- Timestamp:
- Mar 9, 2020, 11:09:41 AM (4 years ago)
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- ADT, arm-eh, ast-experimental, enum, forall-pointer-decay, jacob/cs343-translation, jenkins-sandbox, master, new-ast, new-ast-unique-expr, pthread-emulation, qualifiedEnum
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doc/theses/andrew_beach_MMath/Makefile
r331eacbe r87f572e 5 5 TEXSRC=$(wildcard *.tex) 6 6 BIBSRC=$(wildcard *.bib) 7 STYSRC=$(wildcard *.sty) 8 CLSSRC=$(wildcard *.cls) 7 9 TEXLIB= .:${BUILD}: 8 10 BIBLIB= .:../../bibliography … … 24 26 all: ${DOC} 25 27 26 ${BUILD}/${DOC}: ${TEXSRC} ${BIBSRC} Makefile | ${BUILD}28 ${BUILD}/${DOC}: ${TEXSRC} ${BIBSRC} ${STYSRC} ${CLSSRC} Makefile | ${BUILD} 27 29 ${LATEX} ${BASE} 28 30 ${BIBTEX} ${BUILD}/${BASE} -
doc/theses/andrew_beach_MMath/thesis.tex
r331eacbe r87f572e 1 % uWaterloo Thesis Template for LaTeX 2 % Last Updated June 14, 2017 by Stephen Carr, IST Client Services 3 % FOR ASSISTANCE, please send mail to rt-IST-CSmathsci@ist.uwaterloo.ca 4 5 % Effective October 2006, the University of Waterloo 6 % requires electronic thesis submission. See the uWaterloo thesis regulations at 7 % https://uwaterloo.ca/graduate-studies/thesis. 8 9 % DON'T FORGET TO ADD YOUR OWN NAME AND TITLE in the "hyperref" package 10 % configuration. THIS INFORMATION GETS EMBEDDED IN THE FINAL PDF DOCUMENT. 11 % You can view the information if you view Properties of the PDF document. 12 13 % Many faculties/departments also require one or more printed 14 % copies. This template attempts to satisfy both types of output. 15 % It is based on the standard "book" document class which provides all 16 % necessary sectioning structures and allows multi-part theses. 17 18 % DISCLAIMER 19 % To the best of our knowledge, this template satisfies the current uWaterloo 20 % requirements. However, it is your responsibility to assure that you have met 21 % all requirements of the University and your particular department. 22 % Many thanks for the feedback from many graduates that assisted the 23 % development of this template. 24 25 % ----------------------------------------------------------------------- 26 27 % By default, output is produced that is geared toward generating a PDF 28 % version optimized for viewing on an electronic display, including 29 % hyperlinks within the PDF. 30 31 % E.g. to process a thesis called "mythesis.tex" based on this template, run: 32 33 % pdflatex mythesis -- first pass of the pdflatex processor 34 % bibtex mythesis -- generates bibliography from .bib data file(s) 35 % makeindex -- should be run only if an index is used 36 % pdflatex mythesis -- fixes numbering in cross-references, 37 % pdflatex mythesis -- bibliographic references, glossaries, index, etc. 38 39 % N.B. The "pdftex" program allows graphics in the following formats to be 40 % included with the "\includegraphics" command: PNG, PDF, JPEG, TIFF 41 % Tip 1: Generate your figures and photos in the size you want them to appear 42 % in your thesis, rather than scaling them with \includegraphics options. 43 % Tip 2: Any drawings you do should be in scalable vector graphic formats: 44 % SVG, PNG, WMF, EPS and then converted to PNG or PDF, so they are scalable in 45 % the final PDF as well. 46 % Tip 3: Photographs should be cropped and compressed so as not to be too large. 47 48 % To create a PDF output that is optimized for double-sided printing: 49 % 50 % 1) comment-out the \documentclass statement in the preamble below, and 51 % un-comment the second \documentclass line. 52 % 53 % 2) change the value assigned below to the boolean variable 54 % "PrintVersion" from "false" to "true". 55 56 % --------------------- Start of Document Preamble ----------------------- 57 58 % Specify the document class, default style attributes, and page dimensions 59 % For hyperlinked PDF, suitable for viewing on a computer, use this: 60 \documentclass[letterpaper,12pt,titlepage,oneside,final]{book} 61 62 % For PDF, suitable for double-sided printing, change the PrintVersion 63 % variable below to "true" and use this \documentclass line instead of the one 64 % above: 65 %\documentclass[letterpaper,12pt,titlepage,openright,twoside,final]{book} 66 67 % Some LaTeX commands I define for my own nomenclature. 68 % If you have to, it's better to change nomenclature once here than in a 69 % million places throughout your thesis! 70 \newcommand{\package}[1]{\textbf{#1}} % package names in bold text 71 \newcommand{\cmmd}[1]{\textbackslash\texttt{#1}} % command name in tt font 72 \newcommand{\href}[1]{#1} % does nothing, but defines the command so the 73 % print-optimized version will ignore \href tags (redefined by hyperref pkg). 74 %\newcommand{\texorpdfstring}[2]{#1} % does nothing, but defines the command 75 % Anything defined here may be redefined by packages added below... 76 77 % This package allows if-then-else control structures. 78 \usepackage{ifthen} 79 \newboolean{PrintVersion} 80 \setboolean{PrintVersion}{false} 81 % CHANGE THIS VALUE TO "true" as necessary, to improve printed results for 82 % hard copies by overriding some options of the hyperref package below. 1 % Main tex file for thesis document. 2 \documentclass[digital]{uw-ethesis} 3 4 % Commands used in documenting how to use the template. To remove. 5 \newcommand{\package}[1]{\textbf{#1}} 6 \newcommand{\cmmd}[1]{\textbackslash\texttt{#1}} 7 \newcommand{\href}[1]{#1} 83 8 84 9 % For a nomenclature (optional; available from ctan.org) … … 86 11 % Lots of math symbols and environments 87 12 \usepackage{amsmath,amssymb,amstext} 88 % For including graphics N.B. pdftex graphics driver13 % For including graphics, sets the pdftex graphics driver. 89 14 \usepackage[pdftex]{graphicx} 90 15 91 % I believe the general index function is covered by the glossaries. 92 % \usepackage{makeidx} 93 % \makeindex 94 95 % Hyperlinks make it very easy to navigate an electronic document. 96 % In addition, this is where you should specify the thesis title 97 % and author as they appear in the properties of the PDF document. 98 % Use the "hyperref" package 99 % N.B. HYPERREF MUST BE THE LAST PACKAGE LOADED; ADD ADDITIONAL PKGS ABOVE 100 % N.B. pagebackref=true provides links back from the References to the body 101 % text. This can cause trouble for printing. 102 \usepackage[pdftex,pagebackref=false]{hyperref} % with basic options 103 \hypersetup{ 104 plainpages=false, % needed if Roman numbers in frontpages 105 unicode=false, % non-Latin characters in Acrobat’s bookmarks 106 pdftoolbar=true, % show Acrobat’s toolbar? 107 pdfmenubar=true, % show Acrobat’s menu? 108 pdffitwindow=false, % window fit to page when opened 109 pdfstartview={FitH}, % fits the width of the page to the window 110 pdftitle={uWaterloo\ LaTeX\ Thesis\ Template}, % title: CHANGE THIS TEXT! 111 % pdfauthor={Author}, % author: CHANGE THIS TEXT! and uncomment this line 112 % pdfsubject={Subject}, % subject: CHANGE THIS TEXT! and uncomment this line 113 % pdfkeywords={keyword1} {key2} {key3}, % list of keywords, and uncomment this line if desired 114 pdfnewwindow=true, % links in new window 115 colorlinks=true, % false: boxed links; true: colored links 116 linkcolor=blue, % color of internal links 117 citecolor=green, % color of links to bibliography 118 filecolor=magenta, % color of file links 119 urlcolor=cyan % color of external links 120 } 121 \ifthenelse{\boolean{PrintVersion}}{ 122 % for improved print quality, override some hyperref options 123 \hypersetup{ 124 % colorlinks,% 125 citecolor=black,% 126 filecolor=black,% 127 linkcolor=black,% 128 urlcolor=black} 129 }{} % end of ifthenelse (no else) 130 131 \usepackage[toc,abbreviations]{glossaries-extra} % Exception to the 132 % rule of hyperref being the last add-on package. If glossaries-extra is not 133 % in your LaTeX distribution, get it from CTAN 134 % (http://ctan.org/pkg/glossaries-extra). 135 136 % Setting up the page margins... 137 % uWaterloo thesis requirements specify a minimum of 1 inch (72pt) margin at 138 % the top, bottom, and outside page edges and a 1.125 in. (81pt) gutter 139 % margin (on binding side). While this is not an issue for electronic 140 % viewing, a PDF may be printed, and so we have the same page layout for 141 % both printed and electronic versions, we leave the gutter margin in. 142 % Set margins to minimum permitted by uWaterloo thesis regulations: 143 \setlength{\marginparwidth}{0pt} % width of margin notes 144 % N.B. If margin notes are used, you must adjust \textwidth, \marginparwidth 145 % and \marginparsep so that the space left between the margin notes and page 146 % edge is less than 15 mm (0.6 in.) 147 % Width of space between body text and margin notes. 148 \setlength{\marginparsep}{0pt} 149 \setlength{\evensidemargin}{0.125in} % Adds 1/8 in. to binding side of all 150 % even-numbered pages when the "twoside" printing option is selected 151 \setlength{\oddsidemargin}{0.125in} % Adds 1/8 in. to the left of all pages 152 % when "oneside" printing is selected, and to the left of all odd-numbered 153 % pages when "twoside" printing is selected 154 % Assuming US letter paper (8.5 in. x 11 in.) and side margins as above. 155 \setlength{\textwidth}{6.375in} 156 \raggedbottom 157 158 % The following statement specifies the amount of space between paragraphs. 159 % Other reasonable specifications are \bigskipamount and \smallskipamount. 160 \setlength{\parskip}{\medskipamount} 161 162 % The following statement controls the line spacing. The default 163 % spacing corresponds to good typographic conventions and only slight 164 % changes (e.g., perhaps "1.2"), if any, should be made. 165 \renewcommand{\baselinestretch}{1} % this is the default line space setting 166 167 % By default, each chapter will start on a recto (right-hand side) 168 % page. We also force each section of the front pages to start on 169 % a recto page by inserting \cleardoublepage commands. 170 % In many cases, this will require that the verso page be 171 % blank and, while it should be counted, a page number should not be 172 % printed. The following statements ensure a page number is not 173 % printed on an otherwise blank verso page. 174 \let\origdoublepage\cleardoublepage 175 \newcommand{\clearemptydoublepage}{% 176 \clearpage{\pagestyle{empty}\origdoublepage}} 177 \let\cleardoublepage\clearemptydoublepage 178 179 % Define Glossary terms (This is properly done here, in the preamble. 180 % Could be \input{} from a file...) 16 \usehyperrefpackage[pdftex,pagebackref=false]{ 17 pdftitle={Exception Handling in CFA}, 18 pdfauthor={Andrew James Beach}, 19 pdfsubject={Programming Languages}, 20 pdfkeywords={exceptions,implementation}, 21 } 22 23 % The \phantomsection is used to help the hyperref package create links. 24 25 % Maybe only package that should be loaded after the hyperref package. 26 % From http://ctan.org/pkg/glossaries-extra, extends glossaries which replaces 27 % glossary and builds off of the makeindex system. 28 \usepackage[toc,abbreviations]{glossaries-extra} 29 181 30 % Main glossary entries -- definitions of relevant terminology 182 31 \newglossaryentry{computer} … … 209 58 description={Random vector: a location in n-dimensional Cartesian space, where each dimensional component is determined by a random process} 210 59 } 211 60 61 % Generate the glossaries defined above. 212 62 \makeglossaries 213 63 214 %======================================================================215 % L O G I C A L D O C U M E N T -- the content of your thesis216 %======================================================================217 64 \begin{document} 218 65 219 % For a large document, it is a good idea to divide your thesis220 % into several files, each one containing one chapter.221 % To illustrate this idea, the "front pages" (i.e., title page,222 % declaration, borrowers' page, abstract, acknowledgements,223 % dedication, table of contents, list of tables, list of figures,224 % nomenclature) are contained within the file "uw-ethesis-frontpgs.tex" which225 % is included into the document by the following statement.226 66 %---------------------------------------------------------------------- 227 67 % FRONT MATERIAL … … 232 72 % MAIN BODY 233 73 %---------------------------------------------------------------------- 234 % Because this is a short document, and to reduce the number of files235 % needed for this template, the chapters are not separate236 % documents as suggested above, but you get the idea. If they were237 % separate documents, they would each start with the \chapter command, i.e, do238 % not contain \documentclass or \begin{document} and \end{document} commands.239 74 %====================================================================== 240 75 \chapter{Introduction} … … 327 162 %---------------------------------------------------------------------- 328 163 329 % B I B L I O G R A P H Y 330 % ----------------------- 331 332 % The following statement selects the style to use for references. It controls 333 % the sort order of the entries in the bibliography and also the formatting 334 % for the in-text labels. 335 \bibliographystyle{plain} 336 % This specifies the location of the file containing the bibliographic 337 % information. It assumes you're using BibTeX (if not, why not?). 338 339 % This is needed if the book class is used, to place the anchor in the correct 340 % page, because the bibliography will start on its own page. 164 %---------------------------------------------------------------------- 165 % BIBLIOGRAPHY 166 %---------------------------------------------------------------------- 167 341 168 % Use \clearpage instead if the document class uses the "oneside" argument. 342 169 \cleardoublepage 343 % With hyperref package, enables hyperlinking from the table of contents to344 % bibliography345 170 \phantomsection 346 171 347 % The following statement causes the title "References" to be used for the 348 % bibliography section: 349 \renewcommand*{\bibname}{References} 350 351 % Add the References to the Table of Contents 352 \addcontentsline{toc}{chapter}{\textbf{References}} 353 354 % Tip 5: You can create multiple .bib files to organize your references. Just 355 % list them all in the \bibliogaphy command, separated by commas (no spaces). 172 % Bibliography setup and creation, renamed to References. 173 \addcontentsline{toc}{chapter}{\textbf{\bibname}} 174 \bibliographystyle{plain} 356 175 \bibliography{thesis} 357 176 358 % The following statement causes the specified references to be added to the 359 % bibliography even if they were not cited in the text. The asterisk is a 360 % wildcard that causes all entries in the bibliographic database to be 361 % included (optional). 177 % Include all uncited entries in the bibliography. 362 178 \nocite{*} 363 179 364 % The \appendix statement indicates the beginning of the appendices.180 % Begin the appendix, add a title and table of contents entry. 365 181 \appendix 366 % Add a title page before the appendices and a line in the Table of Contents367 182 \chapter*{APPENDICES} 368 183 \addcontentsline{toc}{chapter}{APPENDICES} -
doc/theses/thierry_delisle_PhD/.gitignore
r331eacbe r87f572e 8 8 9 9 comp_II/build/ 10 comp_II/img/*.fig.bak 10 11 comp_II/comp_II.pdf 11 12 comp_II/comp_II.ps -
doc/theses/thierry_delisle_PhD/comp_II/Makefile
r331eacbe r87f572e 18 18 19 19 FIGURES = ${addsuffix .tex, \ 20 base \ 21 empty \ 22 emptybit \ 23 emptytree \ 24 resize \ 20 25 } 21 26 … … 70 75 mkdir -p ${Build} 71 76 72 %.tex : %.fig ${Build}77 %.tex : img/%.fig ${Build} 73 78 fig2dev -L eepic $< > ${Build}/$@ 74 79 75 %.ps : %.fig | ${Build}80 %.ps : img/%.fig | ${Build} 76 81 fig2dev -L ps $< > ${Build}/$@ 77 82 78 %.pstex : %.fig | ${Build}83 %.pstex : img/%.fig | ${Build} 79 84 fig2dev -L pstex $< > ${Build}/$@ 80 85 fig2dev -L pstex_t -p ${Build}/$@ $< > ${Build}/$@_t -
doc/theses/thierry_delisle_PhD/comp_II/comp_II.tex
r331eacbe r87f572e 1 \documentclass[11pt,fullpage]{article} 1 \documentclass[11pt]{article} 2 \usepackage{fullpage} 2 3 \usepackage[T1]{fontenc} 3 4 \usepackage[utf8]{inputenc} … … 6 7 \usepackage{xcolor} 7 8 \usepackage{graphicx} 8 \usepackage [hidelinks]{hyperref}9 \usepackage{epic,eepic} 9 10 \usepackage{glossaries} 10 11 \usepackage{textcomp} 11 \usepackage{geometry} 12 \usepackage[hidelinks]{hyperref} 13 %\usepackage[margin=1in]{geometry} 14 %\usepackage{float} 12 15 13 16 % cfa macros used in the document … … 51 54 \section{Introduction} 52 55 \subsection{\CFA and the \CFA concurrency package} 53 \CFA\cit is a modern, polymorphic, non-object-oriented, backwards-compatible extension of the C programming language. It aims to add high productivity features while maintaning the predictible performance of C. As such concurrency in \CFA\cit aims to offer simple and safe high-level tools while still allowing performant code. Concurrent code is written in the syncrhonous programming paradigm but uses \glspl{uthrd} in order to achieve the simplicity and maintainability of synchronous programming without sacrificing the efficiency of asynchronous programing. As such the \CFA scheduler is a user-level scheduler that maps \glspl{uthrd} onto \glspl{kthrd}. 54 55 The goal of this research is to produce a scheduler that is simple to use and offers acceptable performance in all cases. Here simplicity does not refer to the API but to how much scheduling concerns programmers need to take into account when using the \CFA concurrency package. Therefore, the main goal of this proposal is as follows : 56 \CFA\cit is a modern, polymorphic, non-object-oriented, backwards-compatible extension of the C programming language. It aims to add high-productivity features while maintaning the predictible performance of C. As such, concurrency in \CFA\cit aims to offer simple and safe high-level tools while still allowing performant code. \CFA concurrrent code is written in the synchronous programming paradigm but uses \glspl{uthrd} in order to achieve the simplicity and maintainability of synchronous programming without sacrificing the efficiency of asynchronous programing. As such, the \CFA \emph{scheduler} is a preemptive user-level scheduler that maps \glspl{uthrd} onto \glspl{kthrd}. 57 58 Scheduling occurs when execution switches from one thread to another, where the second thread is implicitly chosen by the scheduler. This scheduling is an indirect handoff, as opposed to generators and coroutines which explicitly switch to the next generator and coroutine respectively. The cost of switching between two threads for an indirect handoff has two components : the cost of actually context-switching, i.e., changing the relevant registers to move execution from one thread to the other, and the cost of scheduling, i.e., deciding which thread to run next among all the threads ready to run. The first cost is generally constant and fixed, while the scheduling cost can vary based on the system state\footnote{Affecting the context-switch cost is whether it is done in one step, after the scheduling, or in two steps, context-switching to a fixed third-thread before scheduling.}. Adding multiple \glspl{kthrd} does not fundamentally change the scheduler semantics or requirements, it simply adds new correctness requirements, i.e. \textit{linearizability}, and a new dimension to performance: scalability, where scheduling cost now also depends on contention. 59 60 The more threads switch, the more the administrating cost of scheduling becomes noticeable. It is therefore important to build a scheduler with the lowest possible cost and latency. Another important consideration is \emph{fairness}. In principle, scheduling should give the illusion of perfect fairness, where all threads ready to run are running \emph{simultaneously}. While the illusion of simultaneity is easier to reason about, it can break down if the scheduler allows to much unfairness. Therefore, the scheduler should offer as much fairness as needed to guarantee eventual progress, but use unfairness to help performance. In practice, threads must wait in turn but there can be advantages to unfair scheduling, e.g., the express cash register at a grocery store. 61 62 The goal of this research is to produce a scheduler that is simple for programmers to understand and offers good performance. Here understandability does not refer to the API but to how much scheduling concerns programmers need to take into account when writing a \CFA concurrent package. Therefore, the main goal of this proposal is : 56 63 \begin{quote} 57 The \CFA scheduler should be \emph{viable} for anyworkload.64 The \CFA scheduler should be \emph{viable} for \emph{any} workload. 58 65 \end{quote} 59 66 60 This objective includes producing a scheduling strategy with minimal fairness guarantees, creating an abstraction layer over the operating system to handle kernel-threads spinning unnecessarily and hide blocking I/O operations and, writing sufficient library tools to allow developpers to properly use the scheduler. 67 For a general purpose scheduler, it is impossible to produce an optimal algorithm as it would require knowledge of the future behaviour of threads. As such, scheduling performance is generally either defined by the best case scenario, a workload to which the scheduler is tailored, or the worst case scenario, i.e., the scheduler behaves no worst than \emph{X}. For this proposal, the performance is evaluated using the second approach to allow \CFA programmers to rely on scheduling performance. A solution to this impossibility is to allow programmers to write their own scheduler, that is not the subject of this proposal, which considers only the default scheduler. As such, it is important that only programmers with exceptionally high performance requirements should need to write their own scheduler and replace the scheduler in this proposal. 68 69 This objective includes producing a scheduling strategy with sufficient fairness guarantees, creating an abstraction layer over the operating system to handle kernel-threads spinning unnecessarily and hide blocking I/O operations, and writing sufficient library tools to allow developers to indirectly use the scheduler. 61 70 62 71 % =============================================================================== … … 64 73 65 74 \section{Scheduling for \CFA} 66 While the \CFA concurrency package doesn't have any particular scheduling needs beyond those of any concurrency package which uses \glspl{uthrd}, it is important that the default \CFA Scheduler be viable in general. Indeed, since the \CFA Scheduler does not target any specific workloads, it is unrealistic to demand that it use the best scheduling strategy in all cases. However, it should offer a viable ``out of the box'' solution for most scheduling problems so that programmers can quickly write performant concurrent without needed to think about which scheduling strategy is more appropriate for their workload. Indeed, only programmers with exceptionnaly high performance requirements should need to write their own scheduler. More specifically, two broad types of schedulering strategies should be avoided in order to avoid penalizing certain types of workloads : feedback-based and priority schedulers. 75 While the \CFA concurrency package does not have any particular scheduling requirements beyond supporting \glspl{uthrd}. Therefore, the detailed requirements of the \CFA scheduler are : 76 77 \paragraph{Correctness} As with any other concurrent data structure or algorithm, the correctness requirement is paramount. The scheduler cannot allow threads to be dropped from the ready-queue, i.e., scheduled but never run, or be executed multiple times when only being scheduled once. Since \CFA concurrency has no spurious wakeup, this definition of correctness also means the scheduler should have no spurious wakeup. The \CFA scheduler must be correct. 78 79 \paragraph{Performance} The performance of a scheduler can generally be mesured in terms of scheduling cost, scalability and latency. Scheduling cost is the cost to switch from one thread to another, as mentioned above. For simple applications where a single kernel thread does most of the scheduling, it is generally the dominating cost. When adding many kernel threads, scalability can become an issue, effectively increasing the cost of context-switching when contention is high. Finally, a third axis of performance is tail latency. This measurement is related to fairness and mesures how long is needed for a thread to be run once scheduled but evaluated in the worst cases. The \CFA scheduler should offer good performance in all three metrics. 80 81 \paragraph{Fairness} Like performance, this requirements has several aspect : eventual progress, predictability and performance reliablility. As a hard requirement, the \CFA scheduler must guarantee eventual progress, i.e., prevent starvation, otherwise the above mentioned illusion of simultaneous execution is broken and the scheduler becomes much more complext to reason about. Beyond this requirement, performance should be predictible and reliable, which means similar workloads achieve similar performance and programmer intuition is respected. An example of this is : a thread that yields agressively should not run more often then other tasks. While this is intuitive, it does not hold true for many work-stealing or feedback based schedulers. The \CFA scheduler must guarantee eventual progress and should be predictible and offer reliable performance. 82 83 \paragraph{Efficiency} Finally, efficient usage of CPU resources is also an important requirement. This issue is discussed more in depth towards the end of this proposal. It effectively refers to avoiding using CPU power when there are no threads to run, and conversely, use all CPUs available when the workload can benefit from it. Balancing these two states is where the complexity lies. The \CFA scheduler should be efficient with respect to the underlying (shared) computer. 84 85 To achieve these requirements, I can reject two broad types of scheduling strategies : feedback-based and priority schedulers. 67 86 68 87 \subsection{Feedback-Based Schedulers} 69 Many operating systems use schedulers based on fe adback loops in some form, they measure how much CPU a particular thread has used\footnote{Different metrics can be used tohere but it is not relevant to the discussion.} and schedule threads based on this metric. These strategies are sensible for operating systems but rely on two assumptions on the workload :88 Many operating systems use schedulers based on feedback in some form, e.g., measuring how much CPU a particular thread has used\footnote{Different metrics can measured here but it is not relevant to the discussion.} and schedule threads based on this metric. These strategies are sensible for operating systems but rely on two assumptions on the workload : 70 89 71 90 \begin{enumerate} 72 \item Threads live long enough to be scheduled many times.73 \item Cooperation among all threads is not simply infeasible, it is a security risk.91 \item Threads live long enough for useful feedback information to be to gathered. 92 \item Threads belong to multiple users so fairness across threads is insufficient. 74 93 \end{enumerate} 75 94 76 While these two assumptions generally hold for operating systems, they may not for \CFA programs. In fact, \CFA uses \glspl{uthrd} which have the explicit goal of reducing the cost of threading primitives to allow many smaller threads. This can naturally lead to have threads with much shorter lifetime and only being scheduled a few times. Scheduling strategies based on feadback loops cannot be effective in these cases because they will not have the opportunity to measure the metrics that underlay the algorithm. Note that the problem of feadback loop convergence (reacting too slowly to scheduling events) is not specific to short lived threads but can also occur with threads that show drastic changes in scheduling event, e.g., threads running for long periods of time and then suddenly blocking and unblocking quickly and repeatedly.77 78 In the context of operating systems, these concerns can be overshadowed by a more pressing concern : security. When multiple users are involved, it is possible that some users are malevolent and try to exploit the scheduling strategy in order to achieve some nefarious objective. Security concerns mean that more precise and robust fairness metrics must be used . In the case of the \CFA scheduler, every thread runs in the same user-space and are controlled from the same user. It is then possible to safely ignore the possibility that threads are malevolent and assume that all threads will ignore or cooperate with each other. This allows for a much simpler fairness metric and in this proposal ``fairness'' will be considered as equal opportunities to run once scheduled.79 80 Since fe adback is not necessarily feasible within the lifetime of all threads and a simple fairness metric can be used, the scheduling strategy proposed for the \CFA runtime does not user per-threads feedback. Feedback loops in general are not rejected for secondary concerns like idle sleep, but no feedback loopis used to decide which thread to run next.95 While these two assumptions generally hold for operating systems, they may not for user-level threading. Since \CFA has the explicit goal of allowing many smaller threads, this can naturally lead to threads with much shorter lifetime, only being scheduled a few times. Scheduling strategies based on feedback cannot be effective in these cases because they do not have the opportunity to measure the metrics that underlie the algorithm. Note that the problem of feedback convergence (reacting too slowly to scheduling events) is not specific to short lived threads but can also occur with threads that show drastic changes in scheduling, e.g., threads running for long periods of time and then suddenly blocking and unblocking quickly and repeatedly. 96 97 In the context of operating systems, these concerns can be overshadowed by a more pressing concern : security. When multiple users are involved, it is possible that some users are malevolent and try to exploit the scheduling strategy in order to achieve some nefarious objective. Security concerns mean that more precise and robust fairness metrics must be used to guarantee fairness across processes created by users as well as threads created within a process. In the case of the \CFA scheduler, every thread runs in the same user-space and are controlled by the same user. Fairness across users is therefore a given and it is then possible to safely ignore the possibility that threads are malevolent. This approach allows for a much simpler fairness metric and in this proposal ``fairness'' is considered as follows : when multiple threads are cycling through the system, the total ordering of threads being scheduled, i.e., pushed onto the ready-queue, should not differ much from the total ordering of threads being executed, i.e., popped from the ready-queue. 98 99 Since feedback is not necessarily feasible within the lifetime of all threads and a simple fairness metric can be used, the scheduling strategy proposed for the \CFA runtime does not use per-threads feedback. Feedback in general is not rejected for secondary concerns like idle sleep, but no feedback is used to decide which thread to run next. 81 100 82 101 \subsection{Priority Schedulers} 83 Another broad category of schedulers are priority schedulers. In these scheduling strategies threads have priorities and the runtime schedules the threads with the highest priority before scheduling other threads. Threads with equal priority are scheduled using a secondary strategy, often something simple like round-robin or FIFO. These priority mean that, as long as there is a thread with a higher priority that desires to run, a thread with a lower priority will not run. This possible starving of threads can dramatically increase programming complexity since starving threads and priority inversion (prioritising a lower priority thread) can both lead to serious problems, leaving programmers between a rock and a hard place. 84 85 An important observation to make is that threads do not need to have explicit priorities for problems to be possible. Indeed, any system with multiple ready-queues and attempts to exhaust one queue before accessing the other queues, could encounter starvation problems. A popular scheduling strategy that suffers from implicit priorities is work-stealing. Work-stealing is generally presented as follows : 86 87 \begin{itemize} 88 \item Each processor has a list of threads. 89 \end{itemize} 102 Another broad category of schedulers are priority schedulers. In these scheduling strategies, threads have priorities and the runtime schedules the threads with the highest priority before scheduling other threads. Threads with equal priority are scheduled using a secondary strategy, often something simple like round-robin or FIFO. These priority mean that, as long as there is a thread with a higher priority that desires to run, a thread with a lower priority does not run. This possible starving of threads can dramatically increase programming complexity since starving threads and priority inversion (prioritizing a lower priority thread) can both lead to serious problems. 103 104 An important observation to make is that threads do not need to have explicit priorities for problems to occur. Indeed, any system with multiple ready-queues and attempts to exhaust one queue before accessing the other queues, can encounter starvation problems. A popular scheduling strategy that suffers from implicit priorities is work-stealing. Work-stealing is generally presented as follows, each processor has a list of ready threads. 90 105 \begin{enumerate} 91 106 \item Run threads from ``this'' processor's list. … … 93 108 \end{enumerate} 94 109 95 In a loaded system\footnote{A loaded system is a system where threads are being run at the same rate they are scheduled }, if a thread does not yield or block for an extended period of time, threads on the same processor list will starve if no other processors canexhaust their list.110 In a loaded system\footnote{A loaded system is a system where threads are being run at the same rate they are scheduled.}, if a thread does not yield, block or preempt for an extended period of time, threads on the same processor list starve if no other processors exhaust their list. 96 111 97 112 Since priorities can be complex to handle for programmers, the scheduling strategy proposed for the \CFA runtime does not use a strategy with either implicit or explicit thread priorities. 98 113 99 \subsection{Schedulers without feadback or priorities} 100 I claim that the ideal default scheduler for the \CFA runtime is a scheduler that offers good scalability and a simple fairness guarantee that is easy for programmers to reason about. The simplest fairness guarantee is to guarantee FIFO ordering, i.e., threads scheduled first will run first. However, enforcing FIFO ordering generally conflicts with scalability across multiple processors because of the additionnal synchronization. Thankfully, strict FIFO is not needed for scheduling. Since concurrency is inherently non-deterministic, fairness concerns in scheduling are only a problem if a thread repeatedly runs before another thread can run\footnote{This is because the non-determinism means that programmers must already handle ordering problems in order to produce correct code and already must rely on weak guarantees, for example that a specific thread will \emph{eventually} run.}. This need for unfairness to persist before problems occur means that the FIFO fairness guarantee can be significantly relaxed without causing problems. For this proposal, the target guarantee is that the \CFA scheduler guarantees \emph{probable} FIFO ordering, which is defined as follows : 114 \subsection{Schedulers without feedback or priorities} 115 This proposal conjectures that is is possible to construct a default scheduler for the \CFA runtime that offers good scalability and a simple fairness guarantee that is easy for programmers to reason about. The simplest fairness guarantee is FIFO ordering, i.e., threads scheduled first run first. However, enforcing FIFO ordering generally conflicts with scalability across multiple processors because of the additionnal synchronization. Thankfully, strict FIFO is not needed for sufficient fairness. Since concurrency is inherently non-deterministic, fairness concerns in scheduling are only a problem if a thread repeatedly runs before another thread can run. This is because the non-determinism means that programmers must already handle ordering problems in order to produce correct code and already must rely on weak guarantees, for example that a specific thread will \emph{eventually} run. Since some reordering does not break correctness, the FIFO fairness guarantee can be significantly relaxed without causing problems. For this proposal, the target guarantee is that the \CFA scheduler provides \emph{probable} FIFO ordering, which allows reordering but makes it improbable that threads are reordered far from their position in total ordering. 116 117 Scheduling is defined as follows : 101 118 \begin{itemize} 102 \item Given two threads $X$ and $Y$, the odds that thread $X$ runs $N$ times \emph{after} thread $Y$ is scheduled but \emph{before} it is run, decreases exponentially with regard sto $N$.119 \item Given two threads $X$ and $Y$, the odds that thread $X$ runs $N$ times \emph{after} thread $Y$ is scheduled but \emph{before} it is run, decreases exponentially with regard to $N$. 103 120 \end{itemize} 104 121 105 While this is not a strong guarantee, the probability that problems persist for long period of times decreases exponentially, making persisting problems virtually impossible. 106 107 \subsection{Real-Time} 108 While the objective of this proposed scheduler is similar to the objective of real-time scheduling, this proposal is not a proposal for real-time scheduler and as such makes no attempt to offer either soft or hard guarantees on scheduling delays. 122 While this is not a bounded guarantee, the probability that unfairness persist for long periods of times decreases exponentially, making persisting unfairness virtually impossible. 109 123 110 124 % =============================================================================== … … 112 126 \section{Proposal} 113 127 114 \subsection{Ready-Queue} 115 Using trevor's paper\cit as basis, it is simple to build a relaxed FIFO list that is fast and scalable for loaded or overloaded systems. The described queue uses an array of underlying strictly FIFO queue. Pushing new data is done by selecting one of these underlying queues at random, recording a timestamp for the push and pushing to the selected queue. Popping is done by selecting two queues at random and popping from the queue for which the head has the oldest timestamp. In loaded or overloaded systems, it is higly likely that the queues is far from empty, e.i., several tasks are on each of the underlying queues. This means that selecting a queue at random to pop from is higly likely to yield a queue that is not empty. 116 117 When the ready queue is "more empty", i.e., several of the inner queues are empty, selecting a random queue for popping is less likely to yield a valid selection and more attempts need to be made, resulting in a performance degradation. In cases, with few elements on the ready queue and few processors running, performance can be improved by adding information to help processors find which inner queues are used. Preliminary performance tests indicate that with few processors, a bitmask can be used to identify which inner queues are currently in use. This is especially effective in the single-thread case, where the bitmask will always be up-to-date. Furthermore, modern x86 CPUs have a BMI2 extension which allow using the bitmask with very little overhead over directly accessing the readyqueue offerring decent performance even in cases with many empty inner queues. This technique does not solve the problem completely, it randomly attempts to find a block of 64 queues where at least one is used, instead of attempting to find a used queue. For systems with a large number of cores this does not completely solve the problem, but it is a fixed improvement. The size of the blocks are limited by the maximum size atomic instruction can operate on, therefore atomic instructions on large words would increase the 64 queues per block limit. 118 119 \TODO double check the next sentence 120 Preliminary result indicate that the bitmask approach with the BMI2 extension can lead to multi-threaded performance that is contention agnostic in the worst case. 121 This result suggests that the contention penalty and the increase performance for additionnal thread cancel each other exactly. This may indicate that a relatively small reduction in contention may tip the performance into positive scalling even for the worst case. It can be noted that in cases of high-contention, the use of the bitmask to find queues that are not empty is much less reliable. Indeed, if contention on the bitmask is high, it means it probably changes significantly between the moment it is read and the actual operation on the queues it represents. Furthermore, the objective of the bitmask is to avoid probing queues that are empty. Therefore, in cases where the bitmask is highly contented, it may be preferrable to probe queues randomly, either until contention decreases or until a prior prefetch of the bitmask completes. Ideally, the scheduler would be able to observe that the bitmask is highly contented and adjust its behaviour appropriately. However, I am not aware of any mechanism to query whether a cacheline is in cache or to run other instructions until a cacheline is fetch without blocking on the cacheline. As such, an alternative that may have a similar impact would be for each thread to have their own bitmask, which would be updated both after each scheduler action and after a certain number of failed probing. If the bitmask has little contention, the local bitmask will be mostly up-to-date and several threads won't need to contend as much on the global bitmask. If the bitmask has significant contention, then fetching it becomes more expensive and threads may as well probe randomly. This solution claims that probing randomly or against an out-of-date bitmask is equivalent. 122 123 In cases where this is insufficient, another approach is to use a hiearchical data structure. Creating a tree of nodes to reduce contention has been shown to work in similar cases\cit(SNZI: Scalable NonZero Indicators)\footnote{This particular paper seems to be patented in the US. How does that affect \CFA? Can I use it in my work?}. However, this approach may lead to poorer single-threaded performance due to the inherent pointer chasing, as such, it was not considered as the first approach but as a fallback in case the bitmask approach does not satisfy the performance goals. 124 125 Part of this performance relies on contention being low when there are few threads on the readyqueue. However, this can be assumed reliably if the system handles putting idle processors to sleep, which is addressed in section \ref{sleep}. 128 \subsection{Ready-Queue} \label{sec:queue} 129 A simple ready-queue can be built from a FIFO queue, user-threads are pushed onto the queue when they are ready to run and processors (kernel-threads acting as virtual processors) pop the user-threads from the queue and execute them. Using Trevor's paper\cit as basis, it is simple to build a relaxed FIFO list that is fast and scalable for loaded or overloaded systems. The described queue uses an array of underlying strictly FIFO queues as shown in Figure~\ref{fig:base}\footnote{For this section, the number of underlying queues is assumed to be constant, Section~\ref{sec:resize} will discuss resizing the array.}. Pushing new data is done by selecting one of these underlying queues at random, recording a timestamp for the push and pushing to the selected queue. Popping is done by selecting two queues at random and popping from the queue for which the head has the oldest timestamp. A higher number of underlying queues leads to less contention on each queue and therefore better performance. In a loaded system, it is higly likely the queues are non-empty, i.e., several tasks are on each of the underlying queues. This means that selecting a queue at random to pop from is higly likely to yield a queue with available items. In Figure~\ref{fig:base}, ignoring the ellipsis, the chances of getting an empty queue is 2/7 per pick, meaning two randoms pick will yield an item approximately 9 times out of 10. 130 131 \begin{figure} 132 \begin{center} 133 % {\resizebox{0.8\textwidth}{!}{\input{base}}} 134 \input{base} 135 \end{center} 136 \caption{Relaxed FIFO list at the base of the scheduler: an array of strictly FIFO lists. } 137 \label{fig:base} 138 \end{figure} 139 140 \begin{figure} 141 \begin{center} 142 % {\resizebox{0.8\textwidth}{!}{\input{empty}}} 143 \input{empty} 144 \end{center} 145 \caption{``More empty'' state of the queue: the array contains many empty cells.} 146 \label{fig:empty} 147 \end{figure} 148 149 When the ready queue is "more empty", i.e., several of the inner queues are empty, selecting a random queue for popping is less likely to yield a valid selection and more attempts need to be made, resulting in a performance degradation. Figure~\ref{fig:empty} shows an example with fewer elements where the chances of getting an empty queue is 5/7 per pick, meaning two randoms pick will yield an item only half the time. Since the overarching queue is not empty, the pop operation \emph{must} find an element before returning and therefore must retry. Overall performance is therefore influenced by the contention on the underlying queues and pop performance is influenced by the items density. This leads to four performance cases, as depicted in Table~\ref{tab:perfcases}. 150 151 \begin{table} 152 \begin{center} 153 \begin{tabular}{|r|l|l|} 154 \cline{2-3} 155 \multicolumn{1}{r|}{} & \multicolumn{1}{c|}{Many Processors} & \multicolumn{1}{c|}{Few Processors} \\ 156 \hline 157 Many Threads & A: good performance & B: good performance \\ 158 \hline 159 Few Threads & C: poor performance & D: poor performance \\ 160 \hline 161 \end{tabular} 162 \end{center} 163 \caption{Performance of the relaxed FIFO list in different cases. The number of processors (many or few) refers to the number of kernel-threads \emph{actively} attempting to pop user-threads from the queues, not the total number of kernel-threads. The number of threads (many of few) refers to the number of user-threads ready to be run. Many threads means they outnumber processors significantly and most underlying queues have items, few threads means there are barely more threads than processors and most underlying queues are empty. Cases with fewer threads than processors are discussed in Section~\ref{sec:sleep}.} 164 \label{tab:perfcases} 165 \end{table} 166 167 Performance can be improved in case~D (Table~\ref{tab:perfcases}) by adding information to help processors find which inner queues are used. This aims to avoid the cost of retrying the pop operation but does not affect contention on the underlying queues and can incur some management cost for both push and pop operations. 168 169 A bitmask can be used to identify which inner queues are currently in use, as shown in Figure~\ref{fig:emptybit}. This means that processors can often find user-threads in constant time, regardless of how many underlying queues are empty. Furthermore, modern x86 CPUs have an extension (BMI2) which allow using the bitmask with very little overhead compared to a filled readyqueue, offerring decent performance even in cases with many empty inner queues. However, this technique has its limits, with a single word\footnote{Word refers here to however many bits can be written atomicly.} bitmask, the total number of underlying queues in the overarching queue is limited to the number of bits in the word. With a multi-word bitmask, this maximum limit can be increased arbitrarily, but it is not possible to check if the queue is empty by reading the bitmask atomicly. A dense bitmap, i.e., either a single word bitmask or a multi word bitmask where all words are densely packed, also causes additionnal problems in case~C (Table~\ref{tab:perfcases}), which the increased contention on the bitmask both causes new performance degradation and means the accuracy of the bitmask is less reliable due to more hardware threads potentially racing to read and/or update that information. 170 171 \begin{figure} 172 \begin{center} 173 {\resizebox{0.8\textwidth}{!}{\input{emptybit}}} 174 \end{center} 175 \caption{``More empty'' queue with added bitmask to indicate which array cells have items.} 176 \label{fig:emptybit} 177 \end{figure} 178 179 Another approach is to use a hiearchical data structure, for example Figure~\ref{fig:emptytree}. Creating a tree of nodes to reduce contention has been shown to work in similar cases\cit(SNZI: Scalable NonZero Indicators)\footnote{This particular paper seems to be patented in the US. How does that affect \CFA? Can I use it in my work?}. However, this approach may lead to poorer performance in case~B (Table~\ref{tab:perfcases}) due to the inherent pointer chasing cost and already low contention cost in that case. 180 181 \begin{figure} 182 \begin{center} 183 {\resizebox{0.8\textwidth}{!}{\input{emptytree}}} 184 \end{center} 185 \caption{``More empty'' queue with added binary search tree indicate which array cells have items.} 186 \label{fig:emptytree} 187 \end{figure} 188 189 Finally, a third approach is to use dense information, similar to the bitmap, but have each thread keep its own independant copies of it. While this approach can offer good scalability \emph{and} low latency, the livelyness of the information can become a problem. In the simple cases, local copies of which underlying queues are empty can become stale and end-up not be useful when for the pop operation. A more serious problem is that reliable information is necessary for some parts of this algorithm to be correct. As mentionned in this section, processors must know \emph{reliably} whether the list is empty or not to decide if they can return \texttt{NULL} or if they must keep looking during a pop operation. Section~\ref{sec:sleep} discusses an other case where reliable information is required for the algorithm to be correct. 190 191 There is a fundamental tradeoff among these approach. Dense global information about empty underlying queues will help zero-contention cases at the cost of high-contention case. Sparse global information will help high-contention cases but increase latency in zero-contention-cases, to read and ``aggregate'' the information\footnote{Hiearchical structures, e.g., binary search tree, effectively aggregate information but following pointer chains, learning information for each node. Similarly, other sparse schemes would need to read multiple cachelines to acquire all the information needed.}. Finally, dense local information has both the advantages of low latency in zero-contention cases and scalability in high-contention cases, however the information can become stale making it difficult to use to ensure correctness. The fact that these solutions have these fundamental limits suggest that that a more solution that combines these solutions in an interesting ways. The lock discussed in Section~\ref{sec:resize} also allows for solutions that adapt to the number of processors, which couls also prove useful. 126 192 127 193 \paragraph{Objectives and Existing Work} 128 How much scalability is actually needed is highly debatable, libfibre\cit is has compared favorably to other schedulers in webserver tests\cit and uses a single atomic counter in its scheduling algorithm similarly to the proposed bitmask. As such the single atomic instruction on a shared cacheline may be sufficiently performant. 129 130 I have built a prototype of this ready-queue (including the bitmask and BMI2 usage, but not the sharded bitmask) and ran performance experiments on it but it is difficult to compare this prototype to a thread scheduler as the prototype is used as a data-queue. I have also integrated this prototype into the \CFA runtime, but have not yet created performance experiments to compare results. I believe that the bitmask approach is currently one of the larger risks of the proposal, early tests lead me to believe it may work but it is not clear that the contention problem can be overcome. The worst-case scenario is a case where the number of processors and the number of ready threads are similar, yet scheduling events are very frequent. Fewer threads should lead to the Idle Sleep mechanism reducing contention while having many threads ready leads to optimal performance. It is difficult to evaluate the likeliness of this worst-case scenario in real workloads. I believe, frequent scheduling events suggest a more ``bursty'' workload where new work is finely divided among many threads which race to completion. This type of workload would only see a peek of contention close to the end of the work, but no sustained contention. Very fine-grained pipelines are less ``bursty'', these may lead to more sustained contention. However, they could also easily benefit from a direct hand-off strategy which would circumvent the problem entirely. 131 132 \subsection{Dynamic Resizing} 133 The \CFA runtime system currently handles dynamically adding and removing processors from clusters at any time. Since this is part of the existing design, the proposed scheduler must also support this behaviour. However, dynamicly resizing the clusters is considered a rare event associated with setup, teardown and major configuration changes. This assumptions is made both in the design of the proposed scheduler as well as in the original design of the \CFA runtime system. As such, the proposed scheduler must honor the correctness of these behaviour but does not have any performance objectives with regards to resizing a cluster. How long adding or removing processors take and how much this disrupts the performance of other threads is considered a secondary concern since it should be amortized over long period of times. This description effectively matches with te description of a Reader-Writer lock, in frequent but invasive updates among frequent (mostly) read operations. In the case of the Ready-Queue described above, read operations are operations that push or pop from the ready-queue but do not invalidate any references to the ready queue data structures. Writes on the other-hand would add or remove inner queues, invalidating references to the array of inner queues in the process. Therefore, the current proposed approach to this problem is the add a per-cluster Reader Writer lock around the ready queue to prevent restructuring of the ready-queue data structure while threads are being pushed or popped. 194 195 How much scalability is actually needed is highly debatable, libfibre\cit has compared favorably to other schedulers in webserver tests\cit and uses a single atomic counter in its scheduling algorithm similarly to the proposed bitmask. As such, the single atomic instruction on a shared cacheline may be sufficiently performant. 196 197 I have built a prototype of this ready-queue (including the bitmask and BMI2 usage, but not the sharded bitmask) and ran performance experiments on it but it is difficult to compare this prototype to a thread scheduler as the prototype is used as a data-queue. I have also integrated this prototype into the \CFA runtime, but have not yet created performance experiments to compare results. I believe that the bitmask approach is currently one of the larger risks of the proposal, early tests lead me to believe it may work but it is not clear that the contention problem can be overcome. The worst-case scenario is a case where the number of processors and the number of ready threads are similar, yet scheduling events are very frequent. Fewer threads should lead to the Idle Sleep mechanism discussed in Section~\ref{sec:sleep} to reduce contention while having many threads ready leads to optimal performance. It is difficult to evaluate the likeliness of this worst-case scenario in real workloads. I believe, frequent scheduling events suggest a more ``bursty'' workload where new work is finely divided among many threads which race to completion. This type of workload would only see a peek of contention close to the end of the work, but no sustained contention. Very fine-grained pipelines are less ``bursty'', these may lead to more sustained contention. However, they could also easily benefit from a direct hand-off strategy which would circumvent the problem entirely. 198 199 \subsection{Dynamic Resizing} \label{sec:resize} 200 The \CFA runtime system currently handles dynamically adding and removing processors from clusters at any time. Since this is part of the existing design, the proposed scheduler must also support this behaviour. However, dynamicly resizing the clusters is considered a rare event associated with setup, teardown and major configuration changes. This assumptions is made both in the design of the proposed scheduler as well as in the original design of the \CFA runtime system. As such, the proposed scheduler must honor the correctness of these behaviour but does not have any performance objectives with regards to resizing a cluster. How long adding or removing processors take and how much this disrupts the performance of other threads is considered a secondary concern since it should be amortized over long period of times. However, as mentionned in Section~\ref{sec:queue}, contention on the underlying queues can have a direct impact on performance, the number of underlying queues must therefore be adjusted as the number of processors grows or shrinks. Since the underlying queues are stored in a dense array, changing the number of queues requires resizing the array and therefore moving it. This can introduce memory reclamation problems if not done correctly. 201 202 \begin{figure} 203 \begin{center} 204 % {\resizebox{0.8\textwidth}{!}{\input{resize}}} 205 \input{resize} 206 \end{center} 207 \caption{Copy of data structure shown in Figure~\ref{fig:base}. The cells of the array can be modified concurrently but resizing the array, which requires moving it, is not safe to do concurrently. This can also be true of the accompanying data structures used to find non-empty queues.} 208 \label{fig:base2} 209 \end{figure} 210 211 It is important to note that how the array is used in this case. While the array cells are modified by every push and pop operation, the array itself, i.e., the pointer that would change when resized, is only read during these operations. Therefore the use is this pointer can be described as frequent reads and in frequent writes. This description effectively matches with the description of a Reader-Writer lock, infrequent but invasive updates among frequent read operations. In the case of the Ready-Queue described above, read operations are operations that push or pop from the ready-queue but do not invalidate any references to the ready queue data structures. Writes on the other-hand would add or remove inner queues, invalidating references to the array of inner queues in the process. Therefore, the current proposed approach to this problem is the add a per-cluster Reader Writer lock around the ready queue to prevent restructuring of the ready-queue data structure while threads are being pushed or popped. 134 212 135 213 There are possible alternatives to the Reader Writer lock solution. This problem is effectively a memory reclamation problem and as such there is a large body of research on the subject. However, the RWlock solution is simple and can be leveraged to solve other problems (e.g. processor ordering and memory reclamation of threads) which makes it an attractive solution. 136 214 137 215 \paragraph{Objectives and Existing Work} 138 The lock must offer scalability and performance on par with the actual ready-queue in order not to introduce a new bottle 139 140 \subsection{Idle Sleep} \label{s leep}141 As mentionned above, idle sleep is the process of putting processors to sleep while they do not have threads to execute. In this context processors are kernel-threads and sleeping refers to asking the kernel to block a thread. This can be achieved with either thread synchronization operations like pthread\_cond\_wait or using signal operations like sigsuspend.216 The lock must offer scalability and performance on par with the actual ready-queue in order not to introduce a new bottleneck. I have already built a lock that fits the desired requirements and preliminary testing show scalability and performance that exceed the target. As such, I do not consider this lock to be a risk on this project. 217 218 \subsection{Idle Sleep} \label{sec:sleep} 219 As mentionned above, idle sleep is the process of putting processors to sleep while they do not have threads to execute. In this context, processors are kernel-threads and sleeping refers to asking the kernel to block a thread. This can be achieved with either thread synchronization operations like pthread\_cond\_wait or using signal operations like sigsuspend. The goal of putting idle processors to sleep is two-fold, it reduces energy consumption in cases where more idle kernel-threads translate to idle hardware threads, and reduces contention on the ready queue, since the otherwise idle processors generally contend trying to pop items from the queue. Since energy efficiency is a growing concern in many industry sectors\cit, there is not real need to solve the contention problem without using idle sleep. 142 220 143 221 Support for idle sleep broadly involves calling the operating system to block the kernel thread but also handling the race between the sleeping and the waking up, and handling which kernel thread should sleep or wake-up. 144 222 145 When a processor decides to sleep, there is a race that occurs between it signalling that it will go to sleep (so other processors can find sleeping processors) and actually blocking the kernel thread. This is equivalent to the classic problem of missing signals when using condition variables, the ``sleepy'' processor indicates that it will sleep but has not yet gone to sleep, if another processor attempts to wake it up, the waking-up operation may claim nothing needs to be done and the signal will have been missed. In cases where threads are scheduled from processors on the current cluster, loosing signals is not necessarily critical, because at least some processors on the cluster are awake. Individual processors always finish s hceduling threads before looking for new work, which means that the last processor to go to sleep cannot miss threads scheduled from inside the cluster (if they do, that demonstrates the ready-queue is not linearizable). However, this guarantee does not hold if threads are shceduled from outside the cluster, either due to an external event like timers and I/O, or due to a thread migrating from a different cluster. In this case, missed signals can lead to the cluster deadlocking where it should not\footnote{Clusters ``should'' never deadlock, but for this proposal, cases where \CFA users \emph{actually} wrote \CFA code that leads to a deadlock it is considered as a deadlock that ``should'' happen. }. Therefore, it is important that the scheduling of threads include a mechanism where signals \emph{cannot} be missed. For performance reasons, it can be advantageous to have a secondary mechanism that allows signals to be missed in cases where it cannot lead to a deadlock. To be safe, this process must include a ``handshake'' where it is guaranteed that either~: the sleepy processor notices that a thread was scheduled after it signalled its intent to block or code scheduling threads well seethe intent to sleep before scheduling and be able to wake-up the processor. This matter is complicated by the fact that pthread offers few tools to implement this solution and offers no guarantee of ordering of threads waking up for most of these tools.146 147 Another issues is trying to avoid kernel sleeping and waking frequently. A possible partial solution is to order the processors so that the one which most recently went to sleep is woken up. This allows other sleeping processors to reach deeper sleep state (when these are available) while keeping ``hot'' processors warmer. Note that while this generally means organising the processors in a stack, I believe that the unique index provided by the ReaderWriter lock can be reused to strictly order the waking order of processors, causing a LIFO like waking order. While a strict LIFO stack is probably better, using the processor index could pro ove useful and offer a sufficiently LIFO ordering.223 When a processor decides to sleep, there is a race that occurs between it signalling that it will go to sleep (so other processors can find sleeping processors) and actually blocking the kernel thread. This is equivalent to the classic problem of missing signals when using condition variables, the ``sleepy'' processor indicates that it will sleep but has not yet gone to sleep, if another processor attempts to wake it up, the waking-up operation may claim nothing needs to be done and the signal will have been missed. In cases where threads are scheduled from processors on the current cluster, loosing signals is not necessarily critical, because at least some processors on the cluster are awake. Individual processors always finish scheduling threads before looking for new work, which means that the last processor to go to sleep cannot miss threads scheduled from inside the cluster (if they do, that demonstrates the ready-queue is not linearizable). However, this guarantee does not hold if threads are scheduled from outside the cluster, either due to an external event like timers and I/O, or due to a thread migrating from a different cluster. In this case, missed signals can lead to the cluster deadlocking where it should not\footnote{Clusters ``should'' never deadlock, but for this proposal, cases where \CFA users \emph{actually} wrote \CFA code that leads to a deadlock it is considered as a deadlock that ``should'' happen. }. Therefore, it is important that the scheduling of threads include a mechanism where signals \emph{cannot} be missed. For performance reasons, it can be advantageous to have a secondary mechanism that allows signals to be missed in cases where it cannot lead to a deadlock. To be safe, this process must include a ``handshake'' where it is guaranteed that either~: the sleepy processor notices that a thread was scheduled after it signalled its intent to block or code scheduling threads sees the intent to sleep before scheduling and be able to wake-up the processor. This matter is complicated by the fact that pthread offers few tools to implement this solution and offers no guarantee of ordering of threads waking up for most of these tools. 224 225 Another issues is trying to avoid kernel sleeping and waking frequently. A possible partial solution is to order the processors so that the one which most recently went to sleep is woken up. This allows other sleeping processors to reach deeper sleep state (when these are available) while keeping ``hot'' processors warmer. Note that while this generally means organising the processors in a stack, I believe that the unique index provided by the ReaderWriter lock can be reused to strictly order the waking order of processors, causing a LIFO like waking order. While a strict LIFO stack is probably better, using the processor index could prove useful and offer a sufficiently LIFO ordering. 148 226 149 227 Finally, another important aspect of Idle Sleep is when should processors make the decision to sleep and when it is appropriate for sleeping processors to be woken up. Processors that are unnecessarily awake lead to unnecessary contention and power consumption, while too many sleeping processors can lead to sub-optimal throughput. Furthermore, transitions from sleeping to awake and vice-versa also add unnecessary latency. There is already a wealth of research on the subject and I do not plan to implement a novel idea for the Idle Sleep heuristic in this project. … … 153 231 154 232 \paragraph{OS Abstraction} 155 One of the fundamental part of this converting blocking I/O operations into non-blocking ones. This relies on having an underlying asynchronous I/O interface to which to direct the I/O operations. While there exists many different APIs for asynchronous I/O, it is not part of this proposal to create a novel API, simply to use an existing one that is sufficient. uC++ uses the \texttt{select} as its interface, which handles pipes and sockets. It entails significant complexity and has performances problems which make it a less interesting alternative. Another interface which is becoming popular recently\cit is \texttt{epoll}. However, epoll also does not handle file system and seems to have problem to linux pipes and \texttt{TTY}s\cit. A very recent alternative that must still be investigated is \texttt{io\_uring}. It claims to address some of the issues with \texttt{epoll} but is too recent to be confident that it does. Finally, a popular cross-platform alternative is \texttt{libuv}, which offers asynchronous sockets and asynchronous file system operations (among other features). However, as a full-featured library it includes much more than what is needed and could conflict with other features of \CFA unless significant efforts are made to merge them together.233 One of the fundamental part of converting blocking I/O operations into non-blocking ones is having an underlying asynchronous I/O interface to direct the I/O operations. While there exists many different APIs for asynchronous I/O, it is not part of this proposal to create a novel API, simply to use an existing one that is sufficient. uC++ uses the \texttt{select} as its interface, which handles ttys, pipes and sockets, but not disk. It entails significant complexity and is being replaced which make it a less interesting alternative. Another interface which is becoming popular recently\cit is \texttt{epoll}, which is supposed to be cheaper than \texttt{select}. However, epoll also does not handle file system and seems to have problem to linux pipes and \texttt{TTY}s\cit. A very recent alternative that must still be investigated is \texttt{io\_uring}. It claims to address some of the issues with \texttt{epoll} but is too recent to be confident that it does. Finally, a popular cross-platform alternative is \texttt{libuv}, which offers asynchronous sockets and asynchronous file system operations (among other features). However, as a full-featured library it includes much more than what is needed and could conflict with other features of \CFA unless significant efforts are made to merge them together. 156 234 157 235 \paragraph{Event-Engine} … … 159 237 160 238 \paragraph{Interface} 161 Finally, for these components to be available, it is necessary to expose them through a synchronous interface. This can be a novel interface but it is preferrable to attempt to intercept the existing POSIX interface in order to be compatible with existing code. This will allowC programs written using this interface to be transparently converted to \CFA with minimal effeort. Where this is not applicable, a novel interface will be created to fill the gaps.239 Finally, for these components to be available, it is necessary to expose them through a synchronous interface. This can be a novel interface but it is preferrable to attempt to intercept the existing POSIX interface in order to be compatible with existing code. This allows C programs written using this interface to be transparently converted to \CFA with minimal effeort. Where this is not applicable, a novel interface will be created to fill the gaps. 162 240 163 241 … … 171 249 \section{Timeline} 172 250 173 174 \cleardoublepage175 251 176 252 % B I B L I O G R A P H Y 177 253 % ----------------------------- 178 \addcontentsline{toc}{chapter}{Bibliography} 254 \cleardoublepage 255 \phantomsection % allows hyperref to link to the correct page 256 \addcontentsline{toc}{section}{\refname} 179 257 \bibliographystyle{plain} 180 258 \bibliography{pl,local} 259 260 % G L O S S A R Y 261 % ----------------------------- 181 262 \cleardoublepage 182 263 \phantomsection % allows hyperref to link to the correct page 183 184 % G L O S S A R Y 185 % ----------------------------- 186 \addcontentsline{toc}{chapter}{Glossary} 264 \addcontentsline{toc}{section}{Glossary} 187 265 \printglossary 188 \cleardoublepage189 \phantomsection % allows hyperref to link to the correct page190 266 191 267 \end{document} -
doc/user/user.tex
r331eacbe r87f572e 11 11 %% Created On : Wed Apr 6 14:53:29 2016 12 12 %% Last Modified By : Peter A. Buhr 13 %% Last Modified On : Thu Mar 5 12:09:42 202014 %% Update Count : 3 88513 %% Last Modified On : Fri Mar 6 13:34:52 2020 14 %% Update Count : 3924 15 15 %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% 16 16 … … 6621 6621 An array may be filled, resized, or aligned. 6622 6622 \end{description} 6623 \VRef[Table]{t:AllocationVersusCapabilities} shows allocation routines supporting different combinations of storage-management capabilities :6623 \VRef[Table]{t:AllocationVersusCapabilities} shows allocation routines supporting different combinations of storage-management capabilities. 6624 6624 \begin{table} 6625 6625 \centering 6626 \begin{minipage}{0.75\textwidth} 6626 6627 \begin{tabular}{@{}r|l|l|l|l|l@{}} 6627 6628 \multicolumn{1}{c}{}& & \multicolumn{1}{c|}{fill} & resize & alignment & array \\ … … 6631 6632 & ©realloc© & copy & yes & no & no \\ 6632 6633 & ©memalign© & no & no & yes & no \\ 6633 & ©aligned_alloc© & no & no & yes & no \\ 6634 & ©aligned_alloc©\footnote{Same as ©memalign© but size is an integral multiple of alignment, which is universally ignored.} 6635 & no & no & yes & no \\ 6634 6636 & ©posix_memalign© & no & no & yes & no \\ 6637 & ©valloc© & no & no & yes (page size)& no \\ 6638 & ©pvalloc©\footnote{Same as ©valloc© but rounds size to multiple of page size.} 6639 & no & no & yes (page size)& no \\ 6635 6640 \hline 6636 6641 \CFA & ©cmemalign© & yes (0 only) & no & yes & yes \\ 6637 6642 & ©realloc© & copy & yes & yes & no \\ 6638 & ©alloc© & no & no & no & no \\ 6639 & ©alloc© & copy & no/yes & no & yes \\ 6640 & ©alloc© & no/copy/yes & no/yes & no & yes \\ 6641 & ©alloc_set© & no/yes & no & yes & yes \\ 6642 & ©alloc_align© & no/yes & no & yes & yes \\ 6643 & ©alloc_align_set© & no/yes & no & yes & yes \\ 6643 & ©alloc© & no & yes & no & yes \\ 6644 & ©alloc_set© & yes & yes & no & yes \\ 6645 & ©alloc_align© & no & yes & yes & yes \\ 6646 & ©alloc_align_set© & yes & yes & yes & yes \\ 6644 6647 \end{tabular} 6648 \end{minipage} 6645 6649 \caption{Allocation Routines versus Storage-Management Capabilities} 6646 6650 \label{t:AllocationVersusCapabilities} 6647 6651 \end{table} 6652 6653 \CFA memory management extends the type safety of all allocations by using the type of the left-hand-side type to determine the allocation size and return a matching type for the new storage. 6654 Type-safe allocation is provided for all C allocation routines and new \CFA allocation routines, \eg in 6655 \begin{cfa} 6656 int * ip = (int *)malloc( sizeof(int) ); §\C{// C}§ 6657 int * ip = malloc(); §\C{// \CFA type-safe version of C malloc}§ 6658 int * ip = alloc(); §\C{// \CFA type-safe uniform alloc}§ 6659 \end{cfa} 6660 the latter two allocations determine the allocation size from the type of ©p© (©int©) and cast the pointer to the allocated storage to ©int *©. 6661 6662 \CFA memory management extends allocation safety by implicitly honouring all alignment requirements, \eg in 6663 \begin{cfa} 6664 struct S { int i; } __attribute__(( aligned( 128 ) )); // cache-line alignment 6665 S * sp = malloc(); §\C{// honour type alignment}§ 6666 \end{cfa} 6667 the storage allocation is implicitly aligned to 128 rather than the default 16. 6668 The alignment check is performed at compile time so there is no runtime cost. 6648 6669 6649 6670 \CFA memory management extends the resize capability with the notion of \newterm{sticky properties}. … … 6652 6673 Without sticky properties it is dangerous to use ©realloc©, resulting in an idiom of manually performing the reallocation to maintain correctness. 6653 6674 6675 \CFA memory management extends allocation to support constructors for initialization of allocated storage, \eg in 6676 \begin{cfa} 6677 struct S { int i; }; §\C{// cache-line aglinment}§ 6678 void ?{}( S & s, int i ) { s.i = i; } 6679 // assume ?|? operator for printing an S 6680 6681 S & sp = *®new®( 3 ); §\C{// call constructor after allocation}§ 6682 sout | sp.i; 6683 ®delete®( &sp ); 6684 6685 S * spa = ®anew®( 10, 5 ); §\C{// allocate array and initialize each array element}§ 6686 for ( i; 10 ) sout | spa[i] | nonl; 6687 sout | nl; 6688 ®adelete®( 10, spa ); 6689 \end{cfa} 6690 Allocation routines ©new©/©anew© allocate a variable/array and initialize storage using the allocated type's constructor. 6691 Note, the matching deallocation routines ©delete©/©adelete©. 6654 6692 6655 6693 \leavevmode
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