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2 | \documentclass[11pt,epsf,a4wide]{article} |
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3 | \usepackage{../../style/cerco} |
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4 | \usepackage{pdfpages} |
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5 | |
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6 | \usepackage{graphics} |
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7 | |
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8 | % For SLNCS comment above and use |
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9 | % \documentclass{llncs} |
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10 | |
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11 | |
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12 | |
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13 | |
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14 | |
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15 | \RequirePackage[latin1]{inputenc} |
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16 | |
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17 | % Mettre les différents packages et fonctions que l'on utilise |
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18 | \usepackage[english]{babel} |
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19 | \usepackage{amsmath} |
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20 | \usepackage{amsfonts} |
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21 | \usepackage{amssymb} |
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22 | \usepackage{xspace} |
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23 | \usepackage{latexsym} |
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24 | \usepackage{url} |
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25 | \usepackage{xspace} |
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26 | %\usepackage{fancyvrb} |
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27 | \usepackage[all]{xy} |
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28 | %packages pour LNCS |
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29 | %\usepackage{semantic} |
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30 | %\usepackage{cmll} |
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31 | % Packages for RR |
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32 | \usepackage{graphics,color} |
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33 | \RequirePackage[latin1]{inputenc} |
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34 | \usepackage{array} |
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35 | |
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36 | %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% |
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37 | %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% |
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38 | |
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39 | \newenvironment{comment}{{\bf MORE WORK:}} |
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40 | |
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41 | \newenvironment{restate-proposition}[2][{}]{\noindent\textbf{Proposition~{#2}} |
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42 | \;\textbf{#1}\ |
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43 | }{\vskip 1em} |
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44 | |
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45 | \newenvironment{restate-theorem}[2][{}]{\noindent\textbf{Theorem~{#2}}\;\textbf{ |
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46 | #1}\ |
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47 | }{\vskip 1em} |
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48 | |
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49 | \newenvironment{restate-corollary}[2][{}]{\noindent\textbf{Corollary~{#2}} |
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50 | \;\textbf{#1}\ |
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51 | }{\vskip 1em} |
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52 | |
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53 | \newcommand{\myparagraph}[1]{\medskip\noindent\textbf{#1}} |
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54 | |
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55 | \newcommand{\Proofitemb}[1]{\medskip \noindent {\bf #1\;}} |
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56 | \newcommand{\Proofitemfb}[1]{\noindent {\bf #1\;}} |
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57 | \newcommand{\Proofitem}[1]{\medskip \noindent $#1\;$} |
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58 | \newcommand{\Proofitemf}[1]{\noindent $#1\;$} |
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59 | \newcommand{\Defitem}[1]{\smallskip \noindent $#1\;$} |
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60 | \newcommand{\Defitemt}[1]{\smallskip \noindent {\em #1\;}} |
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61 | \newcommand{\Defitemf}[1]{\noindent $#1\;$} |
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62 | |
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63 | |
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64 | %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% |
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65 | %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% |
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66 | |
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67 | |
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68 | |
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69 | %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% |
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70 | %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% |
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71 | |
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72 | \newcommand{\eqdef}{=_{\text{def}}} |
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73 | \newcommand{\concat}{\cdot}%%{\mathbin{+}} |
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74 | \newcommand{\Int}{\mathit{int}} |
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75 | \newcommand{\nat}{\mathit{nat}} |
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76 | \newcommand{\String}{\mathit{string}} |
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77 | \newcommand{\Ident}{\mathit{ident}} |
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78 | \newcommand{\Block}{\mathit{block}} |
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79 | \newcommand{\Signature}{\mathit{signature}} |
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80 | |
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81 | \newcommand{\pc}{\mathit{pc}} |
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82 | \newcommand{\estack}{\mathit{estack}} |
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83 | \newcommand{\Error}{\epsilon} |
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84 | |
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85 | %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% |
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86 | %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% |
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87 | |
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88 | % --------------------------------------------------------------------- % |
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89 | % Proof rule. % |
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90 | % --------------------------------------------------------------------- % |
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91 | |
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92 | \newcommand{\staterule}[3]{% |
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93 | $\begin{array}{@{}l}% |
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94 | \mbox{#1}\\% |
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95 | \begin{array}{c} |
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96 | #2\\ |
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97 | \hline |
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98 | \raisebox{0ex}[2.5ex]{\strut}#3% |
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99 | \end{array} |
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100 | \end{array}$} |
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101 | |
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102 | \newcommand{\GAP}{2ex} |
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103 | |
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104 | \newcommand{\recall}[2]{% |
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105 | $\begin{array}{c} |
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106 | #1\\ |
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107 | \hline |
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108 | \raisebox{0ex}[2.5ex]{\strut}#2% |
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109 | \end{array}$} |
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110 | |
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111 | \newcommand{\hbra}{\noindent\hbox to \textwidth{\leaders\hrule height1.8mm |
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112 | depth-1.5mm\hfill}} |
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113 | \newcommand{\hket}{\noindent\hbox to \textwidth{\leaders\hrule |
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114 | height0.3mm\hfill}} |
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115 | \newcommand{\ratio}{.3} |
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116 | |
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117 | \newenvironment{display}[1]{\begin{tabbing} |
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118 | \hspace{1.5em} \= \hspace{\ratio\linewidth-1.5em} \= \hspace{1.5em} \= \kill |
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119 | \noindent\hbra\\[-.5em] |
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120 | {\ }\textsc{#1}\\[-.8ex] |
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121 | \hbox to \textwidth{\leaders\hrule height1.6mm depth-1.5mm\hfill}\\[-.8ex] |
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122 | }{\\[-.8ex]\hket |
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123 | \end{tabbing}} |
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124 | |
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125 | |
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126 | \newcommand{\sbline}{\hfill\smash[t]{\rule[1.5em]{\textwidth}{0.2ex} |
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127 | \hfill\hspace*{0ex}}} |
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128 | \newcommand{\sline}{\hfill\smash[t]{\rule[1.5em]{\textwidth}{0.1ex} |
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129 | \hfill\hspace*{0ex}}} |
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130 | \newcommand{\sentry}[2]{\>$#1$\>\ \smash[t]{\vrule width 0.2mm height |
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131 | 1.2\baselineskip depth 1.5\baselineskip}\>#2} |
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132 | |
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133 | \newcommand{\entry}[2]{\>$#1$\>\>#2} |
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134 | \newcommand{\clause}[2]{$#1$\>\>#2} |
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135 | \newcommand{\category}[2]{\clause{#1::=}{#2}} |
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136 | \newcommand{\subclause}[1]{\>\>\>#1} |
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137 | \newcommand{\redrule}[3]{$#1$\>\>$#2$\>\>\>#3} |
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138 | |
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139 | %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% |
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140 | %%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%%% |
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141 | |
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142 | % environments |
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143 | |
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144 | % To be commented for LNCS |
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145 | \newtheorem{theorem}{Theorem} |
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146 | \newtheorem{fact}[theorem]{Fact} |
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147 | \newtheorem{definition}[theorem]{Definition} |
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148 | \newtheorem{lemma}[theorem]{Lemma} |
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149 | \newtheorem{corollary}[theorem]{Corollary} |
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150 | \newtheorem{proposition}[theorem]{Proposition} |
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151 | \newtheorem{example}[theorem]{Example} |
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152 | \newtheorem{exercise}[theorem]{Exercise} |
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153 | \newtheorem{remark}[theorem]{Remark} |
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154 | \newtheorem{question}[theorem]{Question} |
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155 | \newtheorem{proviso}[theorem]{Proviso} |
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156 | \newtheorem{conjecture}[theorem]{Conjecture} |
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157 | |
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158 | % proofs |
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159 | |
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160 | \newcommand{\Proof}{\noindent {\sc Proof}. } |
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161 | \newcommand{\Proofhint}{\noindent {\sc Proof hint}. } |
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162 | % To be commented for LNCS |
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163 | \newcommand{\qed}{\hfill${\Box}$} |
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164 | \newcommand{\EndProof}{\qed} |
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165 | |
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166 | % figure environment |
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167 | |
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168 | \newcommand{\Figbar}{{\center \rule{\hsize}{0.3mm}}} |
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169 | %horizontal thiner line for figures |
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170 | \newenvironment{figureplr}[1][t]{\begin{figure}[#1] \Figbar}{\Figbar \end{figure}} |
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171 | %environment for figures |
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172 | % ************Macros for mathematical symbols************* |
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173 | % Style |
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174 | |
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175 | \newcommand{\cl}[1]{{\cal #1}} % \cl{R} to make R calligraphic |
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176 | \newcommand{\la}{\langle} % the brackets for pairing (see also \pair) |
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177 | \newcommand{\ra}{\rangle} |
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178 | |
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179 | \newcommand{\lf}{\lfloor} |
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180 | \newcommand{\rf}{\rfloor} |
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181 | \newcommand{\ul}[1]{\underline{#1}} % to underline |
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182 | \newcommand{\ol}[1]{\overline{#1}} % to overline |
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183 | \newcommand{\ok}{~ok} % well formed context |
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184 | |
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185 | % Syntax |
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186 | |
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187 | \newcommand{\Gives}{\vdash} % in a type judgment |
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188 | \newcommand{\IGives}{\vdash_{I}} % intuitionistic provability |
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189 | \newcommand{\AIGives}{\vdash_{{\it AI}}} %affine-intuitionistic provability |
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190 | \newcommand{\CGives}{\vdash_{C}} % affine-intuitionistic confluent provability |
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191 | |
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192 | |
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193 | \newcommand{\Models}{\mid \! =} % models |
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194 | |
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195 | \newcommand{\emptycxt}{\On} % empty context |
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196 | \newcommand{\subs}[2]{[#1 / #2]} |
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197 | \newcommand{\sub}[2]{[#2 / #1]} % substitution \sub{x}{U} gives [U/x] |
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198 | |
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199 | \newcommand{\Sub}[3]{[#3 / #2]#1} % Substitution with three arguments \Sub{V}{x}{U} |
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200 | |
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201 | \newcommand{\lsub}[2]{#2 / #1} % substitution \lsub{x}{U} gives U/x, to be used in a list. |
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202 | |
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203 | \newcommand{\impl}{\supset} |
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204 | \newcommand{\arrow}{\rightarrow} % right thin arrow |
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205 | \newcommand{\trarrow}{\stackrel{*}{\rightarrow}} % trans closure |
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206 | %\newcommand{\limp}{\makebox[5mm]{\,$- \! {\circ}\,$}} % linear |
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207 | % implication |
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208 | \newcommand{\limp}{\multimap} %linear implication |
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209 | \newcommand{\bang}{\, !} |
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210 | % LNCS |
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211 | %\newcommand{\bang}{\oc} |
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212 | \newcommand{\limpe}[1]{\stackrel{#1}{\multimap}} |
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213 | \newcommand{\hyp}[3]{#1:(#2, #3)} |
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214 | \newcommand{\letm}[3]{{\sf let} \ ! #1 = #2 \ {\sf in} \ #3} % modal let |
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215 | \newcommand{\lets}[3]{{\sf let} \ #1 = #2 \ {\sf in} \ #3} % simple let |
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216 | \newcommand{\letp}[3]{{\sf let} \ \S #1 = #2 \ {\sf in} \ #3} % paragraph let |
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217 | \newcommand{\tertype}{{\bf 1}} |
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218 | \newcommand{\behtype}{{\bf B}} |
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219 | \newcommand{\bt}[1]{{\it BT}(#1)} % Boehm tree |
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220 | \newcommand{\cxt}[1]{#1[~]} % Context with one hole |
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221 | \newcommand{\pr}{\parallel} % parallel || |
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222 | \newcommand{\Nat}{\mathbf{N}} % natural numbers |
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223 | \newcommand{\Natmax}{\mathbf{N}_{{\it max}}} % natural numbers with minus infinity |
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224 | \newcommand{\Rat}{\mathbf{Q}^{+}} % non-negative rationals |
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225 | \newcommand{\Ratmax}{\mathbf{Q}^{+}_{{\it max}}} % non-negative rationals with minus infinity |
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226 | \newcommand{\Alt}{ \mid\!\!\mid } |
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227 | \newcommand{\isum}{\oplus} |
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228 | \newcommand{\csum}{\uplus} %context sum |
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229 | \newcommand{\dpar}{\mid\!\mid} |
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230 | % for the production of a grammar containing \mid |
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231 | \newcommand{\infer}[2]{\begin{array}{c} #1 \\ \hline #2 \end{array}} |
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232 | % to make a centered inference rule |
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233 | |
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234 | % (Meta-)Logic |
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235 | |
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236 | \newcommand{\bool}{{\sf bool}} % boolean values |
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237 | \newcommand{\Or}{\vee} % disjunction |
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238 | \newcommand{\OR}{\bigvee} % big disjunction |
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239 | \newcommand{\AND}{\wedge} % conjunction |
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240 | \newcommand{\ANDD}{\bigwedge} % big conjunction |
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241 | \newcommand{\Arrow}{\Rightarrow} % right double arrow |
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242 | \newcommand{\IFF}{\mbox{~~iff~~}} % iff in roman and with spaces |
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243 | \newcommand{\iffArrow}{\Leftrightarrow} % logical equivalence |
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244 | |
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245 | % Semantics |
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246 | |
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247 | \newcommand{\dl}{[\![} % semantic [[ |
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248 | \newcommand{\dr}{]\!]} % semantic ]] |
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249 | \newcommand{\lam}{{\bf \lambda}} % semantic lambda |
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250 | |
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251 | |
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252 | % The equivalences for this paper |
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253 | |
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254 | % the usual ones |
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255 | \newcommand{\ubis}{\approx^u} % usual labelled weak bis |
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256 | \newcommand{\uabis}{\approx^{u}_{ccs}} % usual labelled weak bis on CCS |
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257 | |
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258 | % the contextual conv sensitive |
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259 | \newcommand{\cbis}{\approx} % convergence sensitive bis |
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260 | \newcommand{\cabis}{\approx_{ccs}} % convergence sensitive bis on CCS |
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261 | |
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262 | % the labelled conv sensitive |
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263 | \newcommand{\lcbis}{\approx^{\ell}} % |
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264 | \newcommand{\lcabis}{\approx^{\ell}_{ccs}} % labelled convergence sensitive bis on CCS |
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265 | \newcommand{\lcbiswrong}{\approx^{\ell \Downarrow}} % |
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266 | |
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267 | |
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268 | |
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269 | |
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270 | |
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271 | |
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272 | |
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273 | \newcommand{\maytest}{=_{\Downarrow}} |
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274 | \newcommand{\musttest}{=_{\Downarrow_{S}}} |
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275 | |
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276 | |
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277 | |
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278 | |
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279 | % Sets |
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280 | |
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281 | \newcommand{\prt}[1]{{\cal P}(#1)} % Parts of a set |
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282 | \newcommand{\finprt}[1]{{\cal P}_{fin}(#1)}% Finite parts |
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283 | \newcommand{\finprtp}[1]{{\cal P}_{fin}^{+}(#1)}% Non-empty Finite parts |
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284 | \newcommand{\union}{\cup} % union |
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285 | \newcommand{\inter}{\cap} % intersection |
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286 | \newcommand{\Union}{\bigcup} % big union |
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287 | \newcommand{\Inter}{\bigcap} % big intersection |
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288 | \newcommand{\cpl}[1]{#1^{c}} % complement |
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289 | \newcommand{\card}{\sharp} % cardinality |
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290 | \newcommand{\minus}{\backslash} % set difference |
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291 | \newcommand{\sequence}[2]{\{#1\}_{#2}} % ex. \sequence{d_n}{n\in \omega} |
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292 | \newcommand{\comp}{\circ} % functional composition |
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293 | %\newcommand{\oset}[1]{\{#1\}} % set enumeration |
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294 | \newcommand{\mset}[1]{\{\! | #1 |\!\}} % pseudo-set notation {| |} |
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295 | |
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296 | % Domains |
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297 | |
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298 | \newcommand{\two}{{\bf O}} % Sierpinski space |
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299 | \newcommand{\join}{\vee} % join |
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300 | \newcommand{\JOIN}{\bigvee} % big join |
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301 | \newcommand{\meet}{\wedge} % meet |
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302 | \newcommand{\MEET}{\bigwedge} % big meet |
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303 | \newcommand{\dcl}{\downarrow} % down closure |
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304 | \newcommand{\ucl}{\uparrow} % up closure |
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305 | \newcommand{\conv}{\downarrow} % synt. conv. pred. (postfix) |
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306 | \newcommand{\diver}{\uparrow} % diverging term |
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307 | \newcommand{\Conv}{\Downarrow} % sem. conv. pred. (postfix) |
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308 | \newcommand{\SConv}{\Downarrow_{S}} % sem. conv. pred. (postfix) |
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309 | \newcommand{\CConv}{\Downarrow_{C}} |
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310 | \newcommand{\Diver}{\Uparrow} % diverging map |
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311 | \newcommand{\cpt}[1]{{\cal K}(#1)} % compacts, write \cpt{D} |
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312 | \newcommand{\ret}{\triangleleft} % retract |
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313 | \newcommand{\nor}{\succeq} |
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314 | \newcommand{\prj}{\underline{\ret}} % projection |
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315 | \newcommand{\parrow}{\rightharpoonup} % partial function space |
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316 | \newcommand{\ub}[1]{{\it UB}(#1)} % upper bounds |
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317 | \newcommand{\mub}[1]{{\it MUB}(#1)} % minimal upper bounds |
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318 | \newcommand{\lift}[1]{(#1)_{\bot}} % lifting |
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319 | \newcommand{\forget}[1]{\underline{#1}} % forgetful translation |
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320 | |
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321 | %\newcommand{\rel}[1]{\;{\cal #1}\;} % infix relation (calligraphic) |
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322 | \newcommand{\rl}[1]{\;{\cal #1}\;} % infix relation |
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323 | \newcommand{\rel}[1]{{\cal #1}} %calligraphic relation with no |
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324 | %extra space |
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325 | \newcommand{\per}[1]{\;#1 \;} |
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326 | \newcommand{\wddagger}{\natural} % weak suspension |
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327 | %\newcommand{\wddagger}{=\!\!\!\!\parallel} % weak suspension |
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328 | % Categories |
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329 | |
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330 | \newcommand{\pair}[2]{\langle #1 , #2 \rangle} % pairing \pair{x}{y}, do not use < >. |
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331 | |
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332 | % ******* Notation for the $\pi$-calculus ********* |
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333 | % Syntax: |
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334 | |
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335 | |
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336 | \newcommand{\fn}[1]{{\it fn}(#1)} % free names |
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337 | \newcommand{\bn}[1]{{\it bn}(#1)} % bound names |
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338 | \newcommand{\names}[1]{{\it n}(#1)} % names |
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339 | \newcommand{\true}{{\sf t}} % true |
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340 | \newcommand{\false}{{\sf f}} % false |
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341 | \newcommand{\pio}{\pi_1} % 1 receptor calculus |
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342 | \newcommand{\pioo}{\pi_{1}^{r}} |
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343 | \newcommand{\piom}{\pi_{1}^{-}} % 1 receptor calculus wo match |
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344 | \newcommand{\pioi}{\pi_{1I}} % 1 receptor I-calculus |
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345 | \newcommand{\pifo}{\pi_{\w{1f}}} % functional calculus |
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346 | \newcommand{\pilo}{\pi_{\w{1l}}} % located calculus |
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347 | \newcommand{\sort}[1]{{\it st}(#1)} % sort |
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348 | \newcommand{\ia}[1]{{\it ia}(#1)} % sort |
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349 | \newcommand{\ite}[3]{{\sf if~} #1 {\sf ~then~} #2 {\sf ~else~} #3} %if then else |
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350 | \newcommand{\casep}[2]{{\sf case}^{\times}(#1, \pair{x}{y}\Arrow#2)} %case on pairs |
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351 | \newcommand{\casel}[3]{{\sf case}^{L}(#1, #2, \s{cons}(x,y)\Arrow#3)} %case on lists |
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352 | \newcommand{\caseb}[3]{{\sf case}^{b}(#1, #2, \s{cons}(x,y)\Arrow#3)} %case on lists |
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353 | \newcommand{\nil}{{\sf nil}} |
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354 | \newcommand{\cons}{{\sf cons}} |
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355 | \newcommand{\idle}[1]{{\it Idle}(#1)} %idle process |
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356 | \newcommand{\conf}[1]{\{ #1 \}} %configuration |
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357 | \newcommand{\link}[2]{#1 \mapsto #2} %likn a ->b |
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358 | \newcommand{\mand}{\mbox{ and }} |
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359 | \newcommand{\dvec}[1]{\tilde{{\bf #1}}} %double vector |
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360 | \newcommand{\erloc}[1]{{\it er}_{l}(#1)} % location erasure |
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361 | \newcommand{\w}[1]{{\it #1}} %To write in math style |
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362 | \newcommand{\vcb}[1]{{\bf #1}} |
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363 | \newcommand{\lc}{\langle\!|} |
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364 | \newcommand{\rc}{|\!\rangle} |
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365 | \newcommand{\obj}[1]{{\it obj}(#1)} |
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366 | \newcommand{\move}[1]{{\sf move}(#1)} |
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367 | \newcommand{\qqs}[2]{\forall\, #1\;\: #2} |
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368 | \newcommand{\qtype}[4]{\forall #1 : #2 . (#4,#3)} |
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369 | \newcommand{\xst}[2]{\exists\, #1\;\: #2} |
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370 | \newcommand{\xstu}[2]{\exists\, ! #1\;\: #2} |
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371 | \newcommand{\dpt}{\,:\,} |
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372 | \newcommand{\cond}[3]{\mathsf{if}\ #1\ \mathsf{then}\ #2\ \mathsf{else}\ #3} |
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373 | \newcommand{\s}[1]{{\sf #1}} % sans-serif |
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374 | \newcommand{\vc}[1]{{\bf #1}} |
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375 | \newcommand{\lnorm}{\lbrack\!\lbrack} |
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376 | \newcommand{\rnorm}{\rbrack\!\rbrack} |
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377 | \newcommand{\sem}[1]{\underline{#1}} |
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378 | \newcommand{\tra}[1]{\langle #1 \rangle} |
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379 | \newcommand{\trb}[1]{[ #1 ]} |
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380 | \newcommand{\squn}{\mathop{\scriptstyle\sqcup}} |
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381 | \newcommand{\lcro}{\langle\!|} |
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382 | \newcommand{\rcro}{|\!\rangle} |
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383 | \newcommand{\semi}[1]{\lcro #1\rcro} |
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384 | \newcommand{\sell}{\,\ell\,} |
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385 | \newcommand{\SDZ}[1]{\marginpar{\textbf{SDZ:} {#1}}} |
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386 | |
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387 | \newcommand{\when}[3]{{\sf when}~#1~{\sf then}~#2~{\sf else}~#3} |
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388 | \newcommand{\wthen}[2]{{\sf when}~#1~{\sf then}~#2~} |
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389 | \newcommand{\welse}[1]{{\sf else}~#1} |
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390 | |
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391 | %Pour la fleche double, il faut rajouter : |
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392 | % \usepackage{mathtools} |
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393 | |
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394 | \newcommand{\act}[1]{\xrightarrow{#1}} %labelled actionlow %high |
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395 | |
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396 | \newcommand{\lact}[1]{\stackrel{#1}{\makebox[5mm]{\,$- \! {\circ}\,$}}} |
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397 | |
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398 | \newcommand{\set}[1]{\{#1\}} |
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399 | \newcommand{\pst}[2]{{\sf pset}(#1,#2)} |
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400 | \newcommand{\st}[2]{{\sf set}(#1,#2)} |
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401 | \newcommand{\wrt}[2]{{\sf w}(#1,#2)} |
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402 | |
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403 | \newcommand{\chtype}[2]{{\it Ch_{#1}(#2)}} |
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404 | \newcommand{\rgtype}[2]{{\it {\sf Reg}_{#1} #2}} |
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405 | |
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406 | \newcommand{\get}[1]{{\sf get}(#1)} |
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407 | |
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408 | %\newcommand{\wact}[1]{\xRightarrow{#1}} %weak labelled action low high |
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409 | |
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410 | %\newcommand{\mact}[1]{\xrightarrow{#1}_{m}} %labelled action low %high |
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411 | |
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412 | %\newcommand{\wmact}[1]{\xRightarrow{#1}_{m}} %weak labelled action low high |
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413 | |
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414 | %\newcommand{\act}[1]{\stackrel{#1}{\rightarrow}} %labelled action low |
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415 | %%%high |
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416 | |
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417 | \newcommand{\acteq}[1]{\stackrel{#1}{\leadsto}} %labelled action low |
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418 | %%%high |
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419 | |
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420 | |
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421 | %\newcommand{\actI}[1]{\stackrel{#1}{\rightarrow_{1}}} %labelled action low |
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422 | \newcommand{\actI}[1]{\xrightarrow{#1}_{1}} |
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423 | |
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424 | %\newcommand{\actII}[1]{\stackrel{#1}{\rightarrow_{2}}} %labelled action low |
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425 | \newcommand{\actII}[1]{\xrightarrow{#1}_{2}} |
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426 | |
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427 | |
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428 | \newcommand{\wact}[1]{\stackrel{#1}{\Rightarrow}} %weak labelled action low high |
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429 | \newcommand{\wactI}[1]{\stackrel{#1}{\Rightarrow_{1}}} %weak labelled action low high |
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430 | \newcommand{\wactII}[1]{\stackrel{#1}{\Rightarrow_{2}}} %weak labelled action low high |
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431 | |
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432 | |
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433 | \newcommand{\mact}[1]{\stackrel{#1}{\rightarrow_{m}}} %labelled action low |
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434 | %high |
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435 | \newcommand{\wmact}[1]{\stackrel{#1}{\Rightarrow_{m}}} %weak labelled action low high |
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436 | |
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437 | %\newcommand{\lact}[1]{\stackrel{#1}{\leftarrow}} |
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438 | \newcommand{\lwact}[1]{\stackrel{#1}{\Leftarrow}} |
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439 | |
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440 | |
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441 | |
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442 | \newcommand{\eval}{\Downarrow} |
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443 | \newcommand{\Eval}[1]{\Downarrow^{#1}} |
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444 | |
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445 | |
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446 | \newcommand{\Z}{{\bf Z}} |
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447 | \newcommand{\Real}{\mathbb{R}^{+}} |
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448 | \newcommand{\Return}{\ensuremath{\mathtt{return}}\xspace} |
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449 | \newcommand{\Stop}{\ensuremath{\mathtt{stop}}\xspace} |
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450 | \newcommand{\Wait}{\ensuremath{\mathtt{wait}}\xspace} |
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451 | \newcommand{\Read}{\ensuremath{\mathtt{read}}\xspace} |
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452 | \newcommand{\Write}{\ensuremath{\mathtt{write}}\xspace} |
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453 | \newcommand{\Yield}{\ensuremath{\mathtt{yield}}\xspace} |
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454 | \newcommand{\Next}{\ensuremath{\mathtt{next}}\xspace} |
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455 | \newcommand{\Load}{\ensuremath{\mathtt{load}}\xspace} |
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456 | \newcommand{\Call}{\ensuremath{\mathtt{call}}\xspace} |
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457 | \newcommand{\Tcall}{\ensuremath{\mathtt{tcall}}\xspace} |
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458 | \newcommand{\Pop}{\ensuremath{\mathtt{pop}}\xspace} |
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459 | \newcommand{\Build}{\ensuremath{\mathtt{build}}\xspace} |
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460 | \newcommand{\Branch}{\ensuremath{\mathtt{branch}}\xspace} |
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461 | \newcommand{\Goto}{\ensuremath{\mathtt{goto}}\xspace} |
---|
462 | |
---|
463 | \newcommand{\hatt}[1]{#1^{+}} |
---|
464 | \newcommand{\Of}{\mathbin{\w{of}}} |
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465 | |
---|
466 | \newcommand{\susp}{\downarrow} |
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467 | \newcommand{\lsusp}{\Downarrow_L} |
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468 | \newcommand{\wsusp}{\Downarrow} |
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469 | \newcommand{\commits}{\searrow} |
---|
470 | |
---|
471 | |
---|
472 | \newcommand{\spi}{S\pi} |
---|
473 | |
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474 | |
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475 | \newcommand{\pres}[2]{#1\triangleright #2} %TCCS else next (alternative) |
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476 | % \newcommand{\pres}[2]{ \lfloor #1 \rfloor (#2)} %TCCS else next |
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477 | \newcommand{\present}[3]{{\sf present} \ #1 \ {\sf do } \ #2 \ {\sf else} \ #3} |
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478 | |
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479 | |
---|
480 | \newcommand{\tick}{{\sf tick}} %tick action |
---|
481 | |
---|
482 | |
---|
483 | |
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484 | \newcommand{\sbis}{\equiv_L} |
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485 | \newcommand{\emit}[2]{\ol{#1}#2} |
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486 | %\newcommand{\present}[4]{#1(#2).#3,#4} |
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487 | \newcommand{\match}[4]{[#1=#2]#3,#4} %pi-equality |
---|
488 | |
---|
489 | \newcommand{\matchv}[4]{[#1 \unrhd #2]#3,#4} |
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490 | |
---|
491 | \newcommand{\new}[2]{\nu #1 \ #2} |
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492 | \newcommand{\outact}[3]{\new{{\bf #1}}{\emit{#2}{#3}}} |
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493 | \newcommand{\real}{\makebox[5mm]{\,$\|\!-$}}% realizability relation |
---|
494 | |
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495 | \newcommand{\regterm}[2]{{\sf reg}_{#1} #2} |
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496 | \newcommand{\thread}[1]{{\sf thread} \ #1} |
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497 | \newcommand{\store}[2]{(#1 \leftarrow #2)} |
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498 | \newcommand{\pstore}[2]{(#1 \Leftarrow #2)} |
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499 | |
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500 | \newcommand{\regtype}[2]{{\sf Reg}_{#1} #2} |
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501 | \newcommand{\uregtype}[3]{{\sf Reg}_{#1}(#2, #3)} |
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502 | \newcommand{\urtype}[2]{{\sf Reg}(#1, #2)} |
---|
503 | |
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504 | \newcommand{\upair}[2]{[#1,#2]} |
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505 | \newcommand{\letb}[3]{\mathsf{let}\;\oc #1 = #2\;\mathsf{in}\;#3} |
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506 | |
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507 | \newcommand{\vlt}[1]{{\cal V}(#1)} |
---|
508 | \newcommand{\prs}[1]{{\cal P}(#1)} |
---|
509 | |
---|
510 | \newcommand{\imp}{{\sf Imp}} %imp language |
---|
511 | \newcommand{\vm}{{\sf Vm}} %virtual machine language |
---|
512 | \newcommand{\mips}{{\sf Mips}} %Mips language |
---|
513 | \newcommand{\C}{{\sf C}} % C language |
---|
514 | \newcommand{\Clight}{{\sf Clight}} %C light language |
---|
515 | \newcommand{\Cminor}{{\sf Cminor}} |
---|
516 | \newcommand{\RTLAbs}{{\sf RTLAbs}} |
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517 | \newcommand{\RTL}{{\sf RTL}} |
---|
518 | \newcommand{\ERTL}{{\sf ERTL}} |
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519 | \newcommand{\LTL}{{\sf LTL}} |
---|
520 | \newcommand{\LIN}{{\sf LIN}} |
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521 | \newcommand{\access}[1]{\stackrel{#1}{\leadsto}} |
---|
522 | \newcommand{\ocaml}{{\sf ocaml}} |
---|
523 | \newcommand{\coq}{{\sf Coq}} |
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524 | \newcommand{\compcert}{{\sf CompCert}} |
---|
525 | %\newcommand{\cerco}{{\sf CerCo}} |
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526 | \newcommand{\cil}{{\sf CIL}} |
---|
527 | \newcommand{\scade}{{\sf Scade}} |
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528 | \newcommand{\absint}{{\sf AbsInt}} |
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529 | \newcommand{\framac}{{\sf Frama-C}} |
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530 | \newcommand{\powerpc}{{\sf PowerPc}} |
---|
531 | \newcommand{\lustre}{{\sf Lustre}} |
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532 | \newcommand{\esterel}{{\sf Esterel}} |
---|
533 | \newcommand{\ml}{{\sf ML}} |
---|
534 | |
---|
535 | \newcommand{\codeex}[1]{\texttt{#1}} % code example |
---|
536 | |
---|
537 | \bibliographystyle{abbrv} |
---|
538 | |
---|
539 | \title{ |
---|
540 | INFORMATION AND COMMUNICATION TECHNOLOGIES\\ |
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541 | (ICT)\\ |
---|
542 | PROGRAMME\\ |
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543 | \vspace*{1cm}Project FP7-ICT-2009-C-243881 \cerco{}} |
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544 | |
---|
545 | \date{ } |
---|
546 | \author{} |
---|
547 | %>>>>>> new commands used in this document |
---|
548 | % \usepackage[nohyperref,nosvn]{mystyle} |
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549 | |
---|
550 | \usepackage{multirow} |
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551 | \newcolumntype{b}{@{}>{{}}} |
---|
552 | \newcolumntype{B}{@{}>{{}}c<{{}}@{}} |
---|
553 | \newcolumntype{h}[1]{@{\hspace{#1}}} |
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554 | \newcolumntype{L}{>{$}l<{$}} |
---|
555 | \newcolumntype{C}{>{$}c<{$}} |
---|
556 | \newcolumntype{R}{>{$}r<{$}} |
---|
557 | \newcolumntype{S}{>{$(}r<{)$}} |
---|
558 | \newcolumntype{n}{@{}} |
---|
559 | \newcommand{\spanr}[2]{\multicolumn{1}{Rn}{\multirow{#1}{*}{(#2)}}} |
---|
560 | \def\nocol{\multicolumn{1}{ncn}{}} |
---|
561 | |
---|
562 | \usepackage[disable, colorinlistoftodos]{todonotes} |
---|
563 | \usepackage{enumerate} |
---|
564 | \usepackage{tikz} |
---|
565 | |
---|
566 | \newcommand{\tern}[3]{#1\mathrel ? #2 : #3} |
---|
567 | \newcommand{\sop}[1]{\s{#1}\ } |
---|
568 | \newcommand{\sbin}[1]{\ \s{#1}\ } |
---|
569 | \newcommand{\Ell}{\mathcal L} |
---|
570 | \newcommand{\alphab}{\boldsymbol\alpha} |
---|
571 | \newcommand{\betab}{\boldsymbol\beta} |
---|
572 | \newcommand{\gramm}{\mathrel{::=}} |
---|
573 | \newcommand{\ass}{\mathrel{:=}} |
---|
574 | |
---|
575 | \renewcommand{\to}[1][]{\stackrel{#1}{\rightarrow}} |
---|
576 | %<<<<<<<<<<<< |
---|
577 | \begin{document} |
---|
578 | % \thispagestyle{empty} |
---|
579 | % |
---|
580 | % \vspace*{-1cm} |
---|
581 | % \begin{center} |
---|
582 | % \includegraphics[width=0.6\textwidth]{../style/cerco_logo.png} |
---|
583 | % \end{center} |
---|
584 | % |
---|
585 | % \begin{minipage}{\textwidth} |
---|
586 | % \maketitle |
---|
587 | % \end{minipage} |
---|
588 | % |
---|
589 | % |
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590 | % \vspace*{0.5cm} |
---|
591 | % \begin{center} |
---|
592 | % \begin{LARGE} |
---|
593 | % \bf |
---|
594 | % Report \\ |
---|
595 | % Dependent Cost Labels |
---|
596 | % \\ |
---|
597 | % \end{LARGE} |
---|
598 | % \end{center} |
---|
599 | % |
---|
600 | % \vspace*{2cm} |
---|
601 | % \begin{center} |
---|
602 | % \begin{large} |
---|
603 | % Version 0.1 |
---|
604 | % \end{large} |
---|
605 | % \end{center} |
---|
606 | % |
---|
607 | % \vspace*{0.5cm} |
---|
608 | % \begin{center} |
---|
609 | % \begin{large} |
---|
610 | % Main Author:\\ |
---|
611 | % Paolo Tranquilli |
---|
612 | % \end{large} |
---|
613 | % \end{center} |
---|
614 | % |
---|
615 | % \vspace*{\fill} |
---|
616 | % \noindent |
---|
617 | % Project Acronym: \cerco{}\\ |
---|
618 | % Project full title: Certified Complexity\\ |
---|
619 | % Proposal/Contract no.: FP7-ICT-2009-C-243881 \cerco{}\\ |
---|
620 | % |
---|
621 | % \clearpage \pagestyle{myheadings} \markright{\cerco{}, FP7-ICT-2009-C-243881} |
---|
622 | % |
---|
623 | % \newpage |
---|
624 | |
---|
625 | \listoftodos |
---|
626 | |
---|
627 | \section{Introduction} |
---|
628 | In~\cite{D2.1} an approach is presented tackling the problem of building a |
---|
629 | compiler from a large fragment of C which is capable of lifting execution cost information |
---|
630 | from the compiled code and present it to the user. This endeavour is central to |
---|
631 | the CerCo project, which strives to produce a mechanically certified version of |
---|
632 | such a compiler. |
---|
633 | |
---|
634 | In rough words, the proposed approach consists in decorating the source code with |
---|
635 | labels at key points, and preserving such labels along the compilation chain. |
---|
636 | Once the final object code is produced, such labels should correspond to parts |
---|
637 | of the compiled code which have constant cost. |
---|
638 | |
---|
639 | There are two properties one asks of the cost estimate. The first, paramount to |
---|
640 | the correctness of the method, is \emph{soundness}: the actual execution cost |
---|
641 | is bounded by the estimate. In the labelling approach, this is guaranteed if every |
---|
642 | loop in the control flow of the compiled code passes through at least one cost |
---|
643 | label. The second, optional but desirable, is \emph{preciseness}: the estimate |
---|
644 | is not in fact an estimate but the actual cost. In the labelling approach, this will |
---|
645 | be true if for every label every possible execution of the compiled code starting |
---|
646 | from such a label yields the same cost. In simple architectures such as the 8051 |
---|
647 | micro-controller this can be guaranteed by having labels at the immediate start of any |
---|
648 | branch of the control flow, and by ensuring no duplicate labels are present. |
---|
649 | |
---|
650 | It should be underlined that the above mentioned requirements must hold when |
---|
651 | executing the code at the |
---|
652 | end of the compilation chain. If one is careful to inject the labels at good |
---|
653 | key places in the source code, one can still think of two main obstacles: |
---|
654 | \begin{itemize} |
---|
655 | \item compilation introduces important changes in the control flow, inserting |
---|
656 | loops or branches: an example in addressing this is the insertion of functions |
---|
657 | in the source code replacing instructions unavailable in the target architecture |
---|
658 | that require loops (e.g.\ multi-word division and generic shift in the 8051 |
---|
659 | architecture) so that the annotation process be sound, or the effort put in |
---|
660 | providing unbranching translations of higher level instructions~ |
---|
661 | \cite{D2.2}; |
---|
662 | \item even if the compiled code \emph{does}, as long as the the syntactic |
---|
663 | control flow graph is concerned, respect the conditions for soundness and |
---|
664 | preciseness, the cost of blocks of instructions might not be independent of context, |
---|
665 | so that different passes through a label might have different costs: this |
---|
666 | becomes a concern if one wishes to apply the approach to more complex architectures, |
---|
667 | for example one with cache or pipelining. |
---|
668 | \end{itemize} |
---|
669 | |
---|
670 | The first point unveils a weakness of the current labelling approach when it |
---|
671 | comes to some common code transformations done along a compilation chain. In |
---|
672 | particular most of \emph{loop optimizations} disrupt such conditions directly. |
---|
673 | For example in what we will call \emph{loop peeling} a first iteration of the |
---|
674 | loop is hoisted ahead of it, possibly to have a different cost than later iterations |
---|
675 | because of further transformation such as dead code elimination or invariant |
---|
676 | code motion. |
---|
677 | |
---|
678 | The second point strikes a difference in favour to existing tools for the |
---|
679 | estimate of execution costs to the detriment of CerCo's approach advocated in~ |
---|
680 | \cite{D2.1}. We will take as example the well known tool \s{aiT}~\cite{absint}, |
---|
681 | based on abstract interpretation: such a tool allows to estimate the |
---|
682 | worst-case execution time (WCET) taking into account advanced aspects of the |
---|
683 | target architecture. \s{aiT}'s ability to |
---|
684 | do close estimates of execution costs even when these depend on the context of |
---|
685 | execution would enhance the effectiveness of CerCo's compiler, either by |
---|
686 | integrating such techniques in its development, or by directly calling this tool |
---|
687 | when ascertaining the cost of blocks of compiled code. A typical case where |
---|
688 | cache analysis yields a difference in the execution cost of a block is in loops: |
---|
689 | the first iteration will usually stumble upon more cache misses than subsequent |
---|
690 | iterations. |
---|
691 | |
---|
692 | If one looks closely, the source of the weakness of the labelling approach as |
---|
693 | presented in~\cite{D2.1} is common to both points: the inability to state |
---|
694 | different costs for different occurrences of labels, where the difference |
---|
695 | might be originated by labels being duplicated along compilation |
---|
696 | or the costs being sensitive to the state of execution. |
---|
697 | |
---|
698 | The preliminary work |
---|
699 | we present here addresses this node, introducing cost labels that are dependent |
---|
700 | on which iteration of its containing loops it occurs in. This is achieved by |
---|
701 | means of \emph{indexed labels}: all cost labels are decorated with formal |
---|
702 | indexes coming from the loops containing such labels. These indexes allow to |
---|
703 | rebuild, even after several steps of loop transformations, |
---|
704 | which iterations of the original loops in the source code a particular label |
---|
705 | occurrence belongs to. During annotation of source code, this information |
---|
706 | is given to the user by means of dependent costs. |
---|
707 | |
---|
708 | We concentrate on integrating the labelling approach with two loop transformations. |
---|
709 | For general information on compiler optimization (and loop optimizations in |
---|
710 | particular) we refer the reader to the vast literature on the subject |
---|
711 | (e.g.~\cite{muchnick,morgan}). |
---|
712 | |
---|
713 | \emph{Loop peeling}, as already mentioned, consists in preceding the loop with |
---|
714 | a copy of its body, appropriately guarded. This is in general useful to trigger |
---|
715 | further optimizations, such as ones relying on execution information which can |
---|
716 | be computed at compile time but which is erased by further iterations of the loop, |
---|
717 | or ones that use the hoisted code to be more effective at eliminating redundant |
---|
718 | code. Integrating this transformation to the labelling approach would also allow |
---|
719 | to integrate the common case of cache analysis explained above: the analysis |
---|
720 | of cache hits and misses in the case of usually benefits from a form of |
---|
721 | \emph{virtual} loop peeling~\cite{cacheprediction}. |
---|
722 | |
---|
723 | \emph{Loop unrolling} consists in repeating several times the body of the loop |
---|
724 | inside the loop itself (inserting appropriate guards, or avoiding them altogether |
---|
725 | if enough information on the loop's guard is available at compile time). This |
---|
726 | can limit the number of (conditional or unconditional) jumps executed by the |
---|
727 | code and trigger further optimizations dealing with pipelining, if applicable |
---|
728 | to the architecture. |
---|
729 | |
---|
730 | While covering only two loop optimizations, the work presented here poses good |
---|
731 | bases to extending the labelling approach to cover more and more of them, as well |
---|
732 | as giving hints as to how integrating in CerCo's compiler advanced cost estimation |
---|
733 | techniques such as cache analysis. Moreover loop peeling has the substantial |
---|
734 | advantage of enhancing other optimizations. Experimentation with CerCo's |
---|
735 | untrusted prototype compiler with constant propagation and |
---|
736 | partial redundancy elimination~\cite{PRE,muchnick} show how loop peeling enhances |
---|
737 | those other optimizations. |
---|
738 | |
---|
739 | \paragraph{Outline.} |
---|
740 | We will present our approach on a minimal toy imperative language, \imp{} |
---|
741 | with gotos, which we present in \autoref{sec:defimp} along with formal |
---|
742 | definition of the loop transformations. This language already |
---|
743 | presents most of the difficulties encountered when dealing with C, so |
---|
744 | we stick to it for the sake of this presentation. In \autoref{sec:labelling} |
---|
745 | we summarize the labelling approach as presented in~\cite{D2.1}. The following |
---|
746 | \autoref{sec:indexedlabels} explains \emph{indexed labels}, our proposal for |
---|
747 | dependent labels which are able to describe precise costs even in the presence |
---|
748 | of the loop transformations we consider. Finally \autoref{sec:conc} goes into |
---|
749 | more details regarding the implementation of indexed labels in CerCo's |
---|
750 | untrusted compiler and speculates on further work on the subject. |
---|
751 | |
---|
752 | |
---|
753 | \section{\imp{} with goto}\label{sec:defimp} |
---|
754 | We briefly outline the toy language which contains all the relevant instructions |
---|
755 | to present the development and the problems it is called to solve. |
---|
756 | |
---|
757 | The version of the minimal imperative language \imp that we will present has, |
---|
758 | with respect to the bare-bone usual version, \s{goto}s and labelled statements. |
---|
759 | Break and continue statements can be added at no particular expense. Its syntax |
---|
760 | and operational semantics is presented in \autoref{fig:minimp}. |
---|
761 | \begin{figureplr} |
---|
762 | $$\begin{gathered} |
---|
763 | \begin{array}{nlBl>(R<)n} |
---|
764 | \multicolumn{4}{C}{\bfseries Syntax}\\ |
---|
765 | \multicolumn{4}{ncn}{ |
---|
766 | \ell,\ldots \hfill \text{(labels)} \hfill x,y,\ldots \hfill |
---|
767 | \text{(identifiers)} |
---|
768 | \hfill e,f,\ldots \hfill \text{(expression)} |
---|
769 | }\\ |
---|
770 | P,S,T,\ldots &\gramm& \s{skip} \mid s;t |
---|
771 | \mid \sop{if}e\sbin{then}s\sbin{else}t |
---|
772 | \mid \sop{while} e \sbin{do} s \mid |
---|
773 | x \ass e |
---|
774 | \\&\mid& |
---|
775 | \ell : s \mid \sop{goto}\ell& \spanr{-2}{statements}\\ |
---|
776 | \\ |
---|
777 | \multicolumn{4}{C}{\bfseries Semantics}\\ |
---|
778 | K,\ldots &\gramm& \s{halt} \mid S \cdot K & continuations |
---|
779 | \end{array} |
---|
780 | \\[15pt] |
---|
781 | \s{find}(\ell,S,K) \ass |
---|
782 | \left\{\begin{array}{lL} |
---|
783 | \bot & if $S=\s{skip},\sop{goto} \ell'$ or $x\ass e$,\\ |
---|
784 | (T, K) & if $S=\ell:T$,\\ |
---|
785 | \s{find}(\ell,T,K) & otherwise, if $S = \ell':T$,\\ |
---|
786 | \s{find}(\ell,T_1,T_2\cdot K) & if defined and $S=T_1;T_2$,\\ |
---|
787 | \s{find}(\ell,T_1,K) & if defined and |
---|
788 | $S=\sop{if}b\sbin{then}T_1\sbin{else}T_2$,\\ |
---|
789 | \s{find}(\ell,T_2,K) & otherwise, if $S=T_1;T_2$ or |
---|
790 | $\sop{if}b\sbin{then}T_1\sbin{else}T_2$,\\ |
---|
791 | \s{find}(\ell,T,S\cdot K) & if $S = \sop{while}b\sbin{do}T$. |
---|
792 | \end{array}\right. |
---|
793 | \\[15pt] |
---|
794 | \begin{array}{lBl} |
---|
795 | (x:=e,K,s) &\to_P& (\s{skip},K,s[v/x]) \qquad\mbox{if }(e,s)\eval v \\ \\ |
---|
796 | |
---|
797 | (S;T,K,s) &\to_P& (S,T\cdot K,s) \\ \\ |
---|
798 | |
---|
799 | (\s{if} \ b \ \s{then} \ S \ \s{else} \ T,K,s) |
---|
800 | &\to_P&\left\{ |
---|
801 | \begin{array}{ll} |
---|
802 | (S,K,s) &\mbox{if }(b,s)\eval v \neq 0 \\ |
---|
803 | (T,K,s) &\mbox{if }(b,s)\eval 0 |
---|
804 | \end{array} |
---|
805 | \right. \\ \\ |
---|
806 | |
---|
807 | |
---|
808 | (\s{while} \ b \ \s{do} \ S ,K,s) |
---|
809 | &\to_P&\left\{ |
---|
810 | \begin{array}{ll} |
---|
811 | (S,\s{while} \ b \ \s{do} \ S \cdot K,s) &\mbox{if }(b,s)\eval v \neq 0 \\ |
---|
812 | (\s{skip},K,s) &\mbox{if }(b,s)\eval 0 |
---|
813 | \end{array} |
---|
814 | \right. \\ \\ |
---|
815 | |
---|
816 | |
---|
817 | (\s{skip},S\cdot K,s) &\to_P&(S,K,s) \\ \\ |
---|
818 | |
---|
819 | (\ell : S, K, s) &\to_P& (S,K,s) \\ \\ |
---|
820 | |
---|
821 | (\sop{goto}\ell,K,s) &\to_P& (\s{find}(\ell,P,\s{halt}),s) \\ \\ |
---|
822 | \end{array} |
---|
823 | \end{gathered}$$ |
---|
824 | |
---|
825 | |
---|
826 | \caption{The syntax and operational semantics of \imp.} |
---|
827 | \label{fig:minimp} |
---|
828 | \end{figureplr} |
---|
829 | The actual |
---|
830 | grammar for expressions is not particularly relevant so we do not give a |
---|
831 | precise one. For the sake of this presentation we also treat boolean and |
---|
832 | arithmetic expressions together (with the usual convention of an expression |
---|
833 | being true iff non-zero). |
---|
834 | |
---|
835 | We will suppose programs are |
---|
836 | \emph{well-labelled}, i.e.\ every label labels at most one occurrence of statement |
---|
837 | in the program, and every \s{goto} points to a label actually present in the program. |
---|
838 | The \s{find} helper functions has the task of not only finding the labelled |
---|
839 | statement in the program, but also building the correct continuation. The continuation |
---|
840 | built by \s{find} replaces the current continuation in the case of a jump. |
---|
841 | |
---|
842 | \paragraph{Further down the compilation chain.} |
---|
843 | As for the rest of the compilation chain, we abstract over it. We just posit |
---|
844 | every language $L$ further down the compilation chain has a suitable notion of |
---|
845 | sequential instructions (with one natural successor), to which we can add our |
---|
846 | own. |
---|
847 | |
---|
848 | \subsection{Loop transformations} |
---|
849 | We call a loop $L$ \emph{single-entry} in $P$ if there is no \s{goto} of $P$ |
---|
850 | outside of $L$ which jumps to within $L$\footnote{This is a reasonable |
---|
851 | aproximation: it considers multi-entry also those loops having an external but |
---|
852 | unreachable \s{goto} jumping to them.}. Many loop optimizations do not preserve |
---|
853 | the semantics of multi-entry loops in general, or are otherwise rendered |
---|
854 | ineffective. Usually compilers implement a multi-entry loop detection which |
---|
855 | avoids those loops from being targeted by optimizations~\cite{muchnick,morgan}. |
---|
856 | |
---|
857 | |
---|
858 | \paragraph{Loop peeling.} |
---|
859 | $$ |
---|
860 | \sop{while}b\sbin{do}S\mapsto |
---|
861 | \sop{if}b\sbin{then} S; \sop{while} b \sbin{do} S[\ell'_i/\ell_i] |
---|
862 | $$ |
---|
863 | where $\ell'_i$ is a fresh label for any $\ell_i$ labelling a statement in $S$. |
---|
864 | This relabelling is safe as to \s{goto}s external to the loop because of the |
---|
865 | single-entry condition. Notice that in case of break and continue statements, |
---|
866 | those should be replaced with \s{goto}s in the peeled body $S$. |
---|
867 | |
---|
868 | \paragraph{Loop unrolling.} |
---|
869 | $$ |
---|
870 | \sop{while}b\sbin{do}S\mapsto |
---|
871 | \sop{while} b \sbin{do} (S ; |
---|
872 | \sop{if} b \sbin{then} (S[\ell^1_i/\ell_i] ; |
---|
873 | \cdots |
---|
874 | \sop{if} b \sbin{then} S[\ell^n_i/\ell_i]) \cdots) |
---|
875 | $$ |
---|
876 | where $\ell^j_i$ are again fresh labels for any $\ell_i$ labelling a statement |
---|
877 | in $S$. This is a willingly naïf version of loop unrolling, which usually |
---|
878 | targets less general loops. The problem it poses to Cerco's labelling approach |
---|
879 | are independent of the cleverness of the actual transformation. |
---|
880 | |
---|
881 | \section{Labelling: a quick sketch of the previous approach}\label{sec:labelling} |
---|
882 | Plainly labelled \imp{} is obtained adding to the code \emph{cost labels} |
---|
883 | (with metavariables $\alpha,\beta,\ldots$), and cost-labelled statements: |
---|
884 | $$S,T\gramm \cdots \mid \alpha: S$$ |
---|
885 | Cost labels allow for some program points to be tracked along the compilation |
---|
886 | chain. For further details we refer to\cite{D2.1}. |
---|
887 | |
---|
888 | With labels the small step semantics turns into a labelled transition |
---|
889 | system and a natural notion of trace (i.e.\ lists of labels) arises. |
---|
890 | Evaluation of statements is enriched with traces, so that rules are like |
---|
891 | $$ |
---|
892 | \begin{array}{lblL} |
---|
893 | (\alpha: S, K,s) &\to[\alpha]_P (S,K,s)\\ |
---|
894 | (\s{skip}, S \cdot K,s) &\to[\varepsilon]_P (S, K, s)\\ |
---|
895 | & \text{etc.} |
---|
896 | \end{array}$$ |
---|
897 | where we identify cost labels $\alpha$ with singleton traces and we use $\varepsilon$ |
---|
898 | for the empty trace. We then write $\to[\lambda]\!\!^*$ for the transitive |
---|
899 | closure of the small step semantics which produces by concatenation the trace |
---|
900 | $\lambda$. |
---|
901 | |
---|
902 | \paragraph{Labelling.} |
---|
903 | Given an \imp{} program $P$ its \emph{labelling} $\alpha:\Ell(P)$ in $\ell-\imp$ |
---|
904 | is |
---|
905 | defined by putting cost labels after every branching, at the start of both |
---|
906 | branches, and a cost label at the beginning of the program. So the relevant |
---|
907 | cases are |
---|
908 | $$\begin{aligned} |
---|
909 | \Ell(\sop{if}e\sbin{then}S\sbin{else}T) &= |
---|
910 | \sop{if}e\sbin{then}\alpha:\Ell(S)\sbin{else}\beta:\Ell(T)\\ |
---|
911 | \Ell(\sop{while}e\sbin{do}S) &= |
---|
912 | (\sop{while}e\sbin{do}\alpha:\Ell(S));\beta:\s{skip} |
---|
913 | \end{aligned}$$ |
---|
914 | where $\alpha,\beta$ are fresh cost labels, and while in other cases the |
---|
915 | definition just passes to sub-statements. |
---|
916 | |
---|
917 | \paragraph{Labels in the rest of the compilation chain.} All languages further |
---|
918 | down the chain get a new sequential statement $\sop{emit}\alpha$ whose effect is |
---|
919 | to be consumed in a labelled transition while keeping the same state. All other |
---|
920 | instructions guard their operational semantics and do not emit cost labels. |
---|
921 | |
---|
922 | Preservation of semantics throughout the compilation process is restated, in |
---|
923 | rough terms, as |
---|
924 | $$\text{starting state of $P$}\to[\lambda]\!\!^*\;\text{halting state} \iff |
---|
925 | \text{starting state of $\mathcal C(P)$} \to[\lambda]\!\!^*\;\text{halting state},$$ |
---|
926 | where $P$ is a program of a language along the compilation chain, starting and |
---|
927 | halting states depend on the language, and $\mathcal C$ is the |
---|
928 | compilation function\footnote{The case of divergent computations needs |
---|
929 | to be addressed too. Also, the requirement can be weakened by demanding some |
---|
930 | sort of equivalence of the traces rather than equality. Both these issues escape |
---|
931 | the scope of this presentation.}. |
---|
932 | |
---|
933 | \paragraph{Instrumentations} |
---|
934 | Let $\mathcal C$ be the whole compilation from $\ell\imp$ to the labelled |
---|
935 | version of some low-level language $L$. Supposing such compilation has not |
---|
936 | introduced any new loop or branching, we have that |
---|
937 | \begin{itemize} |
---|
938 | \item every loop contains at least a cost label (\emph{soundness condition}), |
---|
939 | and |
---|
940 | \item every branching has different labels for the two branches |
---|
941 | (\emph{preciseness condition}). |
---|
942 | \end{itemize} |
---|
943 | With these two conditions, we have that each and every cost label in |
---|
944 | $\mathcal C(P)$ for any $P$ corresponds to a block of sequential instructions, |
---|
945 | to which we can assign a constant \emph{cost}\footnote{This in fact requires the |
---|
946 | machine architecture to be simple enough, or for some form of execution analysis |
---|
947 | to take place.} We can therefore assume a \emph{cost mapping} $\kappa_P$ from |
---|
948 | cost labels to natural numbers, assigning to each cost label $\alpha$ the cost |
---|
949 | of the block containing the single occurrance of $\alpha$. |
---|
950 | |
---|
951 | Given any cost mapping $\kappa$, we can enrich a labelled program so that a |
---|
952 | particular fresh variable (the \emph{cost variable} $c$) keeps track of the |
---|
953 | assumulation of costs during the execution. We call this procedure |
---|
954 | \emph{instrumentation} of the program, and it is defined recursively by |
---|
955 | $$ |
---|
956 | \mathcal I(\alpha:S) = c \ass c + \kappa(\alpha) ; \mathcal I(S) |
---|
957 | $$ |
---|
958 | while in all other cases the definition passes to substatements. |
---|
959 | |
---|
960 | \paragraph{The problem with loop optimizations.} |
---|
961 | Let us take loop peeling, and apply it to the labelling of a program without any |
---|
962 | adjustment: |
---|
963 | $$ |
---|
964 | (\sop{while}e\sbin{do}\alpha:S);\beta:\s{skip} |
---|
965 | \mapsto |
---|
966 | (\sop{if}b\sbin{then} \alpha : S; \sop{while} b \sbin{do} \alpha : |
---|
967 | S[\ell'_i/\ell_i]); |
---|
968 | \beta:\s{skip}$$ |
---|
969 | What happens is that the cost label $\alpha$ is duplicated into two distinct |
---|
970 | occurrences. If these two occurrences correspond to different costs in the |
---|
971 | compiled code, the best the cost mapping can do is to take the maximum of the |
---|
972 | two, preserving soundness (i.e.\ the cost estimate still bounds the actual one) |
---|
973 | but loosing preciseness (i.e.\ the actual cost would be strictly less than its |
---|
974 | estimate). |
---|
975 | |
---|
976 | \section{Indexed labels}\label{sec:indexedlabels} |
---|
977 | This section presents the core of the new approach. In brief points it amounts |
---|
978 | to the following. |
---|
979 | \begin{enumerate}[\bfseries \ref*{sec:indexedlabels}.1.] |
---|
980 | \item\label{en:outline1} |
---|
981 | Enrich cost labels with formal indexes corresponding, at the beginning of |
---|
982 | the process, to which iteration of the loop they belong to. |
---|
983 | \item\label{en:outline2} |
---|
984 | Each time a loop transformation is applied and a cost labels is split in |
---|
985 | different occurrences, each of these will be reindexed so that every time they |
---|
986 | are emitted their position in the original loop will be reconstructed. |
---|
987 | \item\label{en:outline3} |
---|
988 | Along the compilation chain, add alongside the \s{emit} instruction other |
---|
989 | ones updating the indexes, so that iterations of the original loops can be |
---|
990 | rebuilt at the operational semantics level. |
---|
991 | \item\label{en:outline4} |
---|
992 | The machinery computing the cost mapping will still work, but assigning |
---|
993 | costs to indexed cost labels, rather than to cost labels as we wish: however |
---|
994 | \emph{dependent costs} can be calculated, where dependency is on which iteration |
---|
995 | of the containing loops we are in. |
---|
996 | \end{enumerate} |
---|
997 | |
---|
998 | \subsection{Indexing the cost labels}\label{ssec:indlabs} |
---|
999 | \paragraph{Formal indexes and $\iota\ell\imp$.} |
---|
1000 | Let $i_0,i_1,\ldots$ be a sequence of distinguished fresh identifiers that will |
---|
1001 | be used as loop indexes. A \emph{simple expression} is an affine arithmetical |
---|
1002 | expression in one of these indexes, that is $a*i_k+b$ with $a,b,k \in \mathbb N$. |
---|
1003 | Simple expressions $e_1=a_1*i_k+b_1$, $e_2=a2*i_k+b_2$ in the same index can be |
---|
1004 | composed, yielding $e_1\circ e_2\ass (a_1a_2)*i_k + (a_1b2+b_1)$, and this operation |
---|
1005 | has an identity element in $id_k \ass 1*i_k+0$. Constants can be expressed as simple |
---|
1006 | expressions, so that we identify a natural $c$ with $0*i_k+c$. |
---|
1007 | |
---|
1008 | An \emph{indexing} (with metavariables $I$, $J$, \ldots) is a list of |
---|
1009 | transformations of successive formal indexes dictated by simple expressions, |
---|
1010 | that is a mapping% |
---|
1011 | \footnote{Here we restrict each mapping to be a simple expression |
---|
1012 | \emph{on the same index}. This might not be the case if more loop optimizations |
---|
1013 | are accounted for (for example, interchanging two nested loops).} |
---|
1014 | $$i_0\mapsto a_0*i_0+b_0,\dots, i_{k-1} \mapsto a_{k-1}*i_{k-1}+b_{k-1}.$$ |
---|
1015 | |
---|
1016 | An \emph{indexed cost label} (metavariables $\alphab$, $\betab$, \ldots) is |
---|
1017 | the combination of a cost label $\alpha$ and an indexing $I$, written |
---|
1018 | $\alpha\la I\ra$. The cost label underlying an indexed one is called its |
---|
1019 | \emph{atom}. All plain labels can be considered as indexed ones by taking |
---|
1020 | an empty indexing. |
---|
1021 | |
---|
1022 | \imp{} with indexed labels ($\iota\ell\imp$) is defined by adding to $\imp$ |
---|
1023 | statements with indexed labels, and by having loops with formal indexes |
---|
1024 | attached to them: |
---|
1025 | $$S,T,\ldots \gramm \cdots i_k:\sop{while}e\sbin{do}S\mid \alphab : S.$$ |
---|
1026 | Notice than unindexed loops still are in the language: they will correspond |
---|
1027 | to multi-entry loops which are ignored by indexing and optimizations. |
---|
1028 | We will discuss the semantics later. |
---|
1029 | |
---|
1030 | \paragraph{Indexed labelling.} |
---|
1031 | Given an $\imp$ program $P$, in order to index loops and assign indexed labels |
---|
1032 | we must first of all distinguish single-entry loops. We just sketch how it can |
---|
1033 | be computed. |
---|
1034 | |
---|
1035 | A first pass of the program $P$ can easily compute two maps: $\s{loopof}_P$ |
---|
1036 | from each label $\ell$ to the occurrence (i.e.\ the path) of a $\s{while}$ loop |
---|
1037 | containing $\ell$, or the empty path if none exists; and $\s{gotosof}_P$ from |
---|
1038 | a label $\ell$ to the occurrences of \s{goto}s pointing to it. Then the set |
---|
1039 | $\s{multyentry}_P$ of multy-entry loops of $P$ can be computed by adding to it |
---|
1040 | all occurrences $p$ such that there exists a label $\ell$ and an occurrence |
---|
1041 | $q$ with $\s{loopof}_P(\ell)=p$, $q\in \s{gotosof}_P(\ell)$ and $p\not\le q$ |
---|
1042 | (where $\le$ is the prefix relation)\footnote{Possible simplification to this |
---|
1043 | procedure include keeping track just of while loops containing labels and |
---|
1044 | \s{goto}s (rather than paths in the syntactic tree of the program), and making |
---|
1045 | two passes while avoiding building the map to sets $\s{gotosof}$}. |
---|
1046 | |
---|
1047 | Let $Id_k$ be the indexing of length $k$ made from identity simple expressions, |
---|
1048 | i.e.\ the sequence $i_0\mapsto id_0, \ldots , i_{k-1}\mapsto id_{k-1}$. We define the tiered indexed labelling |
---|
1049 | $\Ell^\iota_P (S,k)$ in program $P$ for occurrence $S$ of a statement in $P$ |
---|
1050 | and a natural $k$ by recursion, setting |
---|
1051 | $$ |
---|
1052 | \Ell^\iota_P(S,k)\ass |
---|
1053 | \left\{ |
---|
1054 | \begin{array}{lh{-100pt}l} |
---|
1055 | (i_k:\sop{while}b\sbin{do}\alpha\la Id_{k+1}\ra : \Ell^\iota_P(T,k+1));\beta\la Id_k \ra : \s{skip} |
---|
1056 | \\& \text{if $S=\sop{while}b\sbin{do}T$ and $S\notin \s{multyentry}_P$,}\\[3pt] |
---|
1057 | (\sop{while}b\sbin{do}\alpha\la Id_k \ra : \Ell^\iota_P(T,k));\beta\la Id_k \ra : \s{skip} |
---|
1058 | \\& \text{otherwise, if $S=\sop{while}b\sbin{do}T$,}\\[3pt] |
---|
1059 | \sop{if}b\sbin{then} \alpha\la Id_k \ra : \Ell^\iota_P(T_1,k) \sbin{else} \beta\la Id_k \ra : \Ell^\iota_P(T_2,k) |
---|
1060 | \\&\text{if $S=\sop{if}b\sbin{then}T_1\sbin{else}T_2$,}\\[3pt] |
---|
1061 | \ldots |
---|
1062 | \end{array} |
---|
1063 | \right. |
---|
1064 | $$ |
---|
1065 | where as usual $\alpha$ and $\beta$ are to be fresh cost labels, and other cases just keep |
---|
1066 | making the recursive calls on the substatements. The \emph{indexed labelling} of |
---|
1067 | a program $P$ is then defined as $\alpha\la \ra : \Ell^\iota_P(P,0)$, i.e.\ a |
---|
1068 | further fresh unindexed cost label is added at the start, and we start from level $0$. |
---|
1069 | |
---|
1070 | In other words: each single-entry loop is indexed by $i_k$ where $k$ is the number of |
---|
1071 | other single-entry loops containing this one, and all cost labels under the scope |
---|
1072 | of a single-entry loop indexed by $i_k$ are indexed by all indexes $i_0,\ldots,i_k$, |
---|
1073 | without any transformation. |
---|
1074 | |
---|
1075 | \subsection{Indexed labels and loop transformations} |
---|
1076 | We define the \emph{reindexing} $I \circ (i_k\mapsto a*i_k+b)$ as an operator |
---|
1077 | on indexings by setting |
---|
1078 | \begin{multline*} |
---|
1079 | (i_0\mapsto e_0,\ldots, i_k \mapsto e_k,\ldots,i_n\mapsto e_n) |
---|
1080 | \circ(i_k\mapsto a*i_k+b) |
---|
1081 | \ass\\ |
---|
1082 | i_0\mapsto e_0,\ldots, i_k \mapsto e_k \circ(a*i_k+b),\ldots,i_n\mapsto e_n, |
---|
1083 | \end{multline*} |
---|
1084 | and extending then to indexed labels (by $\alpha\la I\ra\circ(i_k\mapsto e)\ass |
---|
1085 | \alpha\la I\circ (i_k\mapsto e)\ra$) and to statements in $\iota\ell\imp$ |
---|
1086 | (by applying the above transformation to all indexed labels). |
---|
1087 | |
---|
1088 | We can then redefine loop peeling and loop unrolling taking into account indexed labels. |
---|
1089 | It will be possible to apply the transformation only to indexed loops, that is loops |
---|
1090 | that are single-entry. The attentive reader will notice that no assumption is |
---|
1091 | made on the labelling of the statements involved. In particular the transformation |
---|
1092 | can be repeated and composed at will. Also, notice that erasing all labelling |
---|
1093 | information (i.e.\ indexed cost labels and loop indexes) we recover exactly the |
---|
1094 | same transformations presented in \autoref{sec:defimp}. |
---|
1095 | |
---|
1096 | \paragraph{Indexed loop peeling.} |
---|
1097 | |
---|
1098 | $$ |
---|
1099 | i_k:\sop{while}b\sbin{do}S\mapsto |
---|
1100 | \sop{if}b\sbin{then} S\circ (i_k\mapsto 0); i_k : \sop{while} b \sbin{do} S[\ell'_i/\ell_i]\circ(i_k\mapsto i_k + 1) |
---|
1101 | $$ |
---|
1102 | As can be expected, the peeled iteration of the loop gets reindexed as always |
---|
1103 | being the first iteration of the loop, while the iterations of the remaining |
---|
1104 | loop get shifted by $1$. |
---|
1105 | |
---|
1106 | \paragraph{Indexed loop unrolling.} |
---|
1107 | $$ |
---|
1108 | \begin{array}{l} |
---|
1109 | \begin{array}{ncn} |
---|
1110 | i_k:\sop{while}b\sbin{do}S\\ |
---|
1111 | \tikz\node[rotate=-90,inner sep=0pt]{$\mapsto$}; |
---|
1112 | \end{array}\\ |
---|
1113 | i_k:\sop{while} b \sbin{do}\\ |
---|
1114 | \quad (S\circ(i_k\mapsto n*i_k) ;\\ |
---|
1115 | \quad \sop{if} b \sbin{then}\\ |
---|
1116 | \quad\quad (S[\ell^1_i/\ell_i]\circ(i_k\mapsto n*i_k+1) ;\\ |
---|
1117 | \quad\quad\quad \vdots \\ |
---|
1118 | \quad\quad \quad \sop{if} b \sbin{then}\\ |
---|
1119 | \quad \quad \quad \quad S[\ell^n_i/\ell_i]\circ(i_k\mapsto n*i_k+n-1) |
---|
1120 | )\cdots ) |
---|
1121 | \end{array} |
---|
1122 | $$ |
---|
1123 | Again, the reindexing is as can be expected: each copy of the unrolled body |
---|
1124 | has its indexes remapped so that when they are executed the original iteration |
---|
1125 | of the loop to which they correspond can be recovered. |
---|
1126 | |
---|
1127 | \subsection{Semantics and compilation of indexed labels} |
---|
1128 | |
---|
1129 | In order to make sense of loop indexes, one must keep track of their values |
---|
1130 | in the state. A \emph{constant indexing} (metavariables $C,\ldots$) is an |
---|
1131 | indexing which employs only constant simple expressions. The evaluation |
---|
1132 | of an indexing $I$ in a constant indexing $C$, noted $I|_C$, is defined |
---|
1133 | by |
---|
1134 | $$I\circ(i_0\mapsto c_0,\ldots, i_{k-1}\mapsto c_{k-1}) \ass |
---|
1135 | \alphab\circ(i_0\mapsto c_0)\circ\cdots\circ(i_{k-1}\mapsto c_{k-1})$$ |
---|
1136 | (using the definition of ${-}\circ{-}$ given in \autoref{ssec:indlabs}), considering it defined only |
---|
1137 | if the the resulting indexing is a constant one too\footnote{For example |
---|
1138 | $(i_0\mapsto 2*i_0,i_1\mapsto i_1+1)|_{i_0\mapsto 2}$ is undefined, |
---|
1139 | but $(i_0\mapsto 2*i_0,i_1\mapsto 0)|_{i_0\mapsto 2}= |
---|
1140 | i_0\mapsto 4,i_1\mapsto 2$, which is indeed a constant indexing, |
---|
1141 | even if the domain of the original indexing is not covered by the constant one.}. |
---|
1142 | The definition is extended to indexed labels by $\alpha\la I\ra|_C\ass \alpha\la I|_C\ra$. |
---|
1143 | |
---|
1144 | Constant indexings will be used to keep track of the exact iterations of the |
---|
1145 | original code the emitted labels belong to. We thus define two basic actions to |
---|
1146 | update constant indexings: $C[i_k{\uparrow}]$ which increments the value of |
---|
1147 | $i_k$ by one, and $C[i_k{\downarrow}0]$ which resets it to $0$. |
---|
1148 | |
---|
1149 | We are ready to update the definition of the operational semantics of |
---|
1150 | indexed labelled \imp. The emitted cost labels will now be ones indexed by |
---|
1151 | constant indexings. We add a special indexed loop construct for continuations |
---|
1152 | that keeps track of active indexed loop indexes: |
---|
1153 | $$K,\ldots \gramm \cdots | i_k:\sop{while} b \sbin {do} S \sbin{then} K$$ |
---|
1154 | The difference between the regular stack concatenation |
---|
1155 | $i_k:\sop{while}b\sbin{do}S\cdot K$ and the new constructor is that the latter |
---|
1156 | indicates the loop is the active one in which we already are, while the former |
---|
1157 | is a loop that still needs to be started% |
---|
1158 | \footnote{In the presence of \s{continue} and \s{break} statements active |
---|
1159 | loops need to be kept track of in any case.}. |
---|
1160 | The \s{find} function is updated accordingly with the case |
---|
1161 | $$ \s{find}(\ell, i_k:\sop{while}b\sbin{do}S, K) \ass \s{find}(\ell, S, i_k: |
---|
1162 | \sop{while}b\sbin{do}S\sbin{then}K).$$ |
---|
1163 | The state will now be a 4-uple |
---|
1164 | $(S,K,s,C)$ which adds a constant indexing to the 3-uple of regular |
---|
1165 | semantics. The small-step rules for all statements but the |
---|
1166 | cost-labelled ones and the indexed loops remain the same, without |
---|
1167 | touching the $C$ parameter (in particular unindexed loops behave the same |
---|
1168 | as usual). The remaining cases are: |
---|
1169 | $$\begin{aligned} |
---|
1170 | (\alphab : S,K,s,C) &\to[\alphab|_C]_P (S,K,s,C)\\ |
---|
1171 | (i_k:\sop{while}b\sbin{do}S,K,C) &\to[\varepsilon]_P |
---|
1172 | \begin{cases} |
---|
1173 | (S,i_k:\sop{while}b\sbin{do}S\sbin{then} K,s,C[i_k{\downarrow}0]) |
---|
1174 | \\\hskip 125pt \text{if $(b,s)\eval v\neq 0$,}\\ |
---|
1175 | \rlap{(\s{skip}, K, s, C)}\hskip 125pt \text{otherwise} |
---|
1176 | \end{cases}\\ |
---|
1177 | (\s{skip}, i_k:\sop{while}b\sbin{do}S\sbin{then}K,C) &\to[\varepsilon]_P |
---|
1178 | \begin{cases} |
---|
1179 | (S,i_k:\sop{while}b\sbin{do}S\sbin{then} K,s,C[i_k{\uparrow}]) |
---|
1180 | \\\hskip 125pt \text{if $(b,s)\eval v\neq 0$,}\\ |
---|
1181 | \rlap{(\s{skip}, K, s, C)} \hskip 125pt \text{otherwise} |
---|
1182 | \end{cases} |
---|
1183 | \end{aligned}$$ |
---|
1184 | Some explications are in order. |
---|
1185 | \begin{itemize} |
---|
1186 | \item Emitting a label always instantiates it with the current indexing. |
---|
1187 | \item Hitting an indexed loop the first time initializes to 0 the corresponding |
---|
1188 | index; continuing the same loop inrements the index as expected. |
---|
1189 | \item The \s{find} function ignores the current indexing: this is correct |
---|
1190 | under the assumption that all indexed loops are single entry ones, so that |
---|
1191 | when we land inside an indexed loop with a \s{goto}, we are sure that its |
---|
1192 | current index is right. |
---|
1193 | \item The starting state with store $s$ for a program $P$ is |
---|
1194 | $(P,\s{halt},s,(i_0\mapsto 0,\dots,i_{n-1}\mapsto 0)$ where $i_0,\ldots,i_{n-1}$ cover |
---|
1195 | all loop indexes of $P$\footnote{For a program which is the indexed labelling of an |
---|
1196 | \imp{} one this corresponds to the maximum nesting of single-entry loops. We can also |
---|
1197 | avoid computing this value in advance if we define $C[i{\downarrow}0]$ to extend |
---|
1198 | $C$'s domain as needed, so that the starting constant indexing can be the empty one.}. |
---|
1199 | \end{itemize} |
---|
1200 | |
---|
1201 | \paragraph{Compilation.} |
---|
1202 | Further down the compilation chain the loop |
---|
1203 | structure is usually partially or completely lost. We cannot rely on it anymore |
---|
1204 | to ensure keeping track of original source code iterations. We therefore add |
---|
1205 | alongside the \s{emit} instruction two other sequential instructions |
---|
1206 | $\sop{ind_reset}k$ and $\sop{ind_inc}k$ whose sole effect is to reset to |
---|
1207 | 0 (resp.\ increment by 1) the loop index $i_k$, as kept track of in a constant |
---|
1208 | indexing accompanying the state. |
---|
1209 | |
---|
1210 | The first step of compilation from $\iota\ell\imp$ consists in prefixing the |
---|
1211 | translation of an indexed loop $i_k:\s{while}\ b\ \s{do}\ S$ with |
---|
1212 | $\sop{ind_reset}k$ and postfixing the translation of its body $S$ with |
---|
1213 | $\sop{ind_inc}k$. Later on in the compilation chain we just need to propagate |
---|
1214 | the instructions dealing with cost labels. |
---|
1215 | |
---|
1216 | We would like to stress the fact that this machinery is only needed to give a |
---|
1217 | suitable semantics of observables on which preservation proofs can be done. By no |
---|
1218 | means the added instructions and the constant indexing in the state are meant |
---|
1219 | to change the actual (let us say denotational) semantics of the programs. In this |
---|
1220 | regard the two new instruction have a similar role as the \s{emit} one. A |
---|
1221 | forgetful mapping of everything (syntax, states, operational semantics rules) |
---|
1222 | can be defined erasing all occurrences of cost labels and loop indexes, and the |
---|
1223 | result will always be a regular version of the language considered. |
---|
1224 | |
---|
1225 | \paragraph{Stating the preservation of semantics.} In fact, the statement of preservation |
---|
1226 | of semantics does not change at all, if not for considering traces of evaluated |
---|
1227 | indexed cost labels rather than traces of plain ones. |
---|
1228 | |
---|
1229 | |
---|
1230 | \subsection{Dependent costs in the source code}\label{ssec:depcosts} |
---|
1231 | The task of producing dependent costs out of the constant costs of indexed labels |
---|
1232 | is quite technical. Before presenting it here, we would like to point out that |
---|
1233 | the annotations produced by the procedure described in this subsection, even |
---|
1234 | if correct, can be enormous and unreadable. In \autoref{sec:conc}, when we will |
---|
1235 | detail the actual implementation, we will also sketch how we mitigated this |
---|
1236 | problem. |
---|
1237 | |
---|
1238 | Having the result of compiling the indexed labelling $\Ell^\iota(P)$ of an \imp{} |
---|
1239 | program $P$, we can still suppose that a cost mapping can be computed, but |
---|
1240 | from indexed labels to naturals. We want to annotate the source code, so we need |
---|
1241 | a way to express and compute costs of cost labels, i.e.\ group the costs of |
---|
1242 | indexed labels to ones of their atoms. In order to do so we introduce |
---|
1243 | \emph{dependent costs}. Let us suppose that for the sole purpose of annotation, |
---|
1244 | we have available in the language ternary expressions of the form |
---|
1245 | $$\tern e {f_1}{f_2},$$ |
---|
1246 | and that we have access to common operators on integers such as equality, order |
---|
1247 | and modulus. |
---|
1248 | |
---|
1249 | \paragraph{Simple conditions.} |
---|
1250 | First, we need to shift from \emph{transformations} of loop indexes to |
---|
1251 | \emph{conditions} on them. We identify a set of conditions on natural numbers |
---|
1252 | which are able to express the image of any composition of simple expressions. |
---|
1253 | |
---|
1254 | \emph{Simple conditions} are of three possible forms: |
---|
1255 | \begin{itemize} |
---|
1256 | \item equality $i_k=n$ for some natural $n$; |
---|
1257 | \item inequality $i_k\ge n$ for some natural $n$; |
---|
1258 | \item modular equality together with inequality $i_k\bmod a = b\wedge i_k\ge n$ |
---|
1259 | for naturals $a, b, n$. |
---|
1260 | \end{itemize} |
---|
1261 | The always true simple condition is given by $i_k\ge 0$, and similarly we |
---|
1262 | write $i_k\bmod a = b$ as a simple condition for $i_k\bmod a = b\wedge i_k\ge 0$. |
---|
1263 | Given a simple condition $p$ and a constant indexing $C$ we can easily define |
---|
1264 | when $p$ holds for $C$ (written $p\circ C$). A \emph{dependent cost expression} |
---|
1265 | is an expression built solely out of integer constants and ternary expressions |
---|
1266 | with simple conditions at their head. Given a dependent cost expression $e$ where |
---|
1267 | all of the loop indexes appearing in it are in the domain of a constant indexing |
---|
1268 | $C$, we can define the value $e\circ C\in \mathbb N$ by |
---|
1269 | $$n\circ C\ass n,\qquad (\tern p e f)\circ C\ass |
---|
1270 | \begin{cases} |
---|
1271 | e\circ C& \text{if $p\circ C$,}\\ |
---|
1272 | f\circ C& \text{otherwise.} |
---|
1273 | \end{cases}$$ |
---|
1274 | |
---|
1275 | \paragraph{From indexed costs to dependent ones.} |
---|
1276 | Every simple expression $e$ corresponds to a simple condition $p(e)$ which expresses the |
---|
1277 | set of values that can be the result of it. Following is the definition of such |
---|
1278 | relation. We recall that in this development loop indexes are always mapped to |
---|
1279 | simple expressions over the same index. If it was not the case, the condition |
---|
1280 | obtained from an expression should be on the mapped index, not the indeterminate |
---|
1281 | of the simple expression. We leave to further work all generalizations of what |
---|
1282 | we present here. |
---|
1283 | $$ |
---|
1284 | p(a*i_k+b)\ass |
---|
1285 | \begin{cases} |
---|
1286 | i_k = b & \text{if $a = 0$,}\\ |
---|
1287 | i_k \ge b & \text{if $a = 1$,}\\ |
---|
1288 | i_k\bmod a = b \wedge i_k \ge b & \text{otherwise}. |
---|
1289 | \end{cases}$$ |
---|
1290 | Now, suppose we are given a mapping $\kappa$ from indexed labels to natural |
---|
1291 | numbers. We will transform it in a mapping (identified with an abuse of notation |
---|
1292 | with the same symbol $\kappa$) from atoms to \imp{} expressions built with |
---|
1293 | ternary expressions which depend solely on loop indexes. To that end we define |
---|
1294 | an auxiliary function $\kappa^\alpha_L$ parameterized by atoms and words of |
---|
1295 | simple expressions and defined on \emph{sets} of $n$-uples of simple expressions |
---|
1296 | (with $n$ constant across each such set, i.e.\ each set is made of words with |
---|
1297 | the same length). |
---|
1298 | |
---|
1299 | We will employ a bijection between words of simple expressions and indexings, |
---|
1300 | given by\footnote{Lists of simple expressions is in fact how indexings are |
---|
1301 | represented in Cerco's current implementation of the compiler.} |
---|
1302 | $$i_0\mapsto e_0,\ldots,i_{k-1}\mapsto e_{k-1} \cong e_0\cdots e_{k-1}.$$ |
---|
1303 | As usual, $\varepsilon$ denotes the empty word/indexing, and juxtaposition |
---|
1304 | word concatenation. |
---|
1305 | |
---|
1306 | For every set $s$ of $n$-uples of simple expressions, we are in one of the following |
---|
1307 | three exclusive cases: |
---|
1308 | \begin{itemize} |
---|
1309 | \item $S=\emptyset$, or |
---|
1310 | \item $S=\{\varepsilon\}$, or |
---|
1311 | \item there is a simple expression $e$ such that $S$ can be decomposed in |
---|
1312 | $eS'+S''$, with $S'\neq \emptyset$ and none of the words in $S''$ starting with $e$ |
---|
1313 | \end{itemize} |
---|
1314 | where $eS'$ denotes prepending $e$ to all elements of $S'$ and $+$ is disjoint |
---|
1315 | union. This classification can serve as the basis of a definition by recursion |
---|
1316 | on $n+\card S$ where $n$ is the size of tuples in $S$ and $\card S$ is its cardinality. |
---|
1317 | Indeed in the third case in $S'$ the size of tuples decreases strictly (and |
---|
1318 | cardinality does not increase) while for $S''$ the size of tuples remains the same |
---|
1319 | but cardinality strictly decreases. The expression $e$ of the third case will be chosen |
---|
1320 | as minimal for some total order\footnote{The specific order used does not change |
---|
1321 | the correctness of the procedure, but different orders can give more or less |
---|
1322 | readable results. A ``good'' order is the lexicographic one, with $a*i_k+b \le a'*i_k+b'$ |
---|
1323 | if $a<a'$ or $a=a'$ and $b\le b'$.}. |
---|
1324 | |
---|
1325 | Following is the definition of the auxiliary function $\kappa^\alpha_L$, following |
---|
1326 | the recursion scheme presented above. |
---|
1327 | $$ |
---|
1328 | \begin{aligned} |
---|
1329 | \kappa^\alpha_L(\emptyset) &\ass 0\\ |
---|
1330 | \kappa^\alpha_L(\{\varepsilon\}) &\ass \kappa(\alpha\la L\ra) \\ |
---|
1331 | \kappa^\alpha_L(eS'+S'') &\ass \tern{p(e)}{\kappa^\alpha_{Le}(S')}{\kappa^\alpha_L(S'')} |
---|
1332 | \end{aligned} |
---|
1333 | $$ |
---|
1334 | |
---|
1335 | Finally, the wanted dependent cost mapping is defined by |
---|
1336 | $$\kappa(\alpha)\ass\kappa^\alpha_\varepsilon(\{\,L\mid \alpha\la L\ra \text{ appears |
---|
1337 | in the compiled code}\,\}).$$ |
---|
1338 | |
---|
1339 | \paragraph{Indexed instrumentation.} |
---|
1340 | The \emph{indexed instrumentation} generalizes the instrumentation presented in |
---|
1341 | \autoref{sec:labelling}. |
---|
1342 | We described |
---|
1343 | above how cost atoms can be mapped to dependent costs. The instrumentation must |
---|
1344 | also insert code dealing with the loop indexes. As instrumentation is done |
---|
1345 | on the code produced by the labelling phase, all cost labels are indexed by |
---|
1346 | identity indexings. The relevant cases of the recursive definition |
---|
1347 | (supposing $c$ is the cost variable) are then |
---|
1348 | $$ |
---|
1349 | \begin{aligned} |
---|
1350 | \mathcal I^\iota(\alpha\la Id_k\ra:S) &= c\ass c + \kappa(\alpha);\mathcal I^\iota(S)\\ |
---|
1351 | \mathcal I^\iota(i_k : \sop{while}b\sbin{do}S) &= |
---|
1352 | i_k \ass 0; \sop{while}b\sbin{do}(\mathcal I^\iota (S); i_k \ass i_k + 1) |
---|
1353 | \end{aligned} |
---|
1354 | $$ |
---|
1355 | |
---|
1356 | \section{Notes on the implementation and further work}\label{sec:conc} |
---|
1357 | Implementing the indexed label approach in CerCo's untrusted OcaML prototype |
---|
1358 | does not introduce many aspects beyond what has already been presented for the toy |
---|
1359 | language \imp{} with \s{goto}s. \s{Clight}, the fragment of \s{C} which is the |
---|
1360 | source language of CerCo's compilation chain~\cite{D2.1}, has several more fetaures, |
---|
1361 | but few demand changes in the indexed labelled approach. |
---|
1362 | \paragraph{Indexed loops vs index update instructions.} In this presentation |
---|
1363 | we had indexed loops in $\iota\ell\imp$, while we hinted at the fact that later |
---|
1364 | languages in the compilation chain would have specific index update instructions. |
---|
1365 | In CerCo's actual compilation chain from \s{Clight} to 8051 assembly, indexed |
---|
1366 | loops are only in \s{Clight}, while from \s{Cminor} onward all languages have |
---|
1367 | the same three cost-involving instructions: label emitting, index resetting and |
---|
1368 | index incrementing. |
---|
1369 | \paragraph{Loop transformations in the front end.} We decided to implement |
---|
1370 | the two loop transformations in the front end, namely in \s{Clight}. This |
---|
1371 | decision is due to user readability concerns: if costs are to be presented to |
---|
1372 | the programmer, they should depend on structures written by the programmer |
---|
1373 | himself. If loop transformation were performed later it would be harder to |
---|
1374 | relate information on loops in the control flow graph with actual loops |
---|
1375 | written in the source code. Another solution would be however to index |
---|
1376 | loops in the source code and then use these indexes later in the compilation |
---|
1377 | chain to pinpoint explicit loops of the source code: loop indexes can be used |
---|
1378 | to preserve such information just like cost labels. |
---|
1379 | \paragraph{Break and continue statements.} \s{Clight}'s loop flow control |
---|
1380 | statements for breaking and continuing a loop are equivalent to appropriate |
---|
1381 | \s{goto} statements. The only difference is that we are assured that they cannot |
---|
1382 | cause loops to be multi-entry, and that when a transformation such as loop |
---|
1383 | peeling is done, they need to be replaced by actual \s{goto}s (which will happen |
---|
1384 | further down the compilation chain anyway). |
---|
1385 | \paragraph{Function calls.} Every internal function definition has its own |
---|
1386 | space of loop indexes. Executable semantics must thus take into account saving |
---|
1387 | and resetting the constant indexing of current loops upon hitting a function |
---|
1388 | call, and restoring it upon return of control. A peculiarity is that this cannot |
---|
1389 | be attached to actions saving and restoring frames: namely in the case of |
---|
1390 | tail calls the constant indexing needs to be saved whereas the frame does not. |
---|
1391 | \paragraph{Cost-labelled expressions.} In labelled \s{Clight} expressions get |
---|
1392 | cost labels too, because of the presence of ternary conditional expressions |
---|
1393 | (and lazy logical operators, which get translated to ternary expressions too). |
---|
1394 | Adapting the indexed labelled approach to cost-labelled expressions does not |
---|
1395 | pose particular problems. |
---|
1396 | \paragraph{Simplification of dependent costs.} |
---|
1397 | As already mentioned, the blind application of the procedure described in |
---|
1398 | \autoref{ssec:depcosts} produces unwieldy cost annotations. In the implementation |
---|
1399 | several transformations are used to simplify such dependent costs. |
---|
1400 | |
---|
1401 | Disjunctions of simple conditions are closed under all logical operations, and |
---|
1402 | it can be computed whether such a disjunction implies a simple condition |
---|
1403 | or its negation. This can be used to eliminate useless branches of dependent costs, |
---|
1404 | to merge branches that have the same value, and possibly to simplify the third |
---|
1405 | case of simple condition. Examples of the three transformations are respectively |
---|
1406 | \begin{itemize} |
---|
1407 | \item $ |
---|
1408 | \verb+(_i_0 == 0)?+x\verb+:(_i_0 >= 1)?+y\verb+:+z |
---|
1409 | \mapsto |
---|
1410 | \verb+(_i_0 == 0)?+x\verb+:+y, |
---|
1411 | $ |
---|
1412 | \item $ |
---|
1413 | c\texttt{?}x\verb+:(+d\texttt{?}x\texttt{:}y\verb+)+ |
---|
1414 | \mapsto |
---|
1415 | \texttt{(}c\texttt{ || }d\texttt{)?}x\texttt{:}y, |
---|
1416 | $ |
---|
1417 | \item \begin{tabular}[t]{np{\linewidth}n} |
---|
1418 | $\verb+(_i_0 == 0)?+x\verb+:(_i_0 % 2 == 0 && _i_0 >= 2)?+y\verb+:+z |
---|
1419 | \mapsto$ \\\hfill |
---|
1420 | $\verb+(_i_0 == 0)?+x\verb+:(_i_0 % 2 == 0)?+y\verb+:+z. |
---|
1421 | $\end{tabular} |
---|
1422 | \end{itemize} |
---|
1423 | The second transformation tends to accumulate disjunctions, again to the detriment |
---|
1424 | of readability. A further transformation swaps two branches of the ternary |
---|
1425 | expression if the negation of the condition can be expressed with less clauses. |
---|
1426 | An example is |
---|
1427 | $$ \verb+(_i_0 % 3 == 0 || _i_0 % 3 == 1)?+x\verb+:+y \mapsto |
---|
1428 | \verb+(_i_0 % 3 == 2)?+y\verb+:+x. |
---|
1429 | $$ |
---|
1430 | |
---|
1431 | \paragraph{Updates to the frama-C cost plugin.} |
---|
1432 | Cerco's frama-C~\cite{framac} cost plugin\todo{is there a reference for this?}{} |
---|
1433 | has been updated to take into account dependent |
---|
1434 | costs. The frama-c framework explodes ternary expressions to actual branch |
---|
1435 | statements introducing temporaries along the way, which makes the task of |
---|
1436 | analyzing ternary cost expressions rather daunting. It was deemed necessary to provide |
---|
1437 | an option in the compiler to use actual branch statements for cost annotations |
---|
1438 | rather than ternary expressions, so that at least frama-C's use of temporaries in |
---|
1439 | cost annotation be avoided. The cost analysis carried out by the plugin now |
---|
1440 | takes into account such dependent costs. |
---|
1441 | |
---|
1442 | The only limitation (which provided |
---|
1443 | a simplification in the code) is that within a dependent cost |
---|
1444 | simple conditions with modulus on the |
---|
1445 | same loop index should not be modulo different numbers, which corresponds to |
---|
1446 | the reasonable limitation of not applying multiple times loop unrolling to |
---|
1447 | the same loops. |
---|
1448 | \paragraph{Further work.} |
---|
1449 | Indexed labels are for now implemented only in the untrusted OcaML compiler, |
---|
1450 | while they are not present yet in the Matita code. Porting them should pose no |
---|
1451 | significant problem, and then the proof effort should begin. |
---|
1452 | |
---|
1453 | Because most of the executable operational semantics of the languages across the |
---|
1454 | front end and the back end are oblivious to cost labels, it should be expected |
---|
1455 | that the bulk of the semantic preservation proofs that still needs to be done |
---|
1456 | will not get any harder because of indexed labels. The only trickier point |
---|
1457 | would be in the translation of \s{Clight} to \s{Cminor}, where we |
---|
1458 | pass from structured indexed loops to atomic instructions on loop indexes. |
---|
1459 | |
---|
1460 | An invariant which should probably be proved and provably preserved along compilation |
---|
1461 | is the non-overlap of indexings for the same atom. Then, supposing cost |
---|
1462 | correctness for the unindexed approach, the indexed one will just need to |
---|
1463 | add the proof that |
---|
1464 | $$\forall C\text{ constant indexing}.\forall \alpha\la I\ra\text{ appearing in the compiled code}. |
---|
1465 | \kappa(\alpha)\circ (I\circ C) = \kappa(\alpha\la I \ra).$$ |
---|
1466 | $C$ represents a snapshot of loop indexes in the compiled code, while |
---|
1467 | $I\circ C$ is the corresponding snapshot in the source code. Semantics preservation |
---|
1468 | will make sure that when with snapshot $C$ we emit $\alpha\ra I\la$ (that is, |
---|
1469 | we have $\alpha\la I\circ C\ra$ in the trace), this means that $\alpha$ must be |
---|
1470 | emitted in the source code with indexing $I\circ C$, so the cost |
---|
1471 | $\kappa(\alpha)\circ (I\circ C)$ applies. |
---|
1472 | |
---|
1473 | Apart from carrying over the proofs, we would like to extend the approach |
---|
1474 | to more loop transformations. Important examples are loop inversion |
---|
1475 | (where a for loop is reversed, usually to make iterations appear as truly |
---|
1476 | independent) or loop interchange (where two nested loops are swapped, usually |
---|
1477 | to have more loop invariants or to enhance strength reduction). This introduces |
---|
1478 | interesting changes to the approach, where we would have indexings such as |
---|
1479 | $$i_0\mapsto n - i_0\quad\text{or}\quad i_0\mapsto i_1, i_1\mapsto i_0.$$ |
---|
1480 | In particular dependency over actual variables of the code would enter the |
---|
1481 | frame, as indexings would depend on the number of iterations of a well-behaving |
---|
1482 | guarded loop (the $n$ in the first example). |
---|
1483 | |
---|
1484 | Finally, as stated in the introduction, the approach should allow integrating |
---|
1485 | some techniques for cache analysis, a possibility that for now has been put aside |
---|
1486 | as the standard 8051 architecture targeted by the CerCo project has no cache. |
---|
1487 | As possible developments of this line of work without straying from the 8051 |
---|
1488 | architecture, two possibilities present themselves. |
---|
1489 | \begin{enumerate} |
---|
1490 | \item One could extend the development to some 8051 variants, of which some have |
---|
1491 | been produced with a cache. |
---|
1492 | \item One could make the compiler implement its own cache: this cannot |
---|
1493 | apply to RAM accesses of the standard 8051 architecture, as the difference in |
---|
1494 | cost of accessing the two types of RAM is of just 1 clock cycle, which makes |
---|
1495 | any implementation of cache counterproductive. So this possibility should |
---|
1496 | either artificially change the accessing cost of RAM of the model just for the |
---|
1497 | sake of possible future adaptations to other architectures, or otherwise |
---|
1498 | model access to an external memory by means of the serial port.\end{enumerate} |
---|
1499 | |
---|
1500 | |
---|
1501 | % |
---|
1502 | % \newpage |
---|
1503 | % |
---|
1504 | % \includepdf[pages={-}]{plugin.pdf} |
---|
1505 | % |
---|
1506 | % |
---|
1507 | % \newpage |
---|
1508 | % |
---|
1509 | % \includepdf[pages={-}]{fopara.pdf} |
---|
1510 | % |
---|
1511 | % |
---|
1512 | % \newpage |
---|
1513 | % |
---|
1514 | % \includepdf[pages={-}]{tlca.pdf} |
---|
1515 | % |
---|
1516 | % \bibliographystyle{plain} |
---|
1517 | \bibliography{bib} |
---|
1518 | |
---|
1519 | \end{document} |
---|