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linux_dsm_epyc7002/Documentation/memory-barriers.txt
Linus Torvalds 1bf7067c6e Merge branch 'locking-core-for-linus' of git://git.kernel.org/pub/scm/linux/kernel/git/tip/tip 
Pull locking updates from Ingo Molnar:
 "The main changes are:

   - 'qspinlock' support, enabled on x86: queued spinlocks - these are
     now the spinlock variant used by x86 as they outperform ticket
     spinlocks in every category.  (Waiman Long)

   - 'pvqspinlock' support on x86: paravirtualized variant of queued
     spinlocks.  (Waiman Long, Peter Zijlstra)

   - 'qrwlock' support, enabled on x86: queued rwlocks.  Similar to
     queued spinlocks, they are now the variant used by x86:

       CONFIG_ARCH_USE_QUEUED_SPINLOCKS=y
       CONFIG_QUEUED_SPINLOCKS=y
       CONFIG_ARCH_USE_QUEUED_RWLOCKS=y
       CONFIG_QUEUED_RWLOCKS=y

   - various lockdep fixlets

   - various locking primitives cleanups, further WRITE_ONCE()
     propagation"

* 'locking-core-for-linus' of git://git.kernel.org/pub/scm/linux/kernel/git/tip/tip: (24 commits)
  locking/lockdep: Remove hard coded array size dependency
  locking/qrwlock: Don't contend with readers when setting _QW_WAITING
  lockdep: Do not break user-visible string
  locking/arch: Rename set_mb() to smp_store_mb()
  locking/arch: Add WRITE_ONCE() to set_mb()
  rtmutex: Warn if trylock is called from hard/softirq context
  arch: Remove __ARCH_HAVE_CMPXCHG
  locking/rtmutex: Drop usage of __HAVE_ARCH_CMPXCHG
  locking/qrwlock: Rename QUEUE_RWLOCK to QUEUED_RWLOCKS
  locking/pvqspinlock: Rename QUEUED_SPINLOCK to QUEUED_SPINLOCKS
  locking/pvqspinlock: Replace xchg() by the more descriptive set_mb()
  locking/pvqspinlock, x86: Enable PV qspinlock for Xen
  locking/pvqspinlock, x86: Enable PV qspinlock for KVM
  locking/pvqspinlock, x86: Implement the paravirt qspinlock call patching
  locking/pvqspinlock: Implement simple paravirt support for the qspinlock
  locking/qspinlock: Revert to test-and-set on hypervisors
  locking/qspinlock: Use a simple write to grab the lock
  locking/qspinlock: Optimize for smaller NR_CPUS
  locking/qspinlock: Extract out code snippets for the next patch
  locking/qspinlock: Add pending bit
  ...
2015-06-22 14:54:22 -07:00

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============================
LINUX KERNEL MEMORY BARRIERS
============================
By: David Howells <dhowells@redhat.com>
Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Contents:
(*) Abstract memory access model.
- Device operations.
- Guarantees.
(*) What are memory barriers?
- Varieties of memory barrier.
- What may not be assumed about memory barriers?
- Data dependency barriers.
- Control dependencies.
- SMP barrier pairing.
- Examples of memory barrier sequences.
- Read memory barriers vs load speculation.
- Transitivity
(*) Explicit kernel barriers.
- Compiler barrier.
- CPU memory barriers.
- MMIO write barrier.
(*) Implicit kernel memory barriers.
- Locking functions.
- Interrupt disabling functions.
- Sleep and wake-up functions.
- Miscellaneous functions.
(*) Inter-CPU locking barrier effects.
- Locks vs memory accesses.
- Locks vs I/O accesses.
(*) Where are memory barriers needed?
- Interprocessor interaction.
- Atomic operations.
- Accessing devices.
- Interrupts.
(*) Kernel I/O barrier effects.
(*) Assumed minimum execution ordering model.
(*) The effects of the cpu cache.
- Cache coherency.
- Cache coherency vs DMA.
- Cache coherency vs MMIO.
(*) The things CPUs get up to.
- And then there's the Alpha.
(*) Example uses.
- Circular buffers.
(*) References.
============================
ABSTRACT MEMORY ACCESS MODEL
============================
Consider the following abstract model of the system:
: :
: :
: :
+-------+ : +--------+ : +-------+
| | : | | : | |
| | : | | : | |
| CPU 1 |<----->| Memory |<----->| CPU 2 |
| | : | | : | |
| | : | | : | |
+-------+ : +--------+ : +-------+
^ : ^ : ^
| : | : |
| : | : |
| : v : |
| : +--------+ : |
| : | | : |
| : | | : |
+---------->| Device |<----------+
: | | :
: | | :
: +--------+ :
: :
Each CPU executes a program that generates memory access operations. In the
abstract CPU, memory operation ordering is very relaxed, and a CPU may actually
perform the memory operations in any order it likes, provided program causality
appears to be maintained. Similarly, the compiler may also arrange the
instructions it emits in any order it likes, provided it doesn't affect the
apparent operation of the program.
So in the above diagram, the effects of the memory operations performed by a
CPU are perceived by the rest of the system as the operations cross the
interface between the CPU and rest of the system (the dotted lines).
For example, consider the following sequence of events:
CPU 1 CPU 2
=============== ===============
{ A == 1; B == 2 }
A = 3; x = B;
B = 4; y = A;
The set of accesses as seen by the memory system in the middle can be arranged
in 24 different combinations:
STORE A=3, STORE B=4, y=LOAD A->3, x=LOAD B->4
STORE A=3, STORE B=4, x=LOAD B->4, y=LOAD A->3
STORE A=3, y=LOAD A->3, STORE B=4, x=LOAD B->4
STORE A=3, y=LOAD A->3, x=LOAD B->2, STORE B=4
STORE A=3, x=LOAD B->2, STORE B=4, y=LOAD A->3
STORE A=3, x=LOAD B->2, y=LOAD A->3, STORE B=4
STORE B=4, STORE A=3, y=LOAD A->3, x=LOAD B->4
STORE B=4, ...
...
and can thus result in four different combinations of values:
x == 2, y == 1
x == 2, y == 3
x == 4, y == 1
x == 4, y == 3
Furthermore, the stores committed by a CPU to the memory system may not be
perceived by the loads made by another CPU in the same order as the stores were
committed.
As a further example, consider this sequence of events:
CPU 1 CPU 2
=============== ===============
{ A == 1, B == 2, C = 3, P == &A, Q == &C }
B = 4; Q = P;
P = &B D = *Q;
There is an obvious data dependency here, as the value loaded into D depends on
the address retrieved from P by CPU 2. At the end of the sequence, any of the
following results are possible:
(Q == &A) and (D == 1)
(Q == &B) and (D == 2)
(Q == &B) and (D == 4)
Note that CPU 2 will never try and load C into D because the CPU will load P
into Q before issuing the load of *Q.
DEVICE OPERATIONS
-----------------
Some devices present their control interfaces as collections of memory
locations, but the order in which the control registers are accessed is very
important. For instance, imagine an ethernet card with a set of internal
registers that are accessed through an address port register (A) and a data
port register (D). To read internal register 5, the following code might then
be used:
*A = 5;
x = *D;
but this might show up as either of the following two sequences:
STORE *A = 5, x = LOAD *D
x = LOAD *D, STORE *A = 5
the second of which will almost certainly result in a malfunction, since it set
the address _after_ attempting to read the register.
GUARANTEES
----------
There are some minimal guarantees that may be expected of a CPU:
(*) On any given CPU, dependent memory accesses will be issued in order, with
respect to itself. This means that for:
ACCESS_ONCE(Q) = P; smp_read_barrier_depends(); D = ACCESS_ONCE(*Q);
the CPU will issue the following memory operations:
Q = LOAD P, D = LOAD *Q
and always in that order. On most systems, smp_read_barrier_depends()
does nothing, but it is required for DEC Alpha. The ACCESS_ONCE()
is required to prevent compiler mischief. Please note that you
should normally use something like rcu_dereference() instead of
open-coding smp_read_barrier_depends().
(*) Overlapping loads and stores within a particular CPU will appear to be
ordered within that CPU. This means that for:
a = ACCESS_ONCE(*X); ACCESS_ONCE(*X) = b;
the CPU will only issue the following sequence of memory operations:
a = LOAD *X, STORE *X = b
And for:
ACCESS_ONCE(*X) = c; d = ACCESS_ONCE(*X);
the CPU will only issue:
STORE *X = c, d = LOAD *X
(Loads and stores overlap if they are targeted at overlapping pieces of
memory).
And there are a number of things that _must_ or _must_not_ be assumed:
(*) It _must_not_ be assumed that the compiler will do what you want with
memory references that are not protected by ACCESS_ONCE(). Without
ACCESS_ONCE(), the compiler is within its rights to do all sorts
of "creative" transformations, which are covered in the Compiler
Barrier section.
(*) It _must_not_ be assumed that independent loads and stores will be issued
in the order given. This means that for:
X = *A; Y = *B; *D = Z;
we may get any of the following sequences:
X = LOAD *A, Y = LOAD *B, STORE *D = Z
X = LOAD *A, STORE *D = Z, Y = LOAD *B
Y = LOAD *B, X = LOAD *A, STORE *D = Z
Y = LOAD *B, STORE *D = Z, X = LOAD *A
STORE *D = Z, X = LOAD *A, Y = LOAD *B
STORE *D = Z, Y = LOAD *B, X = LOAD *A
(*) It _must_ be assumed that overlapping memory accesses may be merged or
discarded. This means that for:
X = *A; Y = *(A + 4);
we may get any one of the following sequences:
X = LOAD *A; Y = LOAD *(A + 4);
Y = LOAD *(A + 4); X = LOAD *A;
{X, Y} = LOAD {*A, *(A + 4) };
And for:
*A = X; *(A + 4) = Y;
we may get any of:
STORE *A = X; STORE *(A + 4) = Y;
STORE *(A + 4) = Y; STORE *A = X;
STORE {*A, *(A + 4) } = {X, Y};
And there are anti-guarantees:
(*) These guarantees do not apply to bitfields, because compilers often
generate code to modify these using non-atomic read-modify-write
sequences. Do not attempt to use bitfields to synchronize parallel
algorithms.
(*) Even in cases where bitfields are protected by locks, all fields
in a given bitfield must be protected by one lock. If two fields
in a given bitfield are protected by different locks, the compiler's
non-atomic read-modify-write sequences can cause an update to one
field to corrupt the value of an adjacent field.
(*) These guarantees apply only to properly aligned and sized scalar
variables. "Properly sized" currently means variables that are
the same size as "char", "short", "int" and "long". "Properly
aligned" means the natural alignment, thus no constraints for
"char", two-byte alignment for "short", four-byte alignment for
"int", and either four-byte or eight-byte alignment for "long",
on 32-bit and 64-bit systems, respectively. Note that these
guarantees were introduced into the C11 standard, so beware when
using older pre-C11 compilers (for example, gcc 4.6). The portion
of the standard containing this guarantee is Section 3.14, which
defines "memory location" as follows:
memory location
either an object of scalar type, or a maximal sequence
of adjacent bit-fields all having nonzero width
NOTE 1: Two threads of execution can update and access
separate memory locations without interfering with
each other.
NOTE 2: A bit-field and an adjacent non-bit-field member
are in separate memory locations. The same applies
to two bit-fields, if one is declared inside a nested
structure declaration and the other is not, or if the two
are separated by a zero-length bit-field declaration,
or if they are separated by a non-bit-field member
declaration. It is not safe to concurrently update two
bit-fields in the same structure if all members declared
between them are also bit-fields, no matter what the
sizes of those intervening bit-fields happen to be.
=========================
WHAT ARE MEMORY BARRIERS?
=========================
As can be seen above, independent memory operations are effectively performed
in random order, but this can be a problem for CPU-CPU interaction and for I/O.
What is required is some way of intervening to instruct the compiler and the
CPU to restrict the order.
Memory barriers are such interventions. They impose a perceived partial
ordering over the memory operations on either side of the barrier.
Such enforcement is important because the CPUs and other devices in a system
can use a variety of tricks to improve performance, including reordering,
deferral and combination of memory operations; speculative loads; speculative
branch prediction and various types of caching. Memory barriers are used to
override or suppress these tricks, allowing the code to sanely control the
interaction of multiple CPUs and/or devices.
VARIETIES OF MEMORY BARRIER
---------------------------
Memory barriers come in four basic varieties:
(1) Write (or store) memory barriers.
A write memory barrier gives a guarantee that all the STORE operations
specified before the barrier will appear to happen before all the STORE
operations specified after the barrier with respect to the other
components of the system.
A write barrier is a partial ordering on stores only; it is not required
to have any effect on loads.
A CPU can be viewed as committing a sequence of store operations to the
memory system as time progresses. All stores before a write barrier will
occur in the sequence _before_ all the stores after the write barrier.
[!] Note that write barriers should normally be paired with read or data
dependency barriers; see the "SMP barrier pairing" subsection.
(2) Data dependency barriers.
A data dependency barrier is a weaker form of read barrier. In the case
where two loads are performed such that the second depends on the result
of the first (eg: the first load retrieves the address to which the second
load will be directed), a data dependency barrier would be required to
make sure that the target of the second load is updated before the address
obtained by the first load is accessed.
A data dependency barrier is a partial ordering on interdependent loads
only; it is not required to have any effect on stores, independent loads
or overlapping loads.
As mentioned in (1), the other CPUs in the system can be viewed as
committing sequences of stores to the memory system that the CPU being
considered can then perceive. A data dependency barrier issued by the CPU
under consideration guarantees that for any load preceding it, if that
load touches one of a sequence of stores from another CPU, then by the
time the barrier completes, the effects of all the stores prior to that
touched by the load will be perceptible to any loads issued after the data
dependency barrier.
See the "Examples of memory barrier sequences" subsection for diagrams
showing the ordering constraints.
[!] Note that the first load really has to have a _data_ dependency and
not a control dependency. If the address for the second load is dependent
on the first load, but the dependency is through a conditional rather than
actually loading the address itself, then it's a _control_ dependency and
a full read barrier or better is required. See the "Control dependencies"
subsection for more information.
[!] Note that data dependency barriers should normally be paired with
write barriers; see the "SMP barrier pairing" subsection.
(3) Read (or load) memory barriers.
A read barrier is a data dependency barrier plus a guarantee that all the
LOAD operations specified before the barrier will appear to happen before
all the LOAD operations specified after the barrier with respect to the
other components of the system.
A read barrier is a partial ordering on loads only; it is not required to
have any effect on stores.
Read memory barriers imply data dependency barriers, and so can substitute
for them.
[!] Note that read barriers should normally be paired with write barriers;
see the "SMP barrier pairing" subsection.
(4) General memory barriers.
A general memory barrier gives a guarantee that all the LOAD and STORE
operations specified before the barrier will appear to happen before all
the LOAD and STORE operations specified after the barrier with respect to
the other components of the system.
A general memory barrier is a partial ordering over both loads and stores.
General memory barriers imply both read and write memory barriers, and so
can substitute for either.
And a couple of implicit varieties:
(5) ACQUIRE operations.
This acts as a one-way permeable barrier. It guarantees that all memory
operations after the ACQUIRE operation will appear to happen after the
ACQUIRE operation with respect to the other components of the system.
ACQUIRE operations include LOCK operations and smp_load_acquire()
operations.
Memory operations that occur before an ACQUIRE operation may appear to
happen after it completes.
An ACQUIRE operation should almost always be paired with a RELEASE
operation.
(6) RELEASE operations.
This also acts as a one-way permeable barrier. It guarantees that all
memory operations before the RELEASE operation will appear to happen
before the RELEASE operation with respect to the other components of the
system. RELEASE operations include UNLOCK operations and
smp_store_release() operations.
Memory operations that occur after a RELEASE operation may appear to
happen before it completes.
The use of ACQUIRE and RELEASE operations generally precludes the need
for other sorts of memory barrier (but note the exceptions mentioned in
the subsection "MMIO write barrier"). In addition, a RELEASE+ACQUIRE
pair is -not- guaranteed to act as a full memory barrier. However, after
an ACQUIRE on a given variable, all memory accesses preceding any prior
RELEASE on that same variable are guaranteed to be visible. In other
words, within a given variable's critical section, all accesses of all
previous critical sections for that variable are guaranteed to have
completed.
This means that ACQUIRE acts as a minimal "acquire" operation and
RELEASE acts as a minimal "release" operation.
Memory barriers are only required where there's a possibility of interaction
between two CPUs or between a CPU and a device. If it can be guaranteed that
there won't be any such interaction in any particular piece of code, then
memory barriers are unnecessary in that piece of code.
Note that these are the _minimum_ guarantees. Different architectures may give
more substantial guarantees, but they may _not_ be relied upon outside of arch
specific code.
WHAT MAY NOT BE ASSUMED ABOUT MEMORY BARRIERS?
----------------------------------------------
There are certain things that the Linux kernel memory barriers do not guarantee:
(*) There is no guarantee that any of the memory accesses specified before a
memory barrier will be _complete_ by the completion of a memory barrier
instruction; the barrier can be considered to draw a line in that CPU's
access queue that accesses of the appropriate type may not cross.
(*) There is no guarantee that issuing a memory barrier on one CPU will have
any direct effect on another CPU or any other hardware in the system. The
indirect effect will be the order in which the second CPU sees the effects
of the first CPU's accesses occur, but see the next point:
(*) There is no guarantee that a CPU will see the correct order of effects
from a second CPU's accesses, even _if_ the second CPU uses a memory
barrier, unless the first CPU _also_ uses a matching memory barrier (see
the subsection on "SMP Barrier Pairing").
(*) There is no guarantee that some intervening piece of off-the-CPU
hardware[*] will not reorder the memory accesses. CPU cache coherency
mechanisms should propagate the indirect effects of a memory barrier
between CPUs, but might not do so in order.
[*] For information on bus mastering DMA and coherency please read:
Documentation/PCI/pci.txt
Documentation/DMA-API-HOWTO.txt
Documentation/DMA-API.txt
DATA DEPENDENCY BARRIERS
------------------------
The usage requirements of data dependency barriers are a little subtle, and
it's not always obvious that they're needed. To illustrate, consider the
following sequence of events:
CPU 1 CPU 2
=============== ===============
{ A == 1, B == 2, C = 3, P == &A, Q == &C }
B = 4;
<write barrier>
ACCESS_ONCE(P) = &B
Q = ACCESS_ONCE(P);
D = *Q;
There's a clear data dependency here, and it would seem that by the end of the
sequence, Q must be either &A or &B, and that:
(Q == &A) implies (D == 1)
(Q == &B) implies (D == 4)
But! CPU 2's perception of P may be updated _before_ its perception of B, thus
leading to the following situation:
(Q == &B) and (D == 2) ????
Whilst this may seem like a failure of coherency or causality maintenance, it
isn't, and this behaviour can be observed on certain real CPUs (such as the DEC
Alpha).
To deal with this, a data dependency barrier or better must be inserted
between the address load and the data load:
CPU 1 CPU 2
=============== ===============
{ A == 1, B == 2, C = 3, P == &A, Q == &C }
B = 4;
<write barrier>
ACCESS_ONCE(P) = &B
Q = ACCESS_ONCE(P);
<data dependency barrier>
D = *Q;
This enforces the occurrence of one of the two implications, and prevents the
third possibility from arising.
[!] Note that this extremely counterintuitive situation arises most easily on
machines with split caches, so that, for example, one cache bank processes
even-numbered cache lines and the other bank processes odd-numbered cache
lines. The pointer P might be stored in an odd-numbered cache line, and the
variable B might be stored in an even-numbered cache line. Then, if the
even-numbered bank of the reading CPU's cache is extremely busy while the
odd-numbered bank is idle, one can see the new value of the pointer P (&B),
but the old value of the variable B (2).
Another example of where data dependency barriers might be required is where a
number is read from memory and then used to calculate the index for an array
access:
CPU 1 CPU 2
=============== ===============
{ M[0] == 1, M[1] == 2, M[3] = 3, P == 0, Q == 3 }
M[1] = 4;
<write barrier>
ACCESS_ONCE(P) = 1
Q = ACCESS_ONCE(P);
<data dependency barrier>
D = M[Q];
The data dependency barrier is very important to the RCU system,
for example. See rcu_assign_pointer() and rcu_dereference() in
include/linux/rcupdate.h. This permits the current target of an RCU'd
pointer to be replaced with a new modified target, without the replacement
target appearing to be incompletely initialised.
See also the subsection on "Cache Coherency" for a more thorough example.
CONTROL DEPENDENCIES
--------------------
A load-load control dependency requires a full read memory barrier, not
simply a data dependency barrier to make it work correctly. Consider the
following bit of code:
q = ACCESS_ONCE(a);
if (q) {
<data dependency barrier> /* BUG: No data dependency!!! */
p = ACCESS_ONCE(b);
}
This will not have the desired effect because there is no actual data
dependency, but rather a control dependency that the CPU may short-circuit
by attempting to predict the outcome in advance, so that other CPUs see
the load from b as having happened before the load from a. In such a
case what's actually required is:
q = ACCESS_ONCE(a);
if (q) {
<read barrier>
p = ACCESS_ONCE(b);
}
However, stores are not speculated. This means that ordering -is- provided
for load-store control dependencies, as in the following example:
q = READ_ONCE_CTRL(a);
if (q) {
ACCESS_ONCE(b) = p;
}
Control dependencies pair normally with other types of barriers. That
said, please note that READ_ONCE_CTRL() is not optional! Without the
READ_ONCE_CTRL(), the compiler might combine the load from 'a' with
other loads from 'a', and the store to 'b' with other stores to 'b',
with possible highly counterintuitive effects on ordering.
Worse yet, if the compiler is able to prove (say) that the value of
variable 'a' is always non-zero, it would be well within its rights
to optimize the original example by eliminating the "if" statement
as follows:
q = a;
b = p; /* BUG: Compiler and CPU can both reorder!!! */
Finally, the READ_ONCE_CTRL() includes an smp_read_barrier_depends()
that DEC Alpha needs in order to respect control depedencies.
So don't leave out the READ_ONCE_CTRL().
It is tempting to try to enforce ordering on identical stores on both
branches of the "if" statement as follows:
q = READ_ONCE_CTRL(a);
if (q) {
barrier();
ACCESS_ONCE(b) = p;
do_something();
} else {
barrier();
ACCESS_ONCE(b) = p;
do_something_else();
}
Unfortunately, current compilers will transform this as follows at high
optimization levels:
q = READ_ONCE_CTRL(a);
barrier();
ACCESS_ONCE(b) = p; /* BUG: No ordering vs. load from a!!! */
if (q) {
/* ACCESS_ONCE(b) = p; -- moved up, BUG!!! */
do_something();
} else {
/* ACCESS_ONCE(b) = p; -- moved up, BUG!!! */
do_something_else();
}
Now there is no conditional between the load from 'a' and the store to
'b', which means that the CPU is within its rights to reorder them:
The conditional is absolutely required, and must be present in the
assembly code even after all compiler optimizations have been applied.
Therefore, if you need ordering in this example, you need explicit
memory barriers, for example, smp_store_release():
q = ACCESS_ONCE(a);
if (q) {
smp_store_release(&b, p);
do_something();
} else {
smp_store_release(&b, p);
do_something_else();
}
In contrast, without explicit memory barriers, two-legged-if control
ordering is guaranteed only when the stores differ, for example:
q = READ_ONCE_CTRL(a);
if (q) {
ACCESS_ONCE(b) = p;
do_something();
} else {
ACCESS_ONCE(b) = r;
do_something_else();
}
The initial READ_ONCE_CTRL() is still required to prevent the compiler
from proving the value of 'a'.
In addition, you need to be careful what you do with the local variable 'q',
otherwise the compiler might be able to guess the value and again remove
the needed conditional. For example:
q = READ_ONCE_CTRL(a);
if (q % MAX) {
ACCESS_ONCE(b) = p;
do_something();
} else {
ACCESS_ONCE(b) = r;
do_something_else();
}
If MAX is defined to be 1, then the compiler knows that (q % MAX) is
equal to zero, in which case the compiler is within its rights to
transform the above code into the following:
q = READ_ONCE_CTRL(a);
ACCESS_ONCE(b) = p;
do_something_else();
Given this transformation, the CPU is not required to respect the ordering
between the load from variable 'a' and the store to variable 'b'. It is
tempting to add a barrier(), but this does not help. The conditional
is gone, and the barrier won't bring it back. Therefore, if you are
relying on this ordering, you should make sure that MAX is greater than
one, perhaps as follows:
q = READ_ONCE_CTRL(a);
BUILD_BUG_ON(MAX <= 1); /* Order load from a with store to b. */
if (q % MAX) {
ACCESS_ONCE(b) = p;
do_something();
} else {
ACCESS_ONCE(b) = r;
do_something_else();
}
Please note once again that the stores to 'b' differ. If they were
identical, as noted earlier, the compiler could pull this store outside
of the 'if' statement.
You must also be careful not to rely too much on boolean short-circuit
evaluation. Consider this example:
q = READ_ONCE_CTRL(a);
if (a || 1 > 0)
ACCESS_ONCE(b) = 1;
Because the first condition cannot fault and the second condition is
always true, the compiler can transform this example as following,
defeating control dependency:
q = READ_ONCE_CTRL(a);
ACCESS_ONCE(b) = 1;
This example underscores the need to ensure that the compiler cannot
out-guess your code. More generally, although ACCESS_ONCE() does force
the compiler to actually emit code for a given load, it does not force
the compiler to use the results.
Finally, control dependencies do -not- provide transitivity. This is
demonstrated by two related examples, with the initial values of
x and y both being zero:
CPU 0 CPU 1
======================= =======================
r1 = READ_ONCE_CTRL(x); r2 = READ_ONCE_CTRL(y);
if (r1 > 0) if (r2 > 0)
ACCESS_ONCE(y) = 1; ACCESS_ONCE(x) = 1;
assert(!(r1 == 1 && r2 == 1));
The above two-CPU example will never trigger the assert(). However,
if control dependencies guaranteed transitivity (which they do not),
then adding the following CPU would guarantee a related assertion:
CPU 2
=====================
ACCESS_ONCE(x) = 2;
assert(!(r1 == 2 && r2 == 1 && x == 2)); /* FAILS!!! */
But because control dependencies do -not- provide transitivity, the above
assertion can fail after the combined three-CPU example completes. If you
need the three-CPU example to provide ordering, you will need smp_mb()
between the loads and stores in the CPU 0 and CPU 1 code fragments,
that is, just before or just after the "if" statements. Furthermore,
the original two-CPU example is very fragile and should be avoided.
These two examples are the LB and WWC litmus tests from this paper:
http://www.cl.cam.ac.uk/users/pes20/ppc-supplemental/test6.pdf and this
site: https://www.cl.cam.ac.uk/~pes20/ppcmem/index.html.
In summary:
(*) Control dependencies must be headed by READ_ONCE_CTRL().
Or, as a much less preferable alternative, interpose
be headed by READ_ONCE() or an ACCESS_ONCE() read and must
have smp_read_barrier_depends() between this read and the
control-dependent write.
(*) Control dependencies can order prior loads against later stores.
However, they do -not- guarantee any other sort of ordering:
Not prior loads against later loads, nor prior stores against
later anything. If you need these other forms of ordering,
use smp_rmb(), smp_wmb(), or, in the case of prior stores and
later loads, smp_mb().
(*) If both legs of the "if" statement begin with identical stores
to the same variable, a barrier() statement is required at the
beginning of each leg of the "if" statement.
(*) Control dependencies require at least one run-time conditional
between the prior load and the subsequent store, and this
conditional must involve the prior load. If the compiler
is able to optimize the conditional away, it will have also
optimized away the ordering. Careful use of ACCESS_ONCE() can
help to preserve the needed conditional.
(*) Control dependencies require that the compiler avoid reordering the
dependency into nonexistence. Careful use of ACCESS_ONCE() or
barrier() can help to preserve your control dependency. Please
see the Compiler Barrier section for more information.
(*) Control dependencies pair normally with other types of barriers.
(*) Control dependencies do -not- provide transitivity. If you
need transitivity, use smp_mb().
SMP BARRIER PAIRING
-------------------
When dealing with CPU-CPU interactions, certain types of memory barrier should
always be paired. A lack of appropriate pairing is almost certainly an error.
General barriers pair with each other, though they also pair with most
other types of barriers, albeit without transitivity. An acquire barrier
pairs with a release barrier, but both may also pair with other barriers,
including of course general barriers. A write barrier pairs with a data
dependency barrier, a control dependency, an acquire barrier, a release
barrier, a read barrier, or a general barrier. Similarly a read barrier,
control dependency, or a data dependency barrier pairs with a write
barrier, an acquire barrier, a release barrier, or a general barrier:
CPU 1 CPU 2
=============== ===============
ACCESS_ONCE(a) = 1;
<write barrier>
ACCESS_ONCE(b) = 2; x = ACCESS_ONCE(b);
<read barrier>
y = ACCESS_ONCE(a);
Or:
CPU 1 CPU 2
=============== ===============================
a = 1;
<write barrier>
ACCESS_ONCE(b) = &a; x = ACCESS_ONCE(b);
<data dependency barrier>
y = *x;
Or even:
CPU 1 CPU 2
=============== ===============================
r1 = ACCESS_ONCE(y);
<general barrier>
ACCESS_ONCE(y) = 1; if (r2 = ACCESS_ONCE(x)) {
<implicit control dependency>
ACCESS_ONCE(y) = 1;
}
assert(r1 == 0 || r2 == 0);
Basically, the read barrier always has to be there, even though it can be of
the "weaker" type.
[!] Note that the stores before the write barrier would normally be expected to
match the loads after the read barrier or the data dependency barrier, and vice
versa:
CPU 1 CPU 2
=================== ===================
ACCESS_ONCE(a) = 1; }---- --->{ v = ACCESS_ONCE(c);
ACCESS_ONCE(b) = 2; } \ / { w = ACCESS_ONCE(d);
<write barrier> \ <read barrier>
ACCESS_ONCE(c) = 3; } / \ { x = ACCESS_ONCE(a);
ACCESS_ONCE(d) = 4; }---- --->{ y = ACCESS_ONCE(b);
EXAMPLES OF MEMORY BARRIER SEQUENCES
------------------------------------
Firstly, write barriers act as partial orderings on store operations.
Consider the following sequence of events:
CPU 1
=======================
STORE A = 1
STORE B = 2
STORE C = 3
<write barrier>
STORE D = 4
STORE E = 5
This sequence of events is committed to the memory coherence system in an order
that the rest of the system might perceive as the unordered set of { STORE A,
STORE B, STORE C } all occurring before the unordered set of { STORE D, STORE E
}:
+-------+ : :
| | +------+
| |------>| C=3 | } /\
| | : +------+ }----- \ -----> Events perceptible to
| | : | A=1 | } \/ the rest of the system
| | : +------+ }
| CPU 1 | : | B=2 | }
| | +------+ }
| | wwwwwwwwwwwwwwww } <--- At this point the write barrier
| | +------+ } requires all stores prior to the
| | : | E=5 | } barrier to be committed before
| | : +------+ } further stores may take place
| |------>| D=4 | }
| | +------+
+-------+ : :
|
| Sequence in which stores are committed to the
| memory system by CPU 1
V
Secondly, data dependency barriers act as partial orderings on data-dependent
loads. Consider the following sequence of events:
CPU 1 CPU 2
======================= =======================
{ B = 7; X = 9; Y = 8; C = &Y }
STORE A = 1
STORE B = 2
<write barrier>
STORE C = &B LOAD X
STORE D = 4 LOAD C (gets &B)
LOAD *C (reads B)
Without intervention, CPU 2 may perceive the events on CPU 1 in some
effectively random order, despite the write barrier issued by CPU 1:
+-------+ : : : :
| | +------+ +-------+ | Sequence of update
| |------>| B=2 |----- --->| Y->8 | | of perception on
| | : +------+ \ +-------+ | CPU 2
| CPU 1 | : | A=1 | \ --->| C->&Y | V
| | +------+ | +-------+
| | wwwwwwwwwwwwwwww | : :
| | +------+ | : :
| | : | C=&B |--- | : : +-------+
| | : +------+ \ | +-------+ | |
| |------>| D=4 | ----------->| C->&B |------>| |
| | +------+ | +-------+ | |
+-------+ : : | : : | |
| : : | |
| : : | CPU 2 |
| +-------+ | |
Apparently incorrect ---> | | B->7 |------>| |
perception of B (!) | +-------+ | |
| : : | |
| +-------+ | |
The load of X holds ---> \ | X->9 |------>| |
up the maintenance \ +-------+ | |
of coherence of B ----->| B->2 | +-------+
+-------+
: :
In the above example, CPU 2 perceives that B is 7, despite the load of *C
(which would be B) coming after the LOAD of C.
If, however, a data dependency barrier were to be placed between the load of C
and the load of *C (ie: B) on CPU 2:
CPU 1 CPU 2
======================= =======================
{ B = 7; X = 9; Y = 8; C = &Y }
STORE A = 1
STORE B = 2
<write barrier>
STORE C = &B LOAD X
STORE D = 4 LOAD C (gets &B)
<data dependency barrier>
LOAD *C (reads B)
then the following will occur:
+-------+ : : : :
| | +------+ +-------+
| |------>| B=2 |----- --->| Y->8 |
| | : +------+ \ +-------+
| CPU 1 | : | A=1 | \ --->| C->&Y |
| | +------+ | +-------+
| | wwwwwwwwwwwwwwww | : :
| | +------+ | : :
| | : | C=&B |--- | : : +-------+
| | : +------+ \ | +-------+ | |
| |------>| D=4 | ----------->| C->&B |------>| |
| | +------+ | +-------+ | |
+-------+ : : | : : | |
| : : | |
| : : | CPU 2 |
| +-------+ | |
| | X->9 |------>| |
| +-------+ | |
Makes sure all effects ---> \ ddddddddddddddddd | |
prior to the store of C \ +-------+ | |
are perceptible to ----->| B->2 |------>| |
subsequent loads +-------+ | |
: : +-------+
And thirdly, a read barrier acts as a partial order on loads. Consider the
following sequence of events:
CPU 1 CPU 2
======================= =======================
{ A = 0, B = 9 }
STORE A=1
<write barrier>
STORE B=2
LOAD B
LOAD A
Without intervention, CPU 2 may then choose to perceive the events on CPU 1 in
some effectively random order, despite the write barrier issued by CPU 1:
+-------+ : : : :
| | +------+ +-------+
| |------>| A=1 |------ --->| A->0 |
| | +------+ \ +-------+
| CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
| | +------+ | +-------+
| |------>| B=2 |--- | : :
| | +------+ \ | : : +-------+
+-------+ : : \ | +-------+ | |
---------->| B->2 |------>| |
| +-------+ | CPU 2 |
| | A->0 |------>| |
| +-------+ | |
| : : +-------+
\ : :
\ +-------+
---->| A->1 |
+-------+
: :
If, however, a read barrier were to be placed between the load of B and the
load of A on CPU 2:
CPU 1 CPU 2
======================= =======================
{ A = 0, B = 9 }
STORE A=1
<write barrier>
STORE B=2
LOAD B
<read barrier>
LOAD A
then the partial ordering imposed by CPU 1 will be perceived correctly by CPU
2:
+-------+ : : : :
| | +------+ +-------+
| |------>| A=1 |------ --->| A->0 |
| | +------+ \ +-------+
| CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
| | +------+ | +-------+
| |------>| B=2 |--- | : :
| | +------+ \ | : : +-------+
+-------+ : : \ | +-------+ | |
---------->| B->2 |------>| |
| +-------+ | CPU 2 |
| : : | |
| : : | |
At this point the read ----> \ rrrrrrrrrrrrrrrrr | |
barrier causes all effects \ +-------+ | |
prior to the storage of B ---->| A->1 |------>| |
to be perceptible to CPU 2 +-------+ | |
: : +-------+
To illustrate this more completely, consider what could happen if the code
contained a load of A either side of the read barrier:
CPU 1 CPU 2
======================= =======================
{ A = 0, B = 9 }
STORE A=1
<write barrier>
STORE B=2
LOAD B
LOAD A [first load of A]
<read barrier>
LOAD A [second load of A]
Even though the two loads of A both occur after the load of B, they may both
come up with different values:
+-------+ : : : :
| | +------+ +-------+
| |------>| A=1 |------ --->| A->0 |
| | +------+ \ +-------+
| CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
| | +------+ | +-------+
| |------>| B=2 |--- | : :
| | +------+ \ | : : +-------+
+-------+ : : \ | +-------+ | |
---------->| B->2 |------>| |
| +-------+ | CPU 2 |
| : : | |
| : : | |
| +-------+ | |
| | A->0 |------>| 1st |
| +-------+ | |
At this point the read ----> \ rrrrrrrrrrrrrrrrr | |
barrier causes all effects \ +-------+ | |
prior to the storage of B ---->| A->1 |------>| 2nd |
to be perceptible to CPU 2 +-------+ | |
: : +-------+
But it may be that the update to A from CPU 1 becomes perceptible to CPU 2
before the read barrier completes anyway:
+-------+ : : : :
| | +------+ +-------+
| |------>| A=1 |------ --->| A->0 |
| | +------+ \ +-------+
| CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
| | +------+ | +-------+
| |------>| B=2 |--- | : :
| | +------+ \ | : : +-------+
+-------+ : : \ | +-------+ | |
---------->| B->2 |------>| |
| +-------+ | CPU 2 |
| : : | |
\ : : | |
\ +-------+ | |
---->| A->1 |------>| 1st |
+-------+ | |
rrrrrrrrrrrrrrrrr | |
+-------+ | |
| A->1 |------>| 2nd |
+-------+ | |
: : +-------+
The guarantee is that the second load will always come up with A == 1 if the
load of B came up with B == 2. No such guarantee exists for the first load of
A; that may come up with either A == 0 or A == 1.
READ MEMORY BARRIERS VS LOAD SPECULATION
----------------------------------------
Many CPUs speculate with loads: that is they see that they will need to load an
item from memory, and they find a time where they're not using the bus for any
other loads, and so do the load in advance - even though they haven't actually
got to that point in the instruction execution flow yet. This permits the
actual load instruction to potentially complete immediately because the CPU
already has the value to hand.
It may turn out that the CPU didn't actually need the value - perhaps because a
branch circumvented the load - in which case it can discard the value or just
cache it for later use.
Consider:
CPU 1 CPU 2
======================= =======================
LOAD B
DIVIDE } Divide instructions generally
DIVIDE } take a long time to perform
LOAD A
Which might appear as this:
: : +-------+
+-------+ | |
--->| B->2 |------>| |
+-------+ | CPU 2 |
: :DIVIDE | |
+-------+ | |
The CPU being busy doing a ---> --->| A->0 |~~~~ | |
division speculates on the +-------+ ~ | |
LOAD of A : : ~ | |
: :DIVIDE | |
: : ~ | |
Once the divisions are complete --> : : ~-->| |
the CPU can then perform the : : | |
LOAD with immediate effect : : +-------+
Placing a read barrier or a data dependency barrier just before the second
load:
CPU 1 CPU 2
======================= =======================
LOAD B
DIVIDE
DIVIDE
<read barrier>
LOAD A
will force any value speculatively obtained to be reconsidered to an extent
dependent on the type of barrier used. If there was no change made to the
speculated memory location, then the speculated value will just be used:
: : +-------+
+-------+ | |
--->| B->2 |------>| |
+-------+ | CPU 2 |
: :DIVIDE | |
+-------+ | |
The CPU being busy doing a ---> --->| A->0 |~~~~ | |
division speculates on the +-------+ ~ | |
LOAD of A : : ~ | |
: :DIVIDE | |
: : ~ | |
: : ~ | |
rrrrrrrrrrrrrrrr~ | |
: : ~ | |
: : ~-->| |
: : | |
: : +-------+
but if there was an update or an invalidation from another CPU pending, then
the speculation will be cancelled and the value reloaded:
: : +-------+
+-------+ | |
--->| B->2 |------>| |
+-------+ | CPU 2 |
: :DIVIDE | |
+-------+ | |
The CPU being busy doing a ---> --->| A->0 |~~~~ | |
division speculates on the +-------+ ~ | |
LOAD of A : : ~ | |
: :DIVIDE | |
: : ~ | |
: : ~ | |
rrrrrrrrrrrrrrrrr | |
+-------+ | |
The speculation is discarded ---> --->| A->1 |------>| |
and an updated value is +-------+ | |
retrieved : : +-------+
TRANSITIVITY
------------
Transitivity is a deeply intuitive notion about ordering that is not
always provided by real computer systems. The following example
demonstrates transitivity (also called "cumulativity"):
CPU 1 CPU 2 CPU 3
======================= ======================= =======================
{ X = 0, Y = 0 }
STORE X=1 LOAD X STORE Y=1
<general barrier> <general barrier>
LOAD Y LOAD X
Suppose that CPU 2's load from X returns 1 and its load from Y returns 0.
This indicates that CPU 2's load from X in some sense follows CPU 1's
store to X and that CPU 2's load from Y in some sense preceded CPU 3's
store to Y. The question is then "Can CPU 3's load from X return 0?"
Because CPU 2's load from X in some sense came after CPU 1's store, it
is natural to expect that CPU 3's load from X must therefore return 1.
This expectation is an example of transitivity: if a load executing on
CPU A follows a load from the same variable executing on CPU B, then
CPU A's load must either return the same value that CPU B's load did,
or must return some later value.
In the Linux kernel, use of general memory barriers guarantees
transitivity. Therefore, in the above example, if CPU 2's load from X
returns 1 and its load from Y returns 0, then CPU 3's load from X must
also return 1.
However, transitivity is -not- guaranteed for read or write barriers.
For example, suppose that CPU 2's general barrier in the above example
is changed to a read barrier as shown below:
CPU 1 CPU 2 CPU 3
======================= ======================= =======================
{ X = 0, Y = 0 }
STORE X=1 LOAD X STORE Y=1
<read barrier> <general barrier>
LOAD Y LOAD X
This substitution destroys transitivity: in this example, it is perfectly
legal for CPU 2's load from X to return 1, its load from Y to return 0,
and CPU 3's load from X to return 0.
The key point is that although CPU 2's read barrier orders its pair
of loads, it does not guarantee to order CPU 1's store. Therefore, if
this example runs on a system where CPUs 1 and 2 share a store buffer
or a level of cache, CPU 2 might have early access to CPU 1's writes.
General barriers are therefore required to ensure that all CPUs agree
on the combined order of CPU 1's and CPU 2's accesses.
To reiterate, if your code requires transitivity, use general barriers
throughout.
========================
EXPLICIT KERNEL BARRIERS
========================
The Linux kernel has a variety of different barriers that act at different
levels:
(*) Compiler barrier.
(*) CPU memory barriers.
(*) MMIO write barrier.
COMPILER BARRIER
----------------
The Linux kernel has an explicit compiler barrier function that prevents the
compiler from moving the memory accesses either side of it to the other side:
barrier();
This is a general barrier -- there are no read-read or write-write variants
of barrier(). However, ACCESS_ONCE() can be thought of as a weak form
for barrier() that affects only the specific accesses flagged by the
ACCESS_ONCE().
The barrier() function has the following effects:
(*) Prevents the compiler from reordering accesses following the
barrier() to precede any accesses preceding the barrier().
One example use for this property is to ease communication between
interrupt-handler code and the code that was interrupted.
(*) Within a loop, forces the compiler to load the variables used
in that loop's conditional on each pass through that loop.
The ACCESS_ONCE() function can prevent any number of optimizations that,
while perfectly safe in single-threaded code, can be fatal in concurrent
code. Here are some examples of these sorts of optimizations:
(*) The compiler is within its rights to reorder loads and stores
to the same variable, and in some cases, the CPU is within its
rights to reorder loads to the same variable. This means that
the following code:
a[0] = x;
a[1] = x;
Might result in an older value of x stored in a[1] than in a[0].
Prevent both the compiler and the CPU from doing this as follows:
a[0] = ACCESS_ONCE(x);
a[1] = ACCESS_ONCE(x);
In short, ACCESS_ONCE() provides cache coherence for accesses from
multiple CPUs to a single variable.
(*) The compiler is within its rights to merge successive loads from
the same variable. Such merging can cause the compiler to "optimize"
the following code:
while (tmp = a)
do_something_with(tmp);
into the following code, which, although in some sense legitimate
for single-threaded code, is almost certainly not what the developer
intended:
if (tmp = a)
for (;;)
do_something_with(tmp);
Use ACCESS_ONCE() to prevent the compiler from doing this to you:
while (tmp = ACCESS_ONCE(a))
do_something_with(tmp);
(*) The compiler is within its rights to reload a variable, for example,
in cases where high register pressure prevents the compiler from
keeping all data of interest in registers. The compiler might
therefore optimize the variable 'tmp' out of our previous example:
while (tmp = a)
do_something_with(tmp);
This could result in the following code, which is perfectly safe in
single-threaded code, but can be fatal in concurrent code:
while (a)
do_something_with(a);
For example, the optimized version of this code could result in
passing a zero to do_something_with() in the case where the variable
a was modified by some other CPU between the "while" statement and
the call to do_something_with().
Again, use ACCESS_ONCE() to prevent the compiler from doing this:
while (tmp = ACCESS_ONCE(a))
do_something_with(tmp);
Note that if the compiler runs short of registers, it might save
tmp onto the stack. The overhead of this saving and later restoring
is why compilers reload variables. Doing so is perfectly safe for
single-threaded code, so you need to tell the compiler about cases
where it is not safe.
(*) The compiler is within its rights to omit a load entirely if it knows
what the value will be. For example, if the compiler can prove that
the value of variable 'a' is always zero, it can optimize this code:
while (tmp = a)
do_something_with(tmp);
Into this:
do { } while (0);
This transformation is a win for single-threaded code because it gets
rid of a load and a branch. The problem is that the compiler will
carry out its proof assuming that the current CPU is the only one
updating variable 'a'. If variable 'a' is shared, then the compiler's
proof will be erroneous. Use ACCESS_ONCE() to tell the compiler
that it doesn't know as much as it thinks it does:
while (tmp = ACCESS_ONCE(a))
do_something_with(tmp);
But please note that the compiler is also closely watching what you
do with the value after the ACCESS_ONCE(). For example, suppose you
do the following and MAX is a preprocessor macro with the value 1:
while ((tmp = ACCESS_ONCE(a)) % MAX)
do_something_with(tmp);
Then the compiler knows that the result of the "%" operator applied
to MAX will always be zero, again allowing the compiler to optimize
the code into near-nonexistence. (It will still load from the
variable 'a'.)
(*) Similarly, the compiler is within its rights to omit a store entirely
if it knows that the variable already has the value being stored.
Again, the compiler assumes that the current CPU is the only one
storing into the variable, which can cause the compiler to do the
wrong thing for shared variables. For example, suppose you have
the following:
a = 0;
/* Code that does not store to variable a. */
a = 0;
The compiler sees that the value of variable 'a' is already zero, so
it might well omit the second store. This would come as a fatal
surprise if some other CPU might have stored to variable 'a' in the
meantime.
Use ACCESS_ONCE() to prevent the compiler from making this sort of
wrong guess:
ACCESS_ONCE(a) = 0;
/* Code that does not store to variable a. */
ACCESS_ONCE(a) = 0;
(*) The compiler is within its rights to reorder memory accesses unless
you tell it not to. For example, consider the following interaction
between process-level code and an interrupt handler:
void process_level(void)
{
msg = get_message();
flag = true;
}
void interrupt_handler(void)
{
if (flag)
process_message(msg);
}
There is nothing to prevent the compiler from transforming
process_level() to the following, in fact, this might well be a
win for single-threaded code:
void process_level(void)
{
flag = true;
msg = get_message();
}
If the interrupt occurs between these two statement, then
interrupt_handler() might be passed a garbled msg. Use ACCESS_ONCE()
to prevent this as follows:
void process_level(void)
{
ACCESS_ONCE(msg) = get_message();
ACCESS_ONCE(flag) = true;
}
void interrupt_handler(void)
{
if (ACCESS_ONCE(flag))
process_message(ACCESS_ONCE(msg));
}
Note that the ACCESS_ONCE() wrappers in interrupt_handler()
are needed if this interrupt handler can itself be interrupted
by something that also accesses 'flag' and 'msg', for example,
a nested interrupt or an NMI. Otherwise, ACCESS_ONCE() is not
needed in interrupt_handler() other than for documentation purposes.
(Note also that nested interrupts do not typically occur in modern
Linux kernels, in fact, if an interrupt handler returns with
interrupts enabled, you will get a WARN_ONCE() splat.)
You should assume that the compiler can move ACCESS_ONCE() past
code not containing ACCESS_ONCE(), barrier(), or similar primitives.
This effect could also be achieved using barrier(), but ACCESS_ONCE()
is more selective: With ACCESS_ONCE(), the compiler need only forget
the contents of the indicated memory locations, while with barrier()
the compiler must discard the value of all memory locations that
it has currented cached in any machine registers. Of course,
the compiler must also respect the order in which the ACCESS_ONCE()s
occur, though the CPU of course need not do so.
(*) The compiler is within its rights to invent stores to a variable,
as in the following example:
if (a)
b = a;
else
b = 42;
The compiler might save a branch by optimizing this as follows:
b = 42;
if (a)
b = a;
In single-threaded code, this is not only safe, but also saves
a branch. Unfortunately, in concurrent code, this optimization
could cause some other CPU to see a spurious value of 42 -- even
if variable 'a' was never zero -- when loading variable 'b'.
Use ACCESS_ONCE() to prevent this as follows:
if (a)
ACCESS_ONCE(b) = a;
else
ACCESS_ONCE(b) = 42;
The compiler can also invent loads. These are usually less
damaging, but they can result in cache-line bouncing and thus in
poor performance and scalability. Use ACCESS_ONCE() to prevent
invented loads.
(*) For aligned memory locations whose size allows them to be accessed
with a single memory-reference instruction, prevents "load tearing"
and "store tearing," in which a single large access is replaced by
multiple smaller accesses. For example, given an architecture having
16-bit store instructions with 7-bit immediate fields, the compiler
might be tempted to use two 16-bit store-immediate instructions to
implement the following 32-bit store:
p = 0x00010002;
Please note that GCC really does use this sort of optimization,
which is not surprising given that it would likely take more
than two instructions to build the constant and then store it.
This optimization can therefore be a win in single-threaded code.
In fact, a recent bug (since fixed) caused GCC to incorrectly use
this optimization in a volatile store. In the absence of such bugs,
use of ACCESS_ONCE() prevents store tearing in the following example:
ACCESS_ONCE(p) = 0x00010002;
Use of packed structures can also result in load and store tearing,
as in this example:
struct __attribute__((__packed__)) foo {
short a;
int b;
short c;
};
struct foo foo1, foo2;
...
foo2.a = foo1.a;
foo2.b = foo1.b;
foo2.c = foo1.c;
Because there are no ACCESS_ONCE() wrappers and no volatile markings,
the compiler would be well within its rights to implement these three
assignment statements as a pair of 32-bit loads followed by a pair
of 32-bit stores. This would result in load tearing on 'foo1.b'
and store tearing on 'foo2.b'. ACCESS_ONCE() again prevents tearing
in this example:
foo2.a = foo1.a;
ACCESS_ONCE(foo2.b) = ACCESS_ONCE(foo1.b);
foo2.c = foo1.c;
All that aside, it is never necessary to use ACCESS_ONCE() on a variable
that has been marked volatile. For example, because 'jiffies' is marked
volatile, it is never necessary to say ACCESS_ONCE(jiffies). The reason
for this is that ACCESS_ONCE() is implemented as a volatile cast, which
has no effect when its argument is already marked volatile.
Please note that these compiler barriers have no direct effect on the CPU,
which may then reorder things however it wishes.
CPU MEMORY BARRIERS
-------------------
The Linux kernel has eight basic CPU memory barriers:
TYPE MANDATORY SMP CONDITIONAL
=============== ======================= ===========================
GENERAL mb() smp_mb()
WRITE wmb() smp_wmb()
READ rmb() smp_rmb()
DATA DEPENDENCY read_barrier_depends() smp_read_barrier_depends()
All memory barriers except the data dependency barriers imply a compiler
barrier. Data dependencies do not impose any additional compiler ordering.
Aside: In the case of data dependencies, the compiler would be expected to
issue the loads in the correct order (eg. `a[b]` would have to load the value
of b before loading a[b]), however there is no guarantee in the C specification
that the compiler may not speculate the value of b (eg. is equal to 1) and load
a before b (eg. tmp = a[1]; if (b != 1) tmp = a[b]; ). There is also the
problem of a compiler reloading b after having loaded a[b], thus having a newer
copy of b than a[b]. A consensus has not yet been reached about these problems,
however the ACCESS_ONCE macro is a good place to start looking.
SMP memory barriers are reduced to compiler barriers on uniprocessor compiled
systems because it is assumed that a CPU will appear to be self-consistent,
and will order overlapping accesses correctly with respect to itself.
[!] Note that SMP memory barriers _must_ be used to control the ordering of
references to shared memory on SMP systems, though the use of locking instead
is sufficient.
Mandatory barriers should not be used to control SMP effects, since mandatory
barriers unnecessarily impose overhead on UP systems. They may, however, be
used to control MMIO effects on accesses through relaxed memory I/O windows.
These are required even on non-SMP systems as they affect the order in which
memory operations appear to a device by prohibiting both the compiler and the
CPU from reordering them.
There are some more advanced barrier functions:
(*) smp_store_mb(var, value)
This assigns the value to the variable and then inserts a full memory
barrier after it, depending on the function. It isn't guaranteed to
insert anything more than a compiler barrier in a UP compilation.
(*) smp_mb__before_atomic();
(*) smp_mb__after_atomic();
These are for use with atomic (such as add, subtract, increment and
decrement) functions that don't return a value, especially when used for
reference counting. These functions do not imply memory barriers.
These are also used for atomic bitop functions that do not return a
value (such as set_bit and clear_bit).
As an example, consider a piece of code that marks an object as being dead
and then decrements the object's reference count:
obj->dead = 1;
smp_mb__before_atomic();
atomic_dec(&obj->ref_count);
This makes sure that the death mark on the object is perceived to be set
*before* the reference counter is decremented.
See Documentation/atomic_ops.txt for more information. See the "Atomic
operations" subsection for information on where to use these.
(*) dma_wmb();
(*) dma_rmb();
These are for use with consistent memory to guarantee the ordering
of writes or reads of shared memory accessible to both the CPU and a
DMA capable device.
For example, consider a device driver that shares memory with a device
and uses a descriptor status value to indicate if the descriptor belongs
to the device or the CPU, and a doorbell to notify it when new
descriptors are available:
if (desc->status != DEVICE_OWN) {
/* do not read data until we own descriptor */
dma_rmb();
/* read/modify data */
read_data = desc->data;
desc->data = write_data;
/* flush modifications before status update */
dma_wmb();
/* assign ownership */
desc->status = DEVICE_OWN;
/* force memory to sync before notifying device via MMIO */
wmb();
/* notify device of new descriptors */
writel(DESC_NOTIFY, doorbell);
}
The dma_rmb() allows us guarantee the device has released ownership
before we read the data from the descriptor, and the dma_wmb() allows
us to guarantee the data is written to the descriptor before the device
can see it now has ownership. The wmb() is needed to guarantee that the
cache coherent memory writes have completed before attempting a write to
the cache incoherent MMIO region.
See Documentation/DMA-API.txt for more information on consistent memory.
MMIO WRITE BARRIER
------------------
The Linux kernel also has a special barrier for use with memory-mapped I/O
writes:
mmiowb();
This is a variation on the mandatory write barrier that causes writes to weakly
ordered I/O regions to be partially ordered. Its effects may go beyond the
CPU->Hardware interface and actually affect the hardware at some level.
See the subsection "Locks vs I/O accesses" for more information.
===============================
IMPLICIT KERNEL MEMORY BARRIERS
===============================
Some of the other functions in the linux kernel imply memory barriers, amongst
which are locking and scheduling functions.
This specification is a _minimum_ guarantee; any particular architecture may
provide more substantial guarantees, but these may not be relied upon outside
of arch specific code.
ACQUIRING FUNCTIONS
-------------------
The Linux kernel has a number of locking constructs:
(*) spin locks
(*) R/W spin locks
(*) mutexes
(*) semaphores
(*) R/W semaphores
(*) RCU
In all cases there are variants on "ACQUIRE" operations and "RELEASE" operations
for each construct. These operations all imply certain barriers:
(1) ACQUIRE operation implication:
Memory operations issued after the ACQUIRE will be completed after the
ACQUIRE operation has completed.
Memory operations issued before the ACQUIRE may be completed after
the ACQUIRE operation has completed. An smp_mb__before_spinlock(),
combined with a following ACQUIRE, orders prior stores against
subsequent loads and stores. Note that this is weaker than smp_mb()!
The smp_mb__before_spinlock() primitive is free on many architectures.
(2) RELEASE operation implication:
Memory operations issued before the RELEASE will be completed before the
RELEASE operation has completed.
Memory operations issued after the RELEASE may be completed before the
RELEASE operation has completed.
(3) ACQUIRE vs ACQUIRE implication:
All ACQUIRE operations issued before another ACQUIRE operation will be
completed before that ACQUIRE operation.
(4) ACQUIRE vs RELEASE implication:
All ACQUIRE operations issued before a RELEASE operation will be
completed before the RELEASE operation.
(5) Failed conditional ACQUIRE implication:
Certain locking variants of the ACQUIRE operation may fail, either due to
being unable to get the lock immediately, or due to receiving an unblocked
signal whilst asleep waiting for the lock to become available. Failed
locks do not imply any sort of barrier.
[!] Note: one of the consequences of lock ACQUIREs and RELEASEs being only
one-way barriers is that the effects of instructions outside of a critical
section may seep into the inside of the critical section.
An ACQUIRE followed by a RELEASE may not be assumed to be full memory barrier
because it is possible for an access preceding the ACQUIRE to happen after the
ACQUIRE, and an access following the RELEASE to happen before the RELEASE, and
the two accesses can themselves then cross:
*A = a;
ACQUIRE M
RELEASE M
*B = b;
may occur as:
ACQUIRE M, STORE *B, STORE *A, RELEASE M
When the ACQUIRE and RELEASE are a lock acquisition and release,
respectively, this same reordering can occur if the lock's ACQUIRE and
RELEASE are to the same lock variable, but only from the perspective of
another CPU not holding that lock. In short, a ACQUIRE followed by an
RELEASE may -not- be assumed to be a full memory barrier.
Similarly, the reverse case of a RELEASE followed by an ACQUIRE does not
imply a full memory barrier. If it is necessary for a RELEASE-ACQUIRE
pair to produce a full barrier, the ACQUIRE can be followed by an
smp_mb__after_unlock_lock() invocation. This will produce a full barrier
if either (a) the RELEASE and the ACQUIRE are executed by the same
CPU or task, or (b) the RELEASE and ACQUIRE act on the same variable.
The smp_mb__after_unlock_lock() primitive is free on many architectures.
Without smp_mb__after_unlock_lock(), the CPU's execution of the critical
sections corresponding to the RELEASE and the ACQUIRE can cross, so that:
*A = a;
RELEASE M
ACQUIRE N
*B = b;
could occur as:
ACQUIRE N, STORE *B, STORE *A, RELEASE M
It might appear that this reordering could introduce a deadlock.
However, this cannot happen because if such a deadlock threatened,
the RELEASE would simply complete, thereby avoiding the deadlock.
Why does this work?
One key point is that we are only talking about the CPU doing
the reordering, not the compiler. If the compiler (or, for
that matter, the developer) switched the operations, deadlock
-could- occur.
But suppose the CPU reordered the operations. In this case,
the unlock precedes the lock in the assembly code. The CPU
simply elected to try executing the later lock operation first.
If there is a deadlock, this lock operation will simply spin (or
try to sleep, but more on that later). The CPU will eventually
execute the unlock operation (which preceded the lock operation
in the assembly code), which will unravel the potential deadlock,
allowing the lock operation to succeed.
But what if the lock is a sleeplock? In that case, the code will
try to enter the scheduler, where it will eventually encounter
a memory barrier, which will force the earlier unlock operation
to complete, again unraveling the deadlock. There might be
a sleep-unlock race, but the locking primitive needs to resolve
such races properly in any case.
With smp_mb__after_unlock_lock(), the two critical sections cannot overlap.
For example, with the following code, the store to *A will always be
seen by other CPUs before the store to *B:
*A = a;
RELEASE M
ACQUIRE N
smp_mb__after_unlock_lock();
*B = b;
The operations will always occur in one of the following orders:
STORE *A, RELEASE, ACQUIRE, smp_mb__after_unlock_lock(), STORE *B
STORE *A, ACQUIRE, RELEASE, smp_mb__after_unlock_lock(), STORE *B
ACQUIRE, STORE *A, RELEASE, smp_mb__after_unlock_lock(), STORE *B
If the RELEASE and ACQUIRE were instead both operating on the same lock
variable, only the first of these alternatives can occur. In addition,
the more strongly ordered systems may rule out some of the above orders.
But in any case, as noted earlier, the smp_mb__after_unlock_lock()
ensures that the store to *A will always be seen as happening before
the store to *B.
Locks and semaphores may not provide any guarantee of ordering on UP compiled
systems, and so cannot be counted on in such a situation to actually achieve
anything at all - especially with respect to I/O accesses - unless combined
with interrupt disabling operations.
See also the section on "Inter-CPU locking barrier effects".
As an example, consider the following:
*A = a;
*B = b;
ACQUIRE
*C = c;
*D = d;
RELEASE
*E = e;
*F = f;
The following sequence of events is acceptable:
ACQUIRE, {*F,*A}, *E, {*C,*D}, *B, RELEASE
[+] Note that {*F,*A} indicates a combined access.
But none of the following are:
{*F,*A}, *B, ACQUIRE, *C, *D, RELEASE, *E
*A, *B, *C, ACQUIRE, *D, RELEASE, *E, *F
*A, *B, ACQUIRE, *C, RELEASE, *D, *E, *F
*B, ACQUIRE, *C, *D, RELEASE, {*F,*A}, *E
INTERRUPT DISABLING FUNCTIONS
-----------------------------
Functions that disable interrupts (ACQUIRE equivalent) and enable interrupts
(RELEASE equivalent) will act as compiler barriers only. So if memory or I/O
barriers are required in such a situation, they must be provided from some
other means.
SLEEP AND WAKE-UP FUNCTIONS
---------------------------
Sleeping and waking on an event flagged in global data can be viewed as an
interaction between two pieces of data: the task state of the task waiting for
the event and the global data used to indicate the event. To make sure that
these appear to happen in the right order, the primitives to begin the process
of going to sleep, and the primitives to initiate a wake up imply certain
barriers.
Firstly, the sleeper normally follows something like this sequence of events:
for (;;) {
set_current_state(TASK_UNINTERRUPTIBLE);
if (event_indicated)
break;
schedule();
}
A general memory barrier is interpolated automatically by set_current_state()
after it has altered the task state:
CPU 1
===============================
set_current_state();
smp_store_mb();
STORE current->state
<general barrier>
LOAD event_indicated
set_current_state() may be wrapped by:
prepare_to_wait();
prepare_to_wait_exclusive();
which therefore also imply a general memory barrier after setting the state.
The whole sequence above is available in various canned forms, all of which
interpolate the memory barrier in the right place:
wait_event();
wait_event_interruptible();
wait_event_interruptible_exclusive();
wait_event_interruptible_timeout();
wait_event_killable();
wait_event_timeout();
wait_on_bit();
wait_on_bit_lock();
Secondly, code that performs a wake up normally follows something like this:
event_indicated = 1;
wake_up(&event_wait_queue);
or:
event_indicated = 1;
wake_up_process(event_daemon);
A write memory barrier is implied by wake_up() and co. if and only if they wake
something up. The barrier occurs before the task state is cleared, and so sits
between the STORE to indicate the event and the STORE to set TASK_RUNNING:
CPU 1 CPU 2
=============================== ===============================
set_current_state(); STORE event_indicated
smp_store_mb(); wake_up();
STORE current->state <write barrier>
<general barrier> STORE current->state
LOAD event_indicated
To repeat, this write memory barrier is present if and only if something
is actually awakened. To see this, consider the following sequence of
events, where X and Y are both initially zero:
CPU 1 CPU 2
=============================== ===============================
X = 1; STORE event_indicated
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
In contrast, if a wakeup does occur, CPU 2's load from X would be guaranteed
to see 1.
The available waker functions include:
complete();
wake_up();
wake_up_all();
wake_up_bit();
wake_up_interruptible();
wake_up_interruptible_all();
wake_up_interruptible_nr();
wake_up_interruptible_poll();
wake_up_interruptible_sync();
wake_up_interruptible_sync_poll();
wake_up_locked();
wake_up_locked_poll();
wake_up_nr();
wake_up_poll();
wake_up_process();
[!] Note that the memory barriers implied by the sleeper and the waker do _not_
order multiple stores before the wake-up with respect to loads of those stored
values after the sleeper has called set_current_state(). For instance, if the
sleeper does:
set_current_state(TASK_INTERRUPTIBLE);
if (event_indicated)
break;
__set_current_state(TASK_RUNNING);
do_something(my_data);
and the waker does:
my_data = value;
event_indicated = 1;
wake_up(&event_wait_queue);
there's no guarantee that the change to event_indicated will be perceived by
the sleeper as coming after the change to my_data. In such a circumstance, the
code on both sides must interpolate its own memory barriers between the
separate data accesses. Thus the above sleeper ought to do:
set_current_state(TASK_INTERRUPTIBLE);
if (event_indicated) {
smp_rmb();
do_something(my_data);
}
and the waker should do:
my_data = value;
smp_wmb();
event_indicated = 1;
wake_up(&event_wait_queue);
MISCELLANEOUS FUNCTIONS
-----------------------
Other functions that imply barriers:
(*) schedule() and similar imply full memory barriers.
===================================
INTER-CPU ACQUIRING BARRIER EFFECTS
===================================
On SMP systems locking primitives give a more substantial form of barrier: one
that does affect memory access ordering on other CPUs, within the context of
conflict on any particular lock.
ACQUIRES VS MEMORY ACCESSES
---------------------------
Consider the following: the system has a pair of spinlocks (M) and (Q), and
three CPUs; then should the following sequence of events occur:
CPU 1 CPU 2
=============================== ===============================
ACCESS_ONCE(*A) = a; ACCESS_ONCE(*E) = e;
ACQUIRE M ACQUIRE Q
ACCESS_ONCE(*B) = b; ACCESS_ONCE(*F) = f;
ACCESS_ONCE(*C) = c; ACCESS_ONCE(*G) = g;
RELEASE M RELEASE Q
ACCESS_ONCE(*D) = d; ACCESS_ONCE(*H) = h;
Then there is no guarantee as to what order CPU 3 will see the accesses to *A
through *H occur in, other than the constraints imposed by the separate locks
on the separate CPUs. It might, for example, see:
*E, ACQUIRE M, ACQUIRE Q, *G, *C, *F, *A, *B, RELEASE Q, *D, *H, RELEASE M
But it won't see any of:
*B, *C or *D preceding ACQUIRE M
*A, *B or *C following RELEASE M
*F, *G or *H preceding ACQUIRE Q
*E, *F or *G following RELEASE Q
However, if the following occurs:
CPU 1 CPU 2
=============================== ===============================
ACCESS_ONCE(*A) = a;
ACQUIRE M [1]
ACCESS_ONCE(*B) = b;
ACCESS_ONCE(*C) = c;
RELEASE M [1]
ACCESS_ONCE(*D) = d; ACCESS_ONCE(*E) = e;
ACQUIRE M [2]
smp_mb__after_unlock_lock();
ACCESS_ONCE(*F) = f;
ACCESS_ONCE(*G) = g;
RELEASE M [2]
ACCESS_ONCE(*H) = h;
CPU 3 might see:
*E, ACQUIRE M [1], *C, *B, *A, RELEASE M [1],
ACQUIRE M [2], *H, *F, *G, RELEASE M [2], *D
But assuming CPU 1 gets the lock first, CPU 3 won't see any of:
*B, *C, *D, *F, *G or *H preceding ACQUIRE M [1]
*A, *B or *C following RELEASE M [1]
*F, *G or *H preceding ACQUIRE M [2]
*A, *B, *C, *E, *F or *G following RELEASE M [2]
Note that the smp_mb__after_unlock_lock() is critically important
here: Without it CPU 3 might see some of the above orderings.
Without smp_mb__after_unlock_lock(), the accesses are not guaranteed
to be seen in order unless CPU 3 holds lock M.
ACQUIRES VS I/O ACCESSES
------------------------
Under certain circumstances (especially involving NUMA), I/O accesses within
two spinlocked sections on two different CPUs may be seen as interleaved by the
PCI bridge, because the PCI bridge does not necessarily participate in the
cache-coherence protocol, and is therefore incapable of issuing the required
read memory barriers.
For example:
CPU 1 CPU 2
=============================== ===============================
spin_lock(Q)
writel(0, ADDR)
writel(1, DATA);
spin_unlock(Q);
spin_lock(Q);
writel(4, ADDR);
writel(5, DATA);
spin_unlock(Q);
may be seen by the PCI bridge as follows:
STORE *ADDR = 0, STORE *ADDR = 4, STORE *DATA = 1, STORE *DATA = 5
which would probably cause the hardware to malfunction.
What is necessary here is to intervene with an mmiowb() before dropping the
spinlock, for example:
CPU 1 CPU 2
=============================== ===============================
spin_lock(Q)
writel(0, ADDR)
writel(1, DATA);
mmiowb();
spin_unlock(Q);
spin_lock(Q);
writel(4, ADDR);
writel(5, DATA);
mmiowb();
spin_unlock(Q);
this will ensure that the two stores issued on CPU 1 appear at the PCI bridge
before either of the stores issued on CPU 2.
Furthermore, following a store by a load from the same device obviates the need
for the mmiowb(), because the load forces the store to complete before the load
is performed:
CPU 1 CPU 2
=============================== ===============================
spin_lock(Q)
writel(0, ADDR)
a = readl(DATA);
spin_unlock(Q);
spin_lock(Q);
writel(4, ADDR);
b = readl(DATA);
spin_unlock(Q);
See Documentation/DocBook/deviceiobook.tmpl for more information.
=================================
WHERE ARE MEMORY BARRIERS NEEDED?
=================================
Under normal operation, memory operation reordering is generally not going to
be a problem as a single-threaded linear piece of code will still appear to
work correctly, even if it's in an SMP kernel. There are, however, four
circumstances in which reordering definitely _could_ be a problem:
(*) Interprocessor interaction.
(*) Atomic operations.
(*) Accessing devices.
(*) Interrupts.
INTERPROCESSOR INTERACTION
--------------------------
When there's a system with more than one processor, more than one CPU in the
system may be working on the same data set at the same time. This can cause
synchronisation problems, and the usual way of dealing with them is to use
locks. Locks, however, are quite expensive, and so it may be preferable to
operate without the use of a lock if at all possible. In such a case
operations that affect both CPUs may have to be carefully ordered to prevent
a malfunction.
Consider, for example, the R/W semaphore slow path. Here a waiting process is
queued on the semaphore, by virtue of it having a piece of its stack linked to
the semaphore's list of waiting processes:
struct rw_semaphore {
...
spinlock_t lock;
struct list_head waiters;
};
struct rwsem_waiter {
struct list_head list;
struct task_struct *task;
};
To wake up a particular waiter, the up_read() or up_write() functions have to:
(1) read the next pointer from this waiter's record to know as to where the
next waiter record is;
(2) read the pointer to the waiter's task structure;
(3) clear the task pointer to tell the waiter it has been given the semaphore;
(4) call wake_up_process() on the task; and
(5) release the reference held on the waiter's task struct.
In other words, it has to perform this sequence of events:
LOAD waiter->list.next;
LOAD waiter->task;
STORE waiter->task;
CALL wakeup
RELEASE task
and if any of these steps occur out of order, then the whole thing may
malfunction.
Once it has queued itself and dropped the semaphore lock, the waiter does not
get the lock again; it instead just waits for its task pointer to be cleared
before proceeding. Since the record is on the waiter's stack, this means that
if the task pointer is cleared _before_ the next pointer in the list is read,
another CPU might start processing the waiter and might clobber the waiter's
stack before the up*() function has a chance to read the next pointer.
Consider then what might happen to the above sequence of events:
CPU 1 CPU 2
=============================== ===============================
down_xxx()
Queue waiter
Sleep
up_yyy()
LOAD waiter->task;
STORE waiter->task;
Woken up by other event
<preempt>
Resume processing
down_xxx() returns
call foo()
foo() clobbers *waiter
</preempt>
LOAD waiter->list.next;
--- OOPS ---
This could be dealt with using the semaphore lock, but then the down_xxx()
function has to needlessly get the spinlock again after being woken up.
The way to deal with this is to insert a general SMP memory barrier:
LOAD waiter->list.next;
LOAD waiter->task;
smp_mb();
STORE waiter->task;
CALL wakeup
RELEASE task
In this case, the barrier makes a guarantee that all memory accesses before the
barrier will appear to happen before all the memory accesses after the barrier
with respect to the other CPUs on the system. It does _not_ guarantee that all
the memory accesses before the barrier will be complete by the time the barrier
instruction itself is complete.
On a UP system - where this wouldn't be a problem - the smp_mb() is just a
compiler barrier, thus making sure the compiler emits the instructions in the
right order without actually intervening in the CPU. Since there's only one
CPU, that CPU's dependency ordering logic will take care of everything else.
ATOMIC OPERATIONS
-----------------
Whilst they are technically interprocessor interaction considerations, atomic
operations are noted specially as some of them imply full memory barriers and
some don't, but they're very heavily relied on as a group throughout the
kernel.
Any atomic operation that modifies some state in memory and returns information
about the state (old or new) implies an SMP-conditional general memory barrier
(smp_mb()) on each side of the actual operation (with the exception of
explicit lock operations, described later). These include:
xchg();
cmpxchg();
atomic_xchg(); atomic_long_xchg();
atomic_cmpxchg(); atomic_long_cmpxchg();
atomic_inc_return(); atomic_long_inc_return();
atomic_dec_return(); atomic_long_dec_return();
atomic_add_return(); atomic_long_add_return();
atomic_sub_return(); atomic_long_sub_return();
atomic_inc_and_test(); atomic_long_inc_and_test();
atomic_dec_and_test(); atomic_long_dec_and_test();
atomic_sub_and_test(); atomic_long_sub_and_test();
atomic_add_negative(); atomic_long_add_negative();
test_and_set_bit();
test_and_clear_bit();
test_and_change_bit();
/* when succeeds (returns 1) */
atomic_add_unless(); atomic_long_add_unless();
These are used for such things as implementing ACQUIRE-class and RELEASE-class
operations and adjusting reference counters towards object destruction, and as
such the implicit memory barrier effects are necessary.
The following operations are potential problems as they do _not_ imply memory
barriers, but might be used for implementing such things as RELEASE-class
operations:
atomic_set();
set_bit();
clear_bit();
change_bit();
With these the appropriate explicit memory barrier should be used if necessary
(smp_mb__before_atomic() for instance).
The following also do _not_ imply memory barriers, and so may require explicit
memory barriers under some circumstances (smp_mb__before_atomic() for
instance):
atomic_add();
atomic_sub();
atomic_inc();
atomic_dec();
If they're used for statistics generation, then they probably don't need memory
barriers, unless there's a coupling between statistical data.
If they're used for reference counting on an object to control its lifetime,
they probably don't need memory barriers because either the reference count
will be adjusted inside a locked section, or the caller will already hold
sufficient references to make the lock, and thus a memory barrier unnecessary.
If they're used for constructing a lock of some description, then they probably
do need memory barriers as a lock primitive generally has to do things in a
specific order.
Basically, each usage case has to be carefully considered as to whether memory
barriers are needed or not.
The following operations are special locking primitives:
test_and_set_bit_lock();
clear_bit_unlock();
__clear_bit_unlock();
These implement ACQUIRE-class and RELEASE-class operations. These should be used in
preference to other operations when implementing locking primitives, because
their implementations can be optimised on many architectures.
[!] Note that special memory barrier primitives are available for these
situations because on some CPUs the atomic instructions used imply full memory
barriers, and so barrier instructions are superfluous in conjunction with them,
and in such cases the special barrier primitives will be no-ops.
See Documentation/atomic_ops.txt for more information.
ACCESSING DEVICES
-----------------
Many devices can be memory mapped, and so appear to the CPU as if they're just
a set of memory locations. To control such a device, the driver usually has to
make the right memory accesses in exactly the right order.
However, having a clever CPU or a clever compiler creates a potential problem
in that the carefully sequenced accesses in the driver code won't reach the
device in the requisite order if the CPU or the compiler thinks it is more
efficient to reorder, combine or merge accesses - something that would cause
the device to malfunction.
Inside of the Linux kernel, I/O should be done through the appropriate accessor
routines - such as inb() or writel() - which know how to make such accesses
appropriately sequential. Whilst this, for the most part, renders the explicit
use of memory barriers unnecessary, there are a couple of situations where they
might be needed:
(1) On some systems, I/O stores are not strongly ordered across all CPUs, and
so for _all_ general drivers locks should be used and mmiowb() must be
issued prior to unlocking the critical section.
(2) If the accessor functions are used to refer to an I/O memory window with
relaxed memory access properties, then _mandatory_ memory barriers are
required to enforce ordering.
See Documentation/DocBook/deviceiobook.tmpl for more information.
INTERRUPTS
----------
A driver may be interrupted by its own interrupt service routine, and thus the
two parts of the driver may interfere with each other's attempts to control or
access the device.
This may be alleviated - at least in part - by disabling local interrupts (a
form of locking), such that the critical operations are all contained within
the interrupt-disabled section in the driver. Whilst the driver's interrupt
routine is executing, the driver's core may not run on the same CPU, and its
interrupt is not permitted to happen again until the current interrupt has been
handled, thus the interrupt handler does not need to lock against that.
However, consider a driver that was talking to an ethernet card that sports an
address register and a data register. If that driver's core talks to the card
under interrupt-disablement and then the driver's interrupt handler is invoked:
LOCAL IRQ DISABLE
writew(ADDR, 3);
writew(DATA, y);
LOCAL IRQ ENABLE
<interrupt>
writew(ADDR, 4);
q = readw(DATA);
</interrupt>
The store to the data register might happen after the second store to the
address register if ordering rules are sufficiently relaxed:
STORE *ADDR = 3, STORE *ADDR = 4, STORE *DATA = y, q = LOAD *DATA
If ordering rules are relaxed, it must be assumed that accesses done inside an
interrupt disabled section may leak outside of it and may interleave with
accesses performed in an interrupt - and vice versa - unless implicit or
explicit barriers are used.
Normally this won't be a problem because the I/O accesses done inside such
sections will include synchronous load operations on strictly ordered I/O
registers that form implicit I/O barriers. If this isn't sufficient then an
mmiowb() may need to be used explicitly.
A similar situation may occur between an interrupt routine and two routines
running on separate CPUs that communicate with each other. If such a case is
likely, then interrupt-disabling locks should be used to guarantee ordering.
==========================
KERNEL I/O BARRIER EFFECTS
==========================
When accessing I/O memory, drivers should use the appropriate accessor
functions:
(*) inX(), outX():
These are intended to talk to I/O space rather than memory space, but
that's primarily a CPU-specific concept. The i386 and x86_64 processors do
indeed have special I/O space access cycles and instructions, but many
CPUs don't have such a concept.
The PCI bus, amongst others, defines an I/O space concept which - on such
CPUs as i386 and x86_64 - readily maps to the CPU's concept of I/O
space. However, it may also be mapped as a virtual I/O space in the CPU's
memory map, particularly on those CPUs that don't support alternate I/O
spaces.
Accesses to this space may be fully synchronous (as on i386), but
intermediary bridges (such as the PCI host bridge) may not fully honour
that.
They are guaranteed to be fully ordered with respect to each other.
They are not guaranteed to be fully ordered with respect to other types of
memory and I/O operation.
(*) readX(), writeX():
Whether these are guaranteed to be fully ordered and uncombined with
respect to each other on the issuing CPU depends on the characteristics
defined for the memory window through which they're accessing. On later
i386 architecture machines, for example, this is controlled by way of the
MTRR registers.
Ordinarily, these will be guaranteed to be fully ordered and uncombined,
provided they're not accessing a prefetchable device.
However, intermediary hardware (such as a PCI bridge) may indulge in
deferral if it so wishes; to flush a store, a load from the same location
is preferred[*], but a load from the same device or from configuration
space should suffice for PCI.
[*] NOTE! attempting to load from the same location as was written to may
cause a malfunction - consider the 16550 Rx/Tx serial registers for
example.
Used with prefetchable I/O memory, an mmiowb() barrier may be required to
force stores to be ordered.
Please refer to the PCI specification for more information on interactions
between PCI transactions.
(*) readX_relaxed(), writeX_relaxed()
These are similar to readX() and writeX(), but provide weaker memory
ordering guarantees. Specifically, they do not guarantee ordering with
respect to normal memory accesses (e.g. DMA buffers) nor do they guarantee
ordering with respect to LOCK or UNLOCK operations. If the latter is
required, an mmiowb() barrier can be used. Note that relaxed accesses to
the same peripheral are guaranteed to be ordered with respect to each
other.
(*) ioreadX(), iowriteX()
These will perform appropriately for the type of access they're actually
doing, be it inX()/outX() or readX()/writeX().
========================================
ASSUMED MINIMUM EXECUTION ORDERING MODEL
========================================
It has to be assumed that the conceptual CPU is weakly-ordered but that it will
maintain the appearance of program causality with respect to itself. Some CPUs
(such as i386 or x86_64) are more constrained than others (such as powerpc or
frv), and so the most relaxed case (namely DEC Alpha) must be assumed outside
of arch-specific code.
This means that it must be considered that the CPU will execute its instruction
stream in any order it feels like - or even in parallel - provided that if an
instruction in the stream depends on an earlier instruction, then that
earlier instruction must be sufficiently complete[*] before the later
instruction may proceed; in other words: provided that the appearance of
causality is maintained.
[*] Some instructions have more than one effect - such as changing the
condition codes, changing registers or changing memory - and different
instructions may depend on different effects.
A CPU may also discard any instruction sequence that winds up having no
ultimate effect. For example, if two adjacent instructions both load an
immediate value into the same register, the first may be discarded.
Similarly, it has to be assumed that compiler might reorder the instruction
stream in any way it sees fit, again provided the appearance of causality is
maintained.
============================
THE EFFECTS OF THE CPU CACHE
============================
The way cached memory operations are perceived across the system is affected to
a certain extent by the caches that lie between CPUs and memory, and by the
memory coherence system that maintains the consistency of state in the system.
As far as the way a CPU interacts with another part of the system through the
caches goes, the memory system has to include the CPU's caches, and memory
barriers for the most part act at the interface between the CPU and its cache
(memory barriers logically act on the dotted line in the following diagram):
<--- CPU ---> : <----------- Memory ----------->
:
+--------+ +--------+ : +--------+ +-----------+
| | | | : | | | | +--------+
| CPU | | Memory | : | CPU | | | | |
| Core |--->| Access |----->| Cache |<-->| | | |
| | | Queue | : | | | |--->| Memory |
| | | | : | | | | | |
+--------+ +--------+ : +--------+ | | | |
: | Cache | +--------+
: | Coherency |
: | Mechanism | +--------+
+--------+ +--------+ : +--------+ | | | |
| | | | : | | | | | |
| CPU | | Memory | : | CPU | | |--->| Device |
| Core |--->| Access |----->| Cache |<-->| | | |
| | | Queue | : | | | | | |
| | | | : | | | | +--------+
+--------+ +--------+ : +--------+ +-----------+
:
:
Although any particular load or store may not actually appear outside of the
CPU that issued it since it may have been satisfied within the CPU's own cache,
it will still appear as if the full memory access had taken place as far as the
other CPUs are concerned since the cache coherency mechanisms will migrate the
cacheline over to the accessing CPU and propagate the effects upon conflict.
The CPU core may execute instructions in any order it deems fit, provided the
expected program causality appears to be maintained. Some of the instructions
generate load and store operations which then go into the queue of memory
accesses to be performed. The core may place these in the queue in any order
it wishes, and continue execution until it is forced to wait for an instruction
to complete.
What memory barriers are concerned with is controlling the order in which
accesses cross from the CPU side of things to the memory side of things, and
the order in which the effects are perceived to happen by the other observers
in the system.
[!] Memory barriers are _not_ needed within a given CPU, as CPUs always see
their own loads and stores as if they had happened in program order.
[!] MMIO or other device accesses may bypass the cache system. This depends on
the properties of the memory window through which devices are accessed and/or
the use of any special device communication instructions the CPU may have.
CACHE COHERENCY
---------------
Life isn't quite as simple as it may appear above, however: for while the
caches are expected to be coherent, there's no guarantee that that coherency
will be ordered. This means that whilst changes made on one CPU will
eventually become visible on all CPUs, there's no guarantee that they will
become apparent in the same order on those other CPUs.
Consider dealing with a system that has a pair of CPUs (1 & 2), each of which
has a pair of parallel data caches (CPU 1 has A/B, and CPU 2 has C/D):
:
: +--------+
: +---------+ | |
+--------+ : +--->| Cache A |<------->| |
| | : | +---------+ | |
| CPU 1 |<---+ | |
| | : | +---------+ | |
+--------+ : +--->| Cache B |<------->| |
: +---------+ | |
: | Memory |
: +---------+ | System |
+--------+ : +--->| Cache C |<------->| |
| | : | +---------+ | |
| CPU 2 |<---+ | |
| | : | +---------+ | |
+--------+ : +--->| Cache D |<------->| |
: +---------+ | |
: +--------+
:
Imagine the system has the following properties:
(*) an odd-numbered cache line may be in cache A, cache C or it may still be
resident in memory;
(*) an even-numbered cache line may be in cache B, cache D or it may still be
resident in memory;
(*) whilst the CPU core is interrogating one cache, the other cache may be
making use of the bus to access the rest of the system - perhaps to
displace a dirty cacheline or to do a speculative load;
(*) each cache has a queue of operations that need to be applied to that cache
to maintain coherency with the rest of the system;
(*) the coherency queue is not flushed by normal loads to lines already
present in the cache, even though the contents of the queue may
potentially affect those loads.
Imagine, then, that two writes are made on the first CPU, with a write barrier
between them to guarantee that they will appear to reach that CPU's caches in
the requisite order:
CPU 1 CPU 2 COMMENT
=============== =============== =======================================
u == 0, v == 1 and p == &u, q == &u
v = 2;
smp_wmb(); Make sure change to v is visible before
change to p
<A:modify v=2> v is now in cache A exclusively
p = &v;
<B:modify p=&v> p is now in cache B exclusively
The write memory barrier forces the other CPUs in the system to perceive that
the local CPU's caches have apparently been updated in the correct order. But
now imagine that the second CPU wants to read those values:
CPU 1 CPU 2 COMMENT
=============== =============== =======================================
...
q = p;
x = *q;
The above pair of reads may then fail to happen in the expected order, as the
cacheline holding p may get updated in one of the second CPU's caches whilst
the update to the cacheline holding v is delayed in the other of the second
CPU's caches by some other cache event:
CPU 1 CPU 2 COMMENT
=============== =============== =======================================
u == 0, v == 1 and p == &u, q == &u
v = 2;
smp_wmb();
<A:modify v=2> <C:busy>
<C:queue v=2>
p = &v; q = p;
<D:request p>
<B:modify p=&v> <D:commit p=&v>
<D:read p>
x = *q;
<C:read *q> Reads from v before v updated in cache
<C:unbusy>
<C:commit v=2>
Basically, whilst both cachelines will be updated on CPU 2 eventually, there's
no guarantee that, without intervention, the order of update will be the same
as that committed on CPU 1.
To intervene, we need to interpolate a data dependency barrier or a read
barrier between the loads. This will force the cache to commit its coherency
queue before processing any further requests:
CPU 1 CPU 2 COMMENT
=============== =============== =======================================
u == 0, v == 1 and p == &u, q == &u
v = 2;
smp_wmb();
<A:modify v=2> <C:busy>
<C:queue v=2>
p = &v; q = p;
<D:request p>
<B:modify p=&v> <D:commit p=&v>
<D:read p>
smp_read_barrier_depends()
<C:unbusy>
<C:commit v=2>
x = *q;
<C:read *q> Reads from v after v updated in cache
This sort of problem can be encountered on DEC Alpha processors as they have a
split cache that improves performance by making better use of the data bus.
Whilst most CPUs do imply a data dependency barrier on the read when a memory
access depends on a read, not all do, so it may not be relied on.
Other CPUs may also have split caches, but must coordinate between the various
cachelets for normal memory accesses. The semantics of the Alpha removes the
need for coordination in the absence of memory barriers.
CACHE COHERENCY VS DMA
----------------------
Not all systems maintain cache coherency with respect to devices doing DMA. In
such cases, a device attempting DMA may obtain stale data from RAM because
dirty cache lines may be resident in the caches of various CPUs, and may not
have been written back to RAM yet. To deal with this, the appropriate part of
the kernel must flush the overlapping bits of cache on each CPU (and maybe
invalidate them as well).
In addition, the data DMA'd to RAM by a device may be overwritten by dirty
cache lines being written back to RAM from a CPU's cache after the device has
installed its own data, or cache lines present in the CPU's cache may simply
obscure the fact that RAM has been updated, until at such time as the cacheline
is discarded from the CPU's cache and reloaded. To deal with this, the
appropriate part of the kernel must invalidate the overlapping bits of the
cache on each CPU.
See Documentation/cachetlb.txt for more information on cache management.
CACHE COHERENCY VS MMIO
-----------------------
Memory mapped I/O usually takes place through memory locations that are part of
a window in the CPU's memory space that has different properties assigned than
the usual RAM directed window.
Amongst these properties is usually the fact that such accesses bypass the
caching entirely and go directly to the device buses. This means MMIO accesses
may, in effect, overtake accesses to cached memory that were emitted earlier.
A memory barrier isn't sufficient in such a case, but rather the cache must be
flushed between the cached memory write and the MMIO access if the two are in
any way dependent.
=========================
THE THINGS CPUS GET UP TO
=========================
A programmer might take it for granted that the CPU will perform memory
operations in exactly the order specified, so that if the CPU is, for example,
given the following piece of code to execute:
a = ACCESS_ONCE(*A);
ACCESS_ONCE(*B) = b;
c = ACCESS_ONCE(*C);
d = ACCESS_ONCE(*D);
ACCESS_ONCE(*E) = e;
they would then expect that the CPU will complete the memory operation for each
instruction before moving on to the next one, leading to a definite sequence of
operations as seen by external observers in the system:
LOAD *A, STORE *B, LOAD *C, LOAD *D, STORE *E.
Reality is, of course, much messier. With many CPUs and compilers, the above
assumption doesn't hold because:
(*) loads are more likely to need to be completed immediately to permit
execution progress, whereas stores can often be deferred without a
problem;
(*) loads may be done speculatively, and the result discarded should it prove
to have been unnecessary;
(*) loads may be done speculatively, leading to the result having been fetched
at the wrong time in the expected sequence of events;
(*) the order of the memory accesses may be rearranged to promote better use
of the CPU buses and caches;
(*) loads and stores may be combined to improve performance when talking to
memory or I/O hardware that can do batched accesses of adjacent locations,
thus cutting down on transaction setup costs (memory and PCI devices may
both be able to do this); and
(*) the CPU's data cache may affect the ordering, and whilst cache-coherency
mechanisms may alleviate this - once the store has actually hit the cache
- there's no guarantee that the coherency management will be propagated in
order to other CPUs.
So what another CPU, say, might actually observe from the above piece of code
is:
LOAD *A, ..., LOAD {*C,*D}, STORE *E, STORE *B
(Where "LOAD {*C,*D}" is a combined load)
However, it is guaranteed that a CPU will be self-consistent: it will see its
_own_ accesses appear to be correctly ordered, without the need for a memory
barrier. For instance with the following code:
U = ACCESS_ONCE(*A);
ACCESS_ONCE(*A) = V;
ACCESS_ONCE(*A) = W;
X = ACCESS_ONCE(*A);
ACCESS_ONCE(*A) = Y;
Z = ACCESS_ONCE(*A);
and assuming no intervention by an external influence, it can be assumed that
the final result will appear to be:
U == the original value of *A
X == W
Z == Y
*A == Y
The code above may cause the CPU to generate the full sequence of memory
accesses:
U=LOAD *A, STORE *A=V, STORE *A=W, X=LOAD *A, STORE *A=Y, Z=LOAD *A
in that order, but, without intervention, the sequence may have almost any
combination of elements combined or discarded, provided the program's view of
the world remains consistent. Note that ACCESS_ONCE() is -not- optional
in the above example, as there are architectures where a given CPU might
reorder successive loads to the same location. On such architectures,
ACCESS_ONCE() does whatever is necessary to prevent this, for example, on
Itanium the volatile casts used by ACCESS_ONCE() cause GCC to emit the
special ld.acq and st.rel instructions that prevent such reordering.
The compiler may also combine, discard or defer elements of the sequence before
the CPU even sees them.
For instance:
*A = V;
*A = W;
may be reduced to:
*A = W;
since, without either a write barrier or an ACCESS_ONCE(), it can be
assumed that the effect of the storage of V to *A is lost. Similarly:
*A = Y;
Z = *A;
may, without a memory barrier or an ACCESS_ONCE(), be reduced to:
*A = Y;
Z = Y;
and the LOAD operation never appear outside of the CPU.
AND THEN THERE'S THE ALPHA
--------------------------
The DEC Alpha CPU is one of the most relaxed CPUs there is. Not only that,
some versions of the Alpha CPU have a split data cache, permitting them to have
two semantically-related cache lines updated at separate times. This is where
the data dependency barrier really becomes necessary as this synchronises both
caches with the memory coherence system, thus making it seem like pointer
changes vs new data occur in the right order.
The Alpha defines the Linux kernel's memory barrier model.
See the subsection on "Cache Coherency" above.
============
EXAMPLE USES
============
CIRCULAR BUFFERS
----------------
Memory barriers can be used to implement circular buffering without the need
of a lock to serialise the producer with the consumer. See:
Documentation/circular-buffers.txt
for details.
==========
REFERENCES
==========
Alpha AXP Architecture Reference Manual, Second Edition (Sites & Witek,
Digital Press)
Chapter 5.2: Physical Address Space Characteristics
Chapter 5.4: Caches and Write Buffers
Chapter 5.5: Data Sharing
Chapter 5.6: Read/Write Ordering
AMD64 Architecture Programmer's Manual Volume 2: System Programming
Chapter 7.1: Memory-Access Ordering
Chapter 7.4: Buffering and Combining Memory Writes
IA-32 Intel Architecture Software Developer's Manual, Volume 3:
System Programming Guide
Chapter 7.1: Locked Atomic Operations
Chapter 7.2: Memory Ordering
Chapter 7.4: Serializing Instructions
The SPARC Architecture Manual, Version 9
Chapter 8: Memory Models
Appendix D: Formal Specification of the Memory Models
Appendix J: Programming with the Memory Models
UltraSPARC Programmer Reference Manual
Chapter 5: Memory Accesses and Cacheability
Chapter 15: Sparc-V9 Memory Models
UltraSPARC III Cu User's Manual
Chapter 9: Memory Models
UltraSPARC IIIi Processor User's Manual
Chapter 8: Memory Models
UltraSPARC Architecture 2005
Chapter 9: Memory
Appendix D: Formal Specifications of the Memory Models
UltraSPARC T1 Supplement to the UltraSPARC Architecture 2005
Chapter 8: Memory Models
Appendix F: Caches and Cache Coherency
Solaris Internals, Core Kernel Architecture, p63-68:
Chapter 3.3: Hardware Considerations for Locks and
Synchronization
Unix Systems for Modern Architectures, Symmetric Multiprocessing and Caching
for Kernel Programmers:
Chapter 13: Other Memory Models
Intel Itanium Architecture Software Developer's Manual: Volume 1:
Section 2.6: Speculation
Section 4.4: Memory Access
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