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Merge branch 'lkmm-for-mingo' of git://git.kernel.org/pub/scm/linux/kernel/git/paulmck/linux-rcu into locking/core

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Pull various memory-model (LKMM) updates from Paul E. McKenney.

Signed-off-by: Ingo Molnar <mingo@kernel.org>
This commit is contained in:
Ingo Molnar 2019-04-10 09:14:55 +02:00
commit f7c2b7477b
8 changed files with 287 additions and 181 deletions
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17
Documentation/atomic_t.txt
View File

@ -56,6 +56,23 @@ Barriers:
smp_mb__{before,after}_atomic()
TYPES (signed vs unsigned)
-----
While atomic_t, atomic_long_t and atomic64_t use int, long and s64
respectively (for hysterical raisins), the kernel uses -fno-strict-overflow
(which implies -fwrapv) and defines signed overflow to behave like
2s-complement.
Therefore, an explicitly unsigned variant of the atomic ops is strictly
unnecessary and we can simply cast, there is no UB.
There was a bug in UBSAN prior to GCC-8 that would generate UB warnings for
signed types.
With this we also conform to the C/C++ _Atomic behaviour and things like
P1236R1.
SEMANTICS
---------

49
Documentation/translations/ko_KR/memory-barriers.txt
View File

@ -493,10 +493,8 @@ CPU 에게 기대할 수 있는 최소한의 보장사항 몇가지가 있습니
이 타입의 오퍼레이션은 단방향의 투과성 배리어처럼 동작합니다. ACQUIRE
오퍼레이션 뒤의 모든 메모리 오퍼레이션들이 ACQUIRE 오퍼레이션 후에
일어난 것으로 시스템의 나머지 컴포넌트들에 보이게 될 것이 보장됩니다.
LOCK 오퍼레이션과 smp_load_acquire(), smp_cond_acquire() 오퍼레이션도
ACQUIRE 오퍼레이션에 포함됩니다. smp_cond_acquire() 오퍼레이션은 컨트롤
의존성과 smp_rmb() 를 사용해서 ACQUIRE 의 의미적 요구사항(semantic)을
충족시킵니다.
LOCK 오퍼레이션과 smp_load_acquire(), smp_cond_load_acquire() 오퍼레이션도
ACQUIRE 오퍼레이션에 포함됩니다.
ACQUIRE 오퍼레이션 앞의 메모리 오퍼레이션들은 ACQUIRE 오퍼레이션 완료 후에
수행된 것처럼 보일 수 있습니다.
@ -2146,33 +2144,40 @@ set_current_state() 는 다음의 것들로 감싸질 수도 있습니다:
event_indicated = 1;
wake_up_process(event_daemon);
wake_up() 류에 의해 쓰기 메모리 배리어가 내포됩니다. 만약 그것들이 뭔가를
깨운다면요. 이 배리어는 태스크 상태가 지워지기 전에 수행되므로, 이벤트를
알리기 위한 STORE 와 태스크 상태를 TASK_RUNNING 으로 설정하는 STORE 사이에
위치하게 됩니다.
wake_up() 이 무언가를 깨우게 되면, 이 함수는 범용 메모리 배리어를 수행합니다.
이 함수가 아무것도 깨우지 않는다면 메모리 배리어는 수행될 수도, 수행되지 않을
수도 있습니다; 이 경우에 메모리 배리어를 수행할 거라 오해해선 안됩니다. 이
배리어는 태스크 상태가 접근되기 전에 수행되는데, 자세히 말하면 이 이벤트를
알리기 위한 STORE 와 TASK_RUNNING 으로 상태를 쓰는 STORE 사이에 수행됩니다:
CPU 1 CPU 2
CPU 1 (Sleeper) CPU 2 (Waker)
=============================== ===============================
set_current_state(); STORE event_indicated
smp_store_mb(); wake_up();
STORE current->state <쓰기 배리어>
<범용 배리어> STORE current->state
LOAD event_indicated
STORE current->state ...
<범용 배리어> <범용 배리어>
LOAD event_indicated if ((LOAD task->state) & TASK_NORMAL)
STORE task->state
한번더 말합니다만, 이 쓰기 메모리 배리어는 이 코드가 정말로 뭔가를 깨울 때에만
실행됩니다. 이걸 설명하기 위해, X 와 Y 는 모두 0 으로 초기화 되어 있다는 가정
하에 아래의 이벤트 시퀀스를 생각해 봅시다:
여기서 "task" 는 깨어나지는 쓰레드이고 CPU 1 의 "current" 와 같습니다.
반복하지만, wake_up() 이 무언가를 정말 깨운다면 범용 메모리 배리어가 수행될
것이 보장되지만, 그렇지 않다면 그런 보장이 없습니다. 이걸 이해하기 위해, X 와
Y 는 모두 0 으로 초기화 되어 있다는 가정 하에 아래의 이벤트 시퀀스를 생각해
봅시다:
CPU 1 CPU 2
=============================== ===============================
X = 1; STORE event_indicated
X = 1; Y = 1;
smp_mb(); wake_up();
Y = 1; wait_event(wq, Y == 1);
wake_up(); load from Y sees 1, no memory barrier
load from X might see 0
LOAD Y LOAD X
위 예제에서의 경우와 달리 깨우기가 정말로 행해졌다면, CPU 2 의 X 로드는 1 을
본다고 보장될 수 있을 겁니다.
정말로 깨우기가 행해졌다면, 두 로드 중 (최소한) 하나는 1 을 보게 됩니다.
반면에, 실제 깨우기가 행해지지 않았다면, 두 로드 모두 0을 볼 수도 있습니다.
wake_up_process() 는 항상 범용 메모리 배리어를 수행합니다. 이 배리어 역시
태스크 상태가 접근되기 전에 수행됩니다. 특히, 앞의 예제 코드에서 wake_up() 이
wake_up_process() 로 대체된다면 두 로드 중 하나는 1을 볼 것이 보장됩니다.
사용 가능한 깨우기류 함수들로 다음과 같은 것들이 있습니다:
@ -2192,6 +2197,8 @@ wake_up() 류에 의해 쓰기 메모리 배리어가 내포됩니다. 만약
wake_up_poll();
wake_up_process();
메모리 순서규칙 관점에서, 이 함수들은 모두 wake_up() 과 같거나 보다 강한 순서
보장을 제공합니다.
[!] 잠재우는 코드와 깨우는 코드에 내포되는 메모리 배리어들은 깨우기 전에
이루어진 스토어를 잠재우는 코드가 set_current_state() 를 호출한 후에 행하는

287
tools/memory-model/Documentation/explanation.txt
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@ -27,7 +27,7 @@ Explanation of the Linux-Kernel Memory Consistency Model
19. AND THEN THERE WAS ALPHA
20. THE HAPPENS-BEFORE RELATION: hb
21. THE PROPAGATES-BEFORE RELATION: pb
22. RCU RELATIONS: rcu-link, gp, rscs, rcu-fence, and rb
22. RCU RELATIONS: rcu-link, rcu-gp, rcu-rscsi, rcu-fence, and rb
23. LOCKING
24. ODDS AND ENDS
@ -1430,8 +1430,8 @@ they execute means that it cannot have cycles. This requirement is
the content of the LKMM's "propagation" axiom.
RCU RELATIONS: rcu-link, gp, rscs, rcu-fence, and rb
----------------------------------------------------
RCU RELATIONS: rcu-link, rcu-gp, rcu-rscsi, rcu-fence, and rb
-------------------------------------------------------------
RCU (Read-Copy-Update) is a powerful synchronization mechanism. It
rests on two concepts: grace periods and read-side critical sections.
@ -1446,17 +1446,19 @@ As far as memory models are concerned, RCU's main feature is its
Grace-Period Guarantee, which states that a critical section can never
span a full grace period. In more detail, the Guarantee says:
If a critical section starts before a grace period then it
must end before the grace period does. In addition, every
store that propagates to the critical section's CPU before the
end of the critical section must propagate to every CPU before
the end of the grace period.
For any critical section C and any grace period G, at least
one of the following statements must hold:
If a critical section ends after a grace period ends then it
must start after the grace period does. In addition, every
store that propagates to the grace period's CPU before the
start of the grace period must propagate to every CPU before
the start of the critical section.
(1) C ends before G does, and in addition, every store that
propagates to C's CPU before the end of C must propagate to
every CPU before G ends.
(2) G starts before C does, and in addition, every store that
propagates to G's CPU before the start of G must propagate
to every CPU before C starts.
In particular, it is not possible for a critical section to both start
before and end after a grace period.
Here is a simple example of RCU in action:
@ -1483,10 +1485,11 @@ The Grace Period Guarantee tells us that when this code runs, it will
never end with r1 = 1 and r2 = 0. The reasoning is as follows. r1 = 1
means that P0's store to x propagated to P1 before P1 called
synchronize_rcu(), so P0's critical section must have started before
P1's grace period. On the other hand, r2 = 0 means that P0's store to
y, which occurs before the end of the critical section, did not
propagate to P1 before the end of the grace period, violating the
Guarantee.
P1's grace period, contrary to part (2) of the Guarantee. On the
other hand, r2 = 0 means that P0's store to y, which occurs before the
end of the critical section, did not propagate to P1 before the end of
the grace period, contrary to part (1). Together the results violate
the Guarantee.
In the kernel's implementations of RCU, the requirements for stores
to propagate to every CPU are fulfilled by placing strong fences at
@ -1504,11 +1507,11 @@ before" or "ends after" a grace period? Some aspects of the meaning
are pretty obvious, as in the example above, but the details aren't
entirely clear. The LKMM formalizes this notion by means of the
rcu-link relation. rcu-link encompasses a very general notion of
"before": Among other things, X ->rcu-link Z includes cases where X
happens-before or is equal to some event Y which is equal to or comes
before Z in the coherence order. When Y = Z this says that X ->rfe Z
implies X ->rcu-link Z. In addition, when Y = X it says that X ->fr Z
and X ->co Z each imply X ->rcu-link Z.
"before": If E and F are RCU fence events (i.e., rcu_read_lock(),
rcu_read_unlock(), or synchronize_rcu()) then among other things,
E ->rcu-link F includes cases where E is po-before some memory-access
event X, F is po-after some memory-access event Y, and we have any of
X ->rfe Y, X ->co Y, or X ->fr Y.
The formal definition of the rcu-link relation is more than a little
obscure, and we won't give it here. It is closely related to the pb
@ -1516,171 +1519,173 @@ relation, and the details don't matter unless you want to comb through
a somewhat lengthy formal proof. Pretty much all you need to know
about rcu-link is the information in the preceding paragraph.
The LKMM also defines the gp and rscs relations. They bring grace
periods and read-side critical sections into the picture, in the
The LKMM also defines the rcu-gp and rcu-rscsi relations. They bring
grace periods and read-side critical sections into the picture, in the
following way:
E ->gp F means there is a synchronize_rcu() fence event S such
that E ->po S and either S ->po F or S = F. In simple terms,
there is a grace period po-between E and F.
E ->rcu-gp F means that E and F are in fact the same event,
and that event is a synchronize_rcu() fence (i.e., a grace
period).
E ->rscs F means there is a critical section delimited by an
rcu_read_lock() fence L and an rcu_read_unlock() fence U, such
that E ->po U and either L ->po F or L = F. You can think of
this as saying that E and F are in the same critical section
(in fact, it also allows E to be po-before the start of the
critical section and F to be po-after the end).
E ->rcu-rscsi F means that E and F are the rcu_read_unlock()
and rcu_read_lock() fence events delimiting some read-side
critical section. (The 'i' at the end of the name emphasizes
that this relation is "inverted": It links the end of the
critical section to the start.)
If we think of the rcu-link relation as standing for an extended
"before", then X ->gp Y ->rcu-link Z says that X executes before a
grace period which ends before Z executes. (In fact it covers more
than this, because it also includes cases where X executes before a
grace period and some store propagates to Z's CPU before Z executes
but doesn't propagate to some other CPU until after the grace period
ends.) Similarly, X ->rscs Y ->rcu-link Z says that X is part of (or
before the start of) a critical section which starts before Z
executes.
"before", then X ->rcu-gp Y ->rcu-link Z roughly says that X is a
grace period which ends before Z begins. (In fact it covers more than
this, because it also includes cases where some store propagates to
Z's CPU before Z begins but doesn't propagate to some other CPU until
after X ends.) Similarly, X ->rcu-rscsi Y ->rcu-link Z says that X is
the end of a critical section which starts before Z begins.
The LKMM goes on to define the rcu-fence relation as a sequence of gp
and rscs links separated by rcu-link links, in which the number of gp
links is >= the number of rscs links. For example:
The LKMM goes on to define the rcu-fence relation as a sequence of
rcu-gp and rcu-rscsi links separated by rcu-link links, in which the
number of rcu-gp links is >= the number of rcu-rscsi links. For
example:
X ->gp Y ->rcu-link Z ->rscs T ->rcu-link U ->gp V
X ->rcu-gp Y ->rcu-link Z ->rcu-rscsi T ->rcu-link U ->rcu-gp V
would imply that X ->rcu-fence V, because this sequence contains two
gp links and only one rscs link. (It also implies that X ->rcu-fence T
and Z ->rcu-fence V.) On the other hand:
rcu-gp links and one rcu-rscsi link. (It also implies that
X ->rcu-fence T and Z ->rcu-fence V.) On the other hand:
X ->rscs Y ->rcu-link Z ->rscs T ->rcu-link U ->gp V
X ->rcu-rscsi Y ->rcu-link Z ->rcu-rscsi T ->rcu-link U ->rcu-gp V
does not imply X ->rcu-fence V, because the sequence contains only
one gp link but two rscs links.
one rcu-gp link but two rcu-rscsi links.
The rcu-fence relation is important because the Grace Period Guarantee
means that rcu-fence acts kind of like a strong fence. In particular,
if W is a write and we have W ->rcu-fence Z, the Guarantee says that W
will propagate to every CPU before Z executes.
E ->rcu-fence F implies not only that E begins before F ends, but also
that any write po-before E will propagate to every CPU before any
instruction po-after F can execute. (However, it does not imply that
E must execute before F; in fact, each synchronize_rcu() fence event
is linked to itself by rcu-fence as a degenerate case.)
To prove this in full generality requires some intellectual effort.
We'll consider just a very simple case:
W ->gp X ->rcu-link Y ->rscs Z.
G ->rcu-gp W ->rcu-link Z ->rcu-rscsi F.
This formula means that there is a grace period G and a critical
section C such that:
This formula means that G and W are the same event (a grace period),
and there are events X, Y and a read-side critical section C such that:
1. W is po-before G;
1. G = W is po-before or equal to X;
2. X is equal to or po-after G;
2. X comes "before" Y in some sense (including rfe, co and fr);
3. X comes "before" Y in some sense;
2. Y is po-before Z;
4. Y is po-before the end of C;
4. Z is the rcu_read_unlock() event marking the end of C;
5. Z is equal to or po-after the start of C.
5. F is the rcu_read_lock() event marking the start of C.
From 2 - 4 we deduce that the grace period G ends before the critical
section C. Then the second part of the Grace Period Guarantee says
not only that G starts before C does, but also that W (which executes
on G's CPU before G starts) must propagate to every CPU before C
starts. In particular, W propagates to every CPU before Z executes
(or finishes executing, in the case where Z is equal to the
rcu_read_lock() fence event which starts C.) This sort of reasoning
can be expanded to handle all the situations covered by rcu-fence.
From 1 - 4 we deduce that the grace period G ends before the critical
section C. Then part (2) of the Grace Period Guarantee says not only
that G starts before C does, but also that any write which executes on
G's CPU before G starts must propagate to every CPU before C starts.
In particular, the write propagates to every CPU before F finishes
executing and hence before any instruction po-after F can execute.
This sort of reasoning can be extended to handle all the situations
covered by rcu-fence.
Finally, the LKMM defines the RCU-before (rb) relation in terms of
rcu-fence. This is done in essentially the same way as the pb
relation was defined in terms of strong-fence. We will omit the
details; the end result is that E ->rb F implies E must execute before
F, just as E ->pb F does (and for much the same reasons).
details; the end result is that E ->rb F implies E must execute
before F, just as E ->pb F does (and for much the same reasons).
Putting this all together, the LKMM expresses the Grace Period
Guarantee by requiring that the rb relation does not contain a cycle.
Equivalently, this "rcu" axiom requires that there are no events E and
F with E ->rcu-link F ->rcu-fence E. Or to put it a third way, the
axiom requires that there are no cycles consisting of gp and rscs
alternating with rcu-link, where the number of gp links is >= the
number of rscs links.
Equivalently, this "rcu" axiom requires that there are no events E
and F with E ->rcu-link F ->rcu-fence E. Or to put it a third way,
the axiom requires that there are no cycles consisting of rcu-gp and
rcu-rscsi alternating with rcu-link, where the number of rcu-gp links
is >= the number of rcu-rscsi links.
Justifying the axiom isn't easy, but it is in fact a valid
formalization of the Grace Period Guarantee. We won't attempt to go
through the detailed argument, but the following analysis gives a
taste of what is involved. Suppose we have a violation of the first
part of the Guarantee: A critical section starts before a grace
period, and some store propagates to the critical section's CPU before
the end of the critical section but doesn't propagate to some other
CPU until after the end of the grace period.
taste of what is involved. Suppose both parts of the Guarantee are
violated: A critical section starts before a grace period, and some
store propagates to the critical section's CPU before the end of the
critical section but doesn't propagate to some other CPU until after
the end of the grace period.
Putting symbols to these ideas, let L and U be the rcu_read_lock() and
rcu_read_unlock() fence events delimiting the critical section in
question, and let S be the synchronize_rcu() fence event for the grace
period. Saying that the critical section starts before S means there
are events E and F where E is po-after L (which marks the start of the
critical section), E is "before" F in the sense of the rcu-link
relation, and F is po-before the grace period S:
are events Q and R where Q is po-after L (which marks the start of the
critical section), Q is "before" R in the sense used by the rcu-link
relation, and R is po-before the grace period S. Thus we have:
L ->po E ->rcu-link F ->po S.
L ->rcu-link S.
Let W be the store mentioned above, let Z come before the end of the
Let W be the store mentioned above, let Y come before the end of the
critical section and witness that W propagates to the critical
section's CPU by reading from W, and let Y on some arbitrary CPU be a
witness that W has not propagated to that CPU, where Y happens after
section's CPU by reading from W, and let Z on some arbitrary CPU be a
witness that W has not propagated to that CPU, where Z happens after
some event X which is po-after S. Symbolically, this amounts to:
S ->po X ->hb* Y ->fr W ->rf Z ->po U.
S ->po X ->hb* Z ->fr W ->rf Y ->po U.
The fr link from Y to W indicates that W has not propagated to Y's CPU
at the time that Y executes. From this, it can be shown (see the
discussion of the rcu-link relation earlier) that X and Z are related
by rcu-link, yielding:
The fr link from Z to W indicates that W has not propagated to Z's CPU
at the time that Z executes. From this, it can be shown (see the
discussion of the rcu-link relation earlier) that S and U are related
by rcu-link:
S ->po X ->rcu-link Z ->po U.
S ->rcu-link U.
The formulas say that S is po-between F and X, hence F ->gp X. They
also say that Z comes before the end of the critical section and E
comes after its start, hence Z ->rscs E. From all this we obtain:
Since S is a grace period we have S ->rcu-gp S, and since L and U are
the start and end of the critical section C we have U ->rcu-rscsi L.
From this we obtain:
F ->gp X ->rcu-link Z ->rscs E ->rcu-link F,
S ->rcu-gp S ->rcu-link U ->rcu-rscsi L ->rcu-link S,
a forbidden cycle. Thus the "rcu" axiom rules out this violation of
the Grace Period Guarantee.
For something a little more down-to-earth, let's see how the axiom
works out in practice. Consider the RCU code example from above, this
time with statement labels added to the memory access instructions:
time with statement labels added:
int x, y;
P0()
{
rcu_read_lock();
W: WRITE_ONCE(x, 1);
X: WRITE_ONCE(y, 1);
rcu_read_unlock();
L: rcu_read_lock();
X: WRITE_ONCE(x, 1);
Y: WRITE_ONCE(y, 1);
U: rcu_read_unlock();
}
P1()
{
int r1, r2;
Y: r1 = READ_ONCE(x);
synchronize_rcu();
Z: r2 = READ_ONCE(y);
Z: r1 = READ_ONCE(x);
S: synchronize_rcu();
W: r2 = READ_ONCE(y);
}
If r2 = 0 at the end then P0's store at X overwrites the value that
P1's load at Z reads from, so we have Z ->fre X and thus Z ->rcu-link X.
In addition, there is a synchronize_rcu() between Y and Z, so therefore
we have Y ->gp Z.
If r2 = 0 at the end then P0's store at Y overwrites the value that
P1's load at W reads from, so we have W ->fre Y. Since S ->po W and
also Y ->po U, we get S ->rcu-link U. In addition, S ->rcu-gp S
because S is a grace period.
If r1 = 1 at the end then P1's load at Y reads from P0's store at W,
so we have W ->rcu-link Y. In addition, W and X are in the same critical
section, so therefore we have X ->rscs W.
If r1 = 1 at the end then P1's load at Z reads from P0's store at X,
so we have X ->rfe Z. Together with L ->po X and Z ->po S, this
yields L ->rcu-link S. And since L and U are the start and end of a
critical section, we have U ->rcu-rscsi L.
Then X ->rscs W ->rcu-link Y ->gp Z ->rcu-link X is a forbidden cycle,
violating the "rcu" axiom. Hence the outcome is not allowed by the
LKMM, as we would expect.
Then U ->rcu-rscsi L ->rcu-link S ->rcu-gp S ->rcu-link U is a
forbidden cycle, violating the "rcu" axiom. Hence the outcome is not
allowed by the LKMM, as we would expect.
For contrast, let's see what can happen in a more complicated example:
@ -1690,51 +1695,52 @@ For contrast, let's see what can happen in a more complicated example:
{
int r0;
rcu_read_lock();
W: r0 = READ_ONCE(x);
X: WRITE_ONCE(y, 1);
rcu_read_unlock();
L0: rcu_read_lock();
r0 = READ_ONCE(x);
WRITE_ONCE(y, 1);
U0: rcu_read_unlock();
}
P1()
{
int r1;
Y: r1 = READ_ONCE(y);
synchronize_rcu();
Z: WRITE_ONCE(z, 1);
r1 = READ_ONCE(y);
S1: synchronize_rcu();
WRITE_ONCE(z, 1);
}
P2()
{
int r2;
rcu_read_lock();
U: r2 = READ_ONCE(z);
V: WRITE_ONCE(x, 1);
rcu_read_unlock();
L2: rcu_read_lock();
r2 = READ_ONCE(z);
WRITE_ONCE(x, 1);
U2: rcu_read_unlock();
}
If r0 = r1 = r2 = 1 at the end, then similar reasoning to before shows
that W ->rscs X ->rcu-link Y ->gp Z ->rcu-link U ->rscs V ->rcu-link W.
However this cycle is not forbidden, because the sequence of relations
contains fewer instances of gp (one) than of rscs (two). Consequently
the outcome is allowed by the LKMM. The following instruction timing
diagram shows how it might actually occur:
that U0 ->rcu-rscsi L0 ->rcu-link S1 ->rcu-gp S1 ->rcu-link U2 ->rcu-rscsi
L2 ->rcu-link U0. However this cycle is not forbidden, because the
sequence of relations contains fewer instances of rcu-gp (one) than of
rcu-rscsi (two). Consequently the outcome is allowed by the LKMM.
The following instruction timing diagram shows how it might actually
occur:
P0 P1 P2
-------------------- -------------------- --------------------
rcu_read_lock()
X: WRITE_ONCE(y, 1)
Y: r1 = READ_ONCE(y)
WRITE_ONCE(y, 1)
r1 = READ_ONCE(y)
synchronize_rcu() starts
. rcu_read_lock()
. V: WRITE_ONCE(x, 1)
W: r0 = READ_ONCE(x) .
. WRITE_ONCE(x, 1)
r0 = READ_ONCE(x) .
rcu_read_unlock() .
synchronize_rcu() ends
Z: WRITE_ONCE(z, 1)
U: r2 = READ_ONCE(z)
WRITE_ONCE(z, 1)
r2 = READ_ONCE(z)
rcu_read_unlock()
This requires P0 and P2 to execute their loads and stores out of
@ -1744,6 +1750,15 @@ section in P0 both starts before P1's grace period does and ends
before it does, and the critical section in P2 both starts after P1's
grace period does and ends after it does.
Addendum: The LKMM now supports SRCU (Sleepable Read-Copy-Update) in
addition to normal RCU. The ideas involved are much the same as
above, with new relations srcu-gp and srcu-rscsi added to represent
SRCU grace periods and read-side critical sections. There is a
restriction on the srcu-gp and srcu-rscsi links that can appear in an
rcu-fence sequence (the srcu-rscsi links must be paired with srcu-gp
links having the same SRCU domain with proper nesting); the details
are relatively unimportant.
LOCKING
-------

33
tools/memory-model/README
View File

@ -20,13 +20,17 @@ that litmus test to be exercised within the Linux kernel.
REQUIREMENTS
============
Version 7.49 of the "herd7" and "klitmus7" tools must be downloaded
separately:
Version 7.52 or higher of the "herd7" and "klitmus7" tools must be
downloaded separately:
https://github.com/herd/herdtools7
See "herdtools7/INSTALL.md" for installation instructions.
Note that although these tools usually provide backwards compatibility,
this is not absolutely guaranteed. Therefore, if a later version does
not work, please try using the exact version called out above.
==================
BASIC USAGE: HERD7
@ -221,8 +225,29 @@ The Linux-kernel memory model has the following limitations:
additional call_rcu() process to the site of the
emulated rcu-barrier().
e. Sleepable RCU (SRCU) is not modeled. It can be
emulated, but perhaps not simply.
e. Although sleepable RCU (SRCU) is now modeled, there
are some subtle differences between its semantics and
those in the Linux kernel. For example, the kernel
might interpret the following sequence as two partially
overlapping SRCU read-side critical sections:
1 r1 = srcu_read_lock(&my_srcu);
2 do_something_1();
3 r2 = srcu_read_lock(&my_srcu);
4 do_something_2();
5 srcu_read_unlock(&my_srcu, r1);
6 do_something_3();
7 srcu_read_unlock(&my_srcu, r2);
In contrast, LKMM will interpret this as a nested pair of
SRCU read-side critical sections, with the outer critical
section spanning lines 1-7 and the inner critical section
spanning lines 3-5.
This difference would be more of a concern had anyone
identified a reasonable use case for partially overlapping
SRCU read-side critical sections. For more information,
please see: https://paulmck.livejournal.com/40593.html
f. Reader-writer locking is not modeled. It can be
emulated in litmus tests using atomic read-modify-write

35
tools/memory-model/linux-kernel.bell
View File

@ -33,8 +33,14 @@ enum Barriers = 'wmb (*smp_wmb*) ||
'after-unlock-lock (*smp_mb__after_unlock_lock*)
instructions F[Barriers]
(* SRCU *)
enum SRCU = 'srcu-lock || 'srcu-unlock || 'sync-srcu
instructions SRCU[SRCU]
(* All srcu events *)
let Srcu = Srcu-lock | Srcu-unlock | Sync-srcu
(* Compute matching pairs of nested Rcu-lock and Rcu-unlock *)
let matched = let rec
let rcu-rscs = let rec
unmatched-locks = Rcu-lock \ domain(matched)
and unmatched-unlocks = Rcu-unlock \ range(matched)
and unmatched = unmatched-locks | unmatched-unlocks
@ -46,8 +52,27 @@ let matched = let rec
in matched
(* Validate nesting *)
flag ~empty Rcu-lock \ domain(matched) as unbalanced-rcu-locking
flag ~empty Rcu-unlock \ range(matched) as unbalanced-rcu-locking
flag ~empty Rcu-lock \ domain(rcu-rscs) as unbalanced-rcu-locking
flag ~empty Rcu-unlock \ range(rcu-rscs) as unbalanced-rcu-locking
(* Outermost level of nesting only *)
let crit = matched \ (po^-1 ; matched ; po^-1)
(* Compute matching pairs of nested Srcu-lock and Srcu-unlock *)
let srcu-rscs = let rec
unmatched-locks = Srcu-lock \ domain(matched)
and unmatched-unlocks = Srcu-unlock \ range(matched)
and unmatched = unmatched-locks | unmatched-unlocks
and unmatched-po = ([unmatched] ; po ; [unmatched]) & loc
and unmatched-locks-to-unlocks =
([unmatched-locks] ; po ; [unmatched-unlocks]) & loc
and matched = matched | (unmatched-locks-to-unlocks \
(unmatched-po ; unmatched-po))
in matched
(* Validate nesting *)
flag ~empty Srcu-lock \ domain(srcu-rscs) as unbalanced-srcu-locking
flag ~empty Srcu-unlock \ range(srcu-rscs) as unbalanced-srcu-locking
(* Check for use of synchronize_srcu() inside an RCU critical section *)
flag ~empty rcu-rscs & (po ; [Sync-srcu] ; po) as invalid-sleep
(* Validate SRCU dynamic match *)
flag ~empty different-values(srcu-rscs) as srcu-bad-nesting

39
tools/memory-model/linux-kernel.cat
View File

@ -33,7 +33,7 @@ let mb = ([M] ; fencerel(Mb) ; [M]) |
([M] ; po? ; [LKW] ; fencerel(After-spinlock) ; [M]) |
([M] ; po ; [UL] ; (co | po) ; [LKW] ;
fencerel(After-unlock-lock) ; [M])
let gp = po ; [Sync-rcu] ; po?
let gp = po ; [Sync-rcu | Sync-srcu] ; po?
let strong-fence = mb | gp
@ -91,32 +91,47 @@ acyclic pb as propagation
(*******)
(*
* Effect of read-side critical section proceeds from the rcu_read_lock()
* onward on the one hand and from the rcu_read_unlock() backwards on the
* other hand.
* Effects of read-side critical sections proceed from the rcu_read_unlock()
* or srcu_read_unlock() backwards on the one hand, and from the
* rcu_read_lock() or srcu_read_lock() forwards on the other hand.
*
* In the definition of rcu-fence below, the po term at the left-hand side
* of each disjunct and the po? term at the right-hand end have been factored
* out. They have been moved into the definitions of rcu-link and rb.
* This was necessary in order to apply the "& loc" tests correctly.
*)
let rscs = po ; crit^-1 ; po?
let rcu-gp = [Sync-rcu] (* Compare with gp *)
let srcu-gp = [Sync-srcu]
let rcu-rscsi = rcu-rscs^-1
let srcu-rscsi = srcu-rscs^-1
(*
* The synchronize_rcu() strong fence is special in that it can order not
* one but two non-rf relations, but only in conjunction with an RCU
* read-side critical section.
*)
let rcu-link = hb* ; pb* ; prop
let rcu-link = po? ; hb* ; pb* ; prop ; po
(*
* Any sequence containing at least as many grace periods as RCU read-side
* critical sections (joined by rcu-link) acts as a generalized strong fence.
* Likewise for SRCU grace periods and read-side critical sections, provided
* the synchronize_srcu() and srcu_read_[un]lock() calls refer to the same
* struct srcu_struct location.
*)
let rec rcu-fence = gp |
(gp ; rcu-link ; rscs) |
(rscs ; rcu-link ; gp) |
(gp ; rcu-link ; rcu-fence ; rcu-link ; rscs) |
(rscs ; rcu-link ; rcu-fence ; rcu-link ; gp) |
let rec rcu-fence = rcu-gp | srcu-gp |
(rcu-gp ; rcu-link ; rcu-rscsi) |
((srcu-gp ; rcu-link ; srcu-rscsi) & loc) |
(rcu-rscsi ; rcu-link ; rcu-gp) |
((srcu-rscsi ; rcu-link ; srcu-gp) & loc) |
(rcu-gp ; rcu-link ; rcu-fence ; rcu-link ; rcu-rscsi) |
((srcu-gp ; rcu-link ; rcu-fence ; rcu-link ; srcu-rscsi) & loc) |
(rcu-rscsi ; rcu-link ; rcu-fence ; rcu-link ; rcu-gp) |
((srcu-rscsi ; rcu-link ; rcu-fence ; rcu-link ; srcu-gp) & loc) |
(rcu-fence ; rcu-link ; rcu-fence)
(* rb orders instructions just as pb does *)
let rb = prop ; rcu-fence ; hb* ; pb*
let rb = prop ; po ; rcu-fence ; po? ; hb* ; pb*
irreflexive rb as rcu

5
tools/memory-model/linux-kernel.def
View File

@ -47,6 +47,11 @@ rcu_read_unlock() { __fence{rcu-unlock}; }
synchronize_rcu() { __fence{sync-rcu}; }
synchronize_rcu_expedited() { __fence{sync-rcu}; }
// SRCU
srcu_read_lock(X) __srcu{srcu-lock}(X)
srcu_read_unlock(X,Y) { __srcu{srcu-unlock}(X,Y); }
synchronize_srcu(X) { __srcu{sync-srcu}(X); }
// Atomic
atomic_read(X) READ_ONCE(*X)
atomic_set(X,V) { WRITE_ONCE(*X,V); }

3
tools/memory-model/lock.cat
View File

@ -6,9 +6,6 @@
(*
* Generate coherence orders and handle lock operations
*
* Warning: spin_is_locked() crashes herd7 versions strictly before 7.48.
* spin_is_locked() is functional from herd7 version 7.49.
*)
include "cross.cat"