This fixes the tails=true build (at least for x86_64) and eliminates
the need for a frame table in the tails=false build. In the
tails=true build, we still need a frame table on x86(_64) to help
determine whether we've caught a thread executing code to do a tail
call or pop arguments off the stack. However, I've not yet written
the code to actually use this table, and it is only needed to handle
asynchronous unwinds via Thread.getStackTrace.
This is necessary to accomodate classes loaded at runtime which refer
to primitive array types. Otherwise, they won't be included unless
classes in the bootimage refer to them.
When loading a class which extends another class that contained a
field of primitive array type using defineClass in a bootimage=true
build, the VM was unable to find the primitive array class, and
makeArrayClass refused to create one since it should already have
existed.
The problem was that the bootimage=true build uses an empty
Machine::BootstrapClassMap, and resolveArrayClass expected to find the
primitive array classes there. The fix is to check the
Machine::BootLoader map if we can't find it in
Machine::BootstrapClassMap.
Previously, we unwound the stack by following the chain of frame
pointers for normal returns, stack trace creation, and exception
unwinding. On x86, this required reserving EBP/RBP for frame pointer
duties, making it unavailable for general computation and requiring
that it be explicitly saved and restored on entry and exit,
respectively.
On PowerPC, we use an ABI that makes the stack pointer double as a
frame pointer, so it doesn't cost us anything. We've been using the
same convention on ARM, but it doesn't match the native calling
convention, which makes it unusable when we want to call native code
from Java and pass arguments on the stack.
So far, the ARM calling convention mismatch hasn't been an issue
because we've never passed more arguments from Java to native code
than would fit in registers. However, we must now pass an extra
argument (the thread pointer) to e.g. divideLong so it can throw an
exception on divide by zero, which means the last argument must be
passed on the stack. This will clobber the linkage area we've been
using to hold the frame pointer, so we need to stop using it.
One solution would be to use the same convention on ARM as we do on
x86, but this would introduce the same overhead of making a register
unavailable for general use and extra code at method entry and exit.
Instead, this commit removes the need for a frame pointer. Unwinding
involves consulting a map of instruction offsets to frame sizes which
is generated at compile time. This is necessary because stack trace
creation can happen at any time due to Thread.getStackTrace being
called by another thread, and the frame size varies during the
execution of a method.
So far, only x86(_64) is working, and continuations and tail call
optimization are probably broken. More to come.
This rather large commit modifies the VM to use non-local returns to
throw exceptions instead of simply setting Thread::exception and
returning frame-by-frame as it used to. This has several benefits:
* Functions no longer need to check Thread::exception after each call
which might throw an exception (which would be especially tedious
and error-prone now that any function which allocates objects
directly or indirectly might throw an OutOfMemoryError)
* There's no need to audit the code for calls to functions which
previously did not throw exceptions but later do
* Performance should be improved slightly due to both the reduced
need for conditionals and because undwinding now occurs in a single
jump instead of a series of returns
The main disadvantages are:
* Slightly higher overhead for entering and leaving the VM via the
JNI and JDK methods
* Non-local returns can make the code harder to read
* We must be careful to register destructors for stack-allocated
resources with the Thread so they can be called prior to a
non-local return
The non-local return implementation is similar to setjmp/longjmp,
except it uses continuation-passing style to avoid the need for
cooperation from the C/C++ compiler. Native C++ exceptions would have
also been an option, but that would introduce a dependence on
libstdc++, which we're trying to avoid for portability reasons.
Finally, this commit ensures that the VM throws an OutOfMemoryError
instead of aborting when it reaches its memory ceiling. Currently, we
treat the ceiling as a soft limit and temporarily exceed it as
necessary to allow garbage collection and certain internal allocations
to succeed, but refuse to allocate any Java objects until the heap
size drops back below the ceiling.
There is a delay between when we tell the OS to start a thread and
when it actually starts, and during that time a thread might
mistakenly think it was the last to exit, try to shut down the VM, and
then block in joinAll when it finds it wasn't the last one after all.
The solution is to increment Machine::liveCount and add the new thread
to the process tree before starting it -- all while holding
Machine::stateLock for atomicity. This helps guarantee that when
liveCount is one, we can be sure there's really only one thread
running or staged to run.
If we don't do this, the VM will crash when it tries to create a stack
trace for the error because makeObjectArray will return null
immediately when it sees there is a pending exception.
GCC 4.5.1 and later use a naming convention where functions are not
prefixed with an underscore, whereas previous versions added the
underscore. This change was made to ensure compatibility with
Microsoft's compiler. Since GCC 4.5.0 has a serious code generation
bug, we now only support later versions, so it makes sense to assume
the newer convention.
We now check for stack overflow in the JIT build as well as the
interpreted build, throwing a StackOverflowError if the limit
(currently hard-coded to 64KB, but should be easy to make
configurable) is exceeded.
There was an unlikely but dangerous race condition in monitorRelease
such that when a thread released a monitor and then tried to notify
the next thread in line, the latter thread might exit before it can be
notified. This potentially led to a crash as the former thread tried
to acquire and notify the latter thread's private lock after it had
been disposed.
The solution is to do as we do in the interrupt and join cases: call
acquireSystem first and thereby either block the target thread from
exiting until we're done or find that it has already exited, in which
case nothing needs to be done.
I also looked at monitorNotify to see if we have a similar bug there,
but in that case the target thread can't exit without first acquiring
and releasing the monitor, and since we ensure that no thread can
execute monitorNotify without holding the monitor, there's no
potential for a race.
In makeCodeImage, we were passing zero to Promise::Listener::resolve,
which would lead to an assertion error if the address of the code
image was further from the base of the address space (i.e. zero) than
could be spanned by a jump on the target architecture. Since, in this
context, we immediately overwrite the value stored, we may pass
whatever we want to this function (we're only calling it so we can
retrieve the location of the value in the image), and the code image
pointer is a better choice for the above reason.
When trying to create an array class, we try to resolve
java.lang.Object so we can use its vtable in the array class.
However, if Object is missing, we'll try to create and throw a
ClassNotFoundException, which requires creating an array to store the
stack trace, which requires creating an array class, which requires
resolving Object, etc.. This commit short-circuits this process by
telling resolveClass not to create and throw an exception if it can't
find Object.
While doing the above work, I noticed that the implementations of
Classpath::makeThrowable in classpath-avian.cpp and
classpath-openjdk.cpp were identical, so I made makeThrowable a
top-level function.
Finally, I discovered that Thread.setDaemon can only be called before
the target thread has been started, which allowed me to simplify the
code to track daemon threads in the VM.
We have to be careful about how we calculate return addresses on ARM
due to padding introduced by constant pools interspersed with code.
When calculating the offset of code where we're inserting a constant
pool, we want the offset of the end of the pool for jump targets, but
we want the offset just prior to the beginning of the pool (i.e. the
offset of the instruction responsible for jumping past the pool) when
calculating a return address.
The code added to runJavaThread was unecessary and harmful since it
allowed the global daemon thread count to become permanently
out-of-sync with the actual number of daemon threads.
This mainly involves some makefile ugliness to work around bugs in the
native Windows OpenJDK code involving conflicting static and
not-static declarations which GCC 4.0 and later justifiably reject but
MSVC tolerates.
We weren't properly handling the case where a 64-bit value is
multipled with itself in multiplyRR, leading to wrong code. Also,
addCarryCR didn't know how to handle constants more than 8-bits wide.
* add libnet.so and libnio.so to built-in libraries for openjdk-src build
* implement sun.misc.Unsafe.park/unpark
* implement JVM_SetClassSigners/JVM_GetClassSigners
* etc.
The main change here is to use a lazily-populated vector to associate
runtime data with classes instead of referencing them directly from
the class which requires updating immutable references in the heap
image. The other changes employ other strategies to avoid trying to
update immutable references.
Compiling the entire OpenJDK class library into a bootimage revealed
some corner cases which broke the compiler, including synchronization
in a finally block and gotos targeting the first instruction of an
unsynchronized method.
We must call notifyAll on visitLock after setting threadVisitor to
null in case another thread is waiting to do a visit of its own.
Otherwise, the latter thread will wait forever, eventually deadlocking
the whole VM at the next GC since it's in an active state.
If the VM runs out of heap space and the "avian.heap.dump" system
property was specified at startup, the VM will write a heap dump to
the filename indicated by that property. This dump may be analyzed
using e.g. DumpStats.java.
My recent commit to ensure that OS resources are released immediately
upon thread exit introduced a race condition where interrupting or
joining a thread as it exited could lead to attempts to use
already-released resources. This commit adds locking to avoid the
race.
This makes heap dumps more useful since these classes are now refered
to by name instead of number.
This commit also adds a couple of utilities for parsing heap dumps:
PrintDump and DumpStats.
The primary change is to ensure we output a Mach-O file of appropriate
endianness when cross-compiling for an opposite-endian architecture.
Earlier versions of XCode's linker accepted files of either
endianness, reguardless of architecture, but later versions don't,
hence the change.
Previously, loading an arbitrary 32-bit constant required up to four
instructions (128 bytes), since we did so one byte at a time via
immediate-mode operations.
The preferred way to load constants on ARM is via PC-relative
addressing, but this is challenging because immediate memory offsets
are limited to 4096 bytes in either direction. We frequently need to
compile methods which are larger than 4096, or even 8192, bytes, so we
must intersperse code and data if we want to use PC-relative loads
everywhere.
This commit enables pervasive PC-relative loads by handling the
following cases:
1. Method is shorter than 4096 bytes: append data table to end
2. Method is longer than 4096 bytes, but no basic block is longer
than 4096 bytes: insert data tables as necessary after blocks, taking
care to minimize the total number of tables
3. Method is longer than 4096 bytes, and some blocks are longer than
4096 bytes: split large basic blocks and insert data tables as above
Previously, we waited until the next GC to do this, but that can be
too long for workloads which create a lot of short-lived threads but
don't do much allocation.
This requires adding LinkRegister to the list of reserved registers,
since it must be preserved in the thunk code generated by
compileDirectInvoke. An alternative would be to explicitly preserve
it in that special case, but that would complicate the code quite a
bit.
All the tests are passing for openjdk-src builds, but the non-src
openjdk build is crashing and there's trouble loading time zone info
from the embedded java.home directory.
This allows OpenJDK to access time zone data which is normally found
under java.home, but which we must embed in the executable itself to
create a self-contained build. The VM intercepts various file
operations, looking for paths which start with a prefix specified by
the avian.embed.prefix property and redirecting those operations to an
embedded JAR.
For example, if avian.embed.prefix is "/avian-embedded", and code
calls File.exists() with a path of
"/avian-embedded/javahomeJar/foo.txt", the VM looks for a function
named javahomeJar via dlsym, calls the function to find the memory
region containing the embeded JAR, and finally consults the JAR to see
if the file "foo.txt" exists.
sun.misc.Unsafe.getUnsafe expects a null result if the class loader is
the boot classloader and will throw a SecurityException otherwise
(whereas it should really be checking both for null and comparing
against the system classloader). However, just returning null
whenever the loader is the boot loader can cause trouble for embedded
apps which put everything in the boot loader, including application
resources.
Therefore, we only return null if it's the boot loader and we're being
called from Unsafe.getUnsafe.
As described in readme.txt, a standalone OpenJDK build embeds all
libraries, classes, and other files needed at runtime in the resulting
binary, eliminating dependencies on external resources.
Rather than try to support mixing Avian's core classes with those of
an external class library -- which necessitates adding a lot of stub
methods which throw UnsupportedOperationExceptions, among other
comprimises -- we're looking to support such external class libraries
in their unmodified forms. The latter strategy has already proven
successful with OpenJDK's class library. Thus, this commit removes
the stub methods, etc., which not only cleans up the code but avoids
misleading application developers as to what classes and methods
Avian's built-in class library supports.
We now consult the JAVA_HOME environment variable to determine where
to find the system library JARs and SOs. Ultimately, we'll want to
support self-contained build, but this allows Avian to behave like a
conventional libjvm.so.
The main changes in this commit ensure that we don't hold the global
class lock when doing class resolution using application-defined
classloaders. Such classloaders may do their own locking (in fact,
it's almost certain), making deadlock likely when mixed with VM-level
locking in various orders.
Other changes include a fix to avoid overflow when waiting for
extremely long intervals and a GC root stack mapping bug.
The biggest change in this commit is to split the system classloader
into two: one for boot classes (e.g. java.lang.*) and another for
application classes. This is necessary to make OpenJDK's security
checks happy.
The rest of the changes include bugfixes and additional JVM method
implementations in classpath-openjdk.cpp.
Whereas the GNU Classpath port used the strategy of patching Classpath
with core classes from Avian so as to minimize changes to the VM, this
port uses the opposite strategy: abstract and isolate
classpath-specific features in the VM similar to how we abstract away
platform-specific features in system.h. This allows us to use an
unmodified copy of OpenJDK's class library, including its core classes
and augmented by a few VM-specific classes in the "avian" package.
We've been getting away with not doing this so far since our Java
calling convention matches the native calling convention concerning
where the return address is saved, so when our thunk calls native code
it gets saved for us automatically. However, there was still the
danger that a thread would interrupt another thread after the stack
pointer was saved to the thread field but before the native code was
called and try to get a stack trace, at which point it would try to
find the return address relative to that stack pointer and find
garbage instead. This commit ensures that we save the return address
before saving the stack pointer to avoid such a situation.
In order to facilitate making the VM compatible with multiple class
libraries, it's useful to separate the VM-specific representation of
these classes from the library implementations. This commit
introduces VMClass, VMField, and VMMethod for that purpose.
A long time ago, I refactored the class initialization code in the VM,
but did not notice until today that it had caused the
process=interpret build to break on certain recursive initializations.
In particular, we were not always detecting when a thread recursively
tried to initialize a class it was already in the process of
initializing, leading to the mistaken assumption that another thread
was initializing it and that we should wait until it was done, in
which case we would wait forever.
This commit ensures that we always detect recursive initialization and
short-circuit it.
The shiftLeftC function in powerpc.cpp was miscompiling such shifts,
leading to crashes due to illegal instructions and other weirdness due
to instructions that meant something completely different. This
commit fixes that and adds a test to Longs.java to make sure it stays
fixed.
Previously, we risked segfaults by passing negative numbers to memcpy.
This commit also makes arraycopy throw an IndexOutOfBounds exception
instead of an ArrayStoreException if the specified offsets and lengths
would take us outside the bounds of one or both of the arrays, per the
Sun documentation.
If we catch the target thread in a virtual thunk when getting its
stack trace, we must assume its Thread::stack field is garbage and use
the register values instead. Previously, we treated these thunks as
any other native code, leading to crashes when we tried to use the
garbage pointer.
32MB was just slightly too large for PowerPC immediate call instructions
to span, and 16MB matches the JIT executable memory area we use in
compile.cpp.
compileDirectInvoke does some magic to optimize tail calls to native
methods which involves storing the return address (which we'll never
actually return to, since it's a tail call) in a thread-local field so
the thunk function can figure out which native method to look up at
runtime. Since this address will change when the boot image is
loaded, the boot image creation code needs to know about it.
callContinuation failed to call the correct continuation when feeding
it an exception due to a regression introduced with the
Thread.getStackTrace changes.
The new Thread::defaultHeap declaration has increased the offset of all
the fields following it.
This commit also makes vmInvoke_returnAddress global so it can be refered
to from compile.cpp.
It's not safe to use malloc from a signal handler, so we can't
allocate new memory when handling segfaults or Thread.getStackTrace
signals. Instead, we allocate a fixed-size backup heap for each
thread ahead of time and use it if there's no space left in the normal
heap pool. In the rare case that the backup heap isn't large enough,
we fall back to using a preallocated exception without a stack trace
as a last resort.
This function was broken in two different ways:
1. It only checked MyProcessor::thunks, not MyProcessor::bootThunks.
It needs to check both.
2. When checking MyProcessor::thunks, it used fields from
MyProcessor::bootThunks instead of from the same thunk collection.
This fixes both problems.
Implementing Thread.getStackTrace is tricky. A thread may interrupt
another thread at any time to grab a stack trace, including while the
latter is executing Java code, JNI code, helper thunks, VM code, or
while transitioning between any of these.
To create a stack trace we use several context fields associated with
the target thread, including snapshots of the instruction pointer,
stack pointer, and frame pointer. These fields must be current,
accurate, and consistent with each other in order to get a reliable
trace. Otherwise, we risk crashing the VM by trying to walk garbage
stack frames or by misinterpreting the size and/or content of
legitimate frames.
This commit addresses sensitive transition points such as entering the
helper thunks which bridge the transitions from Java to native code
(where we must save the stack and frame registers for use from native
code) and stack unwinding (where we must atomically update the thread
context fields to indicate which frame we are unwinding to). When
grabbing a trace for another thread, we determine what kind of code we
caught the thread executing in and use that information to choose the
thread context values with which to begin the trace. See
MyProcessor::getStackTrace::Visitor::visit for details.
In order to atomically update the thread context fields, we do the
following:
1. Create a temporary "transition" object to serve as a staging area
and populate it with the new field values.
2. Update a transition pointer in the thread object to point to the
object created above. As long as this pointer is non-null,
interrupting threads will use the context values in the staging
object instead of those in the thread object.
3. Update the fields in the thread object.
4. Clear the transition pointer in the thread object.
We use a memory barrier between each of these steps to ensure they are
made visible to other threads in program order. See
MyThread::doTransition for details.
In Mac OS X, if a path contains a space, the path of the main executable
will contain a special URL-encoded character (%20 in this case). This
probably happens when any non-ASCII character is provided.
The fix is to use CFURLCreateStringByReplacingPercentEscapes which
creates a path that the POSIX API likes better.
We were generating code which clobbered the data we were putting into
64-bit volatile fields (and potentially also clobbering the target or
source object in the case of non-static fields) due to misplaced
synchronization code. Reordering this code ensures that both the data
and the target or source survive across calls to synchronization
helper functions.
Previously, the stack frame mapping code (responsible for statically
calculating the map of GC roots for a method's stack frame during JIT
compilation) would assume that the map of GC roots on entry to an
exception handler is the same as on entry to the "try" block which the
handler is attached to. Technically, this is true, but the algorithm
we use does not consider whether a local variable is still "live"
(i.e. will be read later) when calculating the map - only whether we
can expect to find a reference there via normal (non-exceptional)
control flow. This can backfire if, within a "try" block, the stack
location which held an object reference on entry to the block gets
overwritten with a non-reference (i.e. a primitive). If an exception
is later thrown from such a block, we might end up trying to treat
that non-reference as a reference during GC, which will crash the VM.
The ideal way to fix this is to calculate the true interval for which
each value is live and use that to produce the stack frame maps. This
would provide the added benefit of ensuring that the garbage collector
does not visit references which, although still present on the stack,
will not be used again.
However, this commit uses the less invasive strategy of ANDing
together the root maps at each GC point within a "try" block and using
the result as the map on entry to the corresponding exception
handler(s). This should give us safe, if not optimal, results. Later
on, we can refine it as described above.
See commit 8120bee4dc for the original
problem description and solution. That commit and a couple of related
ones had to be reverted when we found they had introduced GC-safety
regressions leading to crashes.
This commit restores the reverted code and fixes the regressions.
We're seeing race conditions which occasionally lead to assertion
failures and thus crashes, so I'm reverting these changes for now:
29309fb414e92674cb738120bee4dc
We don't want to check Thread::waiting until we have re-acquired the
monitor, since another thread might notify us between releasing
Thread::lock and acquiring the monitor.
We need to prefix instructions of the form "mov R,M" with a REX byte
when R is %spl, %bpl, %sil, or %dil. Such moves are unencodable on
32-bit x86, and, because of the order in which we pick registers,
pretty rare on 64-bit systems, which is why this took so long to
notice.
Due to SWT's nasty habit of creating a new object monitor for every
task added to Display.asyncExec, we've found that, on Windows at
least, we tend to run out of OS handles due to the large number of
mutexes we create between garbage collections.
One way to address this might be to trigger a GC when either the
number of monitors created since the last GC exceeds a certain number
or when the total number of monitors in the VM reaches a certain
number. Both of these risk hurting performance, especially if they
force major collections which would otherwise be infrequent. Also,
it's hard to know what the values of such thresholds should be on a
given system.
Instead, we reimplement Java monitors using atomic compare-and-swap
(CAS) and thread-specific native locks for blocking in the case of
contention. This way, we can create an arbitrary number of monitors
without creating any new native locks. The total number of native
locks needed by the VM is bounded instead by the number of live
threads plus a small constant.
Note that if we ever add support for an architecture which does not
support CAS, we'll need to provide a fallback monitor implementation.
We were miscompiling methods which contained getfield, getstatic,
putfield, or putstatic instructions for volatile 64-bit primitives on
32-bit PowerPC due to not noticing that values in registers are clobbered
across function calls.
The solution is to create a separate Compiler::Operand instance for each
object monitor reference before and after the function call to avoid
confusing the compiler. To avoid duplicate entries in the constant pool,
we add code look for and, if found, reuse any existing entry for the same
constant.
Currently, we just set this to /tmp (or the equivalent) since Avian
doesn't really have a home. This avoids a NullPointerException from
javax/xml/parsers/SAXParserFactory.
The latter is cheaper (avoids a state transition and possible memory
allocation) when we just want to know if an exception is thrown
without needing a handle to that exception.
Before allocating a new reference in NewGlobalReference or when
creating a local reference, we look for a previously-allocated
reference pointing to the same object. This is a linear search, but
usually the number of elements in the reference list is small, whereas
the memory, locking, and allocation overhead of creating duplicate
references can be large.
We need to check to see if we caught the thread somewhere in the thunk
code (i.e. about to call a helper function), in which case the stack
and base pointers are valid and may be used to create an accurate
trace.
If another thread succeeds in entering the "exclusive" state while we
use the fast path to transition the current thread to "active", we
must switch back to "idle" temporarily to allow the exclusive thread a
chance to continue, and then retry the transition to "active" via the
slow path.
These paths reduce contention among threads by using atomic operations
and memory barriers instead of mutexes where possible. This is
especially important for JNI calls, since each such call involves two
state transitions: from "active" to "idle" and back.
Previously, we assumed that the "context" parameter to
GetThreadContext was only an output parameter, but it actually uses at
the value of CONTEXT::ContextFlags on entry to decide what parts of
the structure to fill in. We were getting lucking most of the time,
because whatever garbage was on the stack at that location had the
necessary bits set. When we weren't so lucky, we got all zeros for
the register values which sometimes lead to a crash depending on the
state of the thread being examined.
Previously, we only visited frame locations containing values, but
this invited the possibility of reusing the same site for two
locations in some cases.
These warnings are due to GCC being smart enough to do interprocedural
constant propagation but not smart enough to avoid false positives in
all cases when looking for array bounds errors.
We use this utility instead of objcopy to embed data into object files
because it offers more control over e.g. section alignment, which is
important for bootimage builds.
This is necessary because objcopy does not currently allow us to
specify the alignment requirement for the .boot section used to store
the boot image for AOT builds. This may be a problem for Windows as
well, in which case we'll need to add a binaryToPE utility.
uniqueSite also checks, if applicable, to see if the second word of a
value shares the specified site with the first value as its sole site.
Also renamed a couple of fields in Value for clarity.
This allows the assembler to see the operand types of the comparison
and the condition for jumping in the same operation, which is
essential for generating efficient code in cases such as
multiple-precision compare-and-branch.
Previously, we simply removed the element from the sites array, but
this led to problems when the junction sites are shared among multiple
junctions such that the value at a given index is live at one junction
and dead at another.
This implementation does not conform to the Java standard in that
finalize methods are called from whichever thread happens to be garbage
collecting, and that thread may hold locks, whereas the standard
guarantees that finalize will be run from a thread which holds no locks.
Also, an object will never be finalized more than once, even if its
finalize method "rescues" (i.e. makes reachable) the object such that it
might become unreachable a second time and thus a candidate for
finalization once more. It's not clear to me from the standard if this
is OK or not.
Nonwithstanding the above, this implementation is useful for "normal"
finalize methods which simply release resources associated with an
object.
The previous code relied on the invalid assumption that the thread-local
heaps for all threads would have been cleared immediately following a
garbage collection. However, the last thing the garbage collection
function does is run finalizers which may allocate new objects. This
can lead allocate3 to call allocateSmall with a size which is too large
to accomodate, overflowing the heap.
The solution is to iterate until there really is enough room for the
original allocation request.
We now use SingleRead::successor in pickTarget, where we use it to
determine the prefered target site for the successor without requiring
the target to conform to that preference. The previous code made the
preference a hard requirement, which is not desirable or even possible
in general.
The SingleRead::successor field is used (when non-null) to further
constrain the SiteMask in SingleRead::intersect based on reads of
successor values (as in the cases of moves and condensed-addressing
combine and translate instructions).
We now create a unique thunk for each vtable position so as to avoid
relying on using the return address to determine what method is to be
compiled and invoked, since we will not have the correct return address
in the case of a tail call. This required refactoring how executable
memory is allocated in order to keep AOT compilation working. Also, we
must always use the same register to hold the class pointer when
compiling virtual calls, and ensure that the pointer stays there until
the call instruction is executed so we know where to find it in the
thunk.
This helps us support the Java Memory Model without adding a memory
barrier to every object allocation. It's also potentially more
efficient, since we zero out each heap segment all at once instead of
bit-by-bit with each object allocation.
This is important on the 32-bit OS X PowerPC ABI, since the location
of the low 32-bits of a return value change depending on whether the
entire value is 64-bits or not.
bootimage.cpp
This is useful for debugging the compiler, since it allows you to
compile one method in particular if that's where a bug manifests itself.
This mirrors how snapshots are captured and restored for branch events.
Also, switch back to resolving junction sites before capturing fork
snapshots so we capture the true state of each value immediately before
the branch.
This forces the compiler to explicitly move the constant operand into
a register, avoiding the possibility that AX or DX might be unavailable
when the assembler needs them.
