X-Git-Url: https://bilbo.iut-bm.univ-fcomte.fr/and/gitweb/prng_gpu.git/blobdiff_plain/80c5c05c07373d50d1af355c7a25a336a12f7d73..893deb2477f17914ca2f0420009697c0925d5cbe:/prng_gpu.tex?ds=inline diff --git a/prng_gpu.tex b/prng_gpu.tex index 00752c7..32055e7 100644 --- a/prng_gpu.tex +++ b/prng_gpu.tex @@ -1,4 +1,5 @@ -\documentclass{article} +%\documentclass{article} +\documentclass[10pt,journal,letterpaper,compsoc]{IEEEtran} \usepackage[utf8]{inputenc} \usepackage[T1]{fontenc} \usepackage{fullpage} @@ -7,10 +8,15 @@ \usepackage{amscd} \usepackage{moreverb} \usepackage{commath} -\usepackage{algorithm2e} +\usepackage[ruled,vlined]{algorithm2e} \usepackage{listings} \usepackage[standard]{ntheorem} - +\usepackage{algorithmic} +\usepackage{slashbox} +\usepackage{ctable} +\usepackage{cite} +\usepackage{tabularx} +\usepackage{multirow} % Pour mathds : les ensembles IR, IN, etc. \usepackage{dsfont} @@ -20,8 +26,11 @@ \usepackage{graphicx} % Pour faire des sous-figures dans les figures \usepackage{subfigure} +\usepackage{xr-hyper} +\usepackage{hyperref} +\externaldocument[A-]{supplementary} + -\usepackage{color} \newtheorem{notation}{Notation} @@ -34,31 +43,41 @@ \newcommand{\alert}[1]{\begin{color}{blue}\textit{#1}\end{color}} + + \title{Efficient and Cryptographically Secure Generation of Chaotic Pseudorandom Numbers on GPU} \begin{document} \author{Jacques M. Bahi, Rapha\"{e}l Couturier, Christophe -Guyeux, and Pierre-Cyrille Heam\thanks{Authors in alphabetic order}} +Guyeux, and Pierre-Cyrille Héam\thanks{Authors in alphabetic order}} -\maketitle +\IEEEcompsoctitleabstractindextext{ \begin{abstract} In this paper we present a new pseudorandom number generator (PRNG) on -graphics processing units (GPU). This PRNG is based on the so-called chaotic iterations. It -is firstly proven to be chaotic according to the Devaney's formulation. We thus propose an efficient +graphics processing units (GPU). This PRNG is based on the so-called chaotic iterations and +it is thus chaotic according to the Devaney's formulation. We propose an efficient implementation for GPU that successfully passes the {\it BigCrush} tests, deemed to be the hardest battery of tests in TestU01. Experiments show that this PRNG can generate -about 20 billions of random numbers per second on Tesla C1060 and NVidia GTX280 +about 20 billion of random numbers per second on Tesla C1060 and NVidia GTX280 cards. -It is finally established that, under reasonable assumptions, the proposed PRNG can be cryptographically +It is then established that, under reasonable assumptions, the proposed PRNG can be cryptographically secure. +A chaotic version of the Blum-Goldwasser asymmetric key encryption scheme is finally proposed. \end{abstract} +} + +\maketitle + +\IEEEdisplaynotcompsoctitleabstractindextext +\IEEEpeerreviewmaketitle + \section{Introduction} -Randomness is of importance in many fields as scientific simulations or cryptography. +Randomness is of importance in many fields such as scientific simulations or cryptography. ``Random numbers'' can mainly be generated either by a deterministic and reproducible algorithm called a pseudorandom number generator (PRNG), or by a physical non-deterministic process having all the characteristics of a random noise, called a truly random number @@ -66,21 +85,27 @@ generator (TRNG). In this paper, we focus on reproducible generators, useful for instance in Monte-Carlo based simulators or in several cryptographic schemes. These domains need PRNGs that are statistically irreproachable. -On some fields as in numerical simulations, speed is a strong requirement +In some fields such as in numerical simulations, speed is a strong requirement that is usually attained by using parallel architectures. In that case, -a recurrent problem is that a deflate of the statistical qualities is often +a recurrent problem is that a deflation of the statistical qualities is often reported, when the parallelization of a good PRNG is realized. This is why ad-hoc PRNGs for each possible architecture must be found to achieve both speed and randomness. On the other side, speed is not the main requirement in cryptography: the great -need is to define \emph{secure} generators being able to withstand malicious +need is to define \emph{secure} generators able to withstand malicious attacks. Roughly speaking, an attacker should not be able in practice to make the distinction between numbers obtained with the secure generator and a true random -sequence. -Finally, a small part of the community working in this domain focus on a +sequence. However, in an equivalent formulation, he or she should not be +able (in practice) to predict the next bit of the generator, having the knowledge of all the +binary digits that have been already released. ``Being able in practice'' refers here +to the possibility to achieve this attack in polynomial time, and to the exponential growth +of the difficulty of this challenge when the size of the parameters of the PRNG increases. + + +Finally, a small part of the community working in this domain focuses on a third requirement, that is to define chaotic generators. The main idea is to take benefits from a chaotic dynamical system to obtain a -generator that is unpredictable, disordered, sensible to its seed, or in other words chaotic. +generator that is unpredictable, disordered, sensible to its seed, or in other word chaotic. Their desire is to map a given chaotic dynamics into a sequence that seems random and unassailable due to chaos. However, the chaotic maps used as a pattern are defined in the real line @@ -94,7 +119,7 @@ This is why the use of chaos for PRNG still remains marginal and disputable. The authors' opinion is that topological properties of disorder, as they are properly defined in the mathematical theory of chaos, can reinforce the quality of a PRNG. But they are not substitutable for security or statistical perfection. -Indeed, to the authors' point of view, such properties can be useful in the two following situations. On the +Indeed, to the authors' mind, such properties can be useful in the two following situations. On the one hand, a post-treatment based on a chaotic dynamical system can be applied to a PRNG statistically deflective, in order to improve its statistical properties. Such an improvement can be found, for instance, in~\cite{bgw09:ip,bcgr11:ip}. @@ -109,11 +134,18 @@ Let us finish this paragraph by noticing that, in this paper, statistical perfection refers to the ability to pass the whole {\it BigCrush} battery of tests, which is widely considered as the most stringent statistical evaluation of a sequence claimed as random. -This battery can be found into the well-known TestU01 package~\cite{LEcuyerS07}. +This battery can be found in the well-known TestU01 package~\cite{LEcuyerS07}. +More precisely, each time we performed a test on a PRNG, we ran it +twice in order to observe if all $p-$values are inside [0.01, 0.99]. In +fact, we observed that few $p-$values (less than ten) are sometimes +outside this interval but inside [0.001, 0.999], so that is why a +second run allows us to confirm that the values outside are not for +the same test. With this approach all our PRNGs pass the {\it + BigCrush} successfully and all $p-$values are at least once inside +[0.01, 0.99]. Chaos, for its part, refers to the well-established definition of a chaotic dynamical system proposed by Devaney~\cite{Devaney}. - In a previous work~\cite{bgw09:ip,guyeux10} we have proposed a post-treatment on PRNGs making them behave as a chaotic dynamical system. Such a post-treatment leads to a new category of PRNGs. We have shown that proofs of Devaney's chaos can be established for this @@ -130,39 +162,65 @@ applications. Therefore, it is important to be able to generate pseudorandom numbers inside a GPU when a scientific application runs in it. This remark motivates our proposal of a chaotic and statistically perfect PRNG for GPU. Such device -allows us to generated almost 20 billions of pseudorandom numbers per second. -Last, but not least, we show that the proposed post-treatment preserves the +allows us to generate almost 20 billion of pseudorandom numbers per second. +Furthermore, we show that the proposed post-treatment preserves the cryptographical security of the inputted PRNG, when this last has such a property. +Last, but not least, we propose a rewriting of the Blum-Goldwasser asymmetric +key encryption protocol by using the proposed method. + + +{\bf Main contributions.} In this paper a new PRNG using chaotic iteration +is defined. From a theoretical point of view, it is proven that it has fine +topological chaotic properties and that it is cryptographically secured (when +the initial PRNG is also cryptographically secured). From a practical point of +view, experiments point out a very good statistical behavior. An optimized +original implementation of this PRNG is also proposed and experimented. +Pseudorandom numbers are generated at a rate of 20GSamples/s, which is faster +than in~\cite{conf/fpga/ThomasHL09,Marsaglia2003} (and with a better +statistical behavior). Experiments are also provided using BBS as the initial +random generator. The generation speed is significantly weaker. +Note also that an original qualitative comparison between topological chaotic +properties and statistical test is also proposed. + + + The remainder of this paper is organized as follows. In Section~\ref{section:related works} we review some GPU implementations of PRNGs. Section~\ref{section:BASIC RECALLS} gives some basic recalls on the well-known Devaney's formulation of chaos, and on an iteration process called ``chaotic iterations'' on which the post-treatment is based. -Proofs of chaos are given in Section~\ref{sec:pseudorandom}. -Section~\ref{sec:efficient PRNG} presents an efficient -implementation of this chaotic PRNG on a CPU, whereas Section~\ref{sec:efficient PRNG - gpu} describes the GPU implementation. +The proposed PRNG and its proof of chaos are given in Section~\ref{sec:pseudorandom}. +Section~\ref{sec:efficient PRNG} %{The generation of pseudorandom sequence} %illustrates the statistical +%improvement related to the chaotic iteration based post-treatment, for +%our previously released PRNGs and + contains a new efficient +implementation on CPU. + Section~\ref{sec:efficient PRNG + gpu} describes and evaluates theoretically the GPU implementation. Such generators are experimented in Section~\ref{sec:experiments}. We show in Section~\ref{sec:security analysis} that, if the inputted generator is cryptographically secure, then it is the case too for the generator provided by the post-treatment. +A practical +security evaluation is also outlined in Section~\ref{sec:Practicak evaluation}. Such a proof leads to the proposition of a cryptographically secure and -chaotic generator on GPU based on the famous Blum Blum Shum -in Section~\ref{sec:CSGPU}. +chaotic generator on GPU based on the famous Blum Blum Shub +in Section~\ref{sec:CSGPU} and to an improvement of the +Blum-Goldwasser protocol in Sect.~\ref{Blum-Goldwasser}. This research work ends by a conclusion section, in which the contribution is summarized and intended future work is presented. -\section{Related works on GPU based PRNGs} +\section{Related work on GPU based PRNGs} \label{section:related works} -Numerous research works on defining GPU based PRNGs have yet been proposed in the -literature, so that completeness is impossible. +Numerous research works on defining GPU based PRNGs have already been proposed in the +literature, so that exhaustivity is impossible. This is why authors of this document only give reference to the most significant attempts in this domain, from their subjective point of view. The quantity of pseudorandom numbers generated per second is mentioned here @@ -180,7 +238,7 @@ chaos or cryptography in this document. In \cite{ZRKB10}, the authors propose different versions of efficient GPU PRNGs based on Lagged Fibonacci or Hybrid Taus. They have used these PRNGs for Langevin simulations of biomolecules fully implemented on -GPU. Performance of the GPU versions are far better than those obtained with a +GPU. Performances of the GPU versions are far better than those obtained with a CPU, and these PRNGs succeed to pass the {\it BigCrush} battery of TestU01. However the evaluations of the proposed PRNGs are only statistical ones. @@ -192,11 +250,11 @@ the performance of the same PRNGs on different architectures are compared. FPGA appears as the fastest and the most efficient architecture, providing the fastest number of generated pseudorandom numbers per joule. -However, we can notice that authors can ``only'' generate between 11 and 16GSamples/s +However, we notice that authors can ``only'' generate between 11 and 16GSamples/s with a GTX 280 GPU, which should be compared with the results presented in this document. We can remark too that the PRNGs proposed in~\cite{conf/fpga/ThomasHL09} are only -able to pass the {\it Crush} battery, which is very easy compared to the {\it Big Crush} one. +able to pass the {\it Crush} battery, which is far easier than the {\it Big Crush} one. Lastly, Cuda has developed a library for the generation of pseudorandom numbers called Curand~\cite{curand11}. Several PRNGs are implemented, among @@ -206,15 +264,18 @@ their fastest version provides 15GSamples/s on the new Fermi C2050 card. But their PRNGs cannot pass the whole TestU01 battery (only one test is failed). \newline \newline -We can finally remark that, to the best of our knowledge, no GPU implementation have been proven to be chaotic, and the cryptographically secure property is surprisingly never regarded. +We can finally remark that, to the best of our knowledge, no GPU implementation has been proven to be chaotic, and the cryptographically secure property has surprisingly never been considered. \section{Basic Recalls} \label{section:BASIC RECALLS} This section is devoted to basic definitions and terminologies in the fields of -topological chaos and chaotic iterations. -\subsection{Devaney's Chaotic Dynamical Systems} +topological chaos and chaotic iterations. We assume the reader is familiar +with basic notions on topology (see for instance~\cite{Devaney}). + +\subsection{Devaney's Chaotic Dynamical Systems} +\label{subsec:Devaney} In the sequel $S^{n}$ denotes the $n^{th}$ term of a sequence $S$ and $V_{i}$ denotes the $i^{th}$ component of a vector $V$. $f^{k}=f\circ ...\circ f$ is for the $k^{th}$ composition of a function $f$. Finally, the following @@ -225,7 +286,7 @@ Consider a topological space $(\mathcal{X},\tau)$ and a continuous function $f : \mathcal{X} \rightarrow \mathcal{X}$. \begin{definition} -$f$ is said to be \emph{topologically transitive} if, for any pair of open sets +The function $f$ is said to be \emph{topologically transitive} if, for any pair of open sets $U,V \subset \mathcal{X}$, there exists $k>0$ such that $f^k(U) \cap V \neq \varnothing$. \end{definition} @@ -244,7 +305,7 @@ necessarily the same period). \begin{definition}[Devaney's formulation of chaos~\cite{Devaney}] -$f$ is said to be \emph{chaotic} on $(\mathcal{X},\tau)$ if $f$ is regular and +The function $f$ is said to be \emph{chaotic} on $(\mathcal{X},\tau)$ if $f$ is regular and topologically transitive. \end{definition} @@ -252,12 +313,12 @@ The chaos property is strongly linked to the notion of ``sensitivity'', defined on a metric space $(\mathcal{X},d)$ by: \begin{definition} -\label{sensitivity} $f$ has \emph{sensitive dependence on initial conditions} +\label{sensitivity} The function $f$ has \emph{sensitive dependence on initial conditions} if there exists $\delta >0$ such that, for any $x\in \mathcal{X}$ and any neighborhood $V$ of $x$, there exist $y\in V$ and $n > 0$ such that $d\left(f^{n}(x), f^{n}(y)\right) >\delta $. -$\delta$ is called the \emph{constant of sensitivity} of $f$. +The constant $\delta$ is called the \emph{constant of sensitivity} of $f$. \end{definition} Indeed, Banks \emph{et al.} have proven in~\cite{Banks92} that when $f$ is @@ -316,15 +377,15 @@ Let us now recall how to define a suitable metric space where chaotic iterations are continuous. For further explanations, see, e.g., \cite{guyeux10}. Let $\delta $ be the \emph{discrete Boolean metric}, $\delta -(x,y)=0\Leftrightarrow x=y.$ Given a function $f$, define the function: -\begin{equation} +(x,y)=0\Leftrightarrow x=y.$ Given a function $f$, define the function +$F_{f}: \llbracket1;\mathsf{N}\rrbracket\times \mathds{B}^{\mathsf{N}} +\longrightarrow \mathds{B}^{\mathsf{N}}$ +\begin{equation*} \begin{array}{lrll} -F_{f}: & \llbracket1;\mathsf{N}\rrbracket\times \mathds{B}^{\mathsf{N}} & -\longrightarrow & \mathds{B}^{\mathsf{N}} \\ -& (k,E) & \longmapsto & \left( E_{j}.\delta (k,j)+f(E)_{k}.\overline{\delta -(k,j)}\right) _{j\in \llbracket1;\mathsf{N}\rrbracket},% +& (k,E) & \longmapsto & \left( E_{j}.\delta (k,j)+ f(E)_{k}.\overline{\delta +(k,j)}\right) _{j\in \llbracket1;\mathsf{N}\rrbracket}% \end{array}% -\end{equation}% +\end{equation*}% \noindent where + and . are the Boolean addition and product operations. Consider the phase space: \begin{equation} @@ -380,9 +441,9 @@ their distance should increase too. \item In addition, if two systems present the same cells and their respective strategies start with the same terms, then the distance between these two points must be small because the evolution of the two systems will be the same for a -while. Indeed, the two dynamical systems start with the same initial condition, -use the same update function, and as strategies are the same for a while, then -components that are updated are the same too. +while. Indeed, both dynamical systems start with the same initial condition, +use the same update function, and as strategies are the same for a while, furthermore +updated components are the same as well. \end{itemize} The distance presented above follows these recommendations. Indeed, if the floor value $\lfloor d(X,Y)\rfloor $ is equal to $n$, then the systems $E, \check{E}$ @@ -391,7 +452,7 @@ measure of the differences between strategies $S$ and $\check{S}$. More precisely, this floating part is less than $10^{-k}$ if and only if the first $k$ terms of the two strategies are equal. Moreover, if the $k^{th}$ digit is nonzero, then the $k^{th}$ terms of the two strategies are different. -The impact of this choice for a distance will be investigate at the end of the document. +The impact of this choice for a distance will be investigated at the end of the document. Finally, it has been established in \cite{guyeux10} that, @@ -401,7 +462,7 @@ the metric space $(\mathcal{X},d)$. \end{proposition} The chaotic property of $G_f$ has been firstly established for the vectorial -Boolean negation $f(x_1,\hdots, x_\mathsf{N}) = (\overline{x_1},\hdots, \overline{x_\mathsf{N}})$ \cite{guyeux10}. To obtain a characterization, we have secondly +Boolean negation $f_0(x_1,\hdots, x_\mathsf{N}) = (\overline{x_1},\hdots, \overline{x_\mathsf{N}})$ \cite{guyeux10}. To obtain a characterization, we have secondly introduced the notion of asynchronous iteration graph recalled bellow. Let $f$ be a map from $\mathds{B}^\mathsf{N}$ to itself. The @@ -414,8 +475,7 @@ The relation between $\Gamma(f)$ and $G_f$ is clear: there exists a path from $x$ to $x'$ in $\Gamma(f)$ if and only if there exists a strategy $s$ such that the parallel iteration of $G_f$ from the initial point $(s,x)$ reaches the point $x'$. - -We have finally proven in \cite{bcgr11:ip} that, +We have then proven in \cite{bcgr11:ip} that, \begin{theorem} @@ -424,14 +484,33 @@ Let $f:\mathds{B}^\mathsf{N}\to\mathds{B}^\mathsf{N}$. $G_f$ is chaotic (accord if and only if $\Gamma(f)$ is strongly connected. \end{theorem} -This result of chaos has lead us to study the possibility to build a +Finally, we have established in \cite{bcgr11:ip} that, +\begin{theorem} + Let $f: \mathds{B}^{n} \rightarrow \mathds{B}^{n}$, $\Gamma(f)$ its + iteration graph, $\check{M}$ its adjacency + matrix and $M$ + a $n\times n$ matrix defined by + $ + M_{ij} = \frac{1}{n}\check{M}_{ij}$ %\textrm{ + if $i \neq j$ and + $M_{ii} = 1 - \frac{1}{n} \sum\limits_{j=1, j\neq i}^n \check{M}_{ij}$ otherwise. + + If $\Gamma(f)$ is strongly connected, then + the output of the PRNG detailed in Algorithm~\ref{CI Algorithm} follows + a law that tends to the uniform distribution + if and only if $M$ is a double stochastic matrix. +\end{theorem} + + +These results of chaos and uniform distribution have led us to study the possibility of building a pseudorandom number generator (PRNG) based on the chaotic iterations. As $G_f$, defined on the domain $\llbracket 1 ; \mathsf{N} \rrbracket^{\mathds{N}} -\times \mathds{B}^\mathsf{N}$, is build from Boolean networks $f : \mathds{B}^\mathsf{N} +\times \mathds{B}^\mathsf{N}$, is built from Boolean networks $f : \mathds{B}^\mathsf{N} \rightarrow \mathds{B}^\mathsf{N}$, we can preserve the theoretical properties on $G_f$ -during implementations (due to the discrete nature of $f$). It is as if +during implementations (due to the discrete nature of $f$). Indeed, it is as if $\mathds{B}^\mathsf{N}$ represents the memory of the computer whereas $\llbracket 1 ; \mathsf{N} \rrbracket^{\mathds{N}}$ is its input stream (the seeds, for instance, in PRNG, or a physical noise in TRNG). +Let us finally remark that the vectorial negation satisfies the hypotheses of both theorems above. \section{Application to Pseudorandomness} \label{sec:pseudorandom} @@ -440,30 +519,58 @@ $\mathds{B}^\mathsf{N}$ represents the memory of the computer whereas $\llbracke We have proposed in~\cite{bgw09:ip} a new family of generators that receives two PRNGs as inputs. These two generators are mixed with chaotic iterations, -leading thus to a new PRNG that improves the statistical properties of each -generator taken alone. Furthermore, our generator -possesses various chaos properties that none of the generators used as input -present. +leading thus to a new PRNG that +should improve the statistical properties of each +generator taken alone. +Furthermore, the generator obtained in this way possesses various chaos properties that none of the generators used as present input. + + \begin{algorithm}[h!] -%\begin{scriptsize} +\begin{small} \KwIn{a function $f$, an iteration number $b$, an initial configuration $x^0$ ($n$ bits)} \KwOut{a configuration $x$ ($n$ bits)} $x\leftarrow x^0$\; -$k\leftarrow b + \textit{XORshift}(b)$\; +$k\leftarrow b + PRNG_1(b)$\; \For{$i=0,\dots,k$} { -$s\leftarrow{\textit{XORshift}(n)}$\; +$s\leftarrow{PRNG_2(n)}$\; $x\leftarrow{F_f(s,x)}$\; } return $x$\; -%\end{scriptsize} -\caption{PRNG with chaotic functions} +\end{small} +\caption{An arbitrary round of $Old~ CI~ PRNG_f(PRNG_1,PRNG_2)$} \label{CI Algorithm} \end{algorithm} + + + +This generator is synthesized in Algorithm~\ref{CI Algorithm}. +It takes as input: a Boolean function $f$ satisfying Theorem~\ref{Th:Caractérisation des IC chaotiques}; +an integer $b$, ensuring that the number of executed iterations +between two outputs is at least $b$ +and at most $2b+1$; and an initial configuration $x^0$. +It returns the new generated configuration $x$. Internally, it embeds two +inputted generators $PRNG_i(k), i=1,2$, + which must return integers +uniformly distributed +into $\llbracket 1 ; k \rrbracket$. +For instance, these PRNGs can be the \textit{XORshift}~\cite{Marsaglia2003}, +being a category of very fast PRNGs designed by George Marsaglia +that repeatedly uses the transform of exclusive or (XOR, $\oplus$) on a number +with a bit shifted version of it. Such a PRNG, which has a period of +$2^{32}-1=4.29\times10^9$, is summed up in Algorithm~\ref{XORshift}. +This XORshift, or any other reasonable PRNG, is used +in our own generator to compute both the number of iterations between two +outputs (provided by $PRNG_1$) and the strategy elements ($PRNG_2$). + +%This former generator has successively passed various batteries of statistical tests, as the NIST~\cite{bcgr11:ip}, DieHARD~\cite{Marsaglia1996}, and TestU01~\cite{LEcuyerS07} ones. + + \begin{algorithm}[h!] +\begin{small} \KwIn{the internal configuration $z$ (a 32-bit word)} \KwOut{$y$ (a 32-bit word)} $z\leftarrow{z\oplus{(z\ll13)}}$\; @@ -471,49 +578,98 @@ $z\leftarrow{z\oplus{(z\gg17)}}$\; $z\leftarrow{z\oplus{(z\ll5)}}$\; $y\leftarrow{z}$\; return $y$\; -\medskip +\end{small} \caption{An arbitrary round of \textit{XORshift} algorithm} \label{XORshift} \end{algorithm} +\subsection{A ``New CI PRNG''} + +In order to make the Old CI PRNG usable in practice, we have proposed +an adapted version of the chaotic iteration based generator in~\cite{bg10:ip}. +In this ``New CI PRNG'', we prevent a given bit from changing twice between two outputs. +This new generator is designed by the following process. + +First of all, some chaotic iterations have to be done to generate a sequence +$\left(x^n\right)_{n\in\mathds{N}} \in \left(\mathds{B}^{32}\right)^\mathds{N}$ +of Boolean vectors, which are the successive states of the iterated system. +Some of these vectors will be randomly extracted and our pseudorandom bit +flow will be constituted by their components. Such chaotic iterations are +realized as follows. Initial state $x^0 \in \mathds{B}^{32}$ is a Boolean +vector taken as a seed and chaotic strategy $\left(S^n\right)_{n\in\mathds{N}}\in +\llbracket 1, 32 \rrbracket^\mathds{N}$ is +an \emph{irregular decimation} of $PRNG_2$ sequence, as described in +Algorithm~\ref{Chaotic iteration1}. + +Then, at each iteration, only the $S^n$-th component of state $x^n$ is +updated, as follows: $x_i^n = x_i^{n-1}$ if $i \neq S^n$, else $x_i^n = \overline{x_i^{n-1}}$. +Such a procedure is equivalent to achieving chaotic iterations with +the Boolean vectorial negation $f_0$ and some well-chosen strategies. +Finally, some $x^n$ are selected +by a sequence $m^n$ as the pseudorandom bit sequence of our generator. +$(m^n)_{n \in \mathds{N}} \in \mathcal{M}^\mathds{N}$ is computed from $PRNG_1$, where $\mathcal{M}\subset \mathds{N}^*$ is a finite nonempty set of integers. + +The basic design procedure of the New CI generator is summarized in Algorithm~\ref{Chaotic iteration1}. +The internal state is $x$, the output state is $r$. $a$ and $b$ are those computed by the two input +PRNGs. Lastly, the value $g(a)$ is an integer defined as in Eq.~\ref{Formula}. +This function must be chosen such that the outputs of the resulted PRNG are uniform in $\llbracket 0, 2^\mathsf{N}-1 \rrbracket$. Function of \eqref{Formula} achieves this +goal (other candidates and more information can be found in ~\cite{bg10:ip}). +\begin{equation} +\label{Formula} +m^n = g(y^n)= +\left\{ +\begin{array}{l} +0 \text{ if }0 \leqslant{y^n}<{C^0_{32}},\\ +1 \text{ if }{C^0_{32}} \leqslant{y^n}<\sum_{i=0}^1{C^i_{32}},\\ +2 \text{ if }\sum_{i=0}^1{C^i_{32}} \leqslant{y^n}<\sum_{i=0}^2{C^i_{32}},\\ +\vdots~~~~~ ~~\vdots~~~ ~~~~\\ +N \text{ if }\sum_{i=0}^{N-1}{C^i_{32}}\leqslant{y^n}<1.\\ +\end{array} +\right. +\end{equation} - -This generator is synthesized in Algorithm~\ref{CI Algorithm}. -It takes as input: a Boolean function $f$ satisfying Theorem~\ref{Th:Caractérisation des IC chaotiques}; -an integer $b$, ensuring that the number of executed iterations is at least $b$ -and at most $2b+1$; and an initial configuration $x^0$. -It returns the new generated configuration $x$. Internally, it embeds two -\textit{XORshift}$(k)$ PRNGs~\cite{Marsaglia2003} that returns integers -uniformly distributed -into $\llbracket 1 ; k \rrbracket$. -\textit{XORshift} is a category of very fast PRNGs designed by George Marsaglia, -which repeatedly uses the transform of exclusive or (XOR, $\oplus$) on a number -with a bit shifted version of it. This PRNG, which has a period of -$2^{32}-1=4.29\times10^9$, is summed up in Algorithm~\ref{XORshift}. It is used -in our PRNG to compute the strategy length and the strategy elements. - - -We have proven in \cite{bcgr11:ip} that, -\begin{theorem} - Let $f: \mathds{B}^{n} \rightarrow \mathds{B}^{n}$, $\Gamma(f)$ its - iteration graph, $\check{M}$ its adjacency - matrix and $M$ a $n\times n$ matrix defined as in the previous lemma. - If $\Gamma(f)$ is strongly connected, then - the output of the PRNG detailed in Algorithm~\ref{CI Algorithm} follows - a law that tends to the uniform distribution - if and only if $M$ is a double stochastic matrix. -\end{theorem} - -This former generator as successively passed various batteries of statistical tests, as the NIST~\cite{bcgr11:ip}, DieHARD~\cite{Marsaglia1996}, and TestU01~\cite{LEcuyerS07}. +\begin{algorithm} +\textbf{Input:} the internal state $x$ (32 bits)\\ +\textbf{Output:} a state $r$ of 32 bits +\begin{algorithmic}[1] +\FOR{$i=0,\dots,N$} +{ +\STATE$d_i\leftarrow{0}$\; +} +\ENDFOR +\STATE$a\leftarrow{PRNG_1()}$\; +\STATE$k\leftarrow{g(a)}$\; +\WHILE{$i=0,\dots,k$} + +\STATE$b\leftarrow{PRNG_2()~mod~\mathsf{N}}$\; +\STATE$S\leftarrow{b}$\; + \IF{$d_S=0$} + { +\STATE $x_S\leftarrow{ \overline{x_S}}$\; +\STATE $d_S\leftarrow{1}$\; + + } + \ELSIF{$d_S=1$} + { +\STATE $k\leftarrow{ k+1}$\; + }\ENDIF +\ENDWHILE\\ +\STATE $r\leftarrow{x}$\; +\STATE return $r$\; +\medskip +\caption{An arbitrary round of the new CI generator} +\label{Chaotic iteration1} +\end{algorithmic} +\end{algorithm} \subsection{Improving the Speed of the Former Generator} -Instead of updating only one cell at each iteration, we can try to choose a -subset of components and to update them together. Such an attempt leads -to a kind of merger of the two sequences used in Algorithm -\ref{CI Algorithm}. When the updating function is the vectorial negation, +Instead of updating only one cell at each iteration, we now propose to choose a +subset of components and to update them together, for speed improvement. Such a proposition leads +to a kind of merger of the two sequences used in Algorithms +\ref{CI Algorithm} and \ref{Chaotic iteration1}. When the updating function is the vectorial negation, this algorithm can be rewritten as follows: \begin{equation} @@ -526,7 +682,7 @@ x^0 \in \llbracket 0, 2^\mathsf{N}-1 \rrbracket, S \in \llbracket 0, 2^\mathsf{N \label{equation Oplus} \end{equation} where $\oplus$ is for the bitwise exclusive or between two integers. -This rewritten can be understood as follows. The $n-$th term $S^n$ of the +This rewriting can be understood as follows. The $n-$th term $S^n$ of the sequence $S$, which is an integer of $\mathsf{N}$ binary digits, presents the list of cells to update in the state $x^n$ of the system (represented as an integer having $\mathsf{N}$ bits too). More precisely, the $k-$th @@ -550,304 +706,665 @@ where $f$ is the vectorial negation and $\forall n \in \mathds{N}$, $\mathcal{S}^n \subset \llbracket 1, \mathsf{N} \rrbracket$ is such that $k \in \mathcal{S}^n$ if and only if the $k-$th digit in the binary decomposition of $S^n$ is 1. Such chaotic iterations are more general -than the ones presented in Definition \ref{Def:chaotic iterations} for -the fact that, instead of updating only one term at each iteration, +than the ones presented in Definition \ref{Def:chaotic iterations} because, instead of updating only one term at each iteration, we select a subset of components to change. -Obviously, replacing Algorithm~\ref{CI Algorithm} by -Equation~\ref{equation Oplus}, possible when the iteration function is -the vectorial negation, leads to a speed improvement. However, proofs +Obviously, replacing the previous CI PRNG Algorithms by +Equation~\ref{equation Oplus}, which is possible when the iteration function is +the vectorial negation, leads to a speed improvement +(the resulting generator will be referred as ``Xor CI PRNG'' +in what follows). +However, proofs of chaos obtained in~\cite{bg10:ij} have been established only for chaotic iterations of the form presented in Definition -\ref{Def:chaotic iterations}. The question is now to determine whether the +\ref{Def:chaotic iterations}. The question to determine whether the use of more general chaotic iterations to generate pseudorandom numbers -faster, does not deflate their topological chaos properties. - -\subsection{Proofs of Chaos of the General Formulation of the Chaotic Iterations} -\label{deuxième def} -Let us consider the discrete dynamical systems in chaotic iterations having -the general form: - -\begin{equation} -\forall n\in \mathds{N}^{\ast }, \forall i\in -\llbracket1;\mathsf{N}\rrbracket ,x_i^n=\left\{ -\begin{array}{ll} - x_i^{n-1} & \text{ if } i \notin \mathcal{S}^n \\ - \left(f(x^{n-1})\right)_{S^n} & \text{ if }i \in \mathcal{S}^n. -\end{array}\right. -\label{general CIs} -\end{equation} - -In other words, at the $n^{th}$ iteration, only the cells whose id is -contained into the set $S^{n}$ are iterated. - -Let us now rewrite these general chaotic iterations as usual discrete dynamical -system of the form $X^{n+1}=f(X^n)$ on an ad hoc metric space. Such a formulation -is required in order to study the topological behavior of the system. - -Let us introduce the following function: -\begin{equation} -\begin{array}{cccc} - \chi: & \llbracket 1; \mathsf{N} \rrbracket \times \mathcal{P}\left(\llbracket 1; \mathsf{N} \rrbracket\right) & \longrightarrow & \mathds{B}\\ - & (i,X) & \longmapsto & \left\{ \begin{array}{ll} 0 & \textrm{if }i \notin X, \\ 1 & \textrm{if }i \in X, \end{array}\right. -\end{array} -\end{equation} -where $\mathcal{P}\left(X\right)$ is for the powerset of the set $X$, that is, $Y \in \mathcal{P}\left(X\right) \Longleftrightarrow Y \subset X$. - -Given a function $f:\mathds{B}^\mathsf{N} \longrightarrow \mathds{B}^\mathsf{N} $, define the function: -\begin{equation} -\begin{array}{lrll} -F_{f}: & \mathcal{P}\left(\llbracket1;\mathsf{N}\rrbracket \right) \times \mathds{B}^{\mathsf{N}} & -\longrightarrow & \mathds{B}^{\mathsf{N}} \\ -& (P,E) & \longmapsto & \left( E_{j}.\chi (j,P)+f(E)_{j}.\overline{\chi -(j,P)}\right) _{j\in \llbracket1;\mathsf{N}\rrbracket},% -\end{array}% -\end{equation}% -where + and . are the Boolean addition and product operations, and $\overline{x}$ -is the negation of the Boolean $x$. -Consider the phase space: -\begin{equation} -\mathcal{X} = \mathcal{P}\left(\llbracket 1 ; \mathsf{N} \rrbracket\right)^\mathds{N} \times -\mathds{B}^\mathsf{N}, -\end{equation} -\noindent and the map defined on $\mathcal{X}$: -\begin{equation} -G_f\left(S,E\right) = \left(\sigma(S), F_f(i(S),E)\right), \label{Gf} -\end{equation} -\noindent where $\sigma$ is the \emph{shift} function defined by $\sigma -(S^{n})_{n\in \mathds{N}}\in \mathcal{P}\left(\llbracket 1 ; \mathsf{N} \rrbracket\right)^\mathds{N}\longrightarrow (S^{n+1})_{n\in -\mathds{N}}\in \mathcal{P}\left(\llbracket 1 ; \mathsf{N} \rrbracket\right)^\mathds{N}$ and $i$ is the \emph{initial function} -$i:(S^{n})_{n\in \mathds{N}} \in \mathcal{P}\left(\llbracket 1 ; \mathsf{N} \rrbracket\right)^\mathds{N}\longrightarrow S^{0}\in \mathcal{P}\left(\llbracket 1 ; \mathsf{N} \rrbracket\right)$. -Then the general chaotic iterations defined in Equation \ref{general CIs} can -be described by the following discrete dynamical system: -\begin{equation} -\left\{ -\begin{array}{l} -X^0 \in \mathcal{X} \\ -X^{k+1}=G_{f}(X^k).% -\end{array}% -\right. -\end{equation}% - -Another time, a shift function appears as a component of these general chaotic -iterations. - -To study the Devaney's chaos property, a distance between two points -$X = (S,E), Y = (\check{S},\check{E})$ of $\mathcal{X}$ must be defined. -Let us introduce: -\begin{equation} -d(X,Y)=d_{e}(E,\check{E})+d_{s}(S,\check{S}), -\label{nouveau d} -\end{equation} -\noindent where -\begin{equation} -\left\{ -\begin{array}{lll} -\displaystyle{d_{e}(E,\check{E})} & = & \displaystyle{\sum_{k=1}^{\mathsf{N}% -}\delta (E_{k},\check{E}_{k})}\textrm{ is another time the Hamming distance}, \\ -\displaystyle{d_{s}(S,\check{S})} & = & \displaystyle{\dfrac{9}{\mathsf{N}}% -\sum_{k=1}^{\infty }\dfrac{|S^k\Delta {S}^k|}{10^{k}}}.% -\end{array}% -\right. -\end{equation} -where $|X|$ is the cardinality of a set $X$ and $A\Delta B$ is for the symmetric difference, defined for sets A, B as -$A\,\Delta\,B = (A \setminus B) \cup (B \setminus A)$. - - -\begin{proposition} -The function $d$ defined in Eq.~\ref{nouveau d} is a metric on $\mathcal{X}$. -\end{proposition} - -\begin{proof} - $d_e$ is the Hamming distance. We will prove that $d_s$ is a distance -too, thus $d$ will be a distance as sum of two distances. - \begin{itemize} -\item Obviously, $d_s(S,\check{S})\geqslant 0$, and if $S=\check{S}$, then -$d_s(S,\check{S})=0$. Conversely, if $d_s(S,\check{S})=0$, then -$\forall k \in \mathds{N}, |S^k\Delta {S}^k|=0$, and so $\forall k, S^k=\check{S}^k$. - \item $d_s$ is symmetric -($d_s(S,\check{S})=d_s(\check{S},S)$) due to the commutative property -of the symmetric difference. -\item Finally, $|S \Delta S''| = |(S \Delta \varnothing) \Delta S''|= |S \Delta (S'\Delta S') \Delta S''|= |(S \Delta S') \Delta (S' \Delta S'')|\leqslant |S \Delta S'| + |S' \Delta S''|$, -and so for all subsets $S,S',$ and $S''$ of $\llbracket 1, \mathsf{N} \rrbracket$, -we have $d_s(S,S'') \leqslant d_e(S,S')+d_s(S',S'')$, and the triangle -inequality is obtained. - \end{itemize} -\end{proof} - - -Before being able to study the topological behavior of the general -chaotic iterations, we must firstly establish that: - -\begin{proposition} - For all $f:\mathds{B}^\mathsf{N} \longrightarrow \mathds{B}^\mathsf{N} $, the function $G_f$ is continuous on -$\left( \mathcal{X},d\right)$. -\end{proposition} - - -\begin{proof} -We use the sequential continuity. -Let $(S^n,E^n)_{n\in \mathds{N}}$ be a sequence of the phase space $% -\mathcal{X}$, which converges to $(S,E)$. We will prove that $\left( -G_{f}(S^n,E^n)\right) _{n\in \mathds{N}}$ converges to $\left( -G_{f}(S,E)\right) $. Let us remark that for all $n$, $S^n$ is a strategy, -thus, we consider a sequence of strategies (\emph{i.e.}, a sequence of -sequences).\newline -As $d((S^n,E^n);(S,E))$ converges to 0, each distance $d_{e}(E^n,E)$ and $d_{s}(S^n,S)$ converges -to 0. But $d_{e}(E^n,E)$ is an integer, so $\exists n_{0}\in \mathds{N},$ $% -d_{e}(E^n,E)=0$ for any $n\geqslant n_{0}$.\newline -In other words, there exists a threshold $n_{0}\in \mathds{N}$ after which no -cell will change its state: -$\exists n_{0}\in \mathds{N},n\geqslant n_{0}\Rightarrow E^n = E.$ - -In addition, $d_{s}(S^n,S)\longrightarrow 0,$ so $\exists n_{1}\in % -\mathds{N},d_{s}(S^n,S)<10^{-1}$ for all indexes greater than or equal to $% -n_{1}$. This means that for $n\geqslant n_{1}$, all the $S^n$ have the same -first term, which is $S^0$: $\forall n\geqslant n_{1},S_0^n=S_0.$ - -Thus, after the $max(n_{0},n_{1})^{th}$ term, states of $E^n$ and $E$ are -identical and strategies $S^n$ and $S$ start with the same first term.\newline -Consequently, states of $G_{f}(S^n,E^n)$ and $G_{f}(S,E)$ are equal, -so, after the $max(n_0, n_1)^{th}$ term, the distance $d$ between these two points is strictly less than 1.\newline -\noindent We now prove that the distance between $\left( -G_{f}(S^n,E^n)\right) $ and $\left( G_{f}(S,E)\right) $ is convergent to -0. Let $\varepsilon >0$. \medskip -\begin{itemize} -\item If $\varepsilon \geqslant 1$, we see that distance -between $\left( G_{f}(S^n,E^n)\right) $ and $\left( G_{f}(S,E)\right) $ is -strictly less than 1 after the $max(n_{0},n_{1})^{th}$ term (same state). -\medskip -\item If $\varepsilon <1$, then $\exists k\in \mathds{N},10^{-k}\geqslant -\varepsilon > 10^{-(k+1)}$. But $d_{s}(S^n,S)$ converges to 0, so -\begin{equation*} -\exists n_{2}\in \mathds{N},\forall n\geqslant -n_{2},d_{s}(S^n,S)<10^{-(k+2)}, -\end{equation*}% -thus after $n_{2}$, the $k+2$ first terms of $S^n$ and $S$ are equal. -\end{itemize} -\noindent As a consequence, the $k+1$ first entries of the strategies of $% -G_{f}(S^n,E^n)$ and $G_{f}(S,E)$ are the same ($G_{f}$ is a shift of strategies) and due to the definition of $d_{s}$, the floating part of -the distance between $(S^n,E^n)$ and $(S,E)$ is strictly less than $% -10^{-(k+1)}\leqslant \varepsilon $.\bigskip \newline -In conclusion, -$$ -\forall \varepsilon >0,\exists N_{0}=max(n_{0},n_{1},n_{2})\in \mathds{N}% -,\forall n\geqslant N_{0}, - d\left( G_{f}(S^n,E^n);G_{f}(S,E)\right) -\leqslant \varepsilon . -$$ -$G_{f}$ is consequently continuous. -\end{proof} - - -It is now possible to study the topological behavior of the general chaotic -iterations. We will prove that, - -\begin{theorem} -\label{t:chaos des general} - The general chaotic iterations defined on Equation~\ref{general CIs} satisfy -the Devaney's property of chaos. -\end{theorem} - -Let us firstly prove the following lemma. - -\begin{lemma}[Strong transitivity] -\label{strongTrans} - For all couples $X,Y \in \mathcal{X}$ and any neighborhood $V$ of $X$, we can -find $n \in \mathds{N}^*$ and $X' \in V$ such that $G^n(X')=Y$. -\end{lemma} - -\begin{proof} - Let $X=(S,E)$, $\varepsilon>0$, and $k_0 = \lfloor log_{10}(\varepsilon)+1 \rfloor$. -Any point $X'=(S',E')$ such that $E'=E$ and $\forall k \leqslant k_0, S'^k=S^k$, -are in the open ball $\mathcal{B}\left(X,\varepsilon\right)$. Let us define -$\check{X} = \left(\check{S},\check{E}\right)$, where $\check{X}= G^{k_0}(X)$. -We denote by $s\subset \llbracket 1; \mathsf{N} \rrbracket$ the set of coordinates -that are different between $\check{E}$ and the state of $Y$. Thus each point $X'$ of -the form $(S',E')$ where $E'=E$ and $S'$ starts with -$(S^0, S^1, \hdots, S^{k_0},s,\hdots)$, verifies the following properties: -\begin{itemize} - \item $X'$ is in $\mathcal{B}\left(X,\varepsilon\right)$, - \item the state of $G_f^{k_0+1}(X')$ is the state of $Y$. -\end{itemize} -Finally the point $\left(\left(S^0, S^1, \hdots, S^{k_0},s,s^0, s^1, \hdots\right); E\right)$, -where $(s^0,s^1, \hdots)$ is the strategy of $Y$, satisfies the properties -claimed in the lemma. -\end{proof} - -We can now prove the Theorem~\ref{t:chaos des general}... - -\begin{proof}[Theorem~\ref{t:chaos des general}] -Firstly, strong transitivity implies transitivity. - -Let $(S,E) \in\mathcal{X}$ and $\varepsilon >0$. To -prove that $G_f$ is regular, it is sufficient to prove that -there exists a strategy $\tilde S$ such that the distance between -$(\tilde S,E)$ and $(S,E)$ is less than $\varepsilon$, and such that -$(\tilde S,E)$ is a periodic point. - -Let $t_1=\lfloor-\log_{10}(\varepsilon)\rfloor$, and let $E'$ be the -configuration that we obtain from $(S,E)$ after $t_1$ iterations of -$G_f$. As $G_f$ is strongly transitive, there exists a strategy $S'$ -and $t_2\in\mathds{N}$ such -that $E$ is reached from $(S',E')$ after $t_2$ iterations of $G_f$. - -Consider the strategy $\tilde S$ that alternates the first $t_1$ terms -of $S$ and the first $t_2$ terms of $S'$: $$\tilde -S=(S_0,\dots,S_{t_1-1},S'_0,\dots,S'_{t_2-1},S_0,\dots,S_{t_1-1},S'_0,\dots,S'_{t_2-1},S_0,\dots).$$ It -is clear that $(\tilde S,E)$ is obtained from $(\tilde S,E)$ after -$t_1+t_2$ iterations of $G_f$. So $(\tilde S,E)$ is a periodic -point. Since $\tilde S_t=S_t$ for $t0$. \medskip +%% \begin{itemize} +%% \item If $\varepsilon \geqslant 1$, we see that the distance +%% between $\left( G_{f}(S^n,E^n)\right) $ and $\left( G_{f}(S,E)\right) $ is +%% strictly less than 1 after the $max(n_{0},n_{1})^{th}$ term (same state). +%% \medskip +%% \item If $\varepsilon <1$, then $\exists k\in \mathds{N},10^{-k}\geqslant +%% \varepsilon > 10^{-(k+1)}$. But $d_{s}(S^n,S)$ converges to 0, so +%% \begin{equation*} +%% \exists n_{2}\in \mathds{N},\forall n\geqslant +%% n_{2},d_{s}(S^n,S)<10^{-(k+2)}, +%% \end{equation*}% +%% thus after $n_{2}$, the $k+2$ first terms of $S^n$ and $S$ are equal. +%% \end{itemize} +%% \noindent As a consequence, the $k+1$ first entries of the strategies of $% +%% G_{f}(S^n,E^n)$ and $G_{f}(S,E)$ are the same ($G_{f}$ is a shift of strategies) and due to the definition of $d_{s}$, the floating part of +%% the distance between $(S^n,E^n)$ and $(S,E)$ is strictly less than $% +%% 10^{-(k+1)}\leqslant \varepsilon $. + +%% In conclusion, +%% %%RAPH : ici j'ai rajouté une ligne +%% %%TOF : ici j'ai rajouté un commentaire +%% %%TOF : ici aussi +%% $ +%% \forall \varepsilon >0,$ $\exists N_{0}=max(n_{0},n_{1},n_{2})\in \mathds{N} +%% ,$ $\forall n\geqslant N_{0},$ +%% $ d\left( G_{f}(S^n,E^n);G_{f}(S,E)\right) +%% \leqslant \varepsilon . +%% $ +%% $G_{f}$ is consequently continuous. +%% \end{proof} + + +%% It is now possible to study the topological behavior of the general chaotic +%% iterations. We will prove that, + +%% \begin{theorem} +%% \label{t:chaos des general} +%% The general chaotic iterations defined on Equation~\ref{general CIs} satisfy +%% the Devaney's property of chaos. +%% \end{theorem} + +%% Let us firstly prove the following lemma. + +%% \begin{lemma}[Strong transitivity] +%% \label{strongTrans} +%% For all couples $X,Y \in \mathcal{X}$ and any neighborhood $V$ of $X$, we can +%% find $n \in \mathds{N}^*$ and $X' \in V$ such that $G^n(X')=Y$. +%% \end{lemma} + +%% \begin{proof} +%% Let $X=(S,E)$, $\varepsilon>0$, and $k_0 = \lfloor log_{10}(\varepsilon)+1 \rfloor$. +%% Any point $X'=(S',E')$ such that $E'=E$ and $\forall k \leqslant k_0, S'^k=S^k$, +%% are in the open ball $\mathcal{B}\left(X,\varepsilon\right)$. Let us define +%% $\check{X} = \left(\check{S},\check{E}\right)$, where $\check{X}= G^{k_0}(X)$. +%% We denote by $s\subset \llbracket 1; \mathsf{N} \rrbracket$ the set of coordinates +%% that are different between $\check{E}$ and the state of $Y$. Thus each point $X'$ of +%% the form $(S',E')$ where $E'=E$ and $S'$ starts with +%% $(S^0, S^1, \hdots, S^{k_0},s,\hdots)$, verifies the following properties: +%% \begin{itemize} +%% \item $X'$ is in $\mathcal{B}\left(X,\varepsilon\right)$, +%% \item the state of $G_f^{k_0+1}(X')$ is the state of $Y$. +%% \end{itemize} +%% Finally the point $\left(\left(S^0, S^1, \hdots, S^{k_0},s,s^0, s^1, \hdots\right); E\right)$, +%% where $(s^0,s^1, \hdots)$ is the strategy of $Y$, satisfies the properties +%% claimed in the lemma. +%% \end{proof} + +%% We can now prove the Theorem~\ref{t:chaos des general}. + +%% \begin{proof}[Theorem~\ref{t:chaos des general}] +%% Firstly, strong transitivity implies transitivity. + +%% Let $(S,E) \in\mathcal{X}$ and $\varepsilon >0$. To +%% prove that $G_f$ is regular, it is sufficient to prove that +%% there exists a strategy $\tilde S$ such that the distance between +%% $(\tilde S,E)$ and $(S,E)$ is less than $\varepsilon$, and such that +%% $(\tilde S,E)$ is a periodic point. + +%% Let $t_1=\lfloor-\log_{10}(\varepsilon)\rfloor$, and let $E'$ be the +%% configuration that we obtain from $(S,E)$ after $t_1$ iterations of +%% $G_f$. As $G_f$ is strongly transitive, there exists a strategy $S'$ +%% and $t_2\in\mathds{N}$ such +%% that $E$ is reached from $(S',E')$ after $t_2$ iterations of $G_f$. + +%% Consider the strategy $\tilde S$ that alternates the first $t_1$ terms +%% of $S$ and the first $t_2$ terms of $S'$: +%% %%RAPH : j'ai coupé la ligne en 2 +%% $$\tilde +%% S=(S_0,\dots,S_{t_1-1},S'_0,\dots,S'_{t_2-1},S_0,$$$$\dots,S_{t_1-1},S'_0,\dots,S'_{t_2-1},S_0,\dots).$$ It +%% is clear that $(\tilde S,E)$ is obtained from $(\tilde S,E)$ after +%% $t_1+t_2$ iterations of $G_f$. So $(\tilde S,E)$ is a periodic +%% point. Since $\tilde S_t=S_t$ for $t0$ et $\liminf_{n \rightarrow +\infty} d(f^{(n)}(x), f^{(n)}(y))=0$, meaning that their orbits always oscillate as the iterations pass. When a system is compact and contains an uncountable set of such points, it is claimed as chaotic according +%% to Li-Yorke~\cite{Li75,Ruette2001}. A similar property is regarded in the following NIST test~\cite{Nist10}. +%% \begin{itemize} +%% \item \textbf{Runs Test}. To determine whether the number of runs of ones and zeros of various lengths is as expected for a random sequence. In particular, this test determines whether the oscillation between such zeros and ones is too fast or too slow. +%% \end{itemize} +%% \item \textbf{Topological entropy}. The desire to formulate an equivalency of the thermodynamics entropy +%% has emerged both in the topological and statistical fields. Once again, a similar objective has led to two different +%% rewritting of an entropy based disorder: the famous Shannon definition of entropy is approximated in the statistical approach, +%% whereas topological entropy is defined as follows: +%% $x,y \in \mathcal{X}$ are $\varepsilon-$\emph{separated in time $n$} if there exists $k \leqslant n$ such that $d\left(f^{(k)}(x),f^{(k)}(y)\right)>\varepsilon$. Then $(n,\varepsilon)-$separated sets are sets of points that are all $\varepsilon-$separated in time $n$, which +%% leads to the definition of $s_n(\varepsilon,Y)$, being the maximal cardinality of all $(n,\varepsilon)-$separated sets. Using these notations, +%% the topological entropy is defined as follows: $$h_{top}(\mathcal{X},f) = \displaystyle{\lim_{\varepsilon \rightarrow 0} \Big[ \limsup_{n \rightarrow +\infty} \dfrac{1}{n} \log s_n(\varepsilon,\mathcal{X})\Big]}.$$ +%% This value measures the average exponential growth of the number of distinguishable orbit segments. +%% In this sense, it measures the complexity of the topological dynamical system, whereas +%% the Shannon approach comes to mind when defining the following test~\cite{Nist10}: +%% \begin{itemize} +%% \item \textbf{Approximate Entropy Test}. Compare the frequency of the overlapping blocks of two consecutive/adjacent lengths ($m$ and $m+1$) against the expected result for a random sequence. +%% \end{itemize} + +%% \item \textbf{Non-linearity, complexity}. Finally, let us remark that non-linearity and complexity are +%% not only sought in general to obtain chaos, but they are also required for randomness, as illustrated by the two tests below~\cite{Nist10}. +%% \begin{itemize} +%% \item \textbf{Binary Matrix Rank Test}. Check for linear dependence among fixed length substrings of the original sequence. +%% \item \textbf{Linear Complexity Test}. Determine whether or not the sequence is complex enough to be considered random. +%% \end{itemize} +%% \end{itemize} + + +%% We have proven in our previous works~\cite{guyeux12:bc} that chaotic iterations satisfying Theorem~\ref{Th:Caractérisation des IC chaotiques} are, among other +%% things, strongly transitive, topologically mixing, chaotic as defined by Li and Yorke, +%% and that they have a topological entropy and an exponent of Lyapunov both equal to $ln(\mathsf{N})$, +%% where $\mathsf{N}$ is the size of the iterated vector. +%% These topological properties make that we are ground to believe that a generator based on chaotic +%% iterations will probably be able to pass all the existing statistical batteries for pseudorandomness like +%% the NIST one. The following subsections, in which we prove that defective generators have their +%% statistical properties improved by chaotic iterations, show that such an assumption is true. + +%% \subsection{Details of some Existing Generators} + +%% The list of defective PRNGs we will use +%% as inputs for the statistical tests to come is introduced here. + +%% Firstly, the simple linear congruency generators (LCGs) will be used. +%% They are defined by the following recurrence: +%% \begin{equation} +%% x^n = (ax^{n-1} + c)~mod~m, +%% \label{LCG} +%% \end{equation} +%% where $a$, $c$, and $x^0$ must be, among other things, non-negative and inferior to +%% $m$~\cite{LEcuyerS07}. In what follows, 2LCGs and 3LCGs refer to two (resp. three) +%% combinations of such LCGs. For further details, see~\cite{bfg12a:ip,combined_lcg}. + +%% Secondly, the multiple recursive generators (MRGs) which will be used, +%% are based on a linear recurrence of order +%% $k$, modulo $m$~\cite{LEcuyerS07}: +%% \begin{equation} +%% x^n = (a^1x^{n-1}+~...~+a^kx^{n-k})~mod~m . +%% \label{MRG} +%% \end{equation} +%% The combination of two MRGs (referred as 2MRGs) is also used in these experiments. + +%% Generators based on linear recurrences with carry will be regarded too. +%% This family of generators includes the add-with-carry (AWC) generator, based on the recurrence: +%% \begin{equation} +%% \label{AWC} +%% \begin{array}{l} +%% x^n = (x^{n-r} + x^{n-s} + c^{n-1})~mod~m, \\ +%% c^n= (x^{n-r} + x^{n-s} + c^{n-1}) / m, \end{array}\end{equation} +%% the SWB generator, having the recurrence: +%% \begin{equation} +%% \label{SWB} +%% \begin{array}{l} +%% x^n = (x^{n-r} - x^{n-s} - c^{n-1})~mod~m, \\ +%% c^n=\left\{ +%% \begin{array}{l} +%% 1 ~~~~~\text{if}~ (x^{i-r} - x^{i-s} - c^{i-1})<0\\ +%% 0 ~~~~~\text{else},\end{array} \right. \end{array}\end{equation} +%% and the SWC generator, which is based on the following recurrence: +%% \begin{equation} +%% \label{SWC} +%% \begin{array}{l} +%% x^n = (a^1x^{n-1} \oplus ~...~ \oplus a^rx^{n-r} \oplus c^{n-1}) ~ mod ~ 2^w, \\ +%% c^n = (a^1x^{n-1} \oplus ~...~ \oplus a^rx^{n-r} \oplus c^{n-1}) ~ / ~ 2^w. \end{array}\end{equation} + +%% Then the generalized feedback shift register (GFSR) generator has been implemented, that is: +%% \begin{equation} +%% x^n = x^{n-r} \oplus x^{n-k} . +%% \label{GFSR} +%% \end{equation} + + +%% Finally, the nonlinear inversive (INV) generator~\cite{LEcuyerS07} has been studied, which is: + +%% \begin{equation} +%% \label{INV} +%% \begin{array}{l} +%% x^n=\left\{ +%% \begin{array}{ll} +%% (a^1 + a^2 / z^{n-1})~mod~m & \text{if}~ z^{n-1} \neq 0 \\ +%% a^1 & \text{if}~ z^{n-1} = 0 .\end{array} \right. \end{array}\end{equation} + + + +%% \begin{table} +%% \renewcommand{\arraystretch}{1.3} +%% \caption{TestU01 Statistical Test Failures} +%% \label{TestU011} +%% \centering +%% \begin{tabular}{lccccc} +%% \toprule +%% Test name &Tests& Logistic & XORshift & ISAAC\\ +%% Rabbit & 38 &21 &14 &0 \\ +%% Alphabit & 17 &16 &9 &0 \\ +%% Pseudo DieHARD &126 &0 &2 &0 \\ +%% FIPS\_140\_2 &16 &0 &0 &0 \\ +%% SmallCrush &15 &4 &5 &0 \\ +%% Crush &144 &95 &57 &0 \\ +%% Big Crush &160 &125 &55 &0 \\ \hline +%% Failures & &261 &146 &0 \\ +%% \bottomrule +%% \end{tabular} +%% \end{table} + + + +%% \begin{table} +%% \renewcommand{\arraystretch}{1.3} +%% \caption{TestU01 Statistical Test Failures for Old CI algorithms ($\mathsf{N}=4$)} +%% \label{TestU01 for Old CI} +%% \centering +%% \begin{tabular}{lcccc} +%% \toprule +%% \multirow{3}*{Test name} & \multicolumn{4}{c}{Old CI}\\ +%% &Logistic& XORshift& ISAAC&ISAAC \\ +%% &+& +& + & + \\ +%% &Logistic& XORshift& XORshift&ISAAC \\ \cmidrule(r){2-5} +%% Rabbit &7 &2 &0 &0 \\ +%% Alphabit & 3 &0 &0 &0 \\ +%% DieHARD &0 &0 &0 &0 \\ +%% FIPS\_140\_2 &0 &0 &0 &0 \\ +%% SmallCrush &2 &0 &0 &0 \\ +%% Crush &47 &4 &0 &0 \\ +%% Big Crush &79 &3 &0 &0 \\ \hline +%% Failures &138 &9 &0 &0 \\ +%% \bottomrule +%% \end{tabular} +%% \end{table} + + + + + +%% \subsection{Statistical tests} +%% \label{Security analysis} + +%% Three batteries of tests are reputed and regularly used +%% to evaluate the statistical properties of newly designed pseudorandom +%% number generators. These batteries are named DieHard~\cite{Marsaglia1996}, +%% the NIST suite~\cite{ANDREW2008}, and the most stringent one called +%% TestU01~\cite{LEcuyerS07}, which encompasses the two other batteries. + + + +%% \label{Results and discussion} +%% \begin{table*} +%% \renewcommand{\arraystretch}{1.3} +%% \caption{NIST and DieHARD tests suite passing rates for PRNGs without CI} +%% \label{NIST and DieHARD tests suite passing rate the for PRNGs without CI} +%% \centering +%% \begin{tabular}{|l||c|c|c|c|c|c|c|c|c|c|} +%% \hline\hline +%% Types of PRNGs & \multicolumn{2}{c|}{Linear PRNGs} & \multicolumn{4}{c|}{Lagged PRNGs} & \multicolumn{1}{c|}{ICG PRNGs} & \multicolumn{3}{c|}{Mixed PRNGs}\\ \hline +%% \backslashbox{\textbf{$Tests$}} {\textbf{$PRNG$}} & LCG& MRG& AWC & SWB & SWC & GFSR & INV & LCG2& LCG3& MRG2 \\ \hline +%% NIST & 11/15 & 14/15 &\textbf{15/15} & \textbf{15/15} & 14/15 & 14/15 & 14/15 & 14/15& 14/15& 14/15 \\ \hline +%% DieHARD & 16/18 & 16/18 & 15/18 & 16/18 & \textbf{18/18} & 16/18 & 16/18 & 16/18& 16/18& 16/18\\ \hline +%% \end{tabular} +%% \end{table*} + +%% Table~\ref{NIST and DieHARD tests suite passing rate the for PRNGs without CI} shows the +%% results on the two first batteries recalled above, indicating that all the PRNGs presented +%% in the previous section +%% cannot pass all these tests. In other words, the statistical quality of these PRNGs cannot +%% fulfill the up-to-date standards presented previously. We have shown in~\cite{bfg12a:ip} that the use of chaotic +%% iterations can solve this issue. +%% %More precisely, to +%% %illustrate the effects of chaotic iterations on these defective PRNGs, experiments have been divided in three parts~\cite{bfg12a:ip}: +%% %\begin{enumerate} +%% % \item \textbf{Single CIPRNG}: The PRNGs involved in CI computing are of the same category. +%% % \item \textbf{Mixed CIPRNG}: Two different types of PRNGs are mixed during the chaotic iterations process. +%% % \item \textbf{Multiple CIPRNG}: The generator is obtained by repeating the composition of the iteration function as follows: $x^0\in \mathds{B}^{\mathsf{N}}$, and $\forall n\in \mathds{N}^{\ast },\forall i\in \llbracket1;\mathsf{N}\rrbracket, x_i^n=$ +%% %\begin{equation} +%% %\begin{array}{l} +%% %\left\{ +%% %\begin{array}{l} +%% %x_i^{n-1}~~~~~\text{if}~S^n\neq i \\ +%% %\forall j\in \llbracket1;\mathsf{m}\rrbracket,f^m(x^{n-1})_{S^{nm+j}}~\text{if}~S^{nm+j}=i.\end{array} \right. \end{array} +%% %\end{equation} +%% %$m$ is called the \emph{functional power}. +%% %\end{enumerate} +%% % +%% The obtained results are reproduced in Table +%% \ref{NIST and DieHARD tests suite passing rate the for single CIPRNGs}. +%% The scores written in boldface indicate that all the tests have been passed successfully, whereas an +%% asterisk ``*'' means that the considered passing rate has been improved. +%% The improvements are obvious for both the ``Old CI'' and the ``New CI'' generators. +%% Concerning the ``Xor CI PRNG'', the score is less spectacular. Because of a large speed improvement, the statistics +%% are not as good as for the two other versions of these CIPRNGs. +%% However 8 tests have been improved (with no deflation for the other results). + + +%% \begin{table*} +%% \renewcommand{\arraystretch}{1.3} +%% \caption{NIST and DieHARD tests suite passing rates for PRNGs with CI} +%% \label{NIST and DieHARD tests suite passing rate the for single CIPRNGs} +%% \centering +%% \begin{tabular}{|l||c|c|c|c|c|c|c|c|c|c|c|c|} +%% \hline +%% Types of PRNGs & \multicolumn{2}{c|}{Linear PRNGs} & \multicolumn{4}{c|}{Lagged PRNGs} & \multicolumn{1}{c|}{ICG PRNGs} & \multicolumn{3}{c|}{Mixed PRNGs}\\ \hline +%% \backslashbox{\textbf{$Tests$}} {\textbf{$Single~CIPRNG$}} & LCG & MRG & AWC & SWB & SWC & GFSR & INV& LCG2 & LCG3& MRG2 \\ \hline\hline +%% Old CIPRNG\\ \hline \hline +%% NIST & \textbf{15/15} * & \textbf{15/15} * & \textbf{15/15} & \textbf{15/15} & \textbf{15/15} * & \textbf{15/15} * & \textbf{15/15} *& \textbf{15/15} * & \textbf{15/15} * & \textbf{15/15} \\ \hline +%% DieHARD & \textbf{18/18} * & \textbf{18/18} * & \textbf{18/18} * & \textbf{18/18} * & \textbf{18/18} & \textbf{18/18} * & \textbf{18/18} *& \textbf{18/18} * & \textbf{18/18} *& \textbf{18/18} * \\ \hline +%% New CIPRNG\\ \hline \hline +%% NIST & \textbf{15/15} * & \textbf{15/15} * & \textbf{15/15} & \textbf{15/15} & \textbf{15/15} * & \textbf{15/15} * & \textbf{15/15} *& \textbf{15/15} * & \textbf{15/15} * & \textbf{15/15} \\ \hline +%% DieHARD & \textbf{18/18} * & \textbf{18/18} * & \textbf{18/18} * & \textbf{18/18} * & \textbf{18/18} & \textbf{18/18} * & \textbf{18/18} * & \textbf{18/18} * & \textbf{18/18} *& \textbf{18/18} *\\ \hline +%% Xor CIPRNG\\ \hline\hline +%% NIST & 14/15*& \textbf{15/15} * & \textbf{15/15} & \textbf{15/15} & 14/15 & \textbf{15/15} * & 14/15& \textbf{15/15} * & \textbf{15/15} *& \textbf{15/15} \\ \hline +%% DieHARD & 16/18 & 16/18 & 17/18* & \textbf{18/18} * & \textbf{18/18} & \textbf{18/18} * & 16/18 & 16/18 & 16/18& 16/18\\ \hline +%% \end{tabular} +%% \end{table*} + + +%% We have then investigated in~\cite{bfg12a:ip} if it were possible to improve +%% the statistical behavior of the Xor CI version by combining more than one +%% $\oplus$ operation. Results are summarized in Table~\ref{threshold}, illustrating +%% the progressive increasing effects of chaotic iterations, when giving time to chaos to get settled in. +%% Thus rapid and perfect PRNGs, regarding the NIST and DieHARD batteries, can be obtained +%% using chaotic iterations on defective generators. + +%% \begin{table*} +%% \renewcommand{\arraystretch}{1.3} +%% \caption{Number of $\oplus$ operations to pass the whole NIST and DieHARD batteries} +%% \label{threshold} +%% \centering +%% \begin{tabular}{|l||c|c|c|c|c|c|c|c|} +%% \hline +%% Inputted $PRNG$ & LCG & MRG & SWC & GFSR & INV& LCG2 & LCG3 & MRG2 \\ \hline\hline +%% Threshold value $m$& 19 & 7 & 2& 1 & 11& 9& 3& 4\\ \hline\hline +%% \end{tabular} +%% \end{table*} + +%% Finally, the TestU01 battery has been launched on three well-known generators +%% (a logistic map, a simple XORshift, and the cryptographically secure ISAAC, +%% see Table~\ref{TestU011}). These results can be compared with +%% Table~\ref{TestU01 for Old CI}, which gives the scores obtained by the +%% Old CI PRNG that has received these generators. +%% The obvious improvement speaks for itself, and together with the other +%% results recalled in this section, it reinforces the opinion that a strong +%% correlation between topological properties and statistical behavior exists. + + +%% The next subsection will now give a concrete original implementation of the Xor CI PRNG, the +%% fastest generator in the chaotic iteration based family. In the remainder, +%% this generator will be simply referred to as CIPRNG, or ``the proposed PRNG'', if this statement does not +%% raise ambiguity. + + +\section{First Efficient Implementation of a PRNG based on Chaotic Iterations} \label{sec:efficient PRNG} - -Based on the proof presented in the previous section, it is now possible to -improve the speed of the generator formerly presented in~\cite{bgw09:ip,guyeux10}. -The first idea is to consider -that the provided strategy is a pseudorandom Boolean vector obtained by a -given PRNG. -An iteration of the system is simply the bitwise exclusive or between -the last computed state and the current strategy. -Topological properties of disorder exhibited by chaotic -iterations can be inherited by the inputted generator, hoping by doing so to -obtain some statistical improvements while preserving speed. - - -Let us give an example using 16-bits numbers, to clearly understand how the bitwise xor operations -are -done. -Suppose that $x$ and the strategy $S^i$ are given as -binary vectors. -Table~\ref{TableExemple} shows the result of $x \oplus S^i$. - -\begin{table} -$$ -\begin{array}{|cc|cccccccccccccccc|} -\hline -x &=&1&0&1&1&1&0&1&0&1&0&0&1&0&0&1&0\\ -\hline -S^i &=&0&1&1&0&0&1&1&0&1&1&1&0&0&1&1&1\\ -\hline -x \oplus S^i&=&1&1&0&1&1&1&0&0&0&1&1&1&0&1&0&1\\ -\hline - -\hline - \end{array} -$$ -\caption{Example of an arbitrary round of the proposed generator} -\label{TableExemple} -\end{table} - - - -\lstset{language=C,caption={C code of the sequential PRNG based on chaotic iterations},label=algo:seqCIPRNG} +% +%Based on the proof presented in the previous section, it is now possible to +%improve the speed of the generator formerly presented in~\cite{bgw09:ip,guyeux10}. +%The first idea is to consider +%that the provided strategy is a pseudorandom Boolean vector obtained by a +%given PRNG. +%An iteration of the system is simply the bitwise exclusive or between +%the last computed state and the current strategy. +%Topological properties of disorder exhibited by chaotic +%iterations can be inherited by the inputted generator, we hope by doing so to +%obtain some statistical improvements while preserving speed. +% +%%RAPH : j'ai viré tout ca +%% Let us give an example using 16-bits numbers, to clearly understand how the bitwise xor operations +%% are +%% done. +%% Suppose that $x$ and the strategy $S^i$ are given as +%% binary vectors. +%% Table~\ref{TableExemple} shows the result of $x \oplus S^i$. + +%% \begin{table} +%% \begin{scriptsize} +%% $$ +%% \begin{array}{|cc|cccccccccccccccc|} +%% \hline +%% x &=&1&0&1&1&1&0&1&0&1&0&0&1&0&0&1&0\\ +%% \hline +%% S^i &=&0&1&1&0&0&1&1&0&1&1&1&0&0&1&1&1\\ +%% \hline +%% x \oplus S^i&=&1&1&0&1&1&1&0&0&0&1&1&1&0&1&0&1\\ +%% \hline + +%% \hline +%% \end{array} +%% $$ +%% \end{scriptsize} +%% \caption{Example of an arbitrary round of the proposed generator} +%% \label{TableExemple} +%% \end{table} + + + + +\lstset{language=C,caption={C code of the sequential PRNG based on chaotic iterations},label={algo:seqCIPRNG}} +\begin{small} \begin{lstlisting} + unsigned int CIPRNG() { static unsigned int x = 123123123; unsigned long t1 = xorshift(); @@ -862,24 +1379,29 @@ unsigned int CIPRNG() { return x; } \end{lstlisting} +\end{small} +In Listing~\ref{algo:seqCIPRNG} a sequential version of the proposed PRNG based +on chaotic iterations is presented. The xor operator is represented by +\textasciicircum. This function uses three classical 64-bits PRNGs, namely the +\texttt{xorshift}, the \texttt{xor128}, and the +\texttt{xorwow}~\cite{Marsaglia2003}. In the following, we call them ``xor-like +PRNGs''. As each xor-like PRNG uses 64-bits whereas our proposed generator +works with 32-bits, we use the command \texttt{(unsigned int)}, that selects the +32 least significant bits of a given integer, and the code \texttt{(unsigned + int)(t$>>$32)} in order to obtain the 32 most significant bits of \texttt{t}. +Thus producing a pseudorandom number needs 6 xor operations with 6 32-bits numbers +that are provided by 3 64-bits PRNGs. This version successfully passes the +stringent BigCrush battery of tests~\cite{LEcuyerS07}. +At this point, we thus +have defined an efficient and statistically unbiased generator. Its speed is +directly related to the use of linear operations, but for the same reason, +this fast generator cannot be proven as secure. -In Listing~\ref{algo:seqCIPRNG} a sequential version of the proposed PRNG based on chaotic iterations - is presented. The xor operator is represented by \textasciicircum. -This function uses three classical 64-bits PRNGs, namely the \texttt{xorshift}, the -\texttt{xor128}, and the \texttt{xorwow}~\cite{Marsaglia2003}. In the following, we call them -``xor-like PRNGs''. -As -each xor-like PRNG uses 64-bits whereas our proposed generator works with 32-bits, -we use the command \texttt{(unsigned int)}, that selects the 32 least significant bits of a given integer, and the code -\texttt{(unsigned int)(t3$>>$32)} in order to obtain the 32 most significant bits of \texttt{t}. -So producing a pseudorandom number needs 6 xor operations -with 6 32-bits numbers that are provided by 3 64-bits PRNGs. This version successfully passes the -stringent BigCrush battery of tests~\cite{LEcuyerS07}. \section{Efficient PRNGs based on Chaotic Iterations on GPU} \label{sec:efficient PRNG gpu} @@ -890,12 +1412,12 @@ simultaneously. In general, the larger the number of threads is, the more local memory is used, and the less branching instructions are used (if, while, ...), the better the performances on GPU is. Obviously, having these requirements in mind, it is possible to build -a program similar to the one presented in Algorithm +a program similar to the one presented in Listing \ref{algo:seqCIPRNG}, which computes pseudorandom numbers on GPU. To do so, we must firstly recall that in the CUDA~\cite{Nvid10} environment, threads have a local identifier called \texttt{ThreadIdx}, which is relative to the block containing -them. With CUDA parts of the code which are executed by the GPU are +them. Furthermore, in CUDA, parts of the code that are executed by the GPU, are called {\it kernels}. @@ -903,10 +1425,10 @@ called {\it kernels}. It is possible to deduce from the CPU version a quite similar version adapted to GPU. -The simple principle consists to make each thread of the GPU computing the CPU version of our PRNG. +The simple principle consists in making each thread of the GPU computing the CPU version of our PRNG. Of course, the three xor-like PRNGs used in these computations must have different parameters. -In a given thread, these lasts are +In a given thread, these parameters are randomly picked from another PRNGs. The initialization stage is performed by the CPU. To do it, the ISAAC PRNG~\cite{Jenkins96} is used to set all the @@ -919,8 +1441,9 @@ number $x$ that saves the last generated pseudorandom number. Additionally implementation of the xor128, the xorshift, and the xorwow respectively require 4, 5, and 6 unsigned long as internal variables. -\begin{algorithm} +\begin{algorithm} +\begin{small} \KwIn{InternalVarXorLikeArray: array with internal variables of the 3 xor-like PRNGs in global memory\; NumThreads: number of threads\;} @@ -933,14 +1456,16 @@ NumThreads: number of threads\;} } store internal variables in InternalVarXorLikeArray[threadIdx]\; } - +\end{small} \caption{Main kernel of the GPU ``naive'' version of the PRNG based on chaotic iterations} \label{algo:gpu_kernel} \end{algorithm} + + Algorithm~\ref{algo:gpu_kernel} presents a naive implementation of the proposed PRNG on GPU. Due to the available memory in the GPU and the number of threads -used simultenaously, the number of random numbers that a thread can generate +used simultaneously, the number of random numbers that a thread can generate inside a kernel is limited (\emph{i.e.}, the variable \texttt{n} in algorithm~\ref{algo:gpu_kernel}). For instance, if $100,000$ threads are used and if $n=100$\footnote{in fact, we need to add the initial seed (a 32-bits number)}, @@ -951,14 +1476,14 @@ and the pseudorandom numbers generated by our PRNG, is equal to $100,000\ This generator is able to pass the whole BigCrush battery of tests, for all the versions that have been tested depending on their number of threads -(called \texttt{NumThreads} in our algorithm, tested until $10$ millions). +(called \texttt{NumThreads} in our algorithm, tested up to $5$ million). \begin{remark} -The proposed algorithm has the advantage to manipulate independent +The proposed algorithm has the advantage of manipulating independent PRNGs, so this version is easily adaptable on a cluster of computers too. The only thing to ensure is to use a single ISAAC PRNG. To achieve this requirement, a simple solution consists in using a master node for the initialization. This master node computes the initial parameters -for all the differents nodes involves in the computation. +for all the different nodes involved in the computation. \end{remark} \subsection{Improved Version for GPU} @@ -972,45 +1497,47 @@ thread uses the result of which other one, we can use a combination array that contains the indexes of all threads and for which a combination has been performed. -In Algorithm~\ref{algo:gpu_kernel2}, two combination arrays are used. -The variable \texttt{offset} is computed using the value of +In Algorithm~\ref{algo:gpu_kernel2}, two combination arrays are used. The +variable \texttt{offset} is computed using the value of \texttt{combination\_size}. Then we can compute \texttt{o1} and \texttt{o2} -representing the indexes of the other threads whose results are used -by the current one. In this algorithm, we consider that a 64-bits xor-like -PRNG has been chosen, and so its two 32-bits parts are used. +representing the indexes of the other threads whose results are used by the +current one. In this algorithm, we consider that a 32-bits xor-like PRNG has +been chosen. In practice, we use the xor128 proposed in~\cite{Marsaglia2003} in +which unsigned longs (64 bits) have been replaced by unsigned integers (32 +bits). -This version also can pass the whole {\it BigCrush} battery of tests. +This version can also pass the whole {\it BigCrush} battery of tests. \begin{algorithm} - +\begin{small} \KwIn{InternalVarXorLikeArray: array with internal variables of 1 xor-like PRNGs in global memory\; NumThreads: Number of threads\; -tab1, tab2: Arrays containing combinations of size combination\_size\;} +array\_comb1, array\_comb2: Arrays containing combinations of size combination\_size\;} \KwOut{NewNb: array containing random numbers in global memory} \If{threadId is concerned} { retrieve data from InternalVarXorLikeArray[threadId] in local variables including shared memory and x\; offset = threadIdx\%combination\_size\; - o1 = threadIdx-offset+tab1[offset]\; - o2 = threadIdx-offset+tab2[offset]\; + o1 = threadIdx-offset+array\_comb1[offset]\; + o2 = threadIdx-offset+array\_comb2[offset]\; \For{i=1 to n} { t=xor-like()\; - t=t$\oplus$shmem[o1]$\oplus$shmem[o2]\; + t=t\textasciicircum shmem[o1]\textasciicircum shmem[o2]\; shared\_mem[threadId]=t\; - x = x $\oplus$ t\; + x = x\textasciicircum t\; store the new PRNG in NewNb[NumThreads*threadId+i]\; } store internal variables in InternalVarXorLikeArray[threadId]\; } - -\caption{main kernel for the chaotic iterations based PRNG GPU efficient -version} -\label{algo:gpu_kernel2} +\end{small} +\caption{Main kernel for the chaotic iterations based PRNG GPU efficient +version\label{IR}} +\label{algo:gpu_kernel2} \end{algorithm} -\subsection{Theoretical Evaluation of the Improved Version} +\subsection{Chaos Evaluation of the Improved Version} A run of Algorithm~\ref{algo:gpu_kernel2} consists in an operation ($x=x\oplus t$) having the form of Equation~\ref{equation Oplus}, which is equivalent to the iterative @@ -1021,7 +1548,7 @@ and two values previously obtained by two other threads). To be certain that we are in the framework of Theorem~\ref{t:chaos des general}, we must guarantee that this dynamical system iterates on the space $\mathcal{X} = \mathcal{P}\left(\llbracket 1, \mathsf{N} \rrbracket\right)^\mathds{N}\times\mathds{B}^\mathsf{N}$. -The left term $x$ obviously belongs into $\mathds{B}^ \mathsf{N}$. +The left term $x$ obviously belongs to $\mathds{B}^ \mathsf{N}$. To prevent from any flaws of chaotic properties, we must check that the right term (the last $t$), corresponding to the strategies, can possibly be equal to any integer of $\llbracket 1, \mathsf{N} \rrbracket$. @@ -1032,8 +1559,7 @@ last $t$ respects the requirement. Furthermore, it is possible to prove by an immediate mathematical induction that, as the initial $x$ is uniformly distributed (it is provided by a cryptographically secure PRNG), the two other stored values shmem[o1] and shmem[o2] are uniformly distributed too, -(this can be stated by an immediate mathematical -induction), and thus the next $x$ is finally uniformly distributed. +(this is the induction hypothesis), and thus the next $x$ is finally uniformly distributed. Thus Algorithm~\ref{algo:gpu_kernel2} is a concrete realization of the general chaotic iterations presented previously, and for this reason, it satisfies the @@ -1051,7 +1577,7 @@ All the cards have 240 cores. In Figure~\ref{fig:time_xorlike_gpu} we compare the quantity of pseudorandom numbers -generated per second with various xor-like based PRNG. In this figure, the optimized +generated per second with various xor-like based PRNGs. In this figure, the optimized versions use the {\it xor64} described in~\cite{Marsaglia2003}, whereas the naive versions embed the three xor-like PRNGs described in Listing~\ref{algo:seqCIPRNG}. In order to obtain the optimal performances, the storage of pseudorandom numbers @@ -1059,7 +1585,7 @@ into the GPU memory has been removed. This step is time consuming and slows down generation. Moreover this storage is completely useless, in case of applications that consume the pseudorandom numbers directly after generation. We can see that when the number of threads is greater -than approximately 30,000 and lower than 5 millions, the number of pseudorandom numbers generated +than approximately 30,000 and lower than 5 million, the number of pseudorandom numbers generated per second is almost constant. With the naive version, this value ranges from 2.5 to 3GSamples/s. With the optimized version, it is approximately equal to 20GSamples/s. Finally we can remark that both GPU cards are quite similar, but in @@ -1070,7 +1596,7 @@ As a comparison, Listing~\ref{algo:seqCIPRNG} leads to the generation of \begin{figure}[htbp] \begin{center} - \includegraphics[scale=.7]{curve_time_xorlike_gpu.pdf} + \includegraphics[width=\columnwidth]{curve_time_xorlike_gpu.pdf} \end{center} \caption{Quantity of pseudorandom numbers generated per second with the xorlike-based PRNG} \label{fig:time_xorlike_gpu} @@ -1080,17 +1606,16 @@ As a comparison, Listing~\ref{algo:seqCIPRNG} leads to the generation of -In Figure~\ref{fig:time_bbs_gpu} we highlight the performances of the optimized -BBS-based PRNG on GPU. On the Tesla C1060 we -obtain approximately 1.8GSample/s and on the GTX 280 about 1.6GSample/s, which is -obviously slower than the xorlike-based PRNG on GPU. However, we will show in the -next sections that -this new PRNG has a strong level of security, which is necessary paid by a speed -reduction. +In Figure~\ref{fig:time_bbs_gpu} we highlight the performances of the optimized +BBS-based PRNG on GPU. On the Tesla C1060 we obtain approximately 700MSample/s +and on the GTX 280 about 670MSample/s, which is obviously slower than the +xorlike-based PRNG on GPU. However, we will show in the next sections that this +new PRNG has a strong level of security, which is necessarily paid by a speed +reduction. \begin{figure}[htbp] \begin{center} - \includegraphics[scale=.7]{curve_time_bbs_gpu.pdf} + \includegraphics[width=\columnwidth]{curve_time_bbs_gpu.pdf} \end{center} \caption{Quantity of pseudorandom numbers generated per second using the BBS-based PRNG} \label{fig:time_bbs_gpu} @@ -1098,7 +1623,7 @@ reduction. All these experiments allow us to conclude that it is possible to generate a very large quantity of pseudorandom numbers statistically perfect with the xor-like version. -In a certain extend, it is the case too with the secure BBS-based version, the speed deflation being +To a certain extend, it is also the case with the secure BBS-based version, the speed deflation being explained by the fact that the former version has ``only'' chaotic properties and statistical perfection, whereas the latter is also cryptographically secure, as it is shown in the next sections. @@ -1110,34 +1635,52 @@ as it is shown in the next sections. \section{Security Analysis} -\label{sec:security analysis} +This section is dedicated to the security analysis of the + proposed PRNGs, both from a theoretical and from a practical point of view. + +\subsection{Theoretical Proof of Security} +\label{sec:security analysis} + +The standard definition + of {\it indistinguishability} used is the classical one as defined for + instance in~\cite[chapter~3]{Goldreich}. + This property shows that predicting the future results of the PRNG + cannot be done in a reasonable time compared to the generation time. It is important to emphasize that this + is a relative notion between breaking time and the sizes of the + keys/seeds. Of course, if small keys or seeds are chosen, the system can + be broken in practice. But it also means that if the keys/seeds are large + enough, the system is secured. +As a complement, an example of a concrete practical evaluation of security +is outlined in the next subsection. In this section the concatenation of two strings $u$ and $v$ is classically denoted by $uv$. In a cryptographic context, a pseudorandom generator is a deterministic algorithm $G$ transforming strings into strings and such that, for any -seed $w$ of length $N$, $G(w)$ (the output of $G$ on the input $w$) has size -$\ell_G(N)$ with $\ell_G(N)>N$. +seed $s$ of length $m$, $G(s)$ (the output of $G$ on the input $s$) has size +$\ell_G(m)$ with $\ell_G(m)>m$. The notion of {\it secure} PRNGs can now be defined as follows. \begin{definition} A cryptographic PRNG $G$ is secure if for any probabilistic polynomial time algorithm $D$, for any positive polynomial $p$, and for all sufficiently -large $k$'s, -$$| \mathrm{Pr}[D(G(U_k))=1]-Pr[D(U_{\ell_G(k)})=1]|< \frac{1}{p(N)},$$ +large $m$'s, +$$| \mathrm{Pr}[D(G(U_m))=1]-Pr[D(U_{\ell_G(m)})=1]|< \frac{1}{p(m)},$$ where $U_r$ is the uniform distribution over $\{0,1\}^r$ and the -probabilities are taken over $U_N$, $U_{\ell_G(N)}$ as well as over the +probabilities are taken over $U_m$, $U_{\ell_G(m)}$ as well as over the internal coin tosses of $D$. \end{definition} -Intuitively, it means that there is no polynomial time algorithm that can -distinguish a perfect uniform random generator from $G$ with a non -negligible probability. The interested reader is referred -to~\cite[chapter~3]{Goldreich} for more information. Note that it is -quite easily possible to change the function $\ell$ into any polynomial -function $\ell^\prime$ satisfying $\ell^\prime(N)>N)$~\cite[Chapter 3.3]{Goldreich}. +Intuitively, it means that there is no polynomial time algorithm that can +distinguish a perfect uniform random generator from $G$ with a non negligible +probability. An equivalent formulation of this well-known security property +means that it is possible \emph{in practice} to predict the next bit of the +generator, knowing all the previously produced ones. The interested reader is +referred to~\cite[chapter~3]{Goldreich} for more information. Note that it is +quite easily possible to change the function $\ell$ into any polynomial function +$\ell^\prime$ satisfying $\ell^\prime(m)>m)$~\cite[Chapter 3.3]{Goldreich}. The generation schema developed in (\ref{equation Oplus}) is based on a pseudorandom generator. Let $H$ be a cryptographic PRNG. We may assume, @@ -1148,7 +1691,7 @@ strings of length $N$ such that $H(S_0)=S_1 \ldots S_k$ ($H(S_0)$ is the concate the $S_i$'s). The cryptographic PRNG $X$ defined in (\ref{equation Oplus}) is the algorithm mapping any string of length $2N$ $x_0S_0$ into the string $(x_0\oplus S_0 \oplus S_1)(x_0\oplus S_0 \oplus S_1\oplus S_2)\ldots -(x_o\bigoplus_{i=0}^{i=k}S_i)$. Particularly one has $\ell_{X}(2N)=kN=\ell_H(N)$. +(x_o\bigoplus_{i=0}^{i=k}S_i)$. One in particular has $\ell_{X}(2N)=kN=\ell_H(N)$. We claim now that if this PRNG is secure, then the new one is secure too. @@ -1159,7 +1702,7 @@ PRNG too. \end{proposition} \begin{proof} -The proposition is proved by contraposition. Assume that $X$ is not +The proposition is proven by contraposition. Assume that $X$ is not secure. By Definition, there exists a polynomial time probabilistic algorithm $D$, a positive polynomial $p$, such that for all $k_0$ there exists $N\geq \frac{k_0}{2}$ satisfying @@ -1192,8 +1735,10 @@ $y\bigoplus_{i=1}^{i=j} w_i^\prime=y\bigoplus_{i=1}^{i=j} w_i$. It follows, by a direct induction, that $w_i=w_i^\prime$. Furthermore, since $\mathbb{B}^{kN}$ is finite, each $\varphi_y$ is bijective. Therefore, and using (\ref{PCH-1}), one has +$\mathrm{Pr}[D^\prime(U_{kN})=1]=\mathrm{Pr}[D(\varphi_y(U_{kN}))=1]$ and, +therefore, \begin{equation}\label{PCH-2} -\mathrm{Pr}[D^\prime(U_{kN})=1]=\mathrm{Pr}[D(\varphi_y(U_{kN}))=1]=\mathrm{Pr}[D(U_{kN})=1]. +\mathrm{Pr}[D^\prime(U_{kN})=1]=\mathrm{Pr}[D(U_{kN})=1]. \end{equation} Now, using (\ref{PCH-1}) again, one has for every $x$, @@ -1202,7 +1747,7 @@ D^\prime(H(x))=D(\varphi_y(H(x))), \end{equation} where $y$ is randomly generated. By construction, $\varphi_y(H(x))=X(yx)$, thus -\begin{equation}\label{PCH-3} +\begin{equation}%\label{PCH-3} %%RAPH : j'ai viré ce label qui existe déjà, il est 3 ligne avant D^\prime(H(x))=D(yx), \end{equation} where $y$ is randomly generated. @@ -1212,14 +1757,109 @@ It follows that \mathrm{Pr}[D^\prime(H(U_{N}))=1]=\mathrm{Pr}[D(U_{2N})=1]. \end{equation} From (\ref{PCH-2}) and (\ref{PCH-4}), one can deduce that -there exist a polynomial time probabilistic +there exists a polynomial time probabilistic algorithm $D^\prime$, a positive polynomial $p$, such that for all $k_0$ there exists $N\geq \frac{k_0}{2}$ satisfying $$| \mathrm{Pr}[D(H(U_{N}))=1]-\mathrm{Pr}[D(U_{kN}=1]|\geq \frac{1}{p(2N)},$$ -proving that $H$ is not secure, a contradiction. +proving that $H$ is not secure, which is a contradiction. \end{proof} + +\subsection{Practical Security Evaluation} +\label{sec:Practicak evaluation} +This subsection is given in Section~\ref{A-sec:Practicak evaluation} of the annex document. +%%RAF mis en annexe + + +%% Pseudorandom generators based on Eq.~\eqref{equation Oplus} are thus cryptographically secure when +%% they are XORed with an already cryptographically +%% secure PRNG. But, as stated previously, +%% such a property does not mean that, whatever the +%% key size, no attacker can predict the next bit +%% knowing all the previously released ones. +%% However, given a key size, it is possible to +%% measure in practice the minimum duration needed +%% for an attacker to break a cryptographically +%% secure PRNG, if we know the power of his/her +%% machines. Such a concrete security evaluation +%% is related to the $(T,\varepsilon)-$security +%% notion, which is recalled and evaluated in what +%% follows, for the sake of completeness. + +%% Let us firstly recall that, +%% \begin{definition} +%% Let $\mathcal{D} : \mathds{B}^M \longrightarrow \mathds{B}$ be a probabilistic algorithm that runs +%% in time $T$. +%% Let $\varepsilon > 0$. +%% $\mathcal{D}$ is called a $(T,\varepsilon)-$distinguishing attack on pseudorandom +%% generator $G$ if + +%% \begin{flushleft} +%% $\left| Pr[\mathcal{D}(G(k)) = 1 \mid k \in_R \{0,1\}^\ell ]\right.$ +%% \end{flushleft} + +%% \begin{flushright} +%% $ - \left. Pr[\mathcal{D}(s) = 1 \mid s \in_R \mathds{B}^M ]\right| \geqslant \varepsilon,$ +%% \end{flushright} + +%% \noindent where the probability is taken over the internal coin flips of $\mathcal{D}$, and the notation +%% ``$\in_R$'' indicates the process of selecting an element at random and uniformly over the +%% corresponding set. +%% \end{definition} + +%% Let us recall that the running time of a probabilistic algorithm is defined to be the +%% maximum of the expected number of steps needed to produce an output, maximized +%% over all inputs; the expected number is averaged over all coin flips made by the algorithm~\cite{Knuth97}. +%% We are now able to define the notion of cryptographically secure PRNGs: + +%% \begin{definition} +%% A pseudorandom generator is $(T,\varepsilon)-$secure if there exists no $(T,\varepsilon)-$distinguishing attack on this pseudorandom generator. +%% \end{definition} + + + + + + + +%% Suppose now that the PRNG of Eq.~\eqref{equation Oplus} will work during +%% $M=100$ time units, and that during this period, +%% an attacker can realize $10^{12}$ clock cycles. +%% We thus wonder whether, during the PRNG's +%% lifetime, the attacker can distinguish this +%% sequence from a truly random one, with a probability +%% greater than $\varepsilon = 0.2$. +%% We consider that $N$ has 900 bits. + +%% Predicting the next generated bit knowing all the +%% previously released ones by Eq.~\eqref{equation Oplus} is obviously equivalent to predicting the +%% next bit in the BBS generator, which +%% is cryptographically secure. More precisely, it +%% is $(T,\varepsilon)-$secure: no +%% $(T,\varepsilon)-$distinguishing attack can be +%% successfully realized on this PRNG, if~\cite{Fischlin} +%% \begin{equation} +%% T \leqslant \dfrac{L(N)}{6 N (log_2(N))\varepsilon^{-2}M^2}-2^7 N \varepsilon^{-2} M^2 log_2 (8 N \varepsilon^{-1}M) +%% \label{mesureConcrete} +%% \end{equation} +%% where $M$ is the length of the output ($M=100$ in +%% our example), and $L(N)$ is equal to +%% $$ +%% 2.8\times 10^{-3} exp \left(1.9229 \times (N ~ln~ 2)^\frac{1}{3} \times (ln(N~ln~ 2))^\frac{2}{3}\right) +%% $$ +%% is the number of clock cycles to factor a $N-$bit +%% integer. + + + + +%% A direct numerical application shows that this attacker +%% cannot achieve its $(10^{12},0.2)$ distinguishing +%% attack in that context. + + + \section{Cryptographical Applications} \subsection{A Cryptographically Secure PRNG for GPU} @@ -1229,121 +1869,244 @@ It is possible to build a cryptographically secure PRNG based on the previous algorithm (Algorithm~\ref{algo:gpu_kernel2}). Due to Proposition~\ref{cryptopreuve}, it simply consists in replacing the {\it xor-like} PRNG by a cryptographically secure one. -We have chosen the Blum Blum Shum generator~\cite{BBS} (usually denoted by BBS) having the form: -$$x_{n+1}=x_n^2~ mod~ M$$ where $M$ is the product of two prime numbers. These -prime numbers need to be congruent to 3 modulus 4. BBS is +We have chosen the Blum Blum Shub generator~\cite{BBS} (usually denoted by BBS) having the form: +$$x_{n+1}=x_n^2~ mod~ M$$ where $M$ is the product of two prime numbers (these +prime numbers need to be congruent to 3 modulus 4). BBS is known to be very slow and only usable for cryptographic applications. The modulus operation is the most time consuming operation for current GPU cards. So in order to obtain quite reasonable performances, it is -required to use only modulus on 32 bits integer numbers. Consequently -$x_n^2$ need to be less than $2^{32}$ and the number $M$ need to be -less than $2^{16}$. So in practice we can choose prime numbers around -256 that are congruent to 3 modulus 4. With 32 bits numbers, only the +required to use only modulus on 32-bits integer numbers. Consequently +$x_n^2$ need to be lesser than $2^{32}$, and thus the number $M$ must be +lesser than $2^{16}$. So in practice we can choose prime numbers around +256 that are congruent to 3 modulus 4. With 32-bits numbers, only the 4 least significant bits of $x_n$ can be chosen (the maximum number of indistinguishable bits is lesser than or equals to -$log_2(log_2(x_n))$). So to generate a 32 bits number, we need to use -8 times the BBS algorithm with different combinations of $M$. This -approach is not sufficient to pass all the tests of TestU01 because -the fact of having chosen small values of $M$ for the BBS leads to -have a small period. So, in order to add randomness we proceed with +$log_2(log_2(M))$). In other words, to generate a 32-bits number, we need to use +8 times the BBS algorithm with possibly different combinations of $M$. This +approach is not sufficient to be able to pass all the tests of TestU01, +as small values of $M$ for the BBS lead to + small periods. So, in order to add randomness we have proceeded with the followings modifications. \begin{itemize} \item -First we define 16 arrangement arrays instead of 2 (as described in -algorithm \ref{algo:gpu_kernel2}) but only 2 are used at each call of -the PRNG kernels. In practice, the selection of which combinations -arrays will be used is different for all the threads and is determined +Firstly, we define 16 arrangement arrays instead of 2 (as described in +Algorithm \ref{algo:gpu_kernel2}), but only 2 of them are used at each call of +the PRNG kernels. In practice, the selection of combination +arrays to be used is different for all the threads. It is determined by using the three last bits of two internal variables used by BBS. -This approach adds more randomness. In algorithm~\ref{algo:bbs_gpu}, -character \& performs the AND bitwise. So using \&7 with a number -gives the last 3 bits, so it provides a number between 0 and 7. +%This approach adds more randomness. +In Algorithm~\ref{algo:bbs_gpu}, +character \& is for the bitwise AND. Thus using \&7 with a number +gives the last 3 bits, thus providing a number between 0 and 7. \item -Second, after the generation of the 8 BBS numbers for each thread we -have a 32 bits number for which the period is possibly quite small. So -to add randomness, we generate 4 more BBS numbers which allows us to -shift the 32 bits numbers and add upto 6 new bits. This part is -described in algorithm~\ref{algo:bbs_gpu}. In practice, if we call -{\it strategy}, the number representing the strategy, the last 2 bits -of the first new BBS number are used to make a left shift of at least -3 bits. The last 3 bits of the second new BBS number are add to the +Secondly, after the generation of the 8 BBS numbers for each thread, we +have a 32-bits number whose period is possibly quite small. So +to add randomness, we generate 4 more BBS numbers to +shift the 32-bits numbers, and add up to 6 new bits. This improvement is +described in Algorithm~\ref{algo:bbs_gpu}. In practice, the last 2 bits +of the first new BBS number are used to make a left shift of at most +3 bits. The last 3 bits of the second new BBS number are added to the strategy whatever the value of the first left shift. The third and the fourth new BBS numbers are used similarly to apply a new left shift and add 3 new bits. \item -Finally, as we use 8 BBS numbers for each thread, the store of these +Finally, as we use 8 BBS numbers for each thread, the storage of these numbers at the end of the kernel is performed using a rotation. So, internal variable for BBS number 1 is stored in place 2, internal -variable for BBS number 2 is store ind place 3, ... and internal +variable for BBS number 2 is stored in place 3, ..., and finally, internal variable for BBS number 8 is stored in place 1. \end{itemize} - \begin{algorithm} - +\begin{small} \KwIn{InternalVarBBSArray: array with internal variables of the 8 BBS in global memory\; NumThreads: Number of threads\; -tab: 2D Arrays containing 16 combinations (in first dimension) of size combination\_size (in second dimension)\;} +array\_comb: 2D Arrays containing 16 combinations (in first dimension) of size combination\_size (in second dimension)\; +array\_shift[4]=\{0,1,3,7\}\; +} \KwOut{NewNb: array containing random numbers in global memory} \If{threadId is concerned} { retrieve data from InternalVarBBSArray[threadId] in local variables including shared memory and x\; we consider that bbs1 ... bbs8 represent the internal states of the 8 BBS numbers\; offset = threadIdx\%combination\_size\; - o1 = threadIdx-offset+tab[bbs1\&7][offset]\; - o2 = threadIdx-offset+tab[8+bbs2\&7][offset]\; + o1 = threadIdx-offset+array\_comb[bbs1\&7][offset]\; + o2 = threadIdx-offset+array\_comb[8+bbs2\&7][offset]\; \For{i=1 to n} { - t<<=4\; + t$<<$=4\; t|=BBS1(bbs1)\&15\; ...\; - t<<=4\; + t$<<$=4\; t|=BBS8(bbs8)\&15\; - //two new shifts\; - t<<=BBS3(bbs3)\&3\; - t|=BBS1(bbs1)\&7\; - t<<=BBS7(bbs7)\&3\; - t|=BBS2(bbs2)\&7\; - t=t$\oplus$shmem[o1]$\oplus$shmem[o2]\; + \tcp{two new shifts} + shift=BBS3(bbs3)\&3\; + t$<<$=shift\; + t|=BBS1(bbs1)\&array\_shift[shift]\; + shift=BBS7(bbs7)\&3\; + t$<<$=shift\; + t|=BBS2(bbs2)\&array\_shift[shift]\; + t=t\textasciicircum shmem[o1]\textasciicircum shmem[o2]\; shared\_mem[threadId]=t\; - x = x $\oplus$ t\; + x = x\textasciicircum t\; store the new PRNG in NewNb[NumThreads*threadId+i]\; } store internal variables in InternalVarXorLikeArray[threadId] using a rotation\; } - +\end{small} \caption{main kernel for the BBS based PRNG GPU} \label{algo:bbs_gpu} \end{algorithm} -In algorithm~\ref{algo:bbs_gpu}, t<<=4 performs a left shift of 4 bits -on the variable t and stores the result in t. BBS1(bbs1)\&15 selects -the last four bits of the result of BBS1. It should be noticed that -for the two new shifts, we use arbitrarily 4 BBSs that have previously -been used. - +In Algorithm~\ref{algo:bbs_gpu}, $n$ is for the quantity of random numbers that +a thread has to generate. The operation t<<=4 performs a left shift of 4 bits +on the variable $t$ and stores the result in $t$, and $BBS1(bbs1)\&15$ selects +the last four bits of the result of $BBS1$. Thus an operation of the form +$t<<=4; t|=BBS1(bbs1)\&15\;$ realizes in $t$ a left shift of 4 bits, and then +puts the 4 last bits of $BBS1(bbs1)$ in the four last positions of $t$. Let us +remark that the initialization $t$ is not a necessity as we fill it 4 bits by 4 +bits, until having obtained 32-bits. The two last new shifts are realized in +order to enlarge the small periods of the BBS used here, to introduce a kind of +variability. In these operations, we make twice a left shift of $t$ of \emph{at + most} 3 bits, represented by \texttt{shift} in the algorithm, and we put +\emph{exactly} the \texttt{shift} last bits from a BBS into the \texttt{shift} +last bits of $t$. For this, an array named \texttt{array\_shift}, containing the +correspondence between the shift and the number obtained with \texttt{shift} 1 +to make the \texttt{and} operation is used. For example, with a left shift of 0, +we make an and operation with 0, with a left shift of 3, we make an and +operation with 7 (represented by 111 in binary mode). + +It should be noticed that this generator has once more the form $x^{n+1} = x^n \oplus S^n$, +where $S^n$ is referred in this algorithm as $t$: each iteration of this +PRNG ends with $x = x \wedge t$. This $S^n$ is only constituted +by secure bits produced by the BBS generator, and thus, due to +Proposition~\ref{cryptopreuve}, the resulted PRNG is +cryptographically secure. + +As stated before, even if the proposed PRNG is cryptocaphically +secure, it does not mean that such a generator +can be used as described here when attacks are +awaited. The problem is to determine the minimum +time required for an attacker, with a given +computational power, to predict under a probability +lower than 0.5 the $n+1$th bit, knowing the $n$ +previous ones. The proposed GPU generator will be +useful in a security context, at least in some +situations where a secret protected by a pseudorandom +keystream is rapidly obsolete, if this time to +predict the next bit is large enough when compared +to both the generation and transmission times. +It is true that the prime numbers used in the last +section are very small compared to up-to-date +security recommendations. However the attacker has not +access to each BBS, but to the output produced +by Algorithm~\ref{algo:bbs_gpu}, which is far +more complicated than a simple BBS. Indeed, to +determine if this cryptographically secure PRNG +on GPU can be useful in security context with the +proposed parameters, or if it is only a very fast +and statistically perfect generator on GPU, its +$(T,\varepsilon)-$security must be determined, and +a formulation similar to Eq.\eqref{mesureConcrete} +must be established. Authors +hope to achieve this difficult task in a future +work. + + +\subsection{Toward a Cryptographically Secure and Chaotic Asymmetric Cryptosystem} +\label{Blum-Goldwasser} +We finish this research work by giving some thoughts about the use of +the proposed PRNG in an asymmetric cryptosystem. +This first approach will be further investigated in a future work. + +\subsubsection{Recalls of the Blum-Goldwasser Probabilistic Cryptosystem} + +The Blum-Goldwasser cryptosystem is a cryptographically secure asymmetric key encryption algorithm +proposed in 1984~\cite{Blum:1985:EPP:19478.19501}. The encryption algorithm +implements a XOR-based stream cipher using the BBS PRNG, in order to generate +the keystream. Decryption is done by obtaining the initial seed thanks to +the final state of the BBS generator and the secret key, thus leading to the + reconstruction of the keystream. + +The key generation consists in generating two prime numbers $(p,q)$, +randomly and independently of each other, that are + congruent to 3 mod 4, and to compute the modulus $N=pq$. +The public key is $N$, whereas the secret key is the factorization $(p,q)$. + + +Suppose Bob wishes to send a string $m=(m_0, \dots, m_{L-1})$ of $L$ bits to Alice: +\begin{enumerate} +\item Bob picks an integer $r$ randomly in the interval $\llbracket 1,N\rrbracket$ and computes $x_0 = r^2~mod~N$. +\item He uses the BBS to generate the keystream of $L$ pseudorandom bits $(b_0, \dots, b_{L-1})$, as follows. For $i=0$ to $L-1$, +\begin{itemize} +\item $i=0$. +\item While $i \leqslant L-1$: +\begin{itemize} +\item Set $b_i$ equal to the least-significant\footnote{As signaled previously, BBS can securely output up to $\mathsf{N} = \lfloor log(log(N)) \rfloor$ of the least-significant bits of $x_i$ during each round.} bit of $x_i$, +\item $i=i+1$, +\item $x_i = (x_{i-1})^2~mod~N.$ +\end{itemize} +\end{itemize} +\item The ciphertext is computed by XORing the plaintext bits $m$ with the keystream: $ c = (c_0, \dots, c_{L-1}) = m \oplus b$. This ciphertext is $[c, y]$, where $y=x_{0}^{2^{L}}~mod~N.$ +\end{enumerate} -\subsection{A Secure Asymetric Cryptosystem} +When Alice receives $\left[(c_0, \dots, c_{L-1}), y\right]$, she can recover $m$ as follows: +\begin{enumerate} +\item Using the secret key $(p,q)$, she computes $r_p = y^{((p+1)/4)^{L}}~mod~p$ and $r_q = y^{((q+1)/4)^{L}}~mod~q$. +\item The initial seed can be obtained using the following procedure: $x_0=q(q^{-1}~{mod}~p)r_p + p(p^{-1}~{mod}~q)r_q~{mod}~N$. +\item She recomputes the bit-vector $b$ by using BBS and $x_0$. +\item Alice finally computes the plaintext by XORing the keystream with the ciphertext: $ m = c \oplus b$. +\end{enumerate} +\subsubsection{Proposal of a new Asymmetric Cryptosystem Adapted from Blum-Goldwasser} +We propose to adapt the Blum-Goldwasser protocol as follows. +Let $\mathsf{N} = \lfloor log(log(N)) \rfloor$ be the number of bits that can +be obtained securely with the BBS generator using the public key $N$ of Alice. +Alice will pick randomly $S^0$ in $\llbracket 0, 2^{\mathsf{N}-1}\rrbracket$ too, and +her new public key will be $(S^0, N)$. -\section{Conclusion} +To encrypt his message, Bob will compute +%%RAPH : ici, j'ai mis un simple $ +%\begin{equation} +$c = \left(m_0 \oplus (b_0 \oplus S^0), m_1 \oplus (b_0 \oplus b_1 \oplus S^0), \hdots, \right.$ +$ \left. m_{L-1} \oplus (b_0 \oplus b_1 \hdots \oplus b_{L-1} \oplus S^0) \right)$ +%%\end{equation} +instead of $\left(m_0 \oplus b_0, m_1 \oplus b_1, \hdots, m_{L-1} \oplus b_{L-1} \right)$. +The same decryption stage as in Blum-Goldwasser leads to the sequence +$\left(m_0 \oplus S^0, m_1 \oplus S^0, \hdots, m_{L-1} \oplus S^0 \right)$. +Thus, with a simple use of $S^0$, Alice can obtain the plaintext. +By doing so, the proposed generator is used in place of BBS, leading to +the inheritance of all the properties presented in this paper. -In this paper we have presented a new class of PRNGs based on chaotic -iterations. We have proven that these PRNGs are chaotic in the sense of Devaney. -We also propose a PRNG cryptographically secure and its implementation on GPU. +\section{Conclusion} -An efficient implementation on GPU based on a xor-like PRNG allows us to -generate a huge number of pseudorandom numbers per second (about -20Gsamples/s). This PRNG succeeds to pass the hardest batteries of TestU01. -In future work we plan to extend this work for parallel PRNG for clusters or -grid computing. +In this paper, a formerly proposed PRNG based on chaotic iterations +has been generalized to improve its speed. It has been proven to be +chaotic according to Devaney. +Efficient implementations on GPU using xor-like PRNGs as input generators +have shown that a very large quantity of pseudorandom numbers can be generated per second (about +20Gsamples/s), and that these proposed PRNGs succeed to pass the hardest battery in TestU01, +namely the BigCrush. +Furthermore, we have shown that when the inputted generator is cryptographically +secure, then it is the case too for the PRNG we propose, thus leading to +the possibility to develop fast and secure PRNGs using the GPU architecture. +An improvement of the Blum-Goldwasser cryptosystem, making it +behave chaotically, has finally been proposed. + +In future work we plan to extend this research, building a parallel PRNG for clusters or +grid computing. Topological properties of the various proposed generators will be investigated, +and the use of other categories of PRNGs as input will be studied too. The improvement +of Blum-Goldwasser will be deepened. Finally, we +will try to enlarge the quantity of pseudorandom numbers generated per second either +in a simulation context or in a cryptographic one.