 0, to handle
* unbalanced decrements. Unbalanced decrements may
* happen in te following case: bfqq is inserted into
* the current burst list--without incrementing
* bust_size--because of a split, but the current
* burst list is not the burst list bfqq belonged to
* (see comments on the case of a split in
* bfq_set_request).
*/
if (bfqq->bic && bfqq->bfqd->burst_size > 0)
bfqq->bfqd->burst_size--;
}
/*
* bfqq does not exist any longer, so it cannot be woken by
* any other queue, and cannot wake any other queue. Then bfqq
* must be removed from the woken list of its possible waker
* queue, and all queues in the woken list of bfqq must stop
* having a waker queue. Strictly speaking, these updates
* should be performed when bfqq remains with no I/O source
* attached to it, which happens before bfqq gets freed. In
* particular, this happens when the last process associated
* with bfqq exits or gets associated with a different
* queue. However, both events lead to bfqq being freed soon,
* and dangling references would come out only after bfqq gets
* freed. So these updates are done here, as a simple and safe
* way to handle all cases.
*/
/* remove bfqq from woken list */
if (!hlist_unhashed(&bfqq->woken_list_node))
hlist_del_init(&bfqq->woken_list_node);
/* reset waker for all queues in woken list */
hlist_for_each_entry_safe(item, n, &bfqq->woken_list,
woken_list_node) {
item->waker_bfqq = NULL;
bfq_clear_bfqq_has_waker(item);
hlist_del_init(&item->woken_list_node);
}
if (bfqq->bfqd && bfqq->bfqd->last_completed_rq_bfqq == bfqq)
bfqq->bfqd->last_completed_rq_bfqq = NULL;
kmem_cache_free(bfq_pool, bfqq);
#ifdef CONFIG_BFQ_GROUP_IOSCHED
bfqg_and_blkg_put(bfqg);
#endif
}
static void bfq_put_cooperator(struct bfq_queue *bfqq)
{
struct bfq_queue *__bfqq, *next;
/*
* If this queue was scheduled to merge with another queue, be
* sure to drop the reference taken on that queue (and others in
* the merge chain). See bfq_setup_merge and bfq_merge_bfqqs.
*/
__bfqq = bfqq->new_bfqq;
while (__bfqq) {
if (__bfqq == bfqq)
break;
next = __bfqq->new_bfqq;
bfq_put_queue(__bfqq);
__bfqq = next;
}
}
static void bfq_exit_bfqq(struct bfq_data *bfqd, struct bfq_queue *bfqq)
{
if (bfqq == bfqd->in_service_queue) {
__bfq_bfqq_expire(bfqd, bfqq, BFQQE_BUDGET_TIMEOUT);
bfq_schedule_dispatch(bfqd);
}
bfq_log_bfqq(bfqd, bfqq, "exit_bfqq: %p, %d", bfqq, bfqq->ref);
bfq_put_cooperator(bfqq);
bfq_put_queue(bfqq); /* release process reference */
}
static void bfq_exit_icq_bfqq(struct bfq_io_cq *bic, bool is_sync)
{
struct bfq_queue *bfqq = bic_to_bfqq(bic, is_sync);
struct bfq_data *bfqd;
if (bfqq)
bfqd = bfqq->bfqd; /* NULL if scheduler already exited */
if (bfqq && bfqd) {
unsigned long flags;
spin_lock_irqsave(&bfqd->lock, flags);
bfqq->bic = NULL;
bfq_exit_bfqq(bfqd, bfqq);
bic_set_bfqq(bic, NULL, is_sync);
spin_unlock_irqrestore(&bfqd->lock, flags);
}
}
static void bfq_exit_icq(struct io_cq *icq)
{
struct bfq_io_cq *bic = icq_to_bic(icq);
bfq_exit_icq_bfqq(bic, true);
bfq_exit_icq_bfqq(bic, false);
}
/*
* Update the entity prio values; note that the new values will not
* be used until the next (re)activation.
*/
static void
bfq_set_next_ioprio_data(struct bfq_queue *bfqq, struct bfq_io_cq *bic)
{
struct task_struct *tsk = current;
int ioprio_class;
struct bfq_data *bfqd = bfqq->bfqd;
if (!bfqd)
return;
ioprio_class = IOPRIO_PRIO_CLASS(bic->ioprio);
switch (ioprio_class) {
default:
dev_err(bfqq->bfqd->queue->backing_dev_info->dev,
"bfq: bad prio class %d\n", ioprio_class);
/* fall through */
case IOPRIO_CLASS_NONE:
/*
* No prio set, inherit CPU scheduling settings.
*/
bfqq->new_ioprio = task_nice_ioprio(tsk);
bfqq->new_ioprio_class = task_nice_ioclass(tsk);
break;
case IOPRIO_CLASS_RT:
bfqq->new_ioprio = IOPRIO_PRIO_DATA(bic->ioprio);
bfqq->new_ioprio_class = IOPRIO_CLASS_RT;
break;
case IOPRIO_CLASS_BE:
bfqq->new_ioprio = IOPRIO_PRIO_DATA(bic->ioprio);
bfqq->new_ioprio_class = IOPRIO_CLASS_BE;
break;
case IOPRIO_CLASS_IDLE:
bfqq->new_ioprio_class = IOPRIO_CLASS_IDLE;
bfqq->new_ioprio = 7;
break;
}
if (bfqq->new_ioprio >= IOPRIO_BE_NR) {
pr_crit("bfq_set_next_ioprio_data: new_ioprio %d\n",
bfqq->new_ioprio);
bfqq->new_ioprio = IOPRIO_BE_NR;
}
bfqq->entity.new_weight = bfq_ioprio_to_weight(bfqq->new_ioprio);
bfqq->entity.prio_changed = 1;
}
static struct bfq_queue *bfq_get_queue(struct bfq_data *bfqd,
struct bio *bio, bool is_sync,
struct bfq_io_cq *bic);
static void bfq_check_ioprio_change(struct bfq_io_cq *bic, struct bio *bio)
{
struct bfq_data *bfqd = bic_to_bfqd(bic);
struct bfq_queue *bfqq;
int ioprio = bic->icq.ioc->ioprio;
/*
* This condition may trigger on a newly created bic, be sure to
* drop the lock before returning.
*/
if (unlikely(!bfqd) || likely(bic->ioprio == ioprio))
return;
bic->ioprio = ioprio;
bfqq = bic_to_bfqq(bic, false);
if (bfqq) {
/* release process reference on this queue */
bfq_put_queue(bfqq);
bfqq = bfq_get_queue(bfqd, bio, BLK_RW_ASYNC, bic);
bic_set_bfqq(bic, bfqq, false);
}
bfqq = bic_to_bfqq(bic, true);
if (bfqq)
bfq_set_next_ioprio_data(bfqq, bic);
}
static void bfq_init_bfqq(struct bfq_data *bfqd, struct bfq_queue *bfqq,
struct bfq_io_cq *bic, pid_t pid, int is_sync)
{
RB_CLEAR_NODE(&bfqq->entity.rb_node);
INIT_LIST_HEAD(&bfqq->fifo);
INIT_HLIST_NODE(&bfqq->burst_list_node);
INIT_HLIST_NODE(&bfqq->woken_list_node);
INIT_HLIST_HEAD(&bfqq->woken_list);
bfqq->ref = 0;
bfqq->bfqd = bfqd;
if (bic)
bfq_set_next_ioprio_data(bfqq, bic);
if (is_sync) {
/*
* No need to mark as has_short_ttime if in
* idle_class, because no device idling is performed
* for queues in idle class
*/
if (!bfq_class_idle(bfqq))
/* tentatively mark as has_short_ttime */
bfq_mark_bfqq_has_short_ttime(bfqq);
bfq_mark_bfqq_sync(bfqq);
bfq_mark_bfqq_just_created(bfqq);
} else
bfq_clear_bfqq_sync(bfqq);
/* set end request to minus infinity from now */
bfqq->ttime.last_end_request = ktime_get_ns() + 1;
bfq_mark_bfqq_IO_bound(bfqq);
bfqq->pid = pid;
/* Tentative initial value to trade off between thr and lat */
bfqq->max_budget = (2 * bfq_max_budget(bfqd)) / 3;
bfqq->budget_timeout = bfq_smallest_from_now();
bfqq->wr_coeff = 1;
bfqq->last_wr_start_finish = jiffies;
bfqq->wr_start_at_switch_to_srt = bfq_smallest_from_now();
bfqq->split_time = bfq_smallest_from_now();
/*
* To not forget the possibly high bandwidth consumed by a
* process/queue in the recent past,
* bfq_bfqq_softrt_next_start() returns a value at least equal
* to the current value of bfqq->soft_rt_next_start (see
* comments on bfq_bfqq_softrt_next_start). Set
* soft_rt_next_start to now, to mean that bfqq has consumed
* no bandwidth so far.
*/
bfqq->soft_rt_next_start = jiffies;
/* first request is almost certainly seeky */
bfqq->seek_history = 1;
}
static struct bfq_queue **bfq_async_queue_prio(struct bfq_data *bfqd,
struct bfq_group *bfqg,
int ioprio_class, int ioprio)
{
switch (ioprio_class) {
case IOPRIO_CLASS_RT:
return &bfqg->async_bfqq[0][ioprio];
case IOPRIO_CLASS_NONE:
ioprio = IOPRIO_NORM;
/* fall through */
case IOPRIO_CLASS_BE:
return &bfqg->async_bfqq[1][ioprio];
case IOPRIO_CLASS_IDLE:
return &bfqg->async_idle_bfqq;
default:
return NULL;
}
}
static struct bfq_queue *bfq_get_queue(struct bfq_data *bfqd,
struct bio *bio, bool is_sync,
struct bfq_io_cq *bic)
{
const int ioprio = IOPRIO_PRIO_DATA(bic->ioprio);
const int ioprio_class = IOPRIO_PRIO_CLASS(bic->ioprio);
struct bfq_queue **async_bfqq = NULL;
struct bfq_queue *bfqq;
struct bfq_group *bfqg;
rcu_read_lock();
bfqg = bfq_find_set_group(bfqd, __bio_blkcg(bio));
if (!bfqg) {
bfqq = &bfqd->oom_bfqq;
goto out;
}
if (!is_sync) {
async_bfqq = bfq_async_queue_prio(bfqd, bfqg, ioprio_class,
ioprio);
bfqq = *async_bfqq;
if (bfqq)
goto out;
}
bfqq = kmem_cache_alloc_node(bfq_pool,
GFP_NOWAIT | __GFP_ZERO | __GFP_NOWARN,
bfqd->queue->node);
if (bfqq) {
bfq_init_bfqq(bfqd, bfqq, bic, current->pid,
is_sync);
bfq_init_entity(&bfqq->entity, bfqg);
bfq_log_bfqq(bfqd, bfqq, "allocated");
} else {
bfqq = &bfqd->oom_bfqq;
bfq_log_bfqq(bfqd, bfqq, "using oom bfqq");
goto out;
}
/*
* Pin the queue now that it's allocated, scheduler exit will
* prune it.
*/
if (async_bfqq) {
bfqq->ref++; /*
* Extra group reference, w.r.t. sync
* queue. This extra reference is removed
* only if bfqq->bfqg disappears, to
* guarantee that this queue is not freed
* until its group goes away.
*/
bfq_log_bfqq(bfqd, bfqq, "get_queue, bfqq not in async: %p, %d",
bfqq, bfqq->ref);
*async_bfqq = bfqq;
}
out:
bfqq->ref++; /* get a process reference to this queue */
bfq_log_bfqq(bfqd, bfqq, "get_queue, at end: %p, %d", bfqq, bfqq->ref);
rcu_read_unlock();
return bfqq;
}
static void bfq_update_io_thinktime(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
struct bfq_ttime *ttime = &bfqq->ttime;
u64 elapsed = ktime_get_ns() - bfqq->ttime.last_end_request;
elapsed = min_t(u64, elapsed, 2ULL * bfqd->bfq_slice_idle);
ttime->ttime_samples = (7*bfqq->ttime.ttime_samples + 256) / 8;
ttime->ttime_total = div_u64(7*ttime->ttime_total + 256*elapsed, 8);
ttime->ttime_mean = div64_ul(ttime->ttime_total + 128,
ttime->ttime_samples);
}
static void
bfq_update_io_seektime(struct bfq_data *bfqd, struct bfq_queue *bfqq,
struct request *rq)
{
bfqq->seek_history <<= 1;
bfqq->seek_history |= BFQ_RQ_SEEKY(bfqd, bfqq->last_request_pos, rq);
if (bfqq->wr_coeff > 1 &&
bfqq->wr_cur_max_time == bfqd->bfq_wr_rt_max_time &&
BFQQ_TOTALLY_SEEKY(bfqq))
bfq_bfqq_end_wr(bfqq);
}
static void bfq_update_has_short_ttime(struct bfq_data *bfqd,
struct bfq_queue *bfqq,
struct bfq_io_cq *bic)
{
bool has_short_ttime = true, state_changed;
/*
* No need to update has_short_ttime if bfqq is async or in
* idle io prio class, or if bfq_slice_idle is zero, because
* no device idling is performed for bfqq in this case.
*/
if (!bfq_bfqq_sync(bfqq) || bfq_class_idle(bfqq) ||
bfqd->bfq_slice_idle == 0)
return;
/* Idle window just restored, statistics are meaningless. */
if (time_is_after_eq_jiffies(bfqq->split_time +
bfqd->bfq_wr_min_idle_time))
return;
/* Think time is infinite if no process is linked to
* bfqq. Otherwise check average think time to
* decide whether to mark as has_short_ttime
*/
if (atomic_read(&bic->icq.ioc->active_ref) == 0 ||
(bfq_sample_valid(bfqq->ttime.ttime_samples) &&
bfqq->ttime.ttime_mean > bfqd->bfq_slice_idle))
has_short_ttime = false;
state_changed = has_short_ttime != bfq_bfqq_has_short_ttime(bfqq);
if (has_short_ttime)
bfq_mark_bfqq_has_short_ttime(bfqq);
else
bfq_clear_bfqq_has_short_ttime(bfqq);
/*
* Until the base value for the total service time gets
* finally computed for bfqq, the inject limit does depend on
* the think-time state (short|long). In particular, the limit
* is 0 or 1 if the think time is deemed, respectively, as
* short or long (details in the comments in
* bfq_update_inject_limit()). Accordingly, the next
* instructions reset the inject limit if the think-time state
* has changed and the above base value is still to be
* computed.
*
* However, the reset is performed only if more than 100 ms
* have elapsed since the last update of the inject limit, or
* (inclusive) if the change is from short to long think
* time. The reason for this waiting is as follows.
*
* bfqq may have a long think time because of a
* synchronization with some other queue, i.e., because the
* I/O of some other queue may need to be completed for bfqq
* to receive new I/O. Details in the comments on the choice
* of the queue for injection in bfq_select_queue().
*
* As stressed in those comments, if such a synchronization is
* actually in place, then, without injection on bfqq, the
* blocking I/O cannot happen to served while bfqq is in
* service. As a consequence, if bfqq is granted
* I/O-dispatch-plugging, then bfqq remains empty, and no I/O
* is dispatched, until the idle timeout fires. This is likely
* to result in lower bandwidth and higher latencies for bfqq,
* and in a severe loss of total throughput.
*
* On the opposite end, a non-zero inject limit may allow the
* I/O that blocks bfqq to be executed soon, and therefore
* bfqq to receive new I/O soon.
*
* But, if the blocking gets actually eliminated, then the
* next think-time sample for bfqq may be very low. This in
* turn may cause bfqq's think time to be deemed
* short. Without the 100 ms barrier, this new state change
* would cause the body of the next if to be executed
* immediately. But this would set to 0 the inject
* limit. Without injection, the blocking I/O would cause the
* think time of bfqq to become long again, and therefore the
* inject limit to be raised again, and so on. The only effect
* of such a steady oscillation between the two think-time
* states would be to prevent effective injection on bfqq.
*
* In contrast, if the inject limit is not reset during such a
* long time interval as 100 ms, then the number of short
* think time samples can grow significantly before the reset
* is performed. As a consequence, the think time state can
* become stable before the reset. Therefore there will be no
* state change when the 100 ms elapse, and no reset of the
* inject limit. The inject limit remains steadily equal to 1
* both during and after the 100 ms. So injection can be
* performed at all times, and throughput gets boosted.
*
* An inject limit equal to 1 is however in conflict, in
* general, with the fact that the think time of bfqq is
* short, because injection may be likely to delay bfqq's I/O
* (as explained in the comments in
* bfq_update_inject_limit()). But this does not happen in
* this special case, because bfqq's low think time is due to
* an effective handling of a synchronization, through
* injection. In this special case, bfqq's I/O does not get
* delayed by injection; on the contrary, bfqq's I/O is
* brought forward, because it is not blocked for
* milliseconds.
*
* In addition, serving the blocking I/O much sooner, and much
* more frequently than once per I/O-plugging timeout, makes
* it much quicker to detect a waker queue (the concept of
* waker queue is defined in the comments in
* bfq_add_request()). This makes it possible to start sooner
* to boost throughput more effectively, by injecting the I/O
* of the waker queue unconditionally on every
* bfq_dispatch_request().
*
* One last, important benefit of not resetting the inject
* limit before 100 ms is that, during this time interval, the
* base value for the total service time is likely to get
* finally computed for bfqq, freeing the inject limit from
* its relation with the think time.
*/
if (state_changed && bfqq->last_serv_time_ns == 0 &&
(time_is_before_eq_jiffies(bfqq->decrease_time_jif +
msecs_to_jiffies(100)) ||
!has_short_ttime))
bfq_reset_inject_limit(bfqd, bfqq);
}
/*
* Called when a new fs request (rq) is added to bfqq. Check if there's
* something we should do about it.
*/
static void bfq_rq_enqueued(struct bfq_data *bfqd, struct bfq_queue *bfqq,
struct request *rq)
{
if (rq->cmd_flags & REQ_META)
bfqq->meta_pending++;
bfqq->last_request_pos = blk_rq_pos(rq) + blk_rq_sectors(rq);
if (bfqq == bfqd->in_service_queue && bfq_bfqq_wait_request(bfqq)) {
bool small_req = bfqq->queued[rq_is_sync(rq)] == 1 &&
blk_rq_sectors(rq) < 32;
bool budget_timeout = bfq_bfqq_budget_timeout(bfqq);
/*
* There is just this request queued: if
* - the request is small, and
* - we are idling to boost throughput, and
* - the queue is not to be expired,
* then just exit.
*
* In this way, if the device is being idled to wait
* for a new request from the in-service queue, we
* avoid unplugging the device and committing the
* device to serve just a small request. In contrast
* we wait for the block layer to decide when to
* unplug the device: hopefully, new requests will be
* merged to this one quickly, then the device will be
* unplugged and larger requests will be dispatched.
*/
if (small_req && idling_boosts_thr_without_issues(bfqd, bfqq) &&
!budget_timeout)
return;
/*
* A large enough request arrived, or idling is being
* performed to preserve service guarantees, or
* finally the queue is to be expired: in all these
* cases disk idling is to be stopped, so clear
* wait_request flag and reset timer.
*/
bfq_clear_bfqq_wait_request(bfqq);
hrtimer_try_to_cancel(&bfqd->idle_slice_timer);
/*
* The queue is not empty, because a new request just
* arrived. Hence we can safely expire the queue, in
* case of budget timeout, without risking that the
* timestamps of the queue are not updated correctly.
* See [1] for more details.
*/
if (budget_timeout)
bfq_bfqq_expire(bfqd, bfqq, false,
BFQQE_BUDGET_TIMEOUT);
}
}
/* returns true if it causes the idle timer to be disabled */
static bool __bfq_insert_request(struct bfq_data *bfqd, struct request *rq)
{
struct bfq_queue *bfqq = RQ_BFQQ(rq),
*new_bfqq = bfq_setup_cooperator(bfqd, bfqq, rq, true);
bool waiting, idle_timer_disabled = false;
if (new_bfqq) {
/*
* Release the request's reference to the old bfqq
* and make sure one is taken to the shared queue.
*/
new_bfqq->allocated++;
bfqq->allocated--;
new_bfqq->ref++;
/*
* If the bic associated with the process
* issuing this request still points to bfqq
* (and thus has not been already redirected
* to new_bfqq or even some other bfq_queue),
* then complete the merge and redirect it to
* new_bfqq.
*/
if (bic_to_bfqq(RQ_BIC(rq), 1) == bfqq)
bfq_merge_bfqqs(bfqd, RQ_BIC(rq),
bfqq, new_bfqq);
bfq_clear_bfqq_just_created(bfqq);
/*
* rq is about to be enqueued into new_bfqq,
* release rq reference on bfqq
*/
bfq_put_queue(bfqq);
rq->elv.priv[1] = new_bfqq;
bfqq = new_bfqq;
}
bfq_update_io_thinktime(bfqd, bfqq);
bfq_update_has_short_ttime(bfqd, bfqq, RQ_BIC(rq));
bfq_update_io_seektime(bfqd, bfqq, rq);
waiting = bfqq && bfq_bfqq_wait_request(bfqq);
bfq_add_request(rq);
idle_timer_disabled = waiting && !bfq_bfqq_wait_request(bfqq);
rq->fifo_time = ktime_get_ns() + bfqd->bfq_fifo_expire[rq_is_sync(rq)];
list_add_tail(&rq->queuelist, &bfqq->fifo);
bfq_rq_enqueued(bfqd, bfqq, rq);
return idle_timer_disabled;
}
#ifdef CONFIG_BFQ_CGROUP_DEBUG
static void bfq_update_insert_stats(struct request_queue *q,
struct bfq_queue *bfqq,
bool idle_timer_disabled,
unsigned int cmd_flags)
{
if (!bfqq)
return;
/*
* bfqq still exists, because it can disappear only after
* either it is merged with another queue, or the process it
* is associated with exits. But both actions must be taken by
* the same process currently executing this flow of
* instructions.
*
* In addition, the following queue lock guarantees that
* bfqq_group(bfqq) exists as well.
*/
spin_lock_irq(&q->queue_lock);
bfqg_stats_update_io_add(bfqq_group(bfqq), bfqq, cmd_flags);
if (idle_timer_disabled)
bfqg_stats_update_idle_time(bfqq_group(bfqq));
spin_unlock_irq(&q->queue_lock);
}
#else
static inline void bfq_update_insert_stats(struct request_queue *q,
struct bfq_queue *bfqq,
bool idle_timer_disabled,
unsigned int cmd_flags) {}
#endif /* CONFIG_BFQ_CGROUP_DEBUG */
static void bfq_insert_request(struct blk_mq_hw_ctx *hctx, struct request *rq,
bool at_head)
{
struct request_queue *q = hctx->queue;
struct bfq_data *bfqd = q->elevator->elevator_data;
struct bfq_queue *bfqq;
bool idle_timer_disabled = false;
unsigned int cmd_flags;
spin_lock_irq(&bfqd->lock);
if (blk_mq_sched_try_insert_merge(q, rq)) {
spin_unlock_irq(&bfqd->lock);
return;
}
spin_unlock_irq(&bfqd->lock);
blk_mq_sched_request_inserted(rq);
spin_lock_irq(&bfqd->lock);
bfqq = bfq_init_rq(rq);
if (!bfqq || at_head || blk_rq_is_passthrough(rq)) {
if (at_head)
list_add(&rq->queuelist, &bfqd->dispatch);
else
list_add_tail(&rq->queuelist, &bfqd->dispatch);
} else {
idle_timer_disabled = __bfq_insert_request(bfqd, rq);
/*
* Update bfqq, because, if a queue merge has occurred
* in __bfq_insert_request, then rq has been
* redirected into a new queue.
*/
bfqq = RQ_BFQQ(rq);
if (rq_mergeable(rq)) {
elv_rqhash_add(q, rq);
if (!q->last_merge)
q->last_merge = rq;
}
}
/*
* Cache cmd_flags before releasing scheduler lock, because rq
* may disappear afterwards (for example, because of a request
* merge).
*/
cmd_flags = rq->cmd_flags;
spin_unlock_irq(&bfqd->lock);
bfq_update_insert_stats(q, bfqq, idle_timer_disabled,
cmd_flags);
}
static void bfq_insert_requests(struct blk_mq_hw_ctx *hctx,
struct list_head *list, bool at_head)
{
while (!list_empty(list)) {
struct request *rq;
rq = list_first_entry(list, struct request, queuelist);
list_del_init(&rq->queuelist);
bfq_insert_request(hctx, rq, at_head);
}
}
static void bfq_update_hw_tag(struct bfq_data *bfqd)
{
struct bfq_queue *bfqq = bfqd->in_service_queue;
bfqd->max_rq_in_driver = max_t(int, bfqd->max_rq_in_driver,
bfqd->rq_in_driver);
if (bfqd->hw_tag == 1)
return;
/*
* This sample is valid if the number of outstanding requests
* is large enough to allow a queueing behavior. Note that the
* sum is not exact, as it's not taking into account deactivated
* requests.
*/
if (bfqd->rq_in_driver + bfqd->queued <= BFQ_HW_QUEUE_THRESHOLD)
return;
/*
* If active queue hasn't enough requests and can idle, bfq might not
* dispatch sufficient requests to hardware. Don't zero hw_tag in this
* case
*/
if (bfqq && bfq_bfqq_has_short_ttime(bfqq) &&
bfqq->dispatched + bfqq->queued[0] + bfqq->queued[1] <
BFQ_HW_QUEUE_THRESHOLD &&
bfqd->rq_in_driver < BFQ_HW_QUEUE_THRESHOLD)
return;
if (bfqd->hw_tag_samples++ < BFQ_HW_QUEUE_SAMPLES)
return;
bfqd->hw_tag = bfqd->max_rq_in_driver > BFQ_HW_QUEUE_THRESHOLD;
bfqd->max_rq_in_driver = 0;
bfqd->hw_tag_samples = 0;
bfqd->nonrot_with_queueing =
blk_queue_nonrot(bfqd->queue) && bfqd->hw_tag;
}
static void bfq_completed_request(struct bfq_queue *bfqq, struct bfq_data *bfqd)
{
u64 now_ns;
u32 delta_us;
bfq_update_hw_tag(bfqd);
bfqd->rq_in_driver--;
bfqq->dispatched--;
if (!bfqq->dispatched && !bfq_bfqq_busy(bfqq)) {
/*
* Set budget_timeout (which we overload to store the
* time at which the queue remains with no backlog and
* no outstanding request; used by the weight-raising
* mechanism).
*/
bfqq->budget_timeout = jiffies;
bfq_weights_tree_remove(bfqd, bfqq);
}
now_ns = ktime_get_ns();
bfqq->ttime.last_end_request = now_ns;
/*
* Using us instead of ns, to get a reasonable precision in
* computing rate in next check.
*/
delta_us = div_u64(now_ns - bfqd->last_completion, NSEC_PER_USEC);
/*
* If the request took rather long to complete, and, according
* to the maximum request size recorded, this completion latency
* implies that the request was certainly served at a very low
* rate (less than 1M sectors/sec), then the whole observation
* interval that lasts up to this time instant cannot be a
* valid time interval for computing a new peak rate. Invoke
* bfq_update_rate_reset to have the following three steps
* taken:
* - close the observation interval at the last (previous)
* request dispatch or completion
* - compute rate, if possible, for that observation interval
* - reset to zero samples, which will trigger a proper
* re-initialization of the observation interval on next
* dispatch
*/
if (delta_us > BFQ_MIN_TT/NSEC_PER_USEC &&
(bfqd->last_rq_max_size<<BFQ_RATE_SHIFT)/delta_us <
1UL<<(BFQ_RATE_SHIFT - 10))
bfq_update_rate_reset(bfqd, NULL);
bfqd->last_completion = now_ns;
bfqd->last_completed_rq_bfqq = bfqq;
/*
* If we are waiting to discover whether the request pattern
* of the task associated with the queue is actually
* isochronous, and both requisites for this condition to hold
* are now satisfied, then compute soft_rt_next_start (see the
* comments on the function bfq_bfqq_softrt_next_start()). We
* do not compute soft_rt_next_start if bfqq is in interactive
* weight raising (see the comments in bfq_bfqq_expire() for
* an explanation). We schedule this delayed update when bfqq
* expires, if it still has in-flight requests.
*/
if (bfq_bfqq_softrt_update(bfqq) && bfqq->dispatched == 0 &&
RB_EMPTY_ROOT(&bfqq->sort_list) &&
bfqq->wr_coeff != bfqd->bfq_wr_coeff)
bfqq->soft_rt_next_start =
bfq_bfqq_softrt_next_start(bfqd, bfqq);
/*
* If this is the in-service queue, check if it needs to be expired,
* or if we want to idle in case it has no pending requests.
*/
if (bfqd->in_service_queue == bfqq) {
if (bfq_bfqq_must_idle(bfqq)) {
if (bfqq->dispatched == 0)
bfq_arm_slice_timer(bfqd);
/*
* If we get here, we do not expire bfqq, even
* if bfqq was in budget timeout or had no
* more requests (as controlled in the next
* conditional instructions). The reason for
* not expiring bfqq is as follows.
*
* Here bfqq->dispatched > 0 holds, but
* bfq_bfqq_must_idle() returned true. This
* implies that, even if no request arrives
* for bfqq before bfqq->dispatched reaches 0,
* bfqq will, however, not be expired on the
* completion event that causes bfqq->dispatch
* to reach zero. In contrast, on this event,
* bfqq will start enjoying device idling
* (I/O-dispatch plugging).
*
* But, if we expired bfqq here, bfqq would
* not have the chance to enjoy device idling
* when bfqq->dispatched finally reaches
* zero. This would expose bfqq to violation
* of its reserved service guarantees.
*/
return;
} else if (bfq_may_expire_for_budg_timeout(bfqq))
bfq_bfqq_expire(bfqd, bfqq, false,
BFQQE_BUDGET_TIMEOUT);
else if (RB_EMPTY_ROOT(&bfqq->sort_list) &&
(bfqq->dispatched == 0 ||
!bfq_better_to_idle(bfqq)))
bfq_bfqq_expire(bfqd, bfqq, false,
BFQQE_NO_MORE_REQUESTS);
}
if (!bfqd->rq_in_driver)
bfq_schedule_dispatch(bfqd);
}
static void bfq_finish_requeue_request_body(struct bfq_queue *bfqq)
{
bfqq->allocated--;
bfq_put_queue(bfqq);
}
/*
* The processes associated with bfqq may happen to generate their
* cumulative I/O at a lower rate than the rate at which the device
* could serve the same I/O. This is rather probable, e.g., if only
* one process is associated with bfqq and the device is an SSD. It
* results in bfqq becoming often empty while in service. In this
* respect, if BFQ is allowed to switch to another queue when bfqq
* remains empty, then the device goes on being fed with I/O requests,
* and the throughput is not affected. In contrast, if BFQ is not
* allowed to switch to another queue---because bfqq is sync and
* I/O-dispatch needs to be plugged while bfqq is temporarily
* empty---then, during the service of bfqq, there will be frequent
* "service holes", i.e., time intervals during which bfqq gets empty
* and the device can only consume the I/O already queued in its
* hardware queues. During service holes, the device may even get to
* remaining idle. In the end, during the service of bfqq, the device
* is driven at a lower speed than the one it can reach with the kind
* of I/O flowing through bfqq.
*
* To counter this loss of throughput, BFQ implements a "request
* injection mechanism", which tries to fill the above service holes
* with I/O requests taken from other queues. The hard part in this
* mechanism is finding the right amount of I/O to inject, so as to
* both boost throughput and not break bfqq's bandwidth and latency
* guarantees. In this respect, the mechanism maintains a per-queue
* inject limit, computed as below. While bfqq is empty, the injection
* mechanism dispatches extra I/O requests only until the total number
* of I/O requests in flight---i.e., already dispatched but not yet
* completed---remains lower than this limit.
*
* A first definition comes in handy to introduce the algorithm by
* which the inject limit is computed. We define as first request for
* bfqq, an I/O request for bfqq that arrives while bfqq is in
* service, and causes bfqq to switch from empty to non-empty. The
* algorithm updates the limit as a function of the effect of
* injection on the service times of only the first requests of
* bfqq. The reason for this restriction is that these are the
* requests whose service time is affected most, because they are the
* first to arrive after injection possibly occurred.
*
* To evaluate the effect of injection, the algorithm measures the
* "total service time" of first requests. We define as total service
* time of an I/O request, the time that elapses since when the
* request is enqueued into bfqq, to when it is completed. This
* quantity allows the whole effect of injection to be measured. It is
* easy to see why. Suppose that some requests of other queues are
* actually injected while bfqq is empty, and that a new request R
* then arrives for bfqq. If the device does start to serve all or
* part of the injected requests during the service hole, then,
* because of this extra service, it may delay the next invocation of
* the dispatch hook of BFQ. Then, even after R gets eventually
* dispatched, the device may delay the actual service of R if it is
* still busy serving the extra requests, or if it decides to serve,
* before R, some extra request still present in its queues. As a
* conclusion, the cumulative extra delay caused by injection can be
* easily evaluated by just comparing the total service time of first
* requests with and without injection.
*
* The limit-update algorithm works as follows. On the arrival of a
* first request of bfqq, the algorithm measures the total time of the
* request only if one of the three cases below holds, and, for each
* case, it updates the limit as described below:
*
* (1) If there is no in-flight request. This gives a baseline for the
* total service time of the requests of bfqq. If the baseline has
* not been computed yet, then, after computing it, the limit is
* set to 1, to start boosting throughput, and to prepare the
* ground for the next case. If the baseline has already been
* computed, then it is updated, in case it results to be lower
* than the previous value.
*
* (2) If the limit is higher than 0 and there are in-flight
* requests. By comparing the total service time in this case with
* the above baseline, it is possible to know at which extent the
* current value of the limit is inflating the total service
* time. If the inflation is below a certain threshold, then bfqq
* is assumed to be suffering from no perceivable loss of its
* service guarantees, and the limit is even tentatively
* increased. If the inflation is above the threshold, then the
* limit is decreased. Due to the lack of any hysteresis, this
* logic makes the limit oscillate even in steady workload
* conditions. Yet we opted for it, because it is fast in reaching
* the best value for the limit, as a function of the current I/O
* workload. To reduce oscillations, this step is disabled for a
* short time interval after the limit happens to be decreased.
*
* (3) Periodically, after resetting the limit, to make sure that the
* limit eventually drops in case the workload changes. This is
* needed because, after the limit has gone safely up for a
* certain workload, it is impossible to guess whether the
* baseline total service time may have changed, without measuring
* it again without injection. A more effective version of this
* step might be to just sample the baseline, by interrupting
* injection only once, and then to reset/lower the limit only if
* the total service time with the current limit does happen to be
* too large.
*
* More details on each step are provided in the comments on the
* pieces of code that implement these steps: the branch handling the
* transition from empty to non empty in bfq_add_request(), the branch
* handling injection in bfq_select_queue(), and the function
* bfq_choose_bfqq_for_injection(). These comments also explain some
* exceptions, made by the injection mechanism in some special cases.
*/
static void bfq_update_inject_limit(struct bfq_data *bfqd,
struct bfq_queue *bfqq)
{
u64 tot_time_ns = ktime_get_ns() - bfqd->last_empty_occupied_ns;
unsigned int old_limit = bfqq->inject_limit;
if (bfqq->last_serv_time_ns > 0) {
u64 threshold = (bfqq->last_serv_time_ns * 3)>>1;
if (tot_time_ns >= threshold && old_limit > 0) {
bfqq->inject_limit--;
bfqq->decrease_time_jif = jiffies;
} else if (tot_time_ns < threshold &&
old_limit < bfqd->max_rq_in_driver<<1)
bfqq->inject_limit++;
}
/*
* Either we still have to compute the base value for the
* total service time, and there seem to be the right
* conditions to do it, or we can lower the last base value
* computed.
*
* NOTE: (bfqd->rq_in_driver == 1) means that there is no I/O
* request in flight, because this function is in the code
* path that handles the completion of a request of bfqq, and,
* in particular, this function is executed before
* bfqd->rq_in_driver is decremented in such a code path.
*/
if ((bfqq->last_serv_time_ns == 0 && bfqd->rq_in_driver == 1) ||
tot_time_ns < bfqq->last_serv_time_ns) {
bfqq->last_serv_time_ns = tot_time_ns;
/*
* Now we certainly have a base value: make sure we
* start trying injection.
*/
bfqq->inject_limit = max_t(unsigned int, 1, old_limit);
} else if (!bfqd->rqs_injected && bfqd->rq_in_driver == 1)
/*
* No I/O injected and no request still in service in
* the drive: these are the exact conditions for
* computing the base value of the total service time
* for bfqq. So let's update this value, because it is
* rather variable. For example, it varies if the size
* or the spatial locality of the I/O requests in bfqq
* change.
*/
bfqq->last_serv_time_ns = tot_time_ns;
/* update complete, not waiting for any request completion any longer */
bfqd->waited_rq = NULL;
}
/*
* Handle either a requeue or a finish for rq. The things to do are
* the same in both cases: all references to rq are to be dropped. In
* particular, rq is considered completed from the point of view of
* the scheduler.
*/
static void bfq_finish_requeue_request(struct request *rq)
{
struct bfq_queue *bfqq = RQ_BFQQ(rq);
struct bfq_data *bfqd;
/*
* Requeue and finish hooks are invoked in blk-mq without
* checking whether the involved request is actually still
* referenced in the scheduler. To handle this fact, the
* following two checks make this function exit in case of
* spurious invocations, for which there is nothing to do.
*
* First, check whether rq has nothing to do with an elevator.
*/
if (unlikely(!(rq->rq_flags & RQF_ELVPRIV)))
return;
/*
* rq either is not associated with any icq, or is an already
* requeued request that has not (yet) been re-inserted into
* a bfq_queue.
*/
if (!rq->elv.icq || !bfqq)
return;
bfqd = bfqq->bfqd;
if (rq->rq_flags & RQF_STARTED)
bfqg_stats_update_completion(bfqq_group(bfqq),
rq->start_time_ns,
rq->io_start_time_ns,
rq->cmd_flags);
if (likely(rq->rq_flags & RQF_STARTED)) {
unsigned long flags;
spin_lock_irqsave(&bfqd->lock, flags);
if (rq == bfqd->waited_rq)
bfq_update_inject_limit(bfqd, bfqq);
bfq_completed_request(bfqq, bfqd);
bfq_finish_requeue_request_body(bfqq);
spin_unlock_irqrestore(&bfqd->lock, flags);
} else {
/*
* Request rq may be still/already in the scheduler,
* in which case we need to remove it (this should
* never happen in case of requeue). And we cannot
* defer such a check and removal, to avoid
* inconsistencies in the time interval from the end
* of this function to the start of the deferred work.
* This situation seems to occur only in process
* context, as a consequence of a merge. In the
* current version of the code, this implies that the
* lock is held.
*/
if (!RB_EMPTY_NODE(&rq->rb_node)) {
bfq_remove_request(rq->q, rq);
bfqg_stats_update_io_remove(bfqq_group(bfqq),
rq->cmd_flags);
}
bfq_finish_requeue_request_body(bfqq);
}
/*
* Reset private fields. In case of a requeue, this allows
* this function to correctly do nothing if it is spuriously
* invoked again on this same request (see the check at the
* beginning of the function). Probably, a better general
* design would be to prevent blk-mq from invoking the requeue
* or finish hooks of an elevator, for a request that is not
* referred by that elevator.
*
* Resetting the following fields would break the
* request-insertion logic if rq is re-inserted into a bfq
* internal queue, without a re-preparation. Here we assume
* that re-insertions of requeued requests, without
* re-preparation, can happen only for pass_through or at_head
* requests (which are not re-inserted into bfq internal
* queues).
*/
rq->elv.priv[0] = NULL;
rq->elv.priv[1] = NULL;
}
/*
* Returns NULL if a new bfqq should be allocated, or the old bfqq if this
* was the last process referring to that bfqq.
*/
static struct bfq_queue *
bfq_split_bfqq(struct bfq_io_cq *bic, struct bfq_queue *bfqq)
{
bfq_log_bfqq(bfqq->bfqd, bfqq, "splitting queue");
if (bfqq_process_refs(bfqq) == 1) {
bfqq->pid = current->pid;
bfq_clear_bfqq_coop(bfqq);
bfq_clear_bfqq_split_coop(bfqq);
return bfqq;
}
bic_set_bfqq(bic, NULL, 1);
bfq_put_cooperator(bfqq);
bfq_put_queue(bfqq);
return NULL;
}
static struct bfq_queue *bfq_get_bfqq_handle_split(struct bfq_data *bfqd,
struct bfq_io_cq *bic,
struct bio *bio,
bool split, bool is_sync,
bool *new_queue)
{
struct bfq_queue *bfqq = bic_to_bfqq(bic, is_sync);
if (likely(bfqq && bfqq != &bfqd->oom_bfqq))
return bfqq;
if (new_queue)
*new_queue = true;
if (bfqq)
bfq_put_queue(bfqq);
bfqq = bfq_get_queue(bfqd, bio, is_sync, bic);
bic_set_bfqq(bic, bfqq, is_sync);
if (split && is_sync) {
if ((bic->was_in_burst_list && bfqd->large_burst) ||
bic->saved_in_large_burst)
bfq_mark_bfqq_in_large_burst(bfqq);
else {
bfq_clear_bfqq_in_large_burst(bfqq);
if (bic->was_in_burst_list)
/*
* If bfqq was in the current
* burst list before being
* merged, then we have to add
* it back. And we do not need
* to increase burst_size, as
* we did not decrement
* burst_size when we removed
* bfqq from the burst list as
* a consequence of a merge
* (see comments in
* bfq_put_queue). In this
* respect, it would be rather
* costly to know whether the
* current burst list is still
* the same burst list from
* which bfqq was removed on
* the merge. To avoid this
* cost, if bfqq was in a
* burst list, then we add
* bfqq to the current burst
* list without any further
* check. This can cause
* inappropriate insertions,
* but rarely enough to not
* harm the detection of large
* bursts significantly.
*/
hlist_add_head(&bfqq->burst_list_node,
&bfqd->burst_list);
}
bfqq->split_time = jiffies;
}
return bfqq;
}
/*
* Only reset private fields. The actual request preparation will be
* performed by bfq_init_rq, when rq is either inserted or merged. See
* comments on bfq_init_rq for the reason behind this delayed
* preparation.
*/
static void bfq_prepare_request(struct request *rq, struct bio *bio)
{
/*
* Regardless of whether we have an icq attached, we have to
* clear the scheduler pointers, as they might point to
* previously allocated bic/bfqq structs.
*/
rq->elv.priv[0] = rq->elv.priv[1] = NULL;
}
/*
* If needed, init rq, allocate bfq data structures associated with
* rq, and increment reference counters in the destination bfq_queue
* for rq. Return the destination bfq_queue for rq, or NULL is rq is
* not associated with any bfq_queue.
*
* This function is invoked by the functions that perform rq insertion
* or merging. One may have expected the above preparation operations
* to be performed in bfq_prepare_request, and not delayed to when rq
* is inserted or merged. The rationale behind this delayed
* preparation is that, after the prepare_request hook is invoked for
* rq, rq may still be transformed into a request with no icq, i.e., a
* request not associated with any queue. No bfq hook is invoked to
* signal this transformation. As a consequence, should these
* preparation operations be performed when the prepare_request hook
* is invoked, and should rq be transformed one moment later, bfq
* would end up in an inconsistent state, because it would have
* incremented some queue counters for an rq destined to
* transformation, without any chance to correctly lower these
* counters back. In contrast, no transformation can still happen for
* rq after rq has been inserted or merged. So, it is safe to execute
* these preparation operations when rq is finally inserted or merged.
*/
static struct bfq_queue *bfq_init_rq(struct request *rq)
{
struct request_queue *q = rq->q;
struct bio *bio = rq->bio;
struct bfq_data *bfqd = q->elevator->elevator_data;
struct bfq_io_cq *bic;
const int is_sync = rq_is_sync(rq);
struct bfq_queue *bfqq;
bool new_queue = false;
bool bfqq_already_existing = false, split = false;
if (unlikely(!rq->elv.icq))
return NULL;
/*
* Assuming that elv.priv[1] is set only if everything is set
* for this rq. This holds true, because this function is
* invoked only for insertion or merging, and, after such
* events, a request cannot be manipulated any longer before
* being removed from bfq.
*/
if (rq->elv.priv[1])
return rq->elv.priv[1];
bic = icq_to_bic(rq->elv.icq);
bfq_check_ioprio_change(bic, bio);
bfq_bic_update_cgroup(bic, bio);
bfqq = bfq_get_bfqq_handle_split(bfqd, bic, bio, false, is_sync,
&new_queue);
if (likely(!new_queue)) {
/* If the queue was seeky for too long, break it apart. */
if (bfq_bfqq_coop(bfqq) && bfq_bfqq_split_coop(bfqq)) {
bfq_log_bfqq(bfqd, bfqq, "breaking apart bfqq");
/* Update bic before losing reference to bfqq */
if (bfq_bfqq_in_large_burst(bfqq))
bic->saved_in_large_burst = true;
bfqq = bfq_split_bfqq(bic, bfqq);
split = true;
if (!bfqq)
bfqq = bfq_get_bfqq_handle_split(bfqd, bic, bio,
true, is_sync,
NULL);
else
bfqq_already_existing = true;
}
}
bfqq->allocated++;
bfqq->ref++;
bfq_log_bfqq(bfqd, bfqq, "get_request %p: bfqq %p, %d",
rq, bfqq, bfqq->ref);
rq->elv.priv[0] = bic;
rq->elv.priv[1] = bfqq;
/*
* If a bfq_queue has only one process reference, it is owned
* by only this bic: we can then set bfqq->bic = bic. in
* addition, if the queue has also just been split, we have to
* resume its state.
*/
if (likely(bfqq != &bfqd->oom_bfqq) && bfqq_process_refs(bfqq) == 1) {
bfqq->bic = bic;
if (split) {
/*
* The queue has just been split from a shared
* queue: restore the idle window and the
* possible weight raising period.
*/
bfq_bfqq_resume_state(bfqq, bfqd, bic,
bfqq_already_existing);
}
}
/*
* Consider bfqq as possibly belonging to a burst of newly
* created queues only if:
* 1) A burst is actually happening (bfqd->burst_size > 0)
* or
* 2) There is no other active queue. In fact, if, in
* contrast, there are active queues not belonging to the
* possible burst bfqq may belong to, then there is no gain
* in considering bfqq as belonging to a burst, and
* therefore in not weight-raising bfqq. See comments on
* bfq_handle_burst().
*
* This filtering also helps eliminating false positives,
* occurring when bfqq does not belong to an actual large
* burst, but some background task (e.g., a service) happens
* to trigger the creation of new queues very close to when
* bfqq and its possible companion queues are created. See
* comments on bfq_handle_burst() for further details also on
* this issue.
*/
if (unlikely(bfq_bfqq_just_created(bfqq) &&
(bfqd->burst_size > 0 ||
bfq_tot_busy_queues(bfqd) == 0)))
bfq_handle_burst(bfqd, bfqq);
return bfqq;
}
static void bfq_idle_slice_timer_body(struct bfq_queue *bfqq)
{
struct bfq_data *bfqd = bfqq->bfqd;
enum bfqq_expiration reason;
unsigned long flags;
spin_lock_irqsave(&bfqd->lock, flags);
bfq_clear_bfqq_wait_request(bfqq);
if (bfqq != bfqd->in_service_queue) {
spin_unlock_irqrestore(&bfqd->lock, flags);
return;
}
if (bfq_bfqq_budget_timeout(bfqq))
/*
* Also here the queue can be safely expired
* for budget timeout without wasting
* guarantees
*/
reason = BFQQE_BUDGET_TIMEOUT;
else if (bfqq->queued[0] == 0 && bfqq->queued[1] == 0)
/*
* The queue may not be empty upon timer expiration,
* because we may not disable the timer when the
* first request of the in-service queue arrives
* during disk idling.
*/
reason = BFQQE_TOO_IDLE;
else
goto schedule_dispatch;
bfq_bfqq_expire(bfqd, bfqq, true, reason);
schedule_dispatch:
spin_unlock_irqrestore(&bfqd->lock, flags);
bfq_schedule_dispatch(bfqd);
}
/*
* Handler of the expiration of the timer running if the in-service queue
* is idling inside its time slice.
*/
static enum hrtimer_restart bfq_idle_slice_timer(struct hrtimer *timer)
{
struct bfq_data *bfqd = container_of(timer, struct bfq_data,
idle_slice_timer);
struct bfq_queue *bfqq = bfqd->in_service_queue;
/*
* Theoretical race here: the in-service queue can be NULL or
* different from the queue that was idling if a new request
* arrives for the current queue and there is a full dispatch
* cycle that changes the in-service queue. This can hardly
* happen, but in the worst case we just expire a queue too
* early.
*/
if (bfqq)
bfq_idle_slice_timer_body(bfqq);
return HRTIMER_NORESTART;
}
static void __bfq_put_async_bfqq(struct bfq_data *bfqd,
struct bfq_queue **bfqq_ptr)
{
struct bfq_queue *bfqq = *bfqq_ptr;
bfq_log(bfqd, "put_async_bfqq: %p", bfqq);
if (bfqq) {
bfq_bfqq_move(bfqd, bfqq, bfqd->root_group);
bfq_log_bfqq(bfqd, bfqq, "put_async_bfqq: putting %p, %d",
bfqq, bfqq->ref);
bfq_put_queue(bfqq);
*bfqq_ptr = NULL;
}
}
/*
* Release all the bfqg references to its async queues. If we are
* deallocating the group these queues may still contain requests, so
* we reparent them to the root cgroup (i.e., the only one that will
* exist for sure until all the requests on a device are gone).
*/
void bfq_put_async_queues(struct bfq_data *bfqd, struct bfq_group *bfqg)
{
int i, j;
for (i = 0; i < 2; i++)
for (j = 0; j < IOPRIO_BE_NR; j++)
__bfq_put_async_bfqq(bfqd, &bfqg->async_bfqq[i][j]);
__bfq_put_async_bfqq(bfqd, &bfqg->async_idle_bfqq);
}
/*
* See the comments on bfq_limit_depth for the purpose of
* the depths set in the function. Return minimum shallow depth we'll use.
*/
static unsigned int bfq_update_depths(struct bfq_data *bfqd,
struct sbitmap_queue *bt)
{
unsigned int i, j, min_shallow = UINT_MAX;
/*
* In-word depths if no bfq_queue is being weight-raised:
* leaving 25% of tags only for sync reads.
*
* In next formulas, right-shift the value
* (1U<<bt->sb.shift), instead of computing directly
* (1U<<(bt->sb.shift - something)), to be robust against
* any possible value of bt->sb.shift, without having to
* limit 'something'.
*/
/* no more than 50% of tags for async I/O */
bfqd->word_depths[0][0] = max((1U << bt->sb.shift) >> 1, 1U);
/*
* no more than 75% of tags for sync writes (25% extra tags
* w.r.t. async I/O, to prevent async I/O from starving sync
* writes)
*/
bfqd->word_depths[0][1] = max(((1U << bt->sb.shift) * 3) >> 2, 1U);
/*
* In-word depths in case some bfq_queue is being weight-
* raised: leaving ~63% of tags for sync reads. This is the
* highest percentage for which, in our tests, application
* start-up times didn't suffer from any regression due to tag
* shortage.
*/
/* no more than ~18% of tags for async I/O */
bfqd->word_depths[1][0] = max(((1U << bt->sb.shift) * 3) >> 4, 1U);
/* no more than ~37% of tags for sync writes (~20% extra tags) */
bfqd->word_depths[1][1] = max(((1U << bt->sb.shift) * 6) >> 4, 1U);
for (i = 0; i < 2; i++)
for (j = 0; j < 2; j++)
min_shallow = min(min_shallow, bfqd->word_depths[i][j]);
return min_shallow;
}
static void bfq_depth_updated(struct blk_mq_hw_ctx *hctx)
{
struct bfq_data *bfqd = hctx->queue->elevator->elevator_data;
struct blk_mq_tags *tags = hctx->sched_tags;
unsigned int min_shallow;
min_shallow = bfq_update_depths(bfqd, &tags->bitmap_tags);
sbitmap_queue_min_shallow_depth(&tags->bitmap_tags, min_shallow);
}
static int bfq_init_hctx(struct blk_mq_hw_ctx *hctx, unsigned int index)
{
bfq_depth_updated(hctx);
return 0;
}
static void bfq_exit_queue(struct elevator_queue *e)
{
struct bfq_data *bfqd = e->elevator_data;
struct bfq_queue *bfqq, *n;
hrtimer_cancel(&bfqd->idle_slice_timer);
spin_lock_irq(&bfqd->lock);
list_for_each_entry_safe(bfqq, n, &bfqd->idle_list, bfqq_list)
bfq_deactivate_bfqq(bfqd, bfqq, false, false);
spin_unlock_irq(&bfqd->lock);
hrtimer_cancel(&bfqd->idle_slice_timer);
#ifdef CONFIG_BFQ_GROUP_IOSCHED
/* release oom-queue reference to root group */
bfqg_and_blkg_put(bfqd->root_group);
blkcg_deactivate_policy(bfqd->queue, &blkcg_policy_bfq);
#else
spin_lock_irq(&bfqd->lock);
bfq_put_async_queues(bfqd, bfqd->root_group);
kfree(bfqd->root_group);
spin_unlock_irq(&bfqd->lock);
#endif
kfree(bfqd);
}
static void bfq_init_root_group(struct bfq_group *root_group,
struct bfq_data *bfqd)
{
int i;
#ifdef CONFIG_BFQ_GROUP_IOSCHED
root_group->entity.parent = NULL;
root_group->my_entity = NULL;
root_group->bfqd = bfqd;
#endif
root_group->rq_pos_tree = RB_ROOT;
for (i = 0; i < BFQ_IOPRIO_CLASSES; i++)
root_group->sched_data.service_tree[i] = BFQ_SERVICE_TREE_INIT;
root_group->sched_data.bfq_class_idle_last_service = jiffies;
}
static int bfq_init_queue(struct request_queue *q, struct elevator_type *e)
{
struct bfq_data *bfqd;
struct elevator_queue *eq;
eq = elevator_alloc(q, e);
if (!eq)
return -ENOMEM;
bfqd = kzalloc_node(sizeof(*bfqd), GFP_KERNEL, q->node);
if (!bfqd) {
kobject_put(&eq->kobj);
return -ENOMEM;
}
eq->elevator_data = bfqd;
spin_lock_irq(&q->queue_lock);
q->elevator = eq;
spin_unlock_irq(&q->queue_lock);
/*
* Our fallback bfqq if bfq_find_alloc_queue() runs into OOM issues.
* Grab a permanent reference to it, so that the normal code flow
* will not attempt to free it.
*/
bfq_init_bfqq(bfqd, &bfqd->oom_bfqq, NULL, 1, 0);
bfqd->oom_bfqq.ref++;
bfqd->oom_bfqq.new_ioprio = BFQ_DEFAULT_QUEUE_IOPRIO;
bfqd->oom_bfqq.new_ioprio_class = IOPRIO_CLASS_BE;
bfqd->oom_bfqq.entity.new_weight =
bfq_ioprio_to_weight(bfqd->oom_bfqq.new_ioprio);
/* oom_bfqq does not participate to bursts */
bfq_clear_bfqq_just_created(&bfqd->oom_bfqq);
/*
* Trigger weight initialization, according to ioprio, at the
* oom_bfqq's first activation. The oom_bfqq's ioprio and ioprio
* class won't be changed any more.
*/
bfqd->oom_bfqq.entity.prio_changed = 1;
bfqd->queue = q;
INIT_LIST_HEAD(&bfqd->dispatch);
hrtimer_init(&bfqd->idle_slice_timer, CLOCK_MONOTONIC,
HRTIMER_MODE_REL);
bfqd->idle_slice_timer.function = bfq_idle_slice_timer;
bfqd->queue_weights_tree = RB_ROOT_CACHED;
bfqd->num_groups_with_pending_reqs = 0;
INIT_LIST_HEAD(&bfqd->active_list);
INIT_LIST_HEAD(&bfqd->idle_list);
INIT_HLIST_HEAD(&bfqd->burst_list);
bfqd->hw_tag = -1;
bfqd->nonrot_with_queueing = blk_queue_nonrot(bfqd->queue);
bfqd->bfq_max_budget = bfq_default_max_budget;
bfqd->bfq_fifo_expire[0] = bfq_fifo_expire[0];
bfqd->bfq_fifo_expire[1] = bfq_fifo_expire[1];
bfqd->bfq_back_max = bfq_back_max;
bfqd->bfq_back_penalty = bfq_back_penalty;
bfqd->bfq_slice_idle = bfq_slice_idle;
bfqd->bfq_timeout = bfq_timeout;
bfqd->bfq_requests_within_timer = 120;
bfqd->bfq_large_burst_thresh = 8;
bfqd->bfq_burst_interval = msecs_to_jiffies(180);
bfqd->low_latency = true;
/*
* Trade-off between responsiveness and fairness.
*/
bfqd->bfq_wr_coeff = 30;
bfqd->bfq_wr_rt_max_time = msecs_to_jiffies(300);
bfqd->bfq_wr_max_time = 0;
bfqd->bfq_wr_min_idle_time = msecs_to_jiffies(2000);
bfqd->bfq_wr_min_inter_arr_async = msecs_to_jiffies(500);
bfqd->bfq_wr_max_softrt_rate = 7000; /*
* Approximate rate required
* to playback or record a
* high-definition compressed
* video.
*/
bfqd->wr_busy_queues = 0;
/*
* Begin by assuming, optimistically, that the device peak
* rate is equal to 2/3 of the highest reference rate.
*/
bfqd->rate_dur_prod = ref_rate[blk_queue_nonrot(bfqd->queue)] *
ref_wr_duration[blk_queue_nonrot(bfqd->queue)];
bfqd->peak_rate = ref_rate[blk_queue_nonrot(bfqd->queue)] * 2 / 3;
spin_lock_init(&bfqd->lock);
/*
* The invocation of the next bfq_create_group_hierarchy
* function is the head of a chain of function calls
* (bfq_create_group_hierarchy->blkcg_activate_policy->
* blk_mq_freeze_queue) that may lead to the invocation of the
* has_work hook function. For this reason,
* bfq_create_group_hierarchy is invoked only after all
* scheduler data has been initialized, apart from the fields
* that can be initialized only after invoking
* bfq_create_group_hierarchy. This, in particular, enables
* has_work to correctly return false. Of course, to avoid
* other inconsistencies, the blk-mq stack must then refrain
* from invoking further scheduler hooks before this init
* function is finished.
*/
bfqd->root_group = bfq_create_group_hierarchy(bfqd, q->node);
if (!bfqd->root_group)
goto out_free;
bfq_init_root_group(bfqd->root_group, bfqd);
bfq_init_entity(&bfqd->oom_bfqq.entity, bfqd->root_group);
wbt_disable_default(q);
return 0;
out_free:
kfree(bfqd);
kobject_put(&eq->kobj);
return -ENOMEM;
}
static void bfq_slab_kill(void)
{
kmem_cache_destroy(bfq_pool);
}
static int __init bfq_slab_setup(void)
{
bfq_pool = KMEM_CACHE(bfq_queue, 0);
if (!bfq_pool)
return -ENOMEM;
return 0;
}
static ssize_t bfq_var_show(unsigned int var, char *page)
{
return sprintf(page, "%u\n", var);
}
static int bfq_var_store(unsigned long *var, const char *page)
{
unsigned long new_val;
int ret = kstrtoul(page, 10, &new_val);
if (ret)
return ret;
*var = new_val;
return 0;
}
#define SHOW_FUNCTION(__FUNC, __VAR, __CONV) \
static ssize_t __FUNC(struct elevator_queue *e, char *page) \
{ \
struct bfq_data *bfqd = e->elevator_data; \
u64 __data = __VAR; \
if (__CONV == 1) \
__data = jiffies_to_msecs(__data); \
else if (__CONV == 2) \
__data = div_u64(__data, NSEC_PER_MSEC); \
return bfq_var_show(__data, (page)); \
}
SHOW_FUNCTION(bfq_fifo_expire_sync_show, bfqd->bfq_fifo_expire[1], 2);
SHOW_FUNCTION(bfq_fifo_expire_async_show, bfqd->bfq_fifo_expire[0], 2);
SHOW_FUNCTION(bfq_back_seek_max_show, bfqd->bfq_back_max, 0);
SHOW_FUNCTION(bfq_back_seek_penalty_show, bfqd->bfq_back_penalty, 0);
SHOW_FUNCTION(bfq_slice_idle_show, bfqd->bfq_slice_idle, 2);
SHOW_FUNCTION(bfq_max_budget_show, bfqd->bfq_user_max_budget, 0);
SHOW_FUNCTION(bfq_timeout_sync_show, bfqd->bfq_timeout, 1);
SHOW_FUNCTION(bfq_strict_guarantees_show, bfqd->strict_guarantees, 0);
SHOW_FUNCTION(bfq_low_latency_show, bfqd->low_latency, 0);
#undef SHOW_FUNCTION
#define USEC_SHOW_FUNCTION(__FUNC, __VAR) \
static ssize_t __FUNC(struct elevator_queue *e, char *page) \
{ \
struct bfq_data *bfqd = e->elevator_data; \
u64 __data = __VAR; \
__data = div_u64(__data, NSEC_PER_USEC); \
return bfq_var_show(__data, (page)); \
}
USEC_SHOW_FUNCTION(bfq_slice_idle_us_show, bfqd->bfq_slice_idle);
#undef USEC_SHOW_FUNCTION
#define STORE_FUNCTION(__FUNC, __PTR, MIN, MAX, __CONV) \
static ssize_t \
__FUNC(struct elevator_queue *e, const char *page, size_t count) \
{ \
struct bfq_data *bfqd = e->elevator_data; \
unsigned long __data, __min = (MIN), __max = (MAX); \
int ret; \
\
ret = bfq_var_store(&__data, (page)); \
if (ret) \
return ret; \
if (__data < __min) \
__data = __min; \
else if (__data > __max) \
__data = __max; \
if (__CONV == 1) \
*(__PTR) = msecs_to_jiffies(__data); \
else if (__CONV == 2) \
*(__PTR) = (u64)__data * NSEC_PER_MSEC; \
else \
*(__PTR) = __data; \
return count; \
}
STORE_FUNCTION(bfq_fifo_expire_sync_store, &bfqd->bfq_fifo_expire[1], 1,
INT_MAX, 2);
STORE_FUNCTION(bfq_fifo_expire_async_store, &bfqd->bfq_fifo_expire[0], 1,
INT_MAX, 2);
STORE_FUNCTION(bfq_back_seek_max_store, &bfqd->bfq_back_max, 0, INT_MAX, 0);
STORE_FUNCTION(bfq_back_seek_penalty_store, &bfqd->bfq_back_penalty, 1,
INT_MAX, 0);
STORE_FUNCTION(bfq_slice_idle_store, &bfqd->bfq_slice_idle, 0, INT_MAX, 2);
#undef STORE_FUNCTION
#define USEC_STORE_FUNCTION(__FUNC, __PTR, MIN, MAX) \
static ssize_t __FUNC(struct elevator_queue *e, const char *page, size_t count)\
{ \
struct bfq_data *bfqd = e->elevator_data; \
unsigned long __data, __min = (MIN), __max = (MAX); \
int ret; \
\
ret = bfq_var_store(&__data, (page)); \
if (ret) \
return ret; \
if (__data < __min) \
__data = __min; \
else if (__data > __max) \
__data = __max; \
*(__PTR) = (u64)__data * NSEC_PER_USEC; \
return count; \
}
USEC_STORE_FUNCTION(bfq_slice_idle_us_store, &bfqd->bfq_slice_idle, 0,
UINT_MAX);
#undef USEC_STORE_FUNCTION
static ssize_t bfq_max_budget_store(struct elevator_queue *e,
const char *page, size_t count)
{
struct bfq_data *bfqd = e->elevator_data;
unsigned long __data;
int ret;
ret = bfq_var_store(&__data, (page));
if (ret)
return ret;
if (__data == 0)
bfqd->bfq_max_budget = bfq_calc_max_budget(bfqd);
else {
if (__data > INT_MAX)
__data = INT_MAX;
bfqd->bfq_max_budget = __data;
}
bfqd->bfq_user_max_budget = __data;
return count;
}
/*
* Leaving this name to preserve name compatibility with cfq
* parameters, but this timeout is used for both sync and async.
*/
static ssize_t bfq_timeout_sync_store(struct elevator_queue *e,
const char *page, size_t count)
{
struct bfq_data *bfqd = e->elevator_data;
unsigned long __data;
int ret;
ret = bfq_var_store(&__data, (page));
if (ret)
return ret;
if (__data < 1)
__data = 1;
else if (__data > INT_MAX)
__data = INT_MAX;
bfqd->bfq_timeout = msecs_to_jiffies(__data);
if (bfqd->bfq_user_max_budget == 0)
bfqd->bfq_max_budget = bfq_calc_max_budget(bfqd);
return count;
}
static ssize_t bfq_strict_guarantees_store(struct elevator_queue *e,
const char *page, size_t count)
{
struct bfq_data *bfqd = e->elevator_data;
unsigned long __data;
int ret;
ret = bfq_var_store(&__data, (page));
if (ret)
return ret;
if (__data > 1)
__data = 1;
if (!bfqd->strict_guarantees && __data == 1
&& bfqd->bfq_slice_idle < 8 * NSEC_PER_MSEC)
bfqd->bfq_slice_idle = 8 * NSEC_PER_MSEC;
bfqd->strict_guarantees = __data;
return count;
}
static ssize_t bfq_low_latency_store(struct elevator_queue *e,
const char *page, size_t count)
{
struct bfq_data *bfqd = e->elevator_data;
unsigned long __data;
int ret;
ret = bfq_var_store(&__data, (page));
if (ret)
return ret;
if (__data > 1)
__data = 1;
if (__data == 0 && bfqd->low_latency != 0)
bfq_end_wr(bfqd);
bfqd->low_latency = __data;
return count;
}
#define BFQ_ATTR(name) \
__ATTR(name, 0644, bfq_##name##_show, bfq_##name##_store)
static struct elv_fs_entry bfq_attrs[] = {
BFQ_ATTR(fifo_expire_sync),
BFQ_ATTR(fifo_expire_async),
BFQ_ATTR(back_seek_max),
BFQ_ATTR(back_seek_penalty),
BFQ_ATTR(slice_idle),
BFQ_ATTR(slice_idle_us),
BFQ_ATTR(max_budget),
BFQ_ATTR(timeout_sync),
BFQ_ATTR(strict_guarantees),
BFQ_ATTR(low_latency),
__ATTR_NULL
};
static struct elevator_type iosched_bfq_mq = {
.ops = {
.limit_depth = bfq_limit_depth,
.prepare_request = bfq_prepare_request,
.requeue_request = bfq_finish_requeue_request,
.finish_request = bfq_finish_requeue_request,
.exit_icq = bfq_exit_icq,
.insert_requests = bfq_insert_requests,
.dispatch_request = bfq_dispatch_request,
.next_request = elv_rb_latter_request,
.former_request = elv_rb_former_request,
.allow_merge = bfq_allow_bio_merge,
.bio_merge = bfq_bio_merge,
.request_merge = bfq_request_merge,
.requests_merged = bfq_requests_merged,
.request_merged = bfq_request_merged,
.has_work = bfq_has_work,
.depth_updated = bfq_depth_updated,
.init_hctx = bfq_init_hctx,
.init_sched = bfq_init_queue,
.exit_sched = bfq_exit_queue,
},
.icq_size = sizeof(struct bfq_io_cq),
.icq_align = __alignof__(struct bfq_io_cq),
.elevator_attrs = bfq_attrs,
.elevator_name = "bfq",
.elevator_owner = THIS_MODULE,
};
MODULE_ALIAS("bfq-iosched");
static int __init bfq_init(void)
{
int ret;
#ifdef CONFIG_BFQ_GROUP_IOSCHED
ret = blkcg_policy_register(&blkcg_policy_bfq);
if (ret)
return ret;
#endif
ret = -ENOMEM;
if (bfq_slab_setup())
goto err_pol_unreg;
/*
* Times to load large popular applications for the typical
* systems installed on the reference devices (see the
* comments before the definition of the next
* array). Actually, we use slightly lower values, as the
* estimated peak rate tends to be smaller than the actual
* peak rate. The reason for this last fact is that estimates
* are computed over much shorter time intervals than the long
* intervals typically used for benchmarking. Why? First, to
* adapt more quickly to variations. Second, because an I/O
* scheduler cannot rely on a peak-rate-evaluation workload to
* be run for a long time.
*/
ref_wr_duration[0] = msecs_to_jiffies(7000); /* actually 8 sec */
ref_wr_duration[1] = msecs_to_jiffies(2500); /* actually 3 sec */
ret = elv_register(&iosched_bfq_mq);
if (ret)
goto slab_kill;
return 0;
slab_kill:
bfq_slab_kill();
err_pol_unreg:
#ifdef CONFIG_BFQ_GROUP_IOSCHED
blkcg_policy_unregister(&blkcg_policy_bfq);
#endif
return ret;
}
static void __exit bfq_exit(void)
{
elv_unregister(&iosched_bfq_mq);
#ifdef CONFIG_BFQ_GROUP_IOSCHED
blkcg_policy_unregister(&blkcg_policy_bfq);
#endif
bfq_slab_kill();
}
module_init(bfq_init);
module_exit(bfq_exit);
MODULE_AUTHOR("Paolo Valente");
MODULE_LICENSE("GPL");
MODULE_DESCRIPTION("MQ Budget Fair Queueing I/O Scheduler");
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