![Davidlohr Bueso](/assets/img/avatar_default.png)
Signed-off-by: Davidlohr Bueso <dbueso@suse.de> Signed-off-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
3049 lines
110 KiB
Plaintext
3049 lines
110 KiB
Plaintext
============================
|
|
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 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
|
|
in the following example:
|
|
|
|
q = ACCESS_ONCE(a);
|
|
if (q) {
|
|
ACCESS_ONCE(b) = p;
|
|
}
|
|
|
|
Please note that ACCESS_ONCE() is not optional! Without the
|
|
ACCESS_ONCE(), 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!!! */
|
|
|
|
So don't leave out the ACCESS_ONCE().
|
|
|
|
It is tempting to try to enforce ordering on identical stores on both
|
|
branches of the "if" statement as follows:
|
|
|
|
q = ACCESS_ONCE(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 = ACCESS_ONCE(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 = ACCESS_ONCE(a);
|
|
if (q) {
|
|
ACCESS_ONCE(b) = p;
|
|
do_something();
|
|
} else {
|
|
ACCESS_ONCE(b) = r;
|
|
do_something_else();
|
|
}
|
|
|
|
The initial ACCESS_ONCE() 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 = ACCESS_ONCE(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 = ACCESS_ONCE(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 = ACCESS_ONCE(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 = ACCESS_ONCE(a);
|
|
if (a || 1 > 0)
|
|
ACCESS_ONCE(b) = 1;
|
|
|
|
Because the second condition is always true, the compiler can transform
|
|
this example as following, defeating control dependency:
|
|
|
|
q = ACCESS_ONCE(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 = ACCESS_ONCE(x); r2 = ACCESS_ONCE(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.
|
|
|
|
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 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 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, an acquire barrier, a release barrier,
|
|
a read barrier, or a general barrier. Similarly a read barrier 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;
|
|
|
|
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:
|
|
|
|
(*) set_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 he 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 loads against
|
|
subsequent loads and stores and also orders prior stores against
|
|
subsequent 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();
|
|
set_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
|
|
set_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
|