mirror of
https://github.com/torvalds/linux.git
synced 2024-12-13 14:43:03 +00:00
05a5f51ca5
Replace the lkml.org links with lore to better use a single source that's more likely to stay available long-term. Done by bash script: cvt_lkml_to_lore () { tmpfile=$(mktemp ./.cvt_links.XXXXXXX) header=$(echo $1 | sed 's@/lkml/@/lkml/headers/@') wget -qO - $header > $tmpfile if [[ $? == 0 ]] ; then link=$(grep -i '^Message-Id:' $tmpfile | head -1 | \ sed -r -e 's/^\s*Message-Id:\s*<\s*//' -e 's/\s*>\s*$//' -e 's@^@https://lore.kernel.org/r/@') # echo "testlink: $link" if [ -n "$link" ] ; then wget -qO - $link > /dev/null if [[ $? == 0 ]] ; then echo $link fi fi fi rm -f $tmpfile } git grep -P -o "\bhttps?://(?:www.)?lkml.org/lkml[\/\w]+" $@ | while read line ; do echo $line file=$(echo $line | cut -f1 -d':') link=$(echo $line | cut -f2- -d':') newlink=$(cvt_lkml_to_lore $link) if [[ -n "$newlink" ]] ; then sed -i -e "s#\b$link\b#$newlink#" $file fi done Link: https://lore.kernel.org/patchwork/patch/1265849/#1462688 Signed-off-by: Joe Perches <joe@perches.com> Link: https://lore.kernel.org/r/77cdb7f32cfb087955bfc3600b86c40bed5d4104.camel@perches.com Signed-off-by: Jonathan Corbet <corbet@lwn.net>
889 lines
37 KiB
ReStructuredText
889 lines
37 KiB
ReStructuredText
========================
|
||
Deadline Task Scheduling
|
||
========================
|
||
|
||
.. CONTENTS
|
||
|
||
0. WARNING
|
||
1. Overview
|
||
2. Scheduling algorithm
|
||
2.1 Main algorithm
|
||
2.2 Bandwidth reclaiming
|
||
3. Scheduling Real-Time Tasks
|
||
3.1 Definitions
|
||
3.2 Schedulability Analysis for Uniprocessor Systems
|
||
3.3 Schedulability Analysis for Multiprocessor Systems
|
||
3.4 Relationship with SCHED_DEADLINE Parameters
|
||
4. Bandwidth management
|
||
4.1 System-wide settings
|
||
4.2 Task interface
|
||
4.3 Default behavior
|
||
4.4 Behavior of sched_yield()
|
||
5. Tasks CPU affinity
|
||
5.1 SCHED_DEADLINE and cpusets HOWTO
|
||
6. Future plans
|
||
A. Test suite
|
||
B. Minimal main()
|
||
|
||
|
||
0. WARNING
|
||
==========
|
||
|
||
Fiddling with these settings can result in an unpredictable or even unstable
|
||
system behavior. As for -rt (group) scheduling, it is assumed that root users
|
||
know what they're doing.
|
||
|
||
|
||
1. Overview
|
||
===========
|
||
|
||
The SCHED_DEADLINE policy contained inside the sched_dl scheduling class is
|
||
basically an implementation of the Earliest Deadline First (EDF) scheduling
|
||
algorithm, augmented with a mechanism (called Constant Bandwidth Server, CBS)
|
||
that makes it possible to isolate the behavior of tasks between each other.
|
||
|
||
|
||
2. Scheduling algorithm
|
||
=======================
|
||
|
||
2.1 Main algorithm
|
||
------------------
|
||
|
||
SCHED_DEADLINE [18] uses three parameters, named "runtime", "period", and
|
||
"deadline", to schedule tasks. A SCHED_DEADLINE task should receive
|
||
"runtime" microseconds of execution time every "period" microseconds, and
|
||
these "runtime" microseconds are available within "deadline" microseconds
|
||
from the beginning of the period. In order to implement this behavior,
|
||
every time the task wakes up, the scheduler computes a "scheduling deadline"
|
||
consistent with the guarantee (using the CBS[2,3] algorithm). Tasks are then
|
||
scheduled using EDF[1] on these scheduling deadlines (the task with the
|
||
earliest scheduling deadline is selected for execution). Notice that the
|
||
task actually receives "runtime" time units within "deadline" if a proper
|
||
"admission control" strategy (see Section "4. Bandwidth management") is used
|
||
(clearly, if the system is overloaded this guarantee cannot be respected).
|
||
|
||
Summing up, the CBS[2,3] algorithm assigns scheduling deadlines to tasks so
|
||
that each task runs for at most its runtime every period, avoiding any
|
||
interference between different tasks (bandwidth isolation), while the EDF[1]
|
||
algorithm selects the task with the earliest scheduling deadline as the one
|
||
to be executed next. Thanks to this feature, tasks that do not strictly comply
|
||
with the "traditional" real-time task model (see Section 3) can effectively
|
||
use the new policy.
|
||
|
||
In more details, the CBS algorithm assigns scheduling deadlines to
|
||
tasks in the following way:
|
||
|
||
- Each SCHED_DEADLINE task is characterized by the "runtime",
|
||
"deadline", and "period" parameters;
|
||
|
||
- The state of the task is described by a "scheduling deadline", and
|
||
a "remaining runtime". These two parameters are initially set to 0;
|
||
|
||
- When a SCHED_DEADLINE task wakes up (becomes ready for execution),
|
||
the scheduler checks if::
|
||
|
||
remaining runtime runtime
|
||
---------------------------------- > ---------
|
||
scheduling deadline - current time period
|
||
|
||
then, if the scheduling deadline is smaller than the current time, or
|
||
this condition is verified, the scheduling deadline and the
|
||
remaining runtime are re-initialized as
|
||
|
||
scheduling deadline = current time + deadline
|
||
remaining runtime = runtime
|
||
|
||
otherwise, the scheduling deadline and the remaining runtime are
|
||
left unchanged;
|
||
|
||
- When a SCHED_DEADLINE task executes for an amount of time t, its
|
||
remaining runtime is decreased as::
|
||
|
||
remaining runtime = remaining runtime - t
|
||
|
||
(technically, the runtime is decreased at every tick, or when the
|
||
task is descheduled / preempted);
|
||
|
||
- When the remaining runtime becomes less or equal than 0, the task is
|
||
said to be "throttled" (also known as "depleted" in real-time literature)
|
||
and cannot be scheduled until its scheduling deadline. The "replenishment
|
||
time" for this task (see next item) is set to be equal to the current
|
||
value of the scheduling deadline;
|
||
|
||
- When the current time is equal to the replenishment time of a
|
||
throttled task, the scheduling deadline and the remaining runtime are
|
||
updated as::
|
||
|
||
scheduling deadline = scheduling deadline + period
|
||
remaining runtime = remaining runtime + runtime
|
||
|
||
The SCHED_FLAG_DL_OVERRUN flag in sched_attr's sched_flags field allows a task
|
||
to get informed about runtime overruns through the delivery of SIGXCPU
|
||
signals.
|
||
|
||
|
||
2.2 Bandwidth reclaiming
|
||
------------------------
|
||
|
||
Bandwidth reclaiming for deadline tasks is based on the GRUB (Greedy
|
||
Reclamation of Unused Bandwidth) algorithm [15, 16, 17] and it is enabled
|
||
when flag SCHED_FLAG_RECLAIM is set.
|
||
|
||
The following diagram illustrates the state names for tasks handled by GRUB::
|
||
|
||
------------
|
||
(d) | Active |
|
||
------------->| |
|
||
| | Contending |
|
||
| ------------
|
||
| A |
|
||
---------- | |
|
||
| | | |
|
||
| Inactive | |(b) | (a)
|
||
| | | |
|
||
---------- | |
|
||
A | V
|
||
| ------------
|
||
| | Active |
|
||
--------------| Non |
|
||
(c) | Contending |
|
||
------------
|
||
|
||
A task can be in one of the following states:
|
||
|
||
- ActiveContending: if it is ready for execution (or executing);
|
||
|
||
- ActiveNonContending: if it just blocked and has not yet surpassed the 0-lag
|
||
time;
|
||
|
||
- Inactive: if it is blocked and has surpassed the 0-lag time.
|
||
|
||
State transitions:
|
||
|
||
(a) When a task blocks, it does not become immediately inactive since its
|
||
bandwidth cannot be immediately reclaimed without breaking the
|
||
real-time guarantees. It therefore enters a transitional state called
|
||
ActiveNonContending. The scheduler arms the "inactive timer" to fire at
|
||
the 0-lag time, when the task's bandwidth can be reclaimed without
|
||
breaking the real-time guarantees.
|
||
|
||
The 0-lag time for a task entering the ActiveNonContending state is
|
||
computed as::
|
||
|
||
(runtime * dl_period)
|
||
deadline - ---------------------
|
||
dl_runtime
|
||
|
||
where runtime is the remaining runtime, while dl_runtime and dl_period
|
||
are the reservation parameters.
|
||
|
||
(b) If the task wakes up before the inactive timer fires, the task re-enters
|
||
the ActiveContending state and the "inactive timer" is canceled.
|
||
In addition, if the task wakes up on a different runqueue, then
|
||
the task's utilization must be removed from the previous runqueue's active
|
||
utilization and must be added to the new runqueue's active utilization.
|
||
In order to avoid races between a task waking up on a runqueue while the
|
||
"inactive timer" is running on a different CPU, the "dl_non_contending"
|
||
flag is used to indicate that a task is not on a runqueue but is active
|
||
(so, the flag is set when the task blocks and is cleared when the
|
||
"inactive timer" fires or when the task wakes up).
|
||
|
||
(c) When the "inactive timer" fires, the task enters the Inactive state and
|
||
its utilization is removed from the runqueue's active utilization.
|
||
|
||
(d) When an inactive task wakes up, it enters the ActiveContending state and
|
||
its utilization is added to the active utilization of the runqueue where
|
||
it has been enqueued.
|
||
|
||
For each runqueue, the algorithm GRUB keeps track of two different bandwidths:
|
||
|
||
- Active bandwidth (running_bw): this is the sum of the bandwidths of all
|
||
tasks in active state (i.e., ActiveContending or ActiveNonContending);
|
||
|
||
- Total bandwidth (this_bw): this is the sum of all tasks "belonging" to the
|
||
runqueue, including the tasks in Inactive state.
|
||
|
||
|
||
The algorithm reclaims the bandwidth of the tasks in Inactive state.
|
||
It does so by decrementing the runtime of the executing task Ti at a pace equal
|
||
to
|
||
|
||
dq = -max{ Ui / Umax, (1 - Uinact - Uextra) } dt
|
||
|
||
where:
|
||
|
||
- Ui is the bandwidth of task Ti;
|
||
- Umax is the maximum reclaimable utilization (subjected to RT throttling
|
||
limits);
|
||
- Uinact is the (per runqueue) inactive utilization, computed as
|
||
(this_bq - running_bw);
|
||
- Uextra is the (per runqueue) extra reclaimable utilization
|
||
(subjected to RT throttling limits).
|
||
|
||
|
||
Let's now see a trivial example of two deadline tasks with runtime equal
|
||
to 4 and period equal to 8 (i.e., bandwidth equal to 0.5)::
|
||
|
||
A Task T1
|
||
|
|
||
| |
|
||
| |
|
||
|-------- |----
|
||
| | V
|
||
|---|---|---|---|---|---|---|---|--------->t
|
||
0 1 2 3 4 5 6 7 8
|
||
|
||
|
||
A Task T2
|
||
|
|
||
| |
|
||
| |
|
||
| ------------------------|
|
||
| | V
|
||
|---|---|---|---|---|---|---|---|--------->t
|
||
0 1 2 3 4 5 6 7 8
|
||
|
||
|
||
A running_bw
|
||
|
|
||
1 ----------------- ------
|
||
| | |
|
||
0.5- -----------------
|
||
| |
|
||
|---|---|---|---|---|---|---|---|--------->t
|
||
0 1 2 3 4 5 6 7 8
|
||
|
||
|
||
- Time t = 0:
|
||
|
||
Both tasks are ready for execution and therefore in ActiveContending state.
|
||
Suppose Task T1 is the first task to start execution.
|
||
Since there are no inactive tasks, its runtime is decreased as dq = -1 dt.
|
||
|
||
- Time t = 2:
|
||
|
||
Suppose that task T1 blocks
|
||
Task T1 therefore enters the ActiveNonContending state. Since its remaining
|
||
runtime is equal to 2, its 0-lag time is equal to t = 4.
|
||
Task T2 start execution, with runtime still decreased as dq = -1 dt since
|
||
there are no inactive tasks.
|
||
|
||
- Time t = 4:
|
||
|
||
This is the 0-lag time for Task T1. Since it didn't woken up in the
|
||
meantime, it enters the Inactive state. Its bandwidth is removed from
|
||
running_bw.
|
||
Task T2 continues its execution. However, its runtime is now decreased as
|
||
dq = - 0.5 dt because Uinact = 0.5.
|
||
Task T2 therefore reclaims the bandwidth unused by Task T1.
|
||
|
||
- Time t = 8:
|
||
|
||
Task T1 wakes up. It enters the ActiveContending state again, and the
|
||
running_bw is incremented.
|
||
|
||
|
||
2.3 Energy-aware scheduling
|
||
---------------------------
|
||
|
||
When cpufreq's schedutil governor is selected, SCHED_DEADLINE implements the
|
||
GRUB-PA [19] algorithm, reducing the CPU operating frequency to the minimum
|
||
value that still allows to meet the deadlines. This behavior is currently
|
||
implemented only for ARM architectures.
|
||
|
||
A particular care must be taken in case the time needed for changing frequency
|
||
is of the same order of magnitude of the reservation period. In such cases,
|
||
setting a fixed CPU frequency results in a lower amount of deadline misses.
|
||
|
||
|
||
3. Scheduling Real-Time Tasks
|
||
=============================
|
||
|
||
|
||
|
||
.. BIG FAT WARNING ******************************************************
|
||
|
||
.. warning::
|
||
|
||
This section contains a (not-thorough) summary on classical deadline
|
||
scheduling theory, and how it applies to SCHED_DEADLINE.
|
||
The reader can "safely" skip to Section 4 if only interested in seeing
|
||
how the scheduling policy can be used. Anyway, we strongly recommend
|
||
to come back here and continue reading (once the urge for testing is
|
||
satisfied :P) to be sure of fully understanding all technical details.
|
||
|
||
.. ************************************************************************
|
||
|
||
There are no limitations on what kind of task can exploit this new
|
||
scheduling discipline, even if it must be said that it is particularly
|
||
suited for periodic or sporadic real-time tasks that need guarantees on their
|
||
timing behavior, e.g., multimedia, streaming, control applications, etc.
|
||
|
||
3.1 Definitions
|
||
------------------------
|
||
|
||
A typical real-time task is composed of a repetition of computation phases
|
||
(task instances, or jobs) which are activated on a periodic or sporadic
|
||
fashion.
|
||
Each job J_j (where J_j is the j^th job of the task) is characterized by an
|
||
arrival time r_j (the time when the job starts), an amount of computation
|
||
time c_j needed to finish the job, and a job absolute deadline d_j, which
|
||
is the time within which the job should be finished. The maximum execution
|
||
time max{c_j} is called "Worst Case Execution Time" (WCET) for the task.
|
||
A real-time task can be periodic with period P if r_{j+1} = r_j + P, or
|
||
sporadic with minimum inter-arrival time P is r_{j+1} >= r_j + P. Finally,
|
||
d_j = r_j + D, where D is the task's relative deadline.
|
||
Summing up, a real-time task can be described as
|
||
|
||
Task = (WCET, D, P)
|
||
|
||
The utilization of a real-time task is defined as the ratio between its
|
||
WCET and its period (or minimum inter-arrival time), and represents
|
||
the fraction of CPU time needed to execute the task.
|
||
|
||
If the total utilization U=sum(WCET_i/P_i) is larger than M (with M equal
|
||
to the number of CPUs), then the scheduler is unable to respect all the
|
||
deadlines.
|
||
Note that total utilization is defined as the sum of the utilizations
|
||
WCET_i/P_i over all the real-time tasks in the system. When considering
|
||
multiple real-time tasks, the parameters of the i-th task are indicated
|
||
with the "_i" suffix.
|
||
Moreover, if the total utilization is larger than M, then we risk starving
|
||
non- real-time tasks by real-time tasks.
|
||
If, instead, the total utilization is smaller than M, then non real-time
|
||
tasks will not be starved and the system might be able to respect all the
|
||
deadlines.
|
||
As a matter of fact, in this case it is possible to provide an upper bound
|
||
for tardiness (defined as the maximum between 0 and the difference
|
||
between the finishing time of a job and its absolute deadline).
|
||
More precisely, it can be proven that using a global EDF scheduler the
|
||
maximum tardiness of each task is smaller or equal than
|
||
|
||
((M − 1) · WCET_max − WCET_min)/(M − (M − 2) · U_max) + WCET_max
|
||
|
||
where WCET_max = max{WCET_i} is the maximum WCET, WCET_min=min{WCET_i}
|
||
is the minimum WCET, and U_max = max{WCET_i/P_i} is the maximum
|
||
utilization[12].
|
||
|
||
3.2 Schedulability Analysis for Uniprocessor Systems
|
||
----------------------------------------------------
|
||
|
||
If M=1 (uniprocessor system), or in case of partitioned scheduling (each
|
||
real-time task is statically assigned to one and only one CPU), it is
|
||
possible to formally check if all the deadlines are respected.
|
||
If D_i = P_i for all tasks, then EDF is able to respect all the deadlines
|
||
of all the tasks executing on a CPU if and only if the total utilization
|
||
of the tasks running on such a CPU is smaller or equal than 1.
|
||
If D_i != P_i for some task, then it is possible to define the density of
|
||
a task as WCET_i/min{D_i,P_i}, and EDF is able to respect all the deadlines
|
||
of all the tasks running on a CPU if the sum of the densities of the tasks
|
||
running on such a CPU is smaller or equal than 1:
|
||
|
||
sum(WCET_i / min{D_i, P_i}) <= 1
|
||
|
||
It is important to notice that this condition is only sufficient, and not
|
||
necessary: there are task sets that are schedulable, but do not respect the
|
||
condition. For example, consider the task set {Task_1,Task_2} composed by
|
||
Task_1=(50ms,50ms,100ms) and Task_2=(10ms,100ms,100ms).
|
||
EDF is clearly able to schedule the two tasks without missing any deadline
|
||
(Task_1 is scheduled as soon as it is released, and finishes just in time
|
||
to respect its deadline; Task_2 is scheduled immediately after Task_1, hence
|
||
its response time cannot be larger than 50ms + 10ms = 60ms) even if
|
||
|
||
50 / min{50,100} + 10 / min{100, 100} = 50 / 50 + 10 / 100 = 1.1
|
||
|
||
Of course it is possible to test the exact schedulability of tasks with
|
||
D_i != P_i (checking a condition that is both sufficient and necessary),
|
||
but this cannot be done by comparing the total utilization or density with
|
||
a constant. Instead, the so called "processor demand" approach can be used,
|
||
computing the total amount of CPU time h(t) needed by all the tasks to
|
||
respect all of their deadlines in a time interval of size t, and comparing
|
||
such a time with the interval size t. If h(t) is smaller than t (that is,
|
||
the amount of time needed by the tasks in a time interval of size t is
|
||
smaller than the size of the interval) for all the possible values of t, then
|
||
EDF is able to schedule the tasks respecting all of their deadlines. Since
|
||
performing this check for all possible values of t is impossible, it has been
|
||
proven[4,5,6] that it is sufficient to perform the test for values of t
|
||
between 0 and a maximum value L. The cited papers contain all of the
|
||
mathematical details and explain how to compute h(t) and L.
|
||
In any case, this kind of analysis is too complex as well as too
|
||
time-consuming to be performed on-line. Hence, as explained in Section
|
||
4 Linux uses an admission test based on the tasks' utilizations.
|
||
|
||
3.3 Schedulability Analysis for Multiprocessor Systems
|
||
------------------------------------------------------
|
||
|
||
On multiprocessor systems with global EDF scheduling (non partitioned
|
||
systems), a sufficient test for schedulability can not be based on the
|
||
utilizations or densities: it can be shown that even if D_i = P_i task
|
||
sets with utilizations slightly larger than 1 can miss deadlines regardless
|
||
of the number of CPUs.
|
||
|
||
Consider a set {Task_1,...Task_{M+1}} of M+1 tasks on a system with M
|
||
CPUs, with the first task Task_1=(P,P,P) having period, relative deadline
|
||
and WCET equal to P. The remaining M tasks Task_i=(e,P-1,P-1) have an
|
||
arbitrarily small worst case execution time (indicated as "e" here) and a
|
||
period smaller than the one of the first task. Hence, if all the tasks
|
||
activate at the same time t, global EDF schedules these M tasks first
|
||
(because their absolute deadlines are equal to t + P - 1, hence they are
|
||
smaller than the absolute deadline of Task_1, which is t + P). As a
|
||
result, Task_1 can be scheduled only at time t + e, and will finish at
|
||
time t + e + P, after its absolute deadline. The total utilization of the
|
||
task set is U = M · e / (P - 1) + P / P = M · e / (P - 1) + 1, and for small
|
||
values of e this can become very close to 1. This is known as "Dhall's
|
||
effect"[7]. Note: the example in the original paper by Dhall has been
|
||
slightly simplified here (for example, Dhall more correctly computed
|
||
lim_{e->0}U).
|
||
|
||
More complex schedulability tests for global EDF have been developed in
|
||
real-time literature[8,9], but they are not based on a simple comparison
|
||
between total utilization (or density) and a fixed constant. If all tasks
|
||
have D_i = P_i, a sufficient schedulability condition can be expressed in
|
||
a simple way:
|
||
|
||
sum(WCET_i / P_i) <= M - (M - 1) · U_max
|
||
|
||
where U_max = max{WCET_i / P_i}[10]. Notice that for U_max = 1,
|
||
M - (M - 1) · U_max becomes M - M + 1 = 1 and this schedulability condition
|
||
just confirms the Dhall's effect. A more complete survey of the literature
|
||
about schedulability tests for multi-processor real-time scheduling can be
|
||
found in [11].
|
||
|
||
As seen, enforcing that the total utilization is smaller than M does not
|
||
guarantee that global EDF schedules the tasks without missing any deadline
|
||
(in other words, global EDF is not an optimal scheduling algorithm). However,
|
||
a total utilization smaller than M is enough to guarantee that non real-time
|
||
tasks are not starved and that the tardiness of real-time tasks has an upper
|
||
bound[12] (as previously noted). Different bounds on the maximum tardiness
|
||
experienced by real-time tasks have been developed in various papers[13,14],
|
||
but the theoretical result that is important for SCHED_DEADLINE is that if
|
||
the total utilization is smaller or equal than M then the response times of
|
||
the tasks are limited.
|
||
|
||
3.4 Relationship with SCHED_DEADLINE Parameters
|
||
-----------------------------------------------
|
||
|
||
Finally, it is important to understand the relationship between the
|
||
SCHED_DEADLINE scheduling parameters described in Section 2 (runtime,
|
||
deadline and period) and the real-time task parameters (WCET, D, P)
|
||
described in this section. Note that the tasks' temporal constraints are
|
||
represented by its absolute deadlines d_j = r_j + D described above, while
|
||
SCHED_DEADLINE schedules the tasks according to scheduling deadlines (see
|
||
Section 2).
|
||
If an admission test is used to guarantee that the scheduling deadlines
|
||
are respected, then SCHED_DEADLINE can be used to schedule real-time tasks
|
||
guaranteeing that all the jobs' deadlines of a task are respected.
|
||
In order to do this, a task must be scheduled by setting:
|
||
|
||
- runtime >= WCET
|
||
- deadline = D
|
||
- period <= P
|
||
|
||
IOW, if runtime >= WCET and if period is <= P, then the scheduling deadlines
|
||
and the absolute deadlines (d_j) coincide, so a proper admission control
|
||
allows to respect the jobs' absolute deadlines for this task (this is what is
|
||
called "hard schedulability property" and is an extension of Lemma 1 of [2]).
|
||
Notice that if runtime > deadline the admission control will surely reject
|
||
this task, as it is not possible to respect its temporal constraints.
|
||
|
||
References:
|
||
|
||
1 - C. L. Liu and J. W. Layland. Scheduling algorithms for multiprogram-
|
||
ming in a hard-real-time environment. Journal of the Association for
|
||
Computing Machinery, 20(1), 1973.
|
||
2 - L. Abeni , G. Buttazzo. Integrating Multimedia Applications in Hard
|
||
Real-Time Systems. Proceedings of the 19th IEEE Real-time Systems
|
||
Symposium, 1998. http://retis.sssup.it/~giorgio/paps/1998/rtss98-cbs.pdf
|
||
3 - L. Abeni. Server Mechanisms for Multimedia Applications. ReTiS Lab
|
||
Technical Report. http://disi.unitn.it/~abeni/tr-98-01.pdf
|
||
4 - J. Y. Leung and M.L. Merril. A Note on Preemptive Scheduling of
|
||
Periodic, Real-Time Tasks. Information Processing Letters, vol. 11,
|
||
no. 3, pp. 115-118, 1980.
|
||
5 - S. K. Baruah, A. K. Mok and L. E. Rosier. Preemptively Scheduling
|
||
Hard-Real-Time Sporadic Tasks on One Processor. Proceedings of the
|
||
11th IEEE Real-time Systems Symposium, 1990.
|
||
6 - S. K. Baruah, L. E. Rosier and R. R. Howell. Algorithms and Complexity
|
||
Concerning the Preemptive Scheduling of Periodic Real-Time tasks on
|
||
One Processor. Real-Time Systems Journal, vol. 4, no. 2, pp 301-324,
|
||
1990.
|
||
7 - S. J. Dhall and C. L. Liu. On a real-time scheduling problem. Operations
|
||
research, vol. 26, no. 1, pp 127-140, 1978.
|
||
8 - T. Baker. Multiprocessor EDF and Deadline Monotonic Schedulability
|
||
Analysis. Proceedings of the 24th IEEE Real-Time Systems Symposium, 2003.
|
||
9 - T. Baker. An Analysis of EDF Schedulability on a Multiprocessor.
|
||
IEEE Transactions on Parallel and Distributed Systems, vol. 16, no. 8,
|
||
pp 760-768, 2005.
|
||
10 - J. Goossens, S. Funk and S. Baruah, Priority-Driven Scheduling of
|
||
Periodic Task Systems on Multiprocessors. Real-Time Systems Journal,
|
||
vol. 25, no. 2–3, pp. 187–205, 2003.
|
||
11 - R. Davis and A. Burns. A Survey of Hard Real-Time Scheduling for
|
||
Multiprocessor Systems. ACM Computing Surveys, vol. 43, no. 4, 2011.
|
||
http://www-users.cs.york.ac.uk/~robdavis/papers/MPSurveyv5.0.pdf
|
||
12 - U. C. Devi and J. H. Anderson. Tardiness Bounds under Global EDF
|
||
Scheduling on a Multiprocessor. Real-Time Systems Journal, vol. 32,
|
||
no. 2, pp 133-189, 2008.
|
||
13 - P. Valente and G. Lipari. An Upper Bound to the Lateness of Soft
|
||
Real-Time Tasks Scheduled by EDF on Multiprocessors. Proceedings of
|
||
the 26th IEEE Real-Time Systems Symposium, 2005.
|
||
14 - J. Erickson, U. Devi and S. Baruah. Improved tardiness bounds for
|
||
Global EDF. Proceedings of the 22nd Euromicro Conference on
|
||
Real-Time Systems, 2010.
|
||
15 - G. Lipari, S. Baruah, Greedy reclamation of unused bandwidth in
|
||
constant-bandwidth servers, 12th IEEE Euromicro Conference on Real-Time
|
||
Systems, 2000.
|
||
16 - L. Abeni, J. Lelli, C. Scordino, L. Palopoli, Greedy CPU reclaiming for
|
||
SCHED DEADLINE. In Proceedings of the Real-Time Linux Workshop (RTLWS),
|
||
Dusseldorf, Germany, 2014.
|
||
17 - L. Abeni, G. Lipari, A. Parri, Y. Sun, Multicore CPU reclaiming: parallel
|
||
or sequential?. In Proceedings of the 31st Annual ACM Symposium on Applied
|
||
Computing, 2016.
|
||
18 - J. Lelli, C. Scordino, L. Abeni, D. Faggioli, Deadline scheduling in the
|
||
Linux kernel, Software: Practice and Experience, 46(6): 821-839, June
|
||
2016.
|
||
19 - C. Scordino, L. Abeni, J. Lelli, Energy-Aware Real-Time Scheduling in
|
||
the Linux Kernel, 33rd ACM/SIGAPP Symposium On Applied Computing (SAC
|
||
2018), Pau, France, April 2018.
|
||
|
||
|
||
4. Bandwidth management
|
||
=======================
|
||
|
||
As previously mentioned, in order for -deadline scheduling to be
|
||
effective and useful (that is, to be able to provide "runtime" time units
|
||
within "deadline"), it is important to have some method to keep the allocation
|
||
of the available fractions of CPU time to the various tasks under control.
|
||
This is usually called "admission control" and if it is not performed, then
|
||
no guarantee can be given on the actual scheduling of the -deadline tasks.
|
||
|
||
As already stated in Section 3, a necessary condition to be respected to
|
||
correctly schedule a set of real-time tasks is that the total utilization
|
||
is smaller than M. When talking about -deadline tasks, this requires that
|
||
the sum of the ratio between runtime and period for all tasks is smaller
|
||
than M. Notice that the ratio runtime/period is equivalent to the utilization
|
||
of a "traditional" real-time task, and is also often referred to as
|
||
"bandwidth".
|
||
The interface used to control the CPU bandwidth that can be allocated
|
||
to -deadline tasks is similar to the one already used for -rt
|
||
tasks with real-time group scheduling (a.k.a. RT-throttling - see
|
||
Documentation/scheduler/sched-rt-group.rst), and is based on readable/
|
||
writable control files located in procfs (for system wide settings).
|
||
Notice that per-group settings (controlled through cgroupfs) are still not
|
||
defined for -deadline tasks, because more discussion is needed in order to
|
||
figure out how we want to manage SCHED_DEADLINE bandwidth at the task group
|
||
level.
|
||
|
||
A main difference between deadline bandwidth management and RT-throttling
|
||
is that -deadline tasks have bandwidth on their own (while -rt ones don't!),
|
||
and thus we don't need a higher level throttling mechanism to enforce the
|
||
desired bandwidth. In other words, this means that interface parameters are
|
||
only used at admission control time (i.e., when the user calls
|
||
sched_setattr()). Scheduling is then performed considering actual tasks'
|
||
parameters, so that CPU bandwidth is allocated to SCHED_DEADLINE tasks
|
||
respecting their needs in terms of granularity. Therefore, using this simple
|
||
interface we can put a cap on total utilization of -deadline tasks (i.e.,
|
||
\Sum (runtime_i / period_i) < global_dl_utilization_cap).
|
||
|
||
4.1 System wide settings
|
||
------------------------
|
||
|
||
The system wide settings are configured under the /proc virtual file system.
|
||
|
||
For now the -rt knobs are used for -deadline admission control and the
|
||
-deadline runtime is accounted against the -rt runtime. We realize that this
|
||
isn't entirely desirable; however, it is better to have a small interface for
|
||
now, and be able to change it easily later. The ideal situation (see 5.) is to
|
||
run -rt tasks from a -deadline server; in which case the -rt bandwidth is a
|
||
direct subset of dl_bw.
|
||
|
||
This means that, for a root_domain comprising M CPUs, -deadline tasks
|
||
can be created while the sum of their bandwidths stays below:
|
||
|
||
M * (sched_rt_runtime_us / sched_rt_period_us)
|
||
|
||
It is also possible to disable this bandwidth management logic, and
|
||
be thus free of oversubscribing the system up to any arbitrary level.
|
||
This is done by writing -1 in /proc/sys/kernel/sched_rt_runtime_us.
|
||
|
||
|
||
4.2 Task interface
|
||
------------------
|
||
|
||
Specifying a periodic/sporadic task that executes for a given amount of
|
||
runtime at each instance, and that is scheduled according to the urgency of
|
||
its own timing constraints needs, in general, a way of declaring:
|
||
|
||
- a (maximum/typical) instance execution time,
|
||
- a minimum interval between consecutive instances,
|
||
- a time constraint by which each instance must be completed.
|
||
|
||
Therefore:
|
||
|
||
* a new struct sched_attr, containing all the necessary fields is
|
||
provided;
|
||
* the new scheduling related syscalls that manipulate it, i.e.,
|
||
sched_setattr() and sched_getattr() are implemented.
|
||
|
||
For debugging purposes, the leftover runtime and absolute deadline of a
|
||
SCHED_DEADLINE task can be retrieved through /proc/<pid>/sched (entries
|
||
dl.runtime and dl.deadline, both values in ns). A programmatic way to
|
||
retrieve these values from production code is under discussion.
|
||
|
||
|
||
4.3 Default behavior
|
||
---------------------
|
||
|
||
The default value for SCHED_DEADLINE bandwidth is to have rt_runtime equal to
|
||
950000. With rt_period equal to 1000000, by default, it means that -deadline
|
||
tasks can use at most 95%, multiplied by the number of CPUs that compose the
|
||
root_domain, for each root_domain.
|
||
This means that non -deadline tasks will receive at least 5% of the CPU time,
|
||
and that -deadline tasks will receive their runtime with a guaranteed
|
||
worst-case delay respect to the "deadline" parameter. If "deadline" = "period"
|
||
and the cpuset mechanism is used to implement partitioned scheduling (see
|
||
Section 5), then this simple setting of the bandwidth management is able to
|
||
deterministically guarantee that -deadline tasks will receive their runtime
|
||
in a period.
|
||
|
||
Finally, notice that in order not to jeopardize the admission control a
|
||
-deadline task cannot fork.
|
||
|
||
|
||
4.4 Behavior of sched_yield()
|
||
-----------------------------
|
||
|
||
When a SCHED_DEADLINE task calls sched_yield(), it gives up its
|
||
remaining runtime and is immediately throttled, until the next
|
||
period, when its runtime will be replenished (a special flag
|
||
dl_yielded is set and used to handle correctly throttling and runtime
|
||
replenishment after a call to sched_yield()).
|
||
|
||
This behavior of sched_yield() allows the task to wake-up exactly at
|
||
the beginning of the next period. Also, this may be useful in the
|
||
future with bandwidth reclaiming mechanisms, where sched_yield() will
|
||
make the leftoever runtime available for reclamation by other
|
||
SCHED_DEADLINE tasks.
|
||
|
||
|
||
5. Tasks CPU affinity
|
||
=====================
|
||
|
||
-deadline tasks cannot have an affinity mask smaller that the entire
|
||
root_domain they are created on. However, affinities can be specified
|
||
through the cpuset facility (Documentation/admin-guide/cgroup-v1/cpusets.rst).
|
||
|
||
5.1 SCHED_DEADLINE and cpusets HOWTO
|
||
------------------------------------
|
||
|
||
An example of a simple configuration (pin a -deadline task to CPU0)
|
||
follows (rt-app is used to create a -deadline task)::
|
||
|
||
mkdir /dev/cpuset
|
||
mount -t cgroup -o cpuset cpuset /dev/cpuset
|
||
cd /dev/cpuset
|
||
mkdir cpu0
|
||
echo 0 > cpu0/cpuset.cpus
|
||
echo 0 > cpu0/cpuset.mems
|
||
echo 1 > cpuset.cpu_exclusive
|
||
echo 0 > cpuset.sched_load_balance
|
||
echo 1 > cpu0/cpuset.cpu_exclusive
|
||
echo 1 > cpu0/cpuset.mem_exclusive
|
||
echo $$ > cpu0/tasks
|
||
rt-app -t 100000:10000:d:0 -D5 # it is now actually superfluous to specify
|
||
# task affinity
|
||
|
||
6. Future plans
|
||
===============
|
||
|
||
Still missing:
|
||
|
||
- programmatic way to retrieve current runtime and absolute deadline
|
||
- refinements to deadline inheritance, especially regarding the possibility
|
||
of retaining bandwidth isolation among non-interacting tasks. This is
|
||
being studied from both theoretical and practical points of view, and
|
||
hopefully we should be able to produce some demonstrative code soon;
|
||
- (c)group based bandwidth management, and maybe scheduling;
|
||
- access control for non-root users (and related security concerns to
|
||
address), which is the best way to allow unprivileged use of the mechanisms
|
||
and how to prevent non-root users "cheat" the system?
|
||
|
||
As already discussed, we are planning also to merge this work with the EDF
|
||
throttling patches [https://lore.kernel.org/r/cover.1266931410.git.fabio@helm.retis] but we still are in
|
||
the preliminary phases of the merge and we really seek feedback that would
|
||
help us decide on the direction it should take.
|
||
|
||
Appendix A. Test suite
|
||
======================
|
||
|
||
The SCHED_DEADLINE policy can be easily tested using two applications that
|
||
are part of a wider Linux Scheduler validation suite. The suite is
|
||
available as a GitHub repository: https://github.com/scheduler-tools.
|
||
|
||
The first testing application is called rt-app and can be used to
|
||
start multiple threads with specific parameters. rt-app supports
|
||
SCHED_{OTHER,FIFO,RR,DEADLINE} scheduling policies and their related
|
||
parameters (e.g., niceness, priority, runtime/deadline/period). rt-app
|
||
is a valuable tool, as it can be used to synthetically recreate certain
|
||
workloads (maybe mimicking real use-cases) and evaluate how the scheduler
|
||
behaves under such workloads. In this way, results are easily reproducible.
|
||
rt-app is available at: https://github.com/scheduler-tools/rt-app.
|
||
|
||
Thread parameters can be specified from the command line, with something like
|
||
this::
|
||
|
||
# rt-app -t 100000:10000:d -t 150000:20000:f:10 -D5
|
||
|
||
The above creates 2 threads. The first one, scheduled by SCHED_DEADLINE,
|
||
executes for 10ms every 100ms. The second one, scheduled at SCHED_FIFO
|
||
priority 10, executes for 20ms every 150ms. The test will run for a total
|
||
of 5 seconds.
|
||
|
||
More interestingly, configurations can be described with a json file that
|
||
can be passed as input to rt-app with something like this::
|
||
|
||
# rt-app my_config.json
|
||
|
||
The parameters that can be specified with the second method are a superset
|
||
of the command line options. Please refer to rt-app documentation for more
|
||
details (`<rt-app-sources>/doc/*.json`).
|
||
|
||
The second testing application is a modification of schedtool, called
|
||
schedtool-dl, which can be used to setup SCHED_DEADLINE parameters for a
|
||
certain pid/application. schedtool-dl is available at:
|
||
https://github.com/scheduler-tools/schedtool-dl.git.
|
||
|
||
The usage is straightforward::
|
||
|
||
# schedtool -E -t 10000000:100000000 -e ./my_cpuhog_app
|
||
|
||
With this, my_cpuhog_app is put to run inside a SCHED_DEADLINE reservation
|
||
of 10ms every 100ms (note that parameters are expressed in microseconds).
|
||
You can also use schedtool to create a reservation for an already running
|
||
application, given that you know its pid::
|
||
|
||
# schedtool -E -t 10000000:100000000 my_app_pid
|
||
|
||
Appendix B. Minimal main()
|
||
==========================
|
||
|
||
We provide in what follows a simple (ugly) self-contained code snippet
|
||
showing how SCHED_DEADLINE reservations can be created by a real-time
|
||
application developer::
|
||
|
||
#define _GNU_SOURCE
|
||
#include <unistd.h>
|
||
#include <stdio.h>
|
||
#include <stdlib.h>
|
||
#include <string.h>
|
||
#include <time.h>
|
||
#include <linux/unistd.h>
|
||
#include <linux/kernel.h>
|
||
#include <linux/types.h>
|
||
#include <sys/syscall.h>
|
||
#include <pthread.h>
|
||
|
||
#define gettid() syscall(__NR_gettid)
|
||
|
||
#define SCHED_DEADLINE 6
|
||
|
||
/* XXX use the proper syscall numbers */
|
||
#ifdef __x86_64__
|
||
#define __NR_sched_setattr 314
|
||
#define __NR_sched_getattr 315
|
||
#endif
|
||
|
||
#ifdef __i386__
|
||
#define __NR_sched_setattr 351
|
||
#define __NR_sched_getattr 352
|
||
#endif
|
||
|
||
#ifdef __arm__
|
||
#define __NR_sched_setattr 380
|
||
#define __NR_sched_getattr 381
|
||
#endif
|
||
|
||
static volatile int done;
|
||
|
||
struct sched_attr {
|
||
__u32 size;
|
||
|
||
__u32 sched_policy;
|
||
__u64 sched_flags;
|
||
|
||
/* SCHED_NORMAL, SCHED_BATCH */
|
||
__s32 sched_nice;
|
||
|
||
/* SCHED_FIFO, SCHED_RR */
|
||
__u32 sched_priority;
|
||
|
||
/* SCHED_DEADLINE (nsec) */
|
||
__u64 sched_runtime;
|
||
__u64 sched_deadline;
|
||
__u64 sched_period;
|
||
};
|
||
|
||
int sched_setattr(pid_t pid,
|
||
const struct sched_attr *attr,
|
||
unsigned int flags)
|
||
{
|
||
return syscall(__NR_sched_setattr, pid, attr, flags);
|
||
}
|
||
|
||
int sched_getattr(pid_t pid,
|
||
struct sched_attr *attr,
|
||
unsigned int size,
|
||
unsigned int flags)
|
||
{
|
||
return syscall(__NR_sched_getattr, pid, attr, size, flags);
|
||
}
|
||
|
||
void *run_deadline(void *data)
|
||
{
|
||
struct sched_attr attr;
|
||
int x = 0;
|
||
int ret;
|
||
unsigned int flags = 0;
|
||
|
||
printf("deadline thread started [%ld]\n", gettid());
|
||
|
||
attr.size = sizeof(attr);
|
||
attr.sched_flags = 0;
|
||
attr.sched_nice = 0;
|
||
attr.sched_priority = 0;
|
||
|
||
/* This creates a 10ms/30ms reservation */
|
||
attr.sched_policy = SCHED_DEADLINE;
|
||
attr.sched_runtime = 10 * 1000 * 1000;
|
||
attr.sched_period = attr.sched_deadline = 30 * 1000 * 1000;
|
||
|
||
ret = sched_setattr(0, &attr, flags);
|
||
if (ret < 0) {
|
||
done = 0;
|
||
perror("sched_setattr");
|
||
exit(-1);
|
||
}
|
||
|
||
while (!done) {
|
||
x++;
|
||
}
|
||
|
||
printf("deadline thread dies [%ld]\n", gettid());
|
||
return NULL;
|
||
}
|
||
|
||
int main (int argc, char **argv)
|
||
{
|
||
pthread_t thread;
|
||
|
||
printf("main thread [%ld]\n", gettid());
|
||
|
||
pthread_create(&thread, NULL, run_deadline, NULL);
|
||
|
||
sleep(10);
|
||
|
||
done = 1;
|
||
pthread_join(thread, NULL);
|
||
|
||
printf("main dies [%ld]\n", gettid());
|
||
return 0;
|
||
}
|