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<TITLE> Conservative GC Algorithmic Overview </TITLE> | |
<AUTHOR> Hans-J. Boehm, HP Labs (Some of this was written at SGI)</author> | |
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<H1> <I>This is under construction, and may always be.</i> </h1> | |
<H1> Conservative GC Algorithmic Overview </h1> | |
<P> | |
This is a description of the algorithms and data structures used in our | |
conservative garbage collector. I expect the level of detail to increase | |
with time. For a survey of GC algorithms, see for example | |
<A HREF="ftp://ftp.cs.utexas.edu/pub/garbage/gcsurvey.ps"> Paul Wilson's | |
excellent paper</a>. For an overview of the collector interface, | |
see <A HREF="gcinterface.html">here</a>. | |
<P> | |
This description is targeted primarily at someone trying to understand the | |
source code. It specifically refers to variable and function names. | |
It may also be useful for understanding the algorithms at a higher level. | |
<P> | |
The description here assumes that the collector is used in default mode. | |
In particular, we assume that it used as a garbage collector, and not just | |
a leak detector. We initially assume that it is used in stop-the-world, | |
non-incremental mode, though the presence of the incremental collector | |
will be apparent in the design. | |
We assume the default finalization model, but the code affected by that | |
is very localized. | |
<H2> Introduction </h2> | |
The garbage collector uses a modified mark-sweep algorithm. Conceptually | |
it operates roughly in four phases, which are performed occasionally | |
as part of a memory allocation: | |
<OL> | |
<LI> | |
<I>Preparation</i> Each object has an associated mark bit. | |
Clear all mark bits, indicating that all objects | |
are potentially unreachable. | |
<LI> | |
<I>Mark phase</i> Marks all objects that can be reachable via chains of | |
pointers from variables. Often the collector has no real information | |
about the location of pointer variables in the heap, so it | |
views all static data areas, stacks and registers as potentially containing | |
pointers. Any bit patterns that represent addresses inside | |
heap objects managed by the collector are viewed as pointers. | |
Unless the client program has made heap object layout information | |
available to the collector, any heap objects found to be reachable from | |
variables are again scanned similarly. | |
<LI> | |
<I>Sweep phase</i> Scans the heap for inaccessible, and hence unmarked, | |
objects, and returns them to an appropriate free list for reuse. This is | |
not really a separate phase; even in non incremental mode this is operation | |
is usually performed on demand during an allocation that discovers an empty | |
free list. Thus the sweep phase is very unlikely to touch a page that | |
would not have been touched shortly thereafter anyway. | |
<LI> | |
<I>Finalization phase</i> Unreachable objects which had been registered | |
for finalization are enqueued for finalization outside the collector. | |
</ol> | |
<P> | |
The remaining sections describe the memory allocation data structures, | |
and then the last 3 collection phases in more detail. We conclude by | |
outlining some of the additional features implemented in the collector. | |
<H2>Allocation</h2> | |
The collector includes its own memory allocator. The allocator obtains | |
memory from the system in a platform-dependent way. Under UNIX, it | |
uses either <TT>malloc</tt>, <TT>sbrk</tt>, or <TT>mmap</tt>. | |
<P> | |
Most static data used by the allocator, as well as that needed by the | |
rest of the garbage collector is stored inside the | |
<TT>_GC_arrays</tt> structure. | |
This allows the garbage collector to easily ignore the collectors own | |
data structures when it searches for root pointers. Other allocator | |
and collector internal data structures are allocated dynamically | |
with <TT>GC_scratch_alloc</tt>. <TT>GC_scratch_alloc</tt> does not | |
allow for deallocation, and is therefore used only for permanent data | |
structures. | |
<P> | |
The allocator allocates objects of different <I>kinds</i>. | |
Different kinds are handled somewhat differently by certain parts | |
of the garbage collector. Certain kinds are scanned for pointers, | |
others are not. Some may have per-object type descriptors that | |
determine pointer locations. Or a specific kind may correspond | |
to one specific object layout. Two built-in kinds are uncollectable. | |
One (<TT>STUBBORN</tt>) is immutable without special precautions. | |
In spite of that, it is very likely that most C clients of the | |
collector currently | |
use at most two kinds: <TT>NORMAL</tt> and <TT>PTRFREE</tt> objects. | |
The <A HREF="http://gcc.gnu.org/java">gcj</a> runtime also makes | |
heavy use of a kind (allocated with GC_gcj_malloc) that stores | |
type information at a known offset in method tables. | |
<P> | |
The collector uses a two level allocator. A large block is defined to | |
be one larger than half of <TT>HBLKSIZE</tt>, which is a power of 2, | |
typically on the order of the page size. | |
<P> | |
Large block sizes are rounded up to | |
the next multiple of <TT>HBLKSIZE</tt> and then allocated by | |
<TT>GC_allochblk</tt>. Recent versions of the collector | |
use an approximate best fit algorithm by keeping free lists for | |
several large block sizes. | |
The actual | |
implementation of <TT>GC_allochblk</tt> | |
is significantly complicated by black-listing issues | |
(see below). | |
<P> | |
Small blocks are allocated in chunks of size <TT>HBLKSIZE</tt>. | |
Each chunk is | |
dedicated to only one object size and kind. The allocator maintains | |
separate free lists for each size and kind of object. | |
<P> | |
Once a large block is split for use in smaller objects, it can only | |
be used for objects of that size, unless the collector discovers a completely | |
empty chunk. Completely empty chunks are restored to the appropriate | |
large block free list. | |
<P> | |
In order to avoid allocating blocks for too many distinct object sizes, | |
the collector normally does not directly allocate objects of every possible | |
request size. Instead request are rounded up to one of a smaller number | |
of allocated sizes, for which free lists are maintained. The exact | |
allocated sizes are computed on demand, but subject to the constraint | |
that they increase roughly in geometric progression. Thus objects | |
requested early in the execution are likely to be allocated with exactly | |
the requested size, subject to alignment constraints. | |
See <TT>GC_init_size_map</tt> for details. | |
<P> | |
The actual size rounding operation during small object allocation is | |
implemented as a table lookup in <TT>GC_size_map</tt>. | |
<P> | |
Both collector initialization and computation of allocated sizes are | |
handled carefully so that they do not slow down the small object fast | |
allocation path. An attempt to allocate before the collector is initialized, | |
or before the appropriate <TT>GC_size_map</tt> entry is computed, | |
will take the same path as an allocation attempt with an empty free list. | |
This results in a call to the slow path code (<TT>GC_generic_malloc_inner</tt>) | |
which performs the appropriate initialization checks. | |
<P> | |
In non-incremental mode, we make a decision about whether to garbage collect | |
whenever an allocation would otherwise have failed with the current heap size. | |
If the total amount of allocation since the last collection is less than | |
the heap size divided by <TT>GC_free_space_divisor</tt>, we try to | |
expand the heap. Otherwise, we initiate a garbage collection. This ensures | |
that the amount of garbage collection work per allocated byte remains | |
constant. | |
<P> | |
The above is in fact an oversimplification of the real heap expansion | |
and GC triggering heuristic, which adjusts slightly for root size | |
and certain kinds of | |
fragmentation. In particular: | |
<UL> | |
<LI> Programs with a large root set size and | |
little live heap memory will expand the heap to amortize the cost of | |
scanning the roots. | |
<LI> Versions 5.x of the collector actually collect more frequently in | |
nonincremental mode. The large block allocator usually refuses to split | |
large heap blocks once the garbage collection threshold is | |
reached. This often has the effect of collecting well before the | |
heap fills up, thus reducing fragmentation and working set size at the | |
expense of GC time. Versions 6.x choose an intermediate strategy depending | |
on how much large object allocation has taken place in the past. | |
(If the collector is configured to unmap unused pages, versions 6.x | |
use the 5.x strategy.) | |
<LI> In calculating the amount of allocation since the last collection we | |
give partial credit for objects we expect to be explicitly deallocated. | |
Even if all objects are explicitly managed, it is often desirable to collect | |
on rare occasion, since that is our only mechanism for coalescing completely | |
empty chunks. | |
</ul> | |
<P> | |
It has been suggested that this should be adjusted so that we favor | |
expansion if the resulting heap still fits into physical memory. | |
In many cases, that would no doubt help. But it is tricky to do this | |
in a way that remains robust if multiple application are contending | |
for a single pool of physical memory. | |
<H2>Mark phase</h2> | |
At each collection, the collector marks all objects that are | |
possibly reachable from pointer variables. Since it cannot generally | |
tell where pointer variables are located, it scans the following | |
<I>root segments</i> for pointers: | |
<UL> | |
<LI>The registers. Depending on the architecture, this may be done using | |
assembly code, or by calling a <TT>setjmp</tt>-like function which saves | |
register contents on the stack. | |
<LI>The stack(s). In the case of a single-threaded application, | |
on most platforms this | |
is done by scanning the memory between (an approximation of) the current | |
stack pointer and <TT>GC_stackbottom</tt>. (For Itanium, the register stack | |
scanned separately.) The <TT>GC_stackbottom</tt> variable is set in | |
a highly platform-specific way depending on the appropriate configuration | |
information in <TT>gcconfig.h</tt>. Note that the currently active | |
stack needs to be scanned carefully, since callee-save registers of | |
client code may appear inside collector stack frames, which may | |
change during the mark process. This is addressed by scanning | |
some sections of the stack "eagerly", effectively capturing a snapshot | |
at one point in time. | |
<LI>Static data region(s). In the simplest case, this is the region | |
between <TT>DATASTART</tt> and <TT>DATAEND</tt>, as defined in | |
<TT>gcconfig.h</tt>. However, in most cases, this will also involve | |
static data regions associated with dynamic libraries. These are | |
identified by the mostly platform-specific code in <TT>dyn_load.c</tt>. | |
</ul> | |
The marker maintains an explicit stack of memory regions that are known | |
to be accessible, but that have not yet been searched for contained pointers. | |
Each stack entry contains the starting address of the block to be scanned, | |
as well as a descriptor of the block. If no layout information is | |
available for the block, then the descriptor is simply a length. | |
(For other possibilities, see <TT>gc_mark.h</tt>.) | |
<P> | |
At the beginning of the mark phase, all root segments | |
(as described above) are pushed on the | |
stack by <TT>GC_push_roots</tt>. (Registers and eagerly processed | |
stack sections are processed by pushing the referenced objects instead | |
of the stack section itself.) If <TT>ALL_INTERIOR_PTRS</tt> is not | |
defined, then stack roots require special treatment. In this case, the | |
normal marking code ignores interior pointers, but <TT>GC_push_all_stack</tt> | |
explicitly checks for interior pointers and pushes descriptors for target | |
objects. | |
<P> | |
The marker is structured to allow incremental marking. | |
Each call to <TT>GC_mark_some</tt> performs a small amount of | |
work towards marking the heap. | |
It maintains | |
explicit state in the form of <TT>GC_mark_state</tt>, which | |
identifies a particular sub-phase. Some other pieces of state, most | |
notably the mark stack, identify how much work remains to be done | |
in each sub-phase. The normal progression of mark states for | |
a stop-the-world collection is: | |
<OL> | |
<LI> <TT>MS_INVALID</tt> indicating that there may be accessible unmarked | |
objects. In this case <TT>GC_objects_are_marked</tt> will simultaneously | |
be false, so the mark state is advanced to | |
<LI> <TT>MS_PUSH_UNCOLLECTABLE</tt> indicating that it suffices to push | |
uncollectable objects, roots, and then mark everything reachable from them. | |
<TT>Scan_ptr</tt> is advanced through the heap until all uncollectable | |
objects are pushed, and objects reachable from them are marked. | |
At that point, the next call to <TT>GC_mark_some</tt> calls | |
<TT>GC_push_roots</tt> to push the roots. It the advances the | |
mark state to | |
<LI> <TT>MS_ROOTS_PUSHED</tt> asserting that once the mark stack is | |
empty, all reachable objects are marked. Once in this state, we work | |
only on emptying the mark stack. Once this is completed, the state | |
changes to | |
<LI> <TT>MS_NONE</tt> indicating that reachable objects are marked. | |
</ol> | |
The core mark routine <TT>GC_mark_from</tt>, is called | |
repeatedly by several of the sub-phases when the mark stack starts to fill | |
up. It is also called repeatedly in <TT>MS_ROOTS_PUSHED</tt> state | |
to empty the mark stack. | |
The routine is designed to only perform a limited amount of marking at | |
each call, so that it can also be used by the incremental collector. | |
It is fairly carefully tuned, since it usually consumes a large majority | |
of the garbage collection time. | |
<P> | |
The fact that it perform a only a small amount of work per call also | |
allows it to be used as the core routine of the parallel marker. In that | |
case it is normally invoked on thread-private mark stacks instead of the | |
global mark stack. More details can be found in | |
<A HREF="scale.html">scale.html</a> | |
<P> | |
The marker correctly handles mark stack overflows. Whenever the mark stack | |
overflows, the mark state is reset to <TT>MS_INVALID</tt>. | |
Since there are already marked objects in the heap, | |
this eventually forces a complete | |
scan of the heap, searching for pointers, during which any unmarked objects | |
referenced by marked objects are again pushed on the mark stack. This | |
process is repeated until the mark phase completes without a stack overflow. | |
Each time the stack overflows, an attempt is made to grow the mark stack. | |
All pieces of the collector that push regions onto the mark stack have to be | |
careful to ensure forward progress, even in case of repeated mark stack | |
overflows. Every mark attempt results in additional marked objects. | |
<P> | |
Each mark stack entry is processed by examining all candidate pointers | |
in the range described by the entry. If the region has no associated | |
type information, then this typically requires that each 4-byte aligned | |
quantity (8-byte aligned with 64-bit pointers) be considered a candidate | |
pointer. | |
<P> | |
We determine whether a candidate pointer is actually the address of | |
a heap block. This is done in the following steps: | |
<NL> | |
<LI> The candidate pointer is checked against rough heap bounds. | |
These heap bounds are maintained such that all actual heap objects | |
fall between them. In order to facilitate black-listing (see below) | |
we also include address regions that the heap is likely to expand into. | |
Most non-pointers fail this initial test. | |
<LI> The candidate pointer is divided into two pieces; the most significant | |
bits identify a <TT>HBLKSIZE</tt>-sized page in the address space, and | |
the least significant bits specify an offset within that page. | |
(A hardware page may actually consist of multiple such pages. | |
HBLKSIZE is usually the page size divided by a small power of two.) | |
<LI> | |
The page address part of the candidate pointer is looked up in a | |
<A HREF="tree.html">table</a>. | |
Each table entry contains either 0, indicating that the page is not part | |
of the garbage collected heap, a small integer <I>n</i>, indicating | |
that the page is part of large object, starting at least <I>n</i> pages | |
back, or a pointer to a descriptor for the page. In the first case, | |
the candidate pointer i not a true pointer and can be safely ignored. | |
In the last two cases, we can obtain a descriptor for the page containing | |
the beginning of the object. | |
<LI> | |
The starting address of the referenced object is computed. | |
The page descriptor contains the size of the object(s) | |
in that page, the object kind, and the necessary mark bits for those | |
objects. The size information can be used to map the candidate pointer | |
to the object starting address. To accelerate this process, the page header | |
also contains a pointer to a precomputed map of page offsets to displacements | |
from the beginning of an object. The use of this map avoids a | |
potentially slow integer remainder operation in computing the object | |
start address. | |
<LI> | |
The mark bit for the target object is checked and set. If the object | |
was previously unmarked, the object is pushed on the mark stack. | |
The descriptor is read from the page descriptor. (This is computed | |
from information <TT>GC_obj_kinds</tt> when the page is first allocated.) | |
</nl> | |
<P> | |
At the end of the mark phase, mark bits for left-over free lists are cleared, | |
in case a free list was accidentally marked due to a stray pointer. | |
<H2>Sweep phase</h2> | |
At the end of the mark phase, all blocks in the heap are examined. | |
Unmarked large objects are immediately returned to the large object free list. | |
Each small object page is checked to see if all mark bits are clear. | |
If so, the entire page is returned to the large object free list. | |
Small object pages containing some reachable object are queued for later | |
sweeping, unless we determine that the page contains very little free | |
space, in which case it is not examined further. | |
<P> | |
This initial sweep pass touches only block headers, not | |
the blocks themselves. Thus it does not require significant paging, even | |
if large sections of the heap are not in physical memory. | |
<P> | |
Nonempty small object pages are swept when an allocation attempt | |
encounters an empty free list for that object size and kind. | |
Pages for the correct size and kind are repeatedly swept until at | |
least one empty block is found. Sweeping such a page involves | |
scanning the mark bit array in the page header, and building a free | |
list linked through the first words in the objects themselves. | |
This does involve touching the appropriate data page, but in most cases | |
it will be touched only just before it is used for allocation. | |
Hence any paging is essentially unavoidable. | |
<P> | |
Except in the case of pointer-free objects, we maintain the invariant | |
that any object in a small object free list is cleared (except possibly | |
for the link field). Thus it becomes the burden of the small object | |
sweep routine to clear objects. This has the advantage that we can | |
easily recover from accidentally marking a free list, though that could | |
also be handled by other means. The collector currently spends a fair | |
amount of time clearing objects, and this approach should probably be | |
revisited. | |
<P> | |
In most configurations, we use specialized sweep routines to handle common | |
small object sizes. Since we allocate one mark bit per word, it becomes | |
easier to examine the relevant mark bits if the object size divides | |
the word length evenly. We also suitably unroll the inner sweep loop | |
in each case. (It is conceivable that profile-based procedure cloning | |
in the compiler could make this unnecessary and counterproductive. I | |
know of no existing compiler to which this applies.) | |
<P> | |
The sweeping of small object pages could be avoided completely at the expense | |
of examining mark bits directly in the allocator. This would probably | |
be more expensive, since each allocation call would have to reload | |
a large amount of state (e.g. next object address to be swept, position | |
in mark bit table) before it could do its work. The current scheme | |
keeps the allocator simple and allows useful optimizations in the sweeper. | |
<H2>Finalization</h2> | |
Both <TT>GC_register_disappearing_link</tt> and | |
<TT>GC_register_finalizer</tt> add the request to a corresponding hash | |
table. The hash table is allocated out of collected memory, but | |
the reference to the finalizable object is hidden from the collector. | |
Currently finalization requests are processed non-incrementally at the | |
end of a mark cycle. | |
<P> | |
The collector makes an initial pass over the table of finalizable objects, | |
pushing the contents of unmarked objects onto the mark stack. | |
After pushing each object, the marker is invoked to mark all objects | |
reachable from it. The object itself is not explicitly marked. | |
This assures that objects on which a finalizer depends are neither | |
collected nor finalized. | |
<P> | |
If in the process of marking from an object the | |
object itself becomes marked, we have uncovered | |
a cycle involving the object. This usually results in a warning from the | |
collector. Such objects are not finalized, since it may be | |
unsafe to do so. See the more detailed | |
<A HREF="http://www.hpl.hp.com/personal/Hans_Boehm/gc/finalization.html"> discussion of finalization semantics</a>. | |
<P> | |
Any objects remaining unmarked at the end of this process are added to | |
a queue of objects whose finalizers can be run. Depending on collector | |
configuration, finalizers are dequeued and run either implicitly during | |
allocation calls, or explicitly in response to a user request. | |
(Note that the former is unfortunately both the default and not generally safe. | |
If finalizers perform synchronization, it may result in deadlocks. | |
Nontrivial finalizers generally need to perform synchronization, and | |
thus require a different collector configuration.) | |
<P> | |
The collector provides a mechanism for replacing the procedure that is | |
used to mark through objects. This is used both to provide support for | |
Java-style unordered finalization, and to ignore certain kinds of cycles, | |
<I>e.g.</i> those arising from C++ implementations of virtual inheritance. | |
<H2>Generational Collection and Dirty Bits</h2> | |
We basically use the concurrent and generational GC algorithm described in | |
<A HREF="http://www.hpl.hp.com/personal/Hans_Boehm/gc/papers/pldi91.ps.Z">"Mostly Parallel Garbage Collection"</a>, | |
by Boehm, Demers, and Shenker. | |
<P> | |
The most significant modification is that | |
the collector always starts running in the allocating thread. | |
There is no separate garbage collector thread. (If parallel GC is | |
enabled, helper threads may also be woken up.) | |
If an allocation attempt either requests a large object, or encounters | |
an empty small object free list, and notices that there is a collection | |
in progress, it immediately performs a small amount of marking work | |
as described above. | |
<P> | |
This change was made both because we wanted to easily accommodate | |
single-threaded environments, and because a separate GC thread requires | |
very careful control over the scheduler to prevent the mutator from | |
out-running the collector, and hence provoking unneeded heap growth. | |
<P> | |
In incremental mode, the heap is always expanded when we encounter | |
insufficient space for an allocation. Garbage collection is triggered | |
whenever we notice that more than | |
<TT>GC_heap_size</tt>/2 * <TT>GC_free_space_divisor</tt> | |
bytes of allocation have taken place. | |
After <TT>GC_full_freq</tt> minor collections a major collection | |
is started. | |
<P> | |
All collections initially run interrupted until a predetermined | |
amount of time (50 msecs by default) has expired. If this allows | |
the collection to complete entirely, we can avoid correcting | |
for data structure modifications during the collection. If it does | |
not complete, we return control to the mutator, and perform small | |
amounts of additional GC work during those later allocations that | |
cannot be satisfied from small object free lists. When marking completes, | |
the set of modified pages is retrieved, and we mark once again from | |
marked objects on those pages, this time with the mutator stopped. | |
<P> | |
We keep track of modified pages using one of several distinct mechanisms: | |
<OL> | |
<LI> | |
Through explicit mutator cooperation. Currently this requires | |
the use of <TT>GC_malloc_stubborn</tt>, and is rarely used. | |
<LI> | |
(<TT>MPROTECT_VDB</tt>) By write-protecting physical pages and | |
catching write faults. This is | |
implemented for many Unix-like systems and for win32. It is not possible | |
in a few environments. | |
<LI> | |
(<TT>PROC_VDB</tt>) By retrieving dirty bit information from /proc. | |
(Currently only Sun's | |
Solaris supports this. Though this is considerably cleaner, performance | |
may actually be better with mprotect and signals.) | |
<LI> | |
(<TT>PCR_VDB</tt>) By relying on an external dirty bit implementation, in this | |
case the one in Xerox PCR. | |
<LI> | |
(<TT>DEFAULT_VDB</tt>) By treating all pages as dirty. This is the default if | |
none of the other techniques is known to be usable, and | |
<TT>GC_malloc_stubborn</tt> is not used. Practical only for testing, or if | |
the vast majority of objects use <TT>GC_malloc_stubborn</tt>. | |
</ol> | |
<H2>Black-listing</h2> | |
The collector implements <I>black-listing</i> of pages, as described | |
in | |
<A HREF="http://www.acm.org/pubs/citations/proceedings/pldi/155090/p197-boehm/"> | |
Boehm, ``Space Efficient Conservative Collection'', PLDI '93</a>, also available | |
<A HREF="papers/pldi93.ps.Z">here</a>. | |
<P> | |
During the mark phase, the collector tracks ``near misses'', i.e. attempts | |
to follow a ``pointer'' to just outside the garbage-collected heap, or | |
to a currently unallocated page inside the heap. Pages that have been | |
the targets of such near misses are likely to be the targets of | |
misidentified ``pointers'' in the future. To minimize the future | |
damage caused by such misidentifications they will be allocated only to | |
small pointerfree objects. | |
<P> | |
The collector understands two different kinds of black-listing. A | |
page may be black listed for interior pointer references | |
(<TT>GC_add_to_black_list_stack</tt>), if it was the target of a near | |
miss from a location that requires interior pointer recognition, | |
<I>e.g.</i> the stack, or the heap if <TT>GC_all_interior_pointers</tt> | |
is set. In this case, we also avoid allocating large blocks that include | |
this page. | |
<P> | |
If the near miss came from a source that did not require interior | |
pointer recognition, it is black-listed with | |
<TT>GC_add_to_black_list_normal</tt>. | |
A page black-listed in this way may appear inside a large object, | |
so long as it is not the first page of a large object. | |
<P> | |
The <TT>GC_allochblk</tt> routine respects black-listing when assigning | |
a block to a particular object kind and size. It occasionally | |
drops (i.e. allocates and forgets) blocks that are completely black-listed | |
in order to avoid excessively long large block free lists containing | |
only unusable blocks. This would otherwise become an issue | |
if there is low demand for small pointerfree objects. | |
<H2>Thread support</h2> | |
We support several different threading models. Unfortunately Pthreads, | |
the only reasonably well standardized thread model, supports too narrow | |
an interface for conservative garbage collection. There appears to be | |
no completely portable way to allow the collector | |
to coexist with various Pthreads | |
implementations. Hence we currently support only the more | |
common Pthreads implementations. | |
<P> | |
In particular, it is very difficult for the collector to stop all other | |
threads in the system and examine the register contents. This is currently | |
accomplished with very different mechanisms for some Pthreads | |
implementations. The Solaris implementation temporarily disables much | |
of the user-level threads implementation by stopping kernel-level threads | |
("lwp"s). The Linux/HPUX/OSF1 and Irix implementations sends signals to | |
individual Pthreads and has them wait in the signal handler. | |
<P> | |
The Linux and Irix implementations use | |
only documented Pthreads calls, but rely on extensions to their semantics. | |
The Linux implementation <TT>linux_threads.c</tt> relies on only very | |
mild extensions to the pthreads semantics, and already supports a large number | |
of other Unix-like pthreads implementations. Our goal is to make this the | |
only pthread support in the collector. | |
<P> | |
(The Irix implementation is separate only for historical reasons and should | |
clearly be merged. The current Solaris implementation probably performs | |
better in the uniprocessor case, but does not support thread operations in the | |
collector. Hence it cannot support the parallel marker.) | |
<P> | |
All implementations must | |
intercept thread creation and a few other thread-specific calls to allow | |
enumeration of threads and location of thread stacks. This is current | |
accomplished with <TT># define</tt>'s in <TT>gc.h</tt> | |
(really <TT>gc_pthread_redirects.h</tt>), or optionally | |
by using ld's function call wrapping mechanism under Linux. | |
<P> | |
Recent versions of the collector support several facilites to enhance | |
the processor-scalability and thread performance of the collector. | |
These are discussed in more detail <A HREF="scale.html">here</a>. | |
We briefly outline the data approach to thread-local allocation in the | |
next section. | |
<H2>Thread-local allocation</h2> | |
If thread-local allocation is enabled, the collector keeps separate | |
arrays of free lists for each thread. Thread-local allocation | |
is currently only supported on a few platforms. | |
<P> | |
The free list arrays associated | |
with each thread are only used to satisfy requests for objects that | |
are both very small, and belong to one of a small number of well-known | |
kinds. These currently include "normal" and pointer-free objects. | |
Depending onthe configuration, "gcj" objects may also be included. | |
<P> | |
Thread-local free list entries contain either a pointer to the first | |
element of a free list, or they contain a counter of the number of | |
allocation "granules" allocated so far. Initially they contain the | |
value one, i.e. a small counter value. | |
<P> | |
Thread-local allocation allocates directly through the global | |
allocator, if the object is of a size or kind not covered by the | |
local free lists. | |
<P> | |
If there is an appropriate local free list, the allocator checks whether it | |
contains a sufficiently small counter value. If so, the counter is simply | |
incremented by the counter value, and the global allocator is used. | |
In this way, the initial few allocations of a given size bypass the local | |
allocator. A thread that only allocates a handful of objects of a given | |
size will not build up its own free list for that size. This avoids | |
wasting space for unpopular objects sizes or kinds. | |
<P> | |
Once the counter passes a threshold, <TT>GC_malloc_many</tt> is called | |
to allocate roughly <TT>HBLKSIZE</tt> space and put it on the corresponding | |
local free list. Further allocations of that size and kind then use | |
this free list, and no longer need to acquire the allocation lock. | |
The allocation procedure is otherwise similar to the global free lists. | |
The local free lists are also linked using the first word in the object. | |
In most cases this means they require considerably less time. | |
<P> | |
Local free lists are treated buy most of the rest of the collector | |
as though they were in-use reachable data. This requires some care, | |
since pointer-free objects are not normally traced, and hence a special | |
tracing procedure is required to mark all objects on pointer-free and | |
gcj local free lists. | |
<P> | |
On thread exit, any remaining thread-local free list entries are | |
transferred back to the global free list. | |
<P> | |
Note that if the collector is configured for thread-local allocation, | |
GC versions before 7 do not invoke the thread-local allocator by default. | |
<TT>GC_malloc</tt> only uses thread-local allocation in version 7 and later. | |
In earlier versions, <TT>GC_MALLOC</tt> (all caps) may be directed | |
to use thread-local allocation by defining <TT>GC_REDIRECT_TO_LOCAL</tt> | |
and then include <TT>gc_local_alloc.h</tt>. | |
<P> | |
For some more details see <A HREF="scale.html">here</a>, and the | |
technical report entitled | |
<A HREF="http://www.hpl.hp.com/techreports/2000/HPL-2000-165.html"> | |
``Fast Multiprocessor Memory Allocation and Garbage Collection'' | |
</a> | |
<P> | |
<HR> | |
<P> | |
Comments are appreciated. Please send mail to | |
<A HREF="mailto:boehm@acm.org"><TT>boehm@acm.org</tt></a> or | |
<A HREF="mailto:Hans.Boehm@hp.com"><TT>Hans.Boehm@hp.com</tt></a> | |
<P> | |
This is a modified copy of a page written while the author was at SGI. | |
The original was <A HREF="http://reality.sgi.com/boehm/gcdescr.html">here</a>. | |
</body> | |
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