Changeset 23f1065

Timestamp:

Apr 18, 2022, 9:48:33 AM (3 years ago)

Author:

Peter A. Buhr <pabuhr@…>

Branches:

ADT, ast-experimental, master, pthread-emulation, qualifiedEnum

Children:

75cd27b

Parents:

ac4476d

Message:

small wording changes to intro and background chapters, continue proofreading allocator chapter

Location:

doc/theses/mubeen_zulfiqar_MMath

Files:

• allocator.tex (modified) (22 diffs)
• background.tex (modified) (2 diffs)
• intro.tex (modified) (1 diff)

doc/theses/mubeen_zulfiqar_MMath/allocator.tex ¶

@@ -1 +1 @@
 \chapter{Allocator}
 
-This chapter presents a new stand-lone concurrent low-latency memory-allocator ($\approx$1,200 lines of code), called llheap (low-latency heap), for C/\CC programs using kernel threads (1:1 threading), and specialized versions of the allocator for the programming languages \uC and \CFA using user-level threads running over multiple kernel threads (M:N threading).
+This chapter presents a new stand-alone concurrent low-latency memory-allocator ($\approx$1,200 lines of code), called llheap (low-latency heap), for C/\CC programs using kernel threads (1:1 threading), and specialized versions of the allocator for the programming languages \uC and \CFA using user-level threads running over multiple kernel threads (M:N threading).
 The new allocator fulfills the GNU C Library allocator API~\cite{GNUallocAPI}.
 
@@ -12 +12 @@
 hence, llheap's design is willing to use more storage to lower latency.
 This objective is apropos because systems research and industrial applications are striving for low latency and computers have huge amounts of RAM memory.
-Finally, llheap's performance should be comparable with the current best allocators (see performance comparison in \VRef[Chapter]{Performance}).
+Finally, llheap's performance should be comparable with the current best allocators (see performance comparison in \VRef[Chapter]{c:Performance}).
 
 % The objective of llheap's new design was to fulfill following requirements:
@@ -32 +32 @@
 
 \subsection{Allocation Fastpath}
+\label{s:AllocationFastpath}
 
 These designs look at the allocation/free \newterm{fastpath}, \ie when an allocation can immediately return free storage or returned storage is not coalesced.
@@ -113 +114 @@
 Essentially, the serially-reusable problem is a race condition on an unprotected critical section, where the operating system is providing the second thread via the signal handler.
 
-\noindent
 Library @librseq@~\cite{librseq} was used to perform a fast determination of the CPU and to ensure all memory operations complete on one CPU using @librseq@'s restartable sequences, which restart the critical section after undoing its writes, if the critical section is preempted.
 \end{itemize}
@@ -122 +122 @@
 Since KT$_1$ is still executing on CPU$_1$, @librseq@ takes no action because it assumes KT$_1$ is still executing the same critical section.
 Then UT$_1$ is scheduled onto KT$_2$ by the user-level scheduler, and its memory operation continues in parallel with UT$_2$ using references into the heap associated with CPU$_1$, which corrupts CPU$_1$'s heap.
+If @librseq@ had an @rseq_abort@ which:
+\begin{enumerate}
+\item
+Marked the current restartable critical-section as cancelled so it restarts when attempting to commit.
+\item
+Do nothing if there is no current restartable critical section in progress.
+\end{enumerate}
+Then @rseq_abort@ could be called on the backside of a  user-level context-switching.
+A feature similar to this idea might exist for hardware transactional-memory.
 A significant effort was made to make this approach work but its complexity, lack of robustness, and performance costs resulted in its rejection.
-
 
 \paragraph{1:1 model}
@@ -161 +169 @@
 \vspace{5pt}
 \noindent
-The conclusion from this design exercise is: any atomic fence, instruction (lock free), or lock along the allocation fastpath produces significant slowdown.
+The conclusion from this design exercise is: any atomic fence, atomic instruction (lock free), or lock along the allocation fastpath produces significant slowdown.
 For the T:1 and T:H models, locking must exist along the allocation fastpath because the buckets or heaps maybe shared by multiple threads, even when KTs $\le$ N.
 For the T:H=CPU and 1:1 models, locking is eliminated along the allocation fastpath.
@@ -167 +175 @@
 More operating system support is required to make this model viable, but there is still the serially-reusable problem with user-level threading.
 Leaving the 1:1 model with no atomic actions along the fastpath and no special operating-system support required.
-The 1:1 model still has the serially-reusable problem with user-level threading, which is address in \VRef{}, and the greatest potential for heap blowup for certain allocation patterns.
+The 1:1 model still has the serially-reusable problem with user-level threading, which is addressed in \VRef{}, and the greatest potential for heap blowup for certain allocation patterns.
 
 
@@ -227 +235 @@
 no coalescing to minimize latency,
 \item
-local reserved memory (pool) obtained from the operating system using @sbrk@ call,
-\item
-global reserved memory (pool) obtained from the operating system using @mmap@ call to create and reuse heaps needed by threads.
+global heap memory (pool) obtained from the operating system using @mmap@ to create and reuse heaps needed by threads,
+\item
+local reserved memory (pool) per heap obtained from global pool,
+\item
+global reserved memory (pool) obtained from the operating system using @sbrk@ call,
+\item
+optional fast-lookup table for converting allocation requests into bucket sizes,
+\item
+optional statistic-counters table for accumulating counts of allocation operations.
 \end{itemize}
 
@@ -240 +254 @@
 \end{figure}
 
-llheap starts by creating an array of $N$ global heaps from storage obtained by @mmap@, where $N$ is the number of computer cores.
+llheap starts by creating an array of $N$ global heaps from storage obtained using @mmap@, where $N$ is the number of computer cores, that persists for program duration.
 There is a global bump-pointer to the next free heap in the array.
 When this array is exhausted, another array is allocated.
-There is a global top pointer to a heap intrusive link that chain free heaps from terminated threads, where these heaps are reused by new threads.
-When statistics are turned on, there is a global top pointer to a heap intrusive link that chain \emph{all} the heaps, which is traversed to accumulate statistics counters across heaps (see @malloc_stats@ \VRef{}).
-
-When a KT starts, a heap is allocated from the current array for exclusive used by the KT.
+There is a global top pointer for a heap intrusive link to chain free heaps from terminated threads.
+When statistics are turned on, there is a global top pointer for a heap intrusive link to chain \emph{all} the heaps, which is traversed to accumulate statistics counters across heaps using @malloc_stats@.
+
+When a KT starts, a heap is allocated from the current array for exclusive use by the KT.
 When a KT terminates, its heap is chained onto the heap free-list for reuse by a new KT, which prevents unbounded growth of heaps.
 The free heaps is a stack so hot storage is reused first.
-Preserving all heaps created during the program lifetime, solves the storage lifetime problem.
+Preserving all heaps created during the program lifetime, solves the storage lifetime problem, when ownership is used.
 This approach wastes storage if a large number of KTs are created/terminated at program start and then the program continues sequentially.
 llheap can be configured with object ownership, where an object is freed to the heap from which it is allocated, or object no-ownership, where an object is freed to the KT's current heap.
 
 Each heap uses segregated free-buckets that have free objects distributed across 91 different sizes from 16 to 4M.
-The number of buckets used is determined dynamically depending on the crossover point from @sbrk@ to @mmap@ allocation (see @mallopt@ \VRef{}), \ie small objects managed by the program and large objects managed by the operating system.
-Each free bucket of a specific size has following two lists:
+The number of buckets used is determined dynamically depending on the crossover point from @sbrk@ to @mmap@ allocation using @mallopt( M_MMAP_THRESHOLD )@, \ie small objects managed by the program and large objects managed by the operating system.
+Each free bucket of a specific size has the following two lists:
 \begin{itemize}
 \item
 A free stack used solely by the KT heap-owner, so push/pop operations do not require locking.
-The free objects is a stack so hot storage is reused first.
-\item
-For ownership, a shared away-stack for KTs to return storage allocated by other KTs, so push/pop operation require locking.
-The entire ownership stack be removed and become the head of the corresponding free stack, when the free stack is empty.
+The free objects are a stack so hot storage is reused first.
+\item
+For ownership, a shared away-stack for KTs to return storage allocated by other KTs, so push/pop operations require locking.
+When the free stack is empty, the entire ownership stack is removed and becomes the head of the corresponding free stack.
 \end{itemize}
 
 Algorithm~\ref{alg:heapObjectAlloc} shows the allocation outline for an object of size $S$.
 First, the allocation is divided into small (@sbrk@) or large (@mmap@).
+For large allocations, the storage is mapped directly from the operating system.
 For small allocations, $S$ is quantized into a bucket size.
-Quantizing is performed using a binary search, using the ordered bucket array.
+Quantizing is performed using a binary search over the ordered bucket array.
 An optional optimization is fast lookup $O(1)$ for sizes < 64K from a 64K array of type @char@, where each element has an index to the corresponding bucket.
 (Type @char@ restricts the number of bucket sizes to 256.)
-For $S$ > 64K, the binary search is used.
+For $S$ > 64K, a binary search is used.
 Then, the allocation storage is obtained from the following locations (in order), with increasing latency.
 \begin{enumerate}[topsep=0pt,itemsep=0pt,parsep=0pt]
@@ -286 +301 @@
 \end{enumerate}
 
-\begin{algorithm}
+\begin{figure}
+\vspace*{-10pt}
+\begin{algorithm}[H]
+\small
 \caption{Dynamic object allocation of size $S$}\label{alg:heapObjectAlloc}
 \begin{algorithmic}[1]
 \State $\textit{O} \gets \text{NULL}$
-\If {$S < \textit{mmap-threshhold}$}
+\If {$S >= \textit{mmap-threshhold}$}
+        \State $\textit{O} \gets \text{allocate dynamic memory using system call mmap with size S}$
+\Else
         \State $\textit{B} \gets \text{smallest free-bucket} \geq S$
         \If {$\textit{B's free-list is empty}$}
@@ -306 +326 @@
         \EndIf
         \State $\textit{O's owner} \gets \text{B}$
-\Else
-        \State $\textit{O} \gets \text{allocate dynamic memory using system call mmap with size S}$
 \EndIf
 \State $\Return \textit{ O}$
@@ -313 +331 @@
 \end{algorithm}
 
-Algorithm~\ref{alg:heapObjectFree} shows the de-allocation (free) outline for an object at address $A$.
-
-\begin{algorithm}[h]
-\caption{Dynamic object free at address $A$}\label{alg:heapObjectFree}
-%\begin{algorithmic}[1]
-%\State write this algorithm
-%\end{algorithmic}
+\vspace*{-15pt}
+\begin{algorithm}[H]
+\small
+\caption{Dynamic object free at address $A$ with object ownership}\label{alg:heapObjectFreeOwn}
+\begin{algorithmic}[1]
+\If {$\textit{A mapped allocation}$}
+        \State $\text{return A's dynamic memory to system using system call \lstinline{munmap}}$
+\Else
+        \State $\text{B} \gets \textit{O's owner}$
+        \If {$\textit{B is thread-local heap's bucket}$}
+                \State $\text{push A to B's free-list}$
+        \Else
+                \State $\text{push A to B's away-list}$
+        \EndIf
+\EndIf
+\end{algorithmic}
 \end{algorithm}
 
+\vspace*{-15pt}
+\begin{algorithm}[H]
+\small
+\caption{Dynamic object free at address $A$ without object ownership}\label{alg:heapObjectFreeNoOwn}
+\begin{algorithmic}[1]
+\If {$\textit{A mapped allocation}$}
+        \State $\text{return A's dynamic memory to system using system call \lstinline{munmap}}$
+\Else
+        \State $\text{B} \gets \textit{O's owner}$
+        \If {$\textit{B is thread-local heap's bucket}$}
+                \State $\text{push A to B's free-list}$
+        \Else
+                \State $\text{C} \gets \textit{thread local heap's bucket with same size as B}$
+                \State $\text{push A to C's free-list}$
+        \EndIf
+\EndIf
+\end{algorithmic}
+\end{algorithm}
+\end{figure}
+
+Algorithm~\ref{alg:heapObjectFreeOwn} shows the de-allocation (free) outline for an object at address $A$ with ownership.
+First, the address is divided into small (@sbrk@) or large (@mmap@).
+For large allocations, the storage is unmapped back to the operating system.
+For small allocations, the bucket associated with the request size is retrieved.
+If the bucket is local to the thread, the allocation is pushed onto the thread's associated bucket.
+If the bucket is not local to the thread, the allocation is pushed onto the owning thread's associated away stack.
+
+Algorithm~\ref{alg:heapObjectFreeNoOwn} shows the de-allocation (free) outline for an object at address $A$ without ownership.
+The algorithm is the same as for ownership except if the bucket is not local to the thread.
+Then the corresponding bucket of the owner thread is computed for the deallocating thread, and the allocation is pushed onto the deallocating thread's bucket.
+
+Finally, the llheap design funnels \label{p:FunnelRoutine} all allocation/deallocation operations through routines @malloc@/@free@, which are the only routines to directly access and manage the internal data structures of the heap.
+Other allocation operations, \eg @calloc@, @memalign@, and @realloc@, are composed of calls to @malloc@ and possibly @free@, and may manipulate header information after storage is allocated.
+This design simplifies heap-management code during development and maintenance.
+
+
+\subsection{Alignment}
+
+All dynamic memory allocations must have a minimum storage alignment for the contained object(s).
+Often the minimum memory alignment, M, is the bus width (32 or 64-bit) or the largest register (double, long double) or largest atomic instruction (DCAS) or vector data (MMMX).
+In general, the minimum storage alignment is 8/16-byte boundary on 32/64-bit computers.
+For consistency, the object header is normally aligned at this same boundary.
+Larger alignments must be a power of 2, such page alignment (4/8K).
+Any alignment request, N, $\le$ the minimum alignment is handled as a normal allocation with minimal alignment.
+
+For alignments greater than the minimum, the obvious approach for aligning to address @A@ is: compute the next address that is a multiple of @N@ after the current end of the heap, @E@, plus room for the header before @A@ and the size of the allocation after @A@, moving the end of the heap to @E'@.
+\begin{center}
+\input{Alignment1}
+\end{center}
+The storage between @E@ and @H@ is chained onto the appropriate free list for future allocations.
+This approach is also valid within any sufficiently large free block, where @E@ is the start of the free block, and any unused storage before @H@ or after the allocated object becomes free storage.
+In this approach, the aligned address @A@ is the same as the allocated storage address @P@, \ie @P@ $=$ @A@ for all allocation routines, which simplifies deallocation.
+However, if there are a large number of aligned requests, this approach leads to memory fragmentation from the small free areas around the aligned object.
+As well, it does not work for large allocations, where many memory allocators switch from program @sbrk@ to operating-system @mmap@.
+The reason is that @mmap@ only starts on a page boundary, and it is difficult to reuse the storage before the alignment boundary for other requests.
+Finally, this approach is incompatible with allocator designs that funnel allocation requests through @malloc@ as it directly manipulates management information within the allocator to optimize the space/time of a request.
+
+Instead, llheap alignment is accomplished by making a \emph{pessimistically} allocation request for sufficient storage to ensure that \emph{both} the alignment and size request are satisfied, \eg:
+\begin{center}
+\input{Alignment2}
+\end{center}
+The amount of storage necessary is @alignment - M + size@, which ensures there is an address, @A@, after the storage returned from @malloc@, @P@, that is a multiple of @alignment@ followed by sufficient storage for the data object.
+The approach is pessimistic because if @P@ already has the correct alignment @N@, the initial allocation has already requested sufficient space to move to the next multiple of @N@.
+For this special case, there is @alignment - M@ bytes of unused storage after the data object, which subsequently can be used by @realloc@.
+
+Note, the address returned is @A@, which is subsequently returned to @free@.
+However, to correctly free the allocated object, the value @P@ must be computable, since that is the value generated by @malloc@ and returned within @memalign@.
+Hence, there must be a mechanism to detect when @P@ $\neq$ @A@ and how to compute @P@ from @A@.
+
+The llheap approach uses two headers:
+the \emph{original} header associated with a memory allocation from @malloc@, and a \emph{fake} header within this storage before the alignment boundary @A@, which is returned from @memalign@, e.g.:
+\begin{center}
+\input{Alignment2Impl}
+\end{center}
+Since @malloc@ has a minimum alignment of @M@, @P@ $\neq$ @A@ only holds for alignments of @M@ or greater.
+When @P@ $\neq$ @A@, the minimum distance between @P@ and @A@ is @M@ bytes, due to the pessimistic storage-allocation.
+Therefore, there is always room for an @M@-byte fake header before @A@.
+
+The fake header must supply an indicator to distinguish it from a normal header and the location of address @P@ generated by @malloc@.
+This information is encoded as an offset from A to P and the initialize alignment (discussed in \VRef{s:ReallocStickyProperties}).
+To distinguish a fake header from a normal header, the least-significant bit of the alignment is used because the offset participates in multiple calculations, while the alignment is just remembered data.
+\begin{center}
+\input{FakeHeader}
+\end{center}
+
+
+\subsection{\lstinline{realloc} and Sticky Properties}
+\label{s:ReallocStickyProperties}
+
+Allocation routine @realloc@ provides a memory-management pattern for shrinking/enlarging an existing allocation, while maintaining some or all of the object data, rather than performing the following steps manually.
+\begin{flushleft}
+\begin{tabular}{ll}
+\multicolumn{1}{c}{\textbf{realloc pattern}} & \multicolumn{1}{c}{\textbf{manually}} \\
+\begin{lstlisting}
+T * naddr = realloc( oaddr, newSize );
+
+
+
+\end{lstlisting}
+&
+\begin{lstlisting}
+T * naddr = (T *)malloc( newSize ); $\C[2.4in]{// new storage}$
+memcpy( naddr, addr, oldSize );  $\C{// copy old bytes}$
+free( addr );                           $\C{// free old storage}$
+addr = naddr;                           $\C{// change pointer}\CRT$
+\end{lstlisting}
+\end{tabular}
+\end{flushleft}
+The realloc pattern leverages available storage at the end of an allocation due to bucket sizes, possibly eliminating a new allocation and copying.
+This pattern is not used enough to reduce storage management costs.
+In fact, if @oaddr@ is @nullptr@, @realloc@ does a @malloc@, so even the initial @malloc@ can be a @realloc@ for consistency in the pattern.
+
+The hidden problem for this pattern is the effect of zero fill and alignment with respect to reallocation.
+Are these properties transient or persistent (``sticky'')?
+For example, when memory is initially allocated by @calloc@ or @memalign@ with zero fill or alignment properties, respectively, what happens when those allocations are given to @realloc@ to change size.
+That is, if @realloc@ logically extends storage into unused bucket space or allocates new storage to satisfy a size change, are initial allocation properties preserve?
+Currently, allocation properties are not preserved, so subsequent use of @realloc@ storage may cause inefficient execution or errors due to lack of zero fill or alignment.
+This silent problem is unintuitive to programmers and difficult to locate because it is transient.
+To prevent these problems, llheap preserves initial allocation properties for the lifetime of an allocation and the semantics of @realloc@ are augmented to preserve these properties, with additional query routines.
+This change makes the realloc pattern efficient and safe.
+
+
+\subsection{Header}
+
+To preserve allocation properties requires storing additional information with an allocation,
+The only available location is the header, where \VRef[Figure]{f:llheapNormalHeader} shows the llheap storage layout.
+The header has two data field sized appropriately for 32/64-bit alignment requirements.
+The first field is a union of three values:
+\begin{description}
+\item[bucket pointer]
+is for allocated storage and points back to the bucket associated with this storage requests (see \VRef[Figure]{f:llheapStructure} for the fields accessible in a bucket).
+\item[mapped size]
+is for mapped storage and is the storage size for use in unmapping.
+\item[next free block]
+is for free storage and is an intrusive pointer chaining same-size free blocks onto a bucket's free stack.
+\end{description}
+The second field remembers the request size versus the allocation (bucket) size, \eg request 42 bytes which is rounded up to 64 bytes.
+Since programmers think in request sizes rather than allocation sizes, the request size allows better generation of statistics or errors.
+
+\begin{figure}
+\centering
+\input{Header}
+\caption{llheap Normal Header}
+\label{f:llheapNormalHeader}
+\end{figure}
+
+The low-order 3-bits of the first field are \emph{unused} for any stored values, whereas the second field may use all of its bits.
+The 3 unused bits are used to represent mapped allocation, zero filled, and alignment, respectively.
+Note, the alignment bit is not used in the normal header and the zero-filled/mapped bits are not used in the fake header.
+This implementation allows a fast test if any of the lower 3-bits are on (@&@ and compare).
+If no bits are on, it implies a basic allocation, which is handled quickly;
+otherwise, the bits are analysed and appropriate actions are taken for the complex cases.
+Since most allocations are basic, this implementation results in a significant performance gain along the allocation and free fastpath.
+
+
+\section{Statistics and Debugging Modes}
+
+llheap can be built to accumulate fast and largely contention-free allocation statistics to help understand allocation behaviour.
+Incrementing statistic counters must appear on the allocation fastpath.
+As noted, any atomic operation along the fastpath produces a significant increase in allocation costs.
+To make statistics performant enough for use on running systems, each heap has its own set of statistic counters, so heap operations do not require atomic operations.
+
+To locate all statistic counters, heaps are linked together in statistics mode, and this list is locked and traversed to sum all counters across heaps.
+Note, the list is locked to prevent errors traversing an active list;
+the statistics counters are not locked and can flicker during accumulation, which is not an issue with atomic read/write.
+\VRef[Figure]{f:StatiticsOutput} shows an example of statistics output, which covers all allocation operations and information about deallocating storage not owned by a thread.
+No other memory allocator studied provides as comprehensive statistical information.
+Finally, these statistics were invaluable during the development of this thesis for debugging and verifying correctness, and hence, should be equally valuable to application developers.
+
+\begin{figure}
+\begin{lstlisting}
+Heap statistics: (storage request / allocation)
+  malloc >0 calls 2,766; 0 calls 2,064; storage 12,715 / 13,367 bytes
+  aalloc >0 calls 0; 0 calls 0; storage 0 / 0 bytes
+  calloc >0 calls 6; 0 calls 0; storage 1,008 / 1,104 bytes
+  memalign >0 calls 0; 0 calls 0; storage 0 / 0 bytes
+  amemalign >0 calls 0; 0 calls 0; storage 0 / 0 bytes
+  cmemalign >0 calls 0; 0 calls 0; storage 0 / 0 bytes
+  resize >0 calls 0; 0 calls 0; storage 0 / 0 bytes
+  realloc >0 calls 0; 0 calls 0; storage 0 / 0 bytes
+  free !null calls 2,766; null calls 4,064; storage 12,715 / 13,367 bytes
+  away pulls 0; pushes 0; storage 0 / 0 bytes
+  sbrk calls 1; storage 10,485,760 bytes
+  mmap calls 10,000; storage 10,000 / 10,035 bytes
+  munmap calls 10,000; storage 10,000 / 10,035 bytes
+  threads started 4; exited 3
+  heaps new 4; reused 0
+\end{lstlisting}
+\caption{Statistics Output}
+\label{f:StatiticsOutput}
+\end{figure}
+
+llheap can also be built with debug checking, which inserts many asserts along all allocation paths.
+These assertions detect incorrect allocation usage, like double frees, unfreed storage, or memory corruptions because internal values (like header fields) are overwritten.
+These checks are best effort as opposed to complete allocation checking as in @valgrind@.
+Nevertheless, the checks detect many allocation problems.
+There is an unfortunate problem in detecting unfreed storage because some library routines assume their allocations have life-time duration, and hence, do not free their storage.
+For example, @printf@ allocates a 1024 buffer on first call and never deletes this buffer.
+To prevent a false positive for unfreed storage, it is possible to specify an amount of storage that is never freed (see \VRef{}), and it is subtracted from the total allocate/free difference.
+Determining the amount of never-freed storage is annoying, but once done, any warnings of unfreed storage are application related.
+
+Tests indicate only a 30\% performance increase when statistics \emph{and} debugging are enabled, and the latency cost for accumulating statistic is mitigated by limited calls, often only one at the end of the program.
+
+
+\section{User-level Threading Support}
+
+The serially-reusable problem (see \VRef{s:AllocationFastpath}) occurs for kernel threads in the ``T:H model, H = number of CPUs'' model and for user threads in the ``1:1'' model, where llheap uses the ``1:1'' model.
+The solution is to prevent interrupts that can result in CPU or KT change during operations that are logically critical sections.
+Locking these critical sections negates any attempt for a quick fastpath and results in high contention.
+For user-level threading, the serially-reusable problem appears with time slicing for preemptable scheduling, as the signal handler context switches to another user-level thread.
+Without time slicing, a user thread performing a long computation can prevent execution (starve) other threads.
+To prevent starvation for an allocation-active thread, \ie the time slice always triggers in an allocation critical-section for one thread, a thread-local \newterm{rollforward} flag is set in the signal handler when it aborts a time slice.
+The rollforward flag is tested at the end of each allocation funnel routine (see \VPageref{p:FunnelRoutine}), and if set, it is reset and a volunteer yield (context switch) is performed to allow other threads to execute.
+
+llheap uses two techniques to detect when execution is in a allocation operation or routine called from allocation operation, to abort any time slice during this period.
+On the slowpath when executing expensive operations, like @sbrk@ or @mmap@, interrupts are disabled/enabled by setting thread-local flags so the signal handler aborts immediately.
+On the fastpath, disabling/enabling interrupts is too expensive as accessing thread-local storage can be expensive and not thread-safe.
+For example, the ARM processor stores the thread-local pointer in a coprocessor register that cannot perform atomic base-displacement addressing.
+Hence, there is a window between loading the thread-local pointer from the coprocessor register into a normal register and adding the displacement when a time slice can move a thread.
+
+The fast technique defines a special code section and places all non-interruptible routines in this section.
+The linker places all code in this section into a contiguous block of memory, but the order of routines within the block is unspecified.
+Then, the signal handler compares the program counter at the point of interrupt with the the start and end address of the non-interruptible section, and aborts if executing within this section and sets the rollforward flag.
+This technique is fragile because any calls in the non-interruptible code outside of the non-interruptible section (like @sbrk@) must be bracketed with disable/enable interrupts and these calls must be along the slowpath.
+Hence, for correctness, this approach requires inspection of generated assembler code for routines placed in the non-interruptible section.
+This issue is mitigated by the llheap funnel design so only funnel routines and a few statistics routines are placed in the non-interruptible section and their assembler code examined.
+These techniques are used in both the \uC and \CFA versions of llheap, where both of these systems have user-level threading.
+
+
+\section{Bootstrapping}
+
+There are problems bootstrapping a memory allocator.
+\begin{enumerate}
+\item
+Programs can be statically or dynamically linked.
+\item
+The order the linker schedules startup code is poorly supported.
+\item
+Knowing a KT's start and end independently from the KT code is difficult.
+\end{enumerate}
+
+For static linking, the allocator is loaded with the program.
+Hence, allocation calls immediately invoke the allocator operation defined by the loaded allocation library and there is only one memory allocator used in the program.
+This approach allows allocator substitution by placing an allocation library before any other in the linked/load path.
+
+Allocator substitution is similar for dynamic linking, but the problem is that the dynamic loader starts first and needs to perform dynamic allocations \emph{before} the substitution allocator is loaded.
+As a result, the dynamic loader uses a default allocator until the substitution allocator is loaded, after which all allocation operations are handled by the substitution allocator, including from the dynamic loader.
+Hence, some part of the @sbrk@ area may be used by the default allocator and statistics about allocation operations cannot be correct.
+Furthermore, dynamic linking goes through trampolines, so there is an additional cost along the allocator fastpath for all allocation operations.
+Testing showed up to a 5\% performance increase for dynamic linking over static linking, even when using @tls_model("initial-exec")@ so the dynamic loader can obtain tighter binding.
+
+All allocator libraries need to perform startup code to initialize data structures, such as the heap array for llheap.
+The problem is getting initialized done before the first allocator call.
+However, there does not seem to be mechanism to tell either the static or dynamic loader to first perform initialization code before any calls to a loaded library.
+As a result, calls to allocation routines occur without initialization.
+To deal with this problem, it is necessary to put a conditional initialization check along the allocation fastpath to trigger initialization (singleton pattern).
+
+Two other important execution points are program startup and termination, which include prologue or epilogue code to bootstrap a program, which programmers are unaware of.
+For example, dynamic-memory allocations before/after the application starts should not be considered in statistics because the application does not make these calls.
+llheap establishes these two points using routines:
+\begin{lstlisting}
+__attribute__(( constructor( 100 ) )) static void startup( void ) {
+        // clear statistic counters
+        // reset allocUnfreed counter
+}
+__attribute__(( destructor( 100 ) )) static void shutdown( void ) {
+        // sum allocUnfreed for all heaps
+        // subtract global unfreed storage
+        // if allocUnfreed > 0 then print warning message
+}
+\end{lstlisting}
+which use global constructor/destructor priority 100, where the linker calls these routines at program prologue/epilogue in increasing/decreasing order of priority.
+Application programs may only use global constructor/destructor priorities greater than 100.
+Hence, @startup@ is called after the program prologue but before the application starts, and @shutdown@ is called after the program terminates but before the program epilogue.
+By resetting counters in @startup@, prologue allocations are ignored, and checking unfreed storage in @shutdown@ checks only application memory management, ignoring the program epilogue.
+
+While @startup@/@shutdown@ apply to the program KT, a concurrent program creates additional KTs that do not trigger these routines.
+However, it is essential for the allocator to know when each KT is started/terminated.
+One approach is to create a thread-local object with a construct/destructor, which is triggered after a new KT starts and before it terminates, respectively.
+\begin{lstlisting}
+struct ThreadManager {
+        volatile bool pgm_thread;
+        ThreadManager() {} // unusable
+        ~ThreadManager() { if ( pgm_thread ) heapManagerDtor(); }
+};
+static thread_local ThreadManager threadManager;
+\end{lstlisting}
+Unfortunately, thread-local variables are created lazily, \ie on the first dereference of @threadManager@, which then triggers its constructor.
+Therefore, the constructor is useless for knowing when a KT starts because the KT must reference it, and the allocator does not control the application KT.
+Fortunately, the singleton pattern needed for initializing the program KT also triggers KT allocator initialization, which can then reference @pgm_thread@ to call @threadManager@'s constructor, otherwise its destructor is not called.
+Now when a KT terminates, @~ThreadManager@ is called to chained it onto the global-heap free-stack, where @pgm_thread@ is set to true only for the program KT.
+The conditional destructor call prevents closing down the program heap, which must remain available because epilogue code may free more storage.
+
+Finally, there is a recursive problem when the singleton pattern dereferences @pgm_thread@ to initialize the thread-local object, because its initialization calls @atExit@, which immediately calls @malloc@ to obtain storage.
+This recursion is handled with another thread-local flag to prevent double initialization.
+A similar problem exists when the KT terminates and calls member @~ThreadManager@, because immediately afterwards, the terminating KT calls @free@ to deallocate the storage obtained from the @atExit@.
+In the meantime, the terminated heap has been put on the global-heap free-stack, and may be active by a new KT, so the @atExit@ free is handled as a free to another heap and put onto the away list using locking.
+
+For user threading systems, the KTs are controlled by the runtime, and hence, start/end pointers are known and interact directly with the llheap allocator for \uC and \CFA, which eliminates or simplifies several of these problems.
+The following API was created to provide interaction between the language runtime and the allocator.
+\begin{lstlisting}
+void startTask();                       $\C{// KT starts}$
+void finishTask();                      $\C{// KT ends}$
+void startup();                         $\C{// when application code starts}$
+void shutdown();                        $\C{// when application code ends}$
+bool traceHeap();                       $\C{// enable allocation/free printing for debugging}$
+bool traceHeapOn();                     $\C{// start printing allocation/free calls}$
+bool traceHeapOff();                    $\C{// stop printing allocation/free calls}$
+\end{lstlisting}
+This kind of API is necessary to allow concurrent runtime systems to interact with difference memory allocators in a consistent way.
 
 %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%
 
 \section{Added Features and Methods}
-To improve the llheap allocator (FIX ME: cite llheap) interface and make it more user friendly, we added a few more routines to the C allocator.
-Also, we built a \CFA (FIX ME: cite cforall) interface on top of C interface to increase the usability of the allocator.
-
-\subsection{C Interface}
-We added a few more features and routines to the allocator's C interface that can make the allocator more usable to the programmers.
-These features will programmer more control on the dynamic memory allocation.
+
+The C dynamic-allocation API (see \VRef[Figure]{f:CDynamicAllocationAPI}) is neither orthogonal nor complete.
+For example,
+\begin{itemize}
+\item
+It is possible to zero fill or align an allocation but not both.
+\item
+It is \emph{only} possible to zero fill and array allocation.
+\item
+It is not possible to resize a memory allocation without data copying.
+\item
+@realloc@ does not preserve initial allocation properties.
+\end{itemize}
+As a result, programmers must provide these options, which is error prone, resulting in blaming the entire programming language for a poor dynamic-allocation API.
+Furthermore, newer programming languages have better type systems that can provide safer and more powerful APIs for memory allocation.
+
+\begin{figure}
+\begin{lstlisting}
+void * malloc( size_t size );
+void * calloc( size_t nmemb, size_t size );
+void * realloc( void * ptr, size_t size );
+void * reallocarray( void * ptr, size_t nmemb, size_t size );
+void free( void * ptr );
+void * memalign( size_t alignment, size_t size );
+void * valloc( size_t size );
+void * pvalloc( size_t size );
+struct mallinfo mallinfo( void );
+int mallopt( int param, int val );
+int malloc_trim( size_t pad );
+size_t malloc_usable_size( void * ptr );
+void malloc_stats( void );
+int malloc_info( int options, FILE * fp );
+\end{lstlisting}
+\caption{C Dynamic-Allocation API}
+\label{f:CDynamicAllocationAPI}
+\end{figure}
+
+The following presents design and API changes for C, \CC (\uC), and \CFA, all of which are implemented in llheap.
+
 
 \subsection{Out of Memory}
@@ -339 +709 @@
 The alternative is to abort a program when out of memory.
 In theory, notifying the programmer allows recovery;
-in practice, it is almost impossible to gracefully when out of memory, so the cheaper approach of returning @nullptr@ for a zero-sized allocation is chosen.
-
-
-\subsection{\lstinline{void * aalloc( size_t dim, size_t elemSize )}}
+in practice, it is almost impossible to gracefully recover when out of memory, so the cheaper approach of returning @nullptr@ for a zero-sized allocation is chosen for llheap.
+
+
+\subsection{C Interface}
+
+Within the C type-system, it is still possible to increase orthogonality and functionality of the dynamic-memory API to make the allocator more usable for programmers.
+
+\paragraph{\lstinline{void * aalloc( size_t dim, size_t elemSize )}}
 @aalloc@ is an extension of malloc.
 It allows programmer to allocate a dynamic array of objects without calculating the total size of array explicitly.
@@ -358 +732 @@
 On failure, it returns a @NULL@ pointer.
 
-\subsection{\lstinline{void * resize( void * oaddr, size_t size )}}
+\paragraph{\lstinline{void * resize( void * oaddr, size_t size )}}
 @resize@ is an extension of relloc.
 It allows programmer to reuse a currently allocated dynamic object with a new size requirement.
@@ -374 +748 @@
 On failure, it returns a @NULL@ pointer.
 
-\subsection{\lstinline{void * resize( void * oaddr, size_t nalign, size_t size )}}
+\paragraph{\lstinline{void * resize( void * oaddr, size_t nalign, size_t size )}}
 This @resize@ is an extension of the above @resize@ (FIX ME: cite above resize).
 In addition to resizing the size of of an old object, it can also realign the old object to a new alignment requirement.
@@ -392 +766 @@
 On failure, it returns a @NULL@ pointer.
 
-\subsection{\lstinline{void * amemalign( size_t alignment, size_t dim, size_t elemSize )}}
+\paragraph{\lstinline{void * amemalign( size_t alignment, size_t dim, size_t elemSize )}}
 amemalign is a hybrid of memalign and aalloc.
 It allows programmer to allocate an aligned dynamic array of objects without calculating the total size of the array explicitly.
@@ -411 +785 @@
 On failure, it returns a @NULL@ pointer.
 
-\subsection{\lstinline{void * cmemalign( size_t alignment, size_t dim, size_t elemSize )}}
+\paragraph{\lstinline{void * cmemalign( size_t alignment, size_t dim, size_t elemSize )}}
 cmemalign is a hybrid of amemalign and calloc.
 It allows programmer to allocate an aligned dynamic array of objects that is 0 filled.
@@ -431 +805 @@
 On failure, it returns a @NULL@ pointer.
 
-\subsection{\lstinline{size_t malloc_alignment( void * addr )}}
+\paragraph{\lstinline{size_t malloc_alignment( void * addr )}}
 @malloc_alignment@ returns the alignment of a currently allocated dynamic object.
 It allows the programmer in memory management and personal bookkeeping.
@@ -445 +819 @@
 On failure, it return the value of default alignment of the llheap allocator.
 
-\subsection{\lstinline{bool malloc_zero_fill( void * addr )}}
+\paragraph{\lstinline{bool malloc_zero_fill( void * addr )}}
 @malloc_zero_fill@ returns whether a currently allocated dynamic object was initially zero filled at the time of allocation.
 It allows the programmer in memory management and personal bookkeeping.
@@ -459 +833 @@
 On failure, it returns false.
 
-\subsection{\lstinline{size_t malloc_size( void * addr )}}
-@malloc_size@ returns the allocation size of a currently allocated dynamic object.
+\paragraph{\lstinline{size_t malloc_size( void * addr )}}
+@malloc_size@ returns the request size of a currently allocated dynamic object.
 It allows the programmer in memory management and personal bookkeeping.
 It helps the programmer in verifying the alignment of a dynamic object especially in a scenario similar to producer-consumer where a producer allocates a dynamic object and the consumer needs to assure that the dynamic object was allocated with the required size.
@@ -474 +848 @@
 @addr@: the address of the currently allocated dynamic object.
 \end{itemize}
-@malloc_size@ returns the allocation size of the given dynamic object.
+@malloc_size@ returns the request size of the given dynamic object.
 On failure, it return zero.
 
-\subsection{\lstinline{void * realloc( void * oaddr, size_t nalign, size_t size )}}
+
+\subsection{\CC Interface}
+
+\paragraph{\lstinline{void * realloc( void * oaddr, size_t nalign, size_t size )}}
 This @realloc@ is an extension of the default @realloc@ (FIX ME: cite default @realloc@).
 In addition to reallocating an old object and preserving the data in old object, it can also realign the old object to a new alignment requirement.
@@ -495 +872 @@
 On failure, it returns a @NULL@ pointer.
 
-\subsection{\CFA Malloc Interface}
+
+\subsection{\CFA Interface}
 We added some routines to the @malloc@ interface of \CFA.
 These routines can only be used in \CFA and not in our stand-alone llheap allocator as these routines use some features that are only provided by \CFA and not by C.

doc/theses/mubeen_zulfiqar_MMath/background.tex ¶

@@ -36 +36 @@
 The management data starts with fixed-sized information in the static-data memory that references components in the dynamic-allocation memory.
 The \newterm{storage data} is composed of allocated and freed objects, and \newterm{reserved memory}.
-Allocated objects (white) are variable sized, and allocated and maintained by the program;
+Allocated objects (light grey) are variable sized, and allocated and maintained by the program;
 \ie only the program knows the location of allocated storage, not the memory allocator.
 \begin{figure}[h]
@@ -44 +44 @@
 \label{f:AllocatorComponents}
 \end{figure}
-Freed objects (light grey) represent memory deallocated by the program, which are linked into one or more lists facilitating easy location of new allocations.
+Freed objects (white) represent memory deallocated by the program, which are linked into one or more lists facilitating easy location of new allocations.
 Often the free list is chained internally so it does not consume additional storage, \ie the link fields are placed at known locations in the unused memory blocks.
 Reserved memory (dark grey) is one or more blocks of memory obtained from the operating system but not yet allocated to the program;

doc/theses/mubeen_zulfiqar_MMath/intro.tex ¶

@@ -124 +124 @@
 
 \item
-Provide complete, fast, and contention-free allocation statistics to help understand program behaviour:
+Provide complete, fast, and contention-free allocation statistics to help understand allocation behaviour:
 \begin{itemize}
 \item