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doc-src/ind-defs.tex

author | clasohm |

Thu, 29 Jun 1995 12:28:27 +0200 | |

changeset 1163 | c080ff36d24e |

parent 597 | ebf373c17ee2 |

child 1197 | ae58cd15e802 |

permissions | -rw-r--r-- |

changed 'chol' labels to 'hol'; added a few parentheses

\documentstyle[a4,proof,iman,extra,12pt]{llncs} \newif\ifCADE \CADEfalse \title{A Fixedpoint Approach to Implementing\\ (Co)Inductive Definitions\thanks{J. Grundy and S. Thompson made detailed comments; the referees were also helpful. Research funded by SERC grants GR/G53279, GR/H40570 and by the ESPRIT Project 6453 `Types'.}} \author{Lawrence C. Paulson\\{\tt lcp@cl.cam.ac.uk}} \institute{Computer Laboratory, University of Cambridge, England} \date{\today} \setcounter{secnumdepth}{2} \setcounter{tocdepth}{2} \newcommand\sbs{\subseteq} \let\To=\Rightarrow \newcommand\pow{{\cal P}} %%%\let\pow=\wp \newcommand\RepFun{\hbox{\tt RepFun}} \newcommand\cons{\hbox{\tt cons}} \def\succ{\hbox{\tt succ}} \newcommand\split{\hbox{\tt split}} \newcommand\fst{\hbox{\tt fst}} \newcommand\snd{\hbox{\tt snd}} \newcommand\converse{\hbox{\tt converse}} \newcommand\domain{\hbox{\tt domain}} \newcommand\range{\hbox{\tt range}} \newcommand\field{\hbox{\tt field}} \newcommand\lfp{\hbox{\tt lfp}} \newcommand\gfp{\hbox{\tt gfp}} \newcommand\id{\hbox{\tt id}} \newcommand\trans{\hbox{\tt trans}} \newcommand\wf{\hbox{\tt wf}} \newcommand\nat{\hbox{\tt nat}} \newcommand\rank{\hbox{\tt rank}} \newcommand\univ{\hbox{\tt univ}} \newcommand\Vrec{\hbox{\tt Vrec}} \newcommand\Inl{\hbox{\tt Inl}} \newcommand\Inr{\hbox{\tt Inr}} \newcommand\case{\hbox{\tt case}} \newcommand\lst{\hbox{\tt list}} \newcommand\Nil{\hbox{\tt Nil}} \newcommand\Cons{\hbox{\tt Cons}} \newcommand\lstcase{\hbox{\tt list\_case}} \newcommand\lstrec{\hbox{\tt list\_rec}} \newcommand\length{\hbox{\tt length}} \newcommand\listn{\hbox{\tt listn}} \newcommand\acc{\hbox{\tt acc}} \newcommand\primrec{\hbox{\tt primrec}} \newcommand\SC{\hbox{\tt SC}} \newcommand\CONST{\hbox{\tt CONST}} \newcommand\PROJ{\hbox{\tt PROJ}} \newcommand\COMP{\hbox{\tt COMP}} \newcommand\PREC{\hbox{\tt PREC}} \newcommand\quniv{\hbox{\tt quniv}} \newcommand\llist{\hbox{\tt llist}} \newcommand\LNil{\hbox{\tt LNil}} \newcommand\LCons{\hbox{\tt LCons}} \newcommand\lconst{\hbox{\tt lconst}} \newcommand\lleq{\hbox{\tt lleq}} \newcommand\map{\hbox{\tt map}} \newcommand\term{\hbox{\tt term}} \newcommand\Apply{\hbox{\tt Apply}} \newcommand\termcase{\hbox{\tt term\_case}} \newcommand\rev{\hbox{\tt rev}} \newcommand\reflect{\hbox{\tt reflect}} \newcommand\tree{\hbox{\tt tree}} \newcommand\forest{\hbox{\tt forest}} \newcommand\Part{\hbox{\tt Part}} \newcommand\TF{\hbox{\tt tree\_forest}} \newcommand\Tcons{\hbox{\tt Tcons}} \newcommand\Fcons{\hbox{\tt Fcons}} \newcommand\Fnil{\hbox{\tt Fnil}} \newcommand\TFcase{\hbox{\tt TF\_case}} \newcommand\Fin{\hbox{\tt Fin}} \newcommand\QInl{\hbox{\tt QInl}} \newcommand\QInr{\hbox{\tt QInr}} \newcommand\qsplit{\hbox{\tt qsplit}} \newcommand\qcase{\hbox{\tt qcase}} \newcommand\Con{\hbox{\tt Con}} \newcommand\data{\hbox{\tt data}} \binperiod %%%treat . like a binary operator \begin{document} \pagestyle{empty} \begin{titlepage} \maketitle \begin{abstract} This paper presents a fixedpoint approach to inductive definitions. Instead of using a syntactic test such as `strictly positive,' the approach lets definitions involve any operators that have been proved monotone. It is conceptually simple, which has allowed the easy implementation of mutual recursion and other conveniences. It also handles coinductive definitions: simply replace the least fixedpoint by a greatest fixedpoint. This represents the first automated support for coinductive definitions. The method has been implemented in two of Isabelle's logics, ZF set theory and higher-order logic. It should be applicable to any logic in which the Knaster-Tarski Theorem can be proved. Examples include lists of $n$ elements, the accessible part of a relation and the set of primitive recursive functions. One example of a coinductive definition is bisimulations for lazy lists. \ifCADE\else Recursive datatypes are examined in detail, as well as one example of a {\bf codatatype}: lazy lists. The appendices are simple user's manuals for this Isabelle package.\fi \end{abstract} % \bigskip\centerline{Copyright \copyright{} \number\year{} by Lawrence C. Paulson} \thispagestyle{empty} \end{titlepage} \tableofcontents\cleardoublepage\pagestyle{plain} \section{Introduction} Several theorem provers provide commands for formalizing recursive data structures, like lists and trees. Examples include Boyer and Moore's shell principle~\cite{bm79} and Melham's recursive type package for the Cambridge HOL system~\cite{melham89}. Such data structures are called {\bf datatypes} below, by analogy with {\tt datatype} definitions in Standard~ML\@. A datatype is but one example of an {\bf inductive definition}. This specifies the least set closed under given rules~\cite{aczel77}. The collection of theorems in a logic is inductively defined. A structural operational semantics~\cite{hennessy90} is an inductive definition of a reduction or evaluation relation on programs. A few theorem provers provide commands for formalizing inductive definitions; these include Coq~\cite{paulin92} and again the HOL system~\cite{camilleri92}. The dual notion is that of a {\bf coinductive definition}. This specifies the greatest set closed under given rules. Important examples include using bisimulation relations to formalize equivalence of processes~\cite{milner89} or lazy functional programs~\cite{abramsky90}. Other examples include lazy lists and other infinite data structures; these are called {\bf codatatypes} below. Not all inductive definitions are meaningful. {\bf Monotone} inductive definitions are a large, well-behaved class. Monotonicity can be enforced by syntactic conditions such as `strictly positive,' but this could lead to monotone definitions being rejected on the grounds of their syntactic form. More flexible is to formalize monotonicity within the logic and allow users to prove it. This paper describes a package based on a fixedpoint approach. Least fixedpoints yield inductive definitions; greatest fixedpoints yield coinductive definitions. The package has several advantages: \begin{itemize} \item It allows reference to any operators that have been proved monotone. Thus it accepts all provably monotone inductive definitions, including iterated definitions. \item It accepts a wide class of datatype definitions, including those with infinite branching. \item It handles coinductive and codatatype definitions. Most of the discussion below applies equally to inductive and coinductive definitions, and most of the code is shared. To my knowledge, this is the only package supporting coinductive definitions. \item Definitions may be mutually recursive. \end{itemize} The package has been implemented in Isabelle~\cite{isabelle-intro} using ZF set theory \cite{paulson-set-I,paulson-set-II}; part of it has since been ported to Isabelle's higher-order logic. However, the fixedpoint approach is independent of Isabelle. The recursion equations are specified as introduction rules for the mutually recursive sets. The package transforms these rules into a mapping over sets, and attempts to prove that the mapping is monotonic and well-typed. If successful, the package makes fixedpoint definitions and proves the introduction, elimination and (co)induction rules. The package consists of several Standard ML functors~\cite{paulson91}; it accepts its argument and returns its result as ML structures.\footnote{This use of ML modules is not essential; the package could also be implemented as a function on records.} Most datatype packages equip the new datatype with some means of expressing recursive functions. This is the main omission from my package. Its fixedpoint operators define only recursive sets. To define recursive functions, the Isabelle/ZF theory provides well-founded recursion and other logical tools~\cite{paulson-set-II}. {\bf Outline.} Section~2 introduces the least and greatest fixedpoint operators. Section~3 discusses the form of introduction rules, mutual recursion and other points common to inductive and coinductive definitions. Section~4 discusses induction and coinduction rules separately. Section~5 presents several examples, including a coinductive definition. Section~6 describes datatype definitions. Section~7 presents related work. Section~8 draws brief conclusions. \ifCADE\else The appendices are simple user's manuals for this Isabelle package.\fi Most of the definitions and theorems shown below have been generated by the package. I have renamed some variables to improve readability. \section{Fixedpoint operators} In set theory, the least and greatest fixedpoint operators are defined as follows: \begin{eqnarray*} \lfp(D,h) & \equiv & \inter\{X\sbs D. h(X)\sbs X\} \\ \gfp(D,h) & \equiv & \union\{X\sbs D. X\sbs h(X)\} \end{eqnarray*} Let $D$ be a set. Say that $h$ is {\bf bounded by}~$D$ if $h(D)\sbs D$, and {\bf monotone below~$D$} if $h(A)\sbs h(B)$ for all $A$ and $B$ such that $A\sbs B\sbs D$. If $h$ is bounded by~$D$ and monotone then both operators yield fixedpoints: \begin{eqnarray*} \lfp(D,h) & = & h(\lfp(D,h)) \\ \gfp(D,h) & = & h(\gfp(D,h)) \end{eqnarray*} These equations are instances of the Knaster-Tarski Theorem, which states that every monotonic function over a complete lattice has a fixedpoint~\cite{davey&priestley}. It is obvious from their definitions that $\lfp$ must be the least fixedpoint, and $\gfp$ the greatest. This fixedpoint theory is simple. The Knaster-Tarski Theorem is easy to prove. Showing monotonicity of~$h$ is trivial, in typical cases. We must also exhibit a bounding set~$D$ for~$h$. Frequently this is trivial, as when a set of `theorems' is (co)inductively defined over some previously existing set of `formulae.' Isabelle/ZF provides a suitable bounding set for finitely branching (co)datatype definitions; see~\S\ref{univ-sec} below. Bounding sets are also called {\bf domains}. The powerset operator is monotone, but by Cantor's Theorem there is no set~$A$ such that $A=\pow(A)$. We cannot put $A=\lfp(D,\pow)$ because there is no suitable domain~$D$. But \S\ref{acc-sec} demonstrates that~$\pow$ is still useful in inductive definitions. \section{Elements of an inductive or coinductive definition}\label{basic-sec} Consider a (co)inductive definition of the sets $R_1$, \ldots,~$R_n$, in mutual recursion. They will be constructed from domains $D_1$, \ldots,~$D_n$, respectively. The construction yields not $R_i\sbs D_i$ but $R_i\sbs D_1+\cdots+D_n$, where $R_i$ is contained in the image of~$D_i$ under an injection. Reasons for this are discussed elsewhere~\cite[\S4.5]{paulson-set-II}. The definition may involve arbitrary parameters $\vec{p}=p_1$, \ldots,~$p_k$. Each recursive set then has the form $R_i(\vec{p})$. The parameters must be identical every time they occur within a definition. This would appear to be a serious restriction compared with other systems such as Coq~\cite{paulin92}. For instance, we cannot define the lists of $n$ elements as the set $\listn(A,n)$ using rules where the parameter~$n$ varies. Section~\ref{listn-sec} describes how to express this set using the inductive definition package. To avoid clutter below, the recursive sets are shown as simply $R_i$ instead of $R_i(\vec{p})$. \subsection{The form of the introduction rules}\label{intro-sec} The body of the definition consists of the desired introduction rules, specified as strings. The conclusion of each rule must have the form $t\in R_i$, where $t$ is any term. Premises typically have the same form, but they can have the more general form $t\in M(R_i)$ or express arbitrary side-conditions. The premise $t\in M(R_i)$ is permitted if $M$ is a monotonic operator on sets, satisfying the rule \[ \infer{M(A)\sbs M(B)}{A\sbs B} \] The user must supply the package with monotonicity rules for all such premises. The ability to introduce new monotone operators makes the approach flexible. A suitable choice of~$M$ and~$t$ can express a lot. The powerset operator $\pow$ is monotone, and the premise $t\in\pow(R)$ expresses $t\sbs R$; see \S\ref{acc-sec} for an example. The `list of' operator is monotone, as is easily proved by induction. The premise $t\in\lst(R)$ avoids having to encode the effect of~$\lst(R)$ using mutual recursion; see \S\ref{primrec-sec} and also my earlier paper~\cite[\S4.4]{paulson-set-II}. Introduction rules may also contain {\bf side-conditions}. These are premises consisting of arbitrary formulae not mentioning the recursive sets. Side-conditions typically involve type-checking. One example is the premise $a\in A$ in the following rule from the definition of lists: \[ \infer{\Cons(a,l)\in\lst(A)}{a\in A & l\in\lst(A)} \] \subsection{The fixedpoint definitions} The package translates the list of desired introduction rules into a fixedpoint definition. Consider, as a running example, the finite powerset operator $\Fin(A)$: the set of all finite subsets of~$A$. It can be defined as the least set closed under the rules \[ \emptyset\in\Fin(A) \qquad \infer{\{a\}\un b\in\Fin(A)}{a\in A & b\in\Fin(A)} \] The domain in a (co)inductive definition must be some existing set closed under the rules. A suitable domain for $\Fin(A)$ is $\pow(A)$, the set of all subsets of~$A$. The package generates the definition \begin{eqnarray*} \Fin(A) & \equiv & \lfp(\pow(A), \; \begin{array}[t]{r@{\,}l} \lambda X. \{z\in\pow(A). & z=\emptyset \disj{} \\ &(\exists a\,b. z=\{a\}\un b\conj a\in A\conj b\in X)\}) \end{array} \end{eqnarray*} The contribution of each rule to the definition of $\Fin(A)$ should be obvious. A coinductive definition is similar but uses $\gfp$ instead of~$\lfp$. The package must prove that the fixedpoint operator is applied to a monotonic function. If the introduction rules have the form described above, and if the package is supplied a monotonicity theorem for every $t\in M(R_i)$ premise, then this proof is trivial.\footnote{Due to the presence of logical connectives in the fixedpoint's body, the monotonicity proof requires some unusual rules. These state that the connectives $\conj$, $\disj$ and $\exists$ preserve monotonicity with respect to the partial ordering on unary predicates given by $P\sqsubseteq Q$ if and only if $\forall x.P(x)\imp Q(x)$.} The package returns its result as an ML structure, which consists of named components; we may regard it as a record. The result structure contains the definitions of the recursive sets as a theorem list called {\tt defs}. It also contains, as the theorem {\tt unfold}, a fixedpoint equation such as \begin{eqnarray*} \Fin(A) & = & \begin{array}[t]{r@{\,}l} \{z\in\pow(A). & z=\emptyset \disj{} \\ &(\exists a\,b. z=\{a\}\un b\conj a\in A\conj b\in \Fin(A))\} \end{array} \end{eqnarray*} It also contains, as the theorem {\tt dom\_subset}, an inclusion such as $\Fin(A)\sbs\pow(A)$. \subsection{Mutual recursion} \label{mutual-sec} In a mutually recursive definition, the domain of the fixedpoint construction is the disjoint sum of the domain~$D_i$ of each~$R_i$, for $i=1$, \ldots,~$n$. The package uses the injections of the binary disjoint sum, typically $\Inl$ and~$\Inr$, to express injections $h_{1n}$, \ldots, $h_{nn}$ for the $n$-ary disjoint sum $D_1+\cdots+D_n$. As discussed elsewhere \cite[\S4.5]{paulson-set-II}, Isabelle/ZF defines the operator $\Part$ to support mutual recursion. The set $\Part(A,h)$ contains those elements of~$A$ having the form~$h(z)$: \begin{eqnarray*} \Part(A,h) & \equiv & \{x\in A. \exists z. x=h(z)\}. \end{eqnarray*} For mutually recursive sets $R_1$, \ldots,~$R_n$ with $n>1$, the package makes $n+1$ definitions. The first defines a set $R$ using a fixedpoint operator. The remaining $n$ definitions have the form \begin{eqnarray*} R_i & \equiv & \Part(R,h_{in}), \qquad i=1,\ldots, n. \end{eqnarray*} It follows that $R=R_1\un\cdots\un R_n$, where the $R_i$ are pairwise disjoint. \subsection{Proving the introduction rules} The user supplies the package with the desired form of the introduction rules. Once it has derived the theorem {\tt unfold}, it attempts to prove those rules. From the user's point of view, this is the trickiest stage; the proofs often fail. The task is to show that the domain $D_1+\cdots+D_n$ of the combined set $R_1\un\cdots\un R_n$ is closed under all the introduction rules. This essentially involves replacing each~$R_i$ by $D_1+\cdots+D_n$ in each of the introduction rules and attempting to prove the result. Consider the $\Fin(A)$ example. After substituting $\pow(A)$ for $\Fin(A)$ in the rules, the package must prove \[ \emptyset\in\pow(A) \qquad \infer{\{a\}\un b\in\pow(A)}{a\in A & b\in\pow(A)} \] Such proofs can be regarded as type-checking the definition.\footnote{The Isabelle/HOL version does not require these proofs, as HOL has implicit type-checking.} The user supplies the package with type-checking rules to apply. Usually these are general purpose rules from the ZF theory. They could however be rules specifically proved for a particular inductive definition; sometimes this is the easiest way to get the definition through! The result structure contains the introduction rules as the theorem list {\tt intrs}. \subsection{The case analysis rule} The elimination rule, called {\tt elim}, performs case analysis. There is one case for each introduction rule. The elimination rule for $\Fin(A)$ is \[ \infer{Q}{x\in\Fin(A) & \infer*{Q}{[x=\emptyset]} & \infer*{Q}{[x=\{a\}\un b & a\in A &b\in\Fin(A)]_{a,b}} } \] The subscripted variables $a$ and~$b$ above the third premise are eigenvariables, subject to the usual `not free in \ldots' proviso. The rule states that if $x\in\Fin(A)$ then either $x=\emptyset$ or else $x=\{a\}\un b$ for some $a\in A$ and $b\in\Fin(A)$; it is a simple consequence of {\tt unfold}. The package also returns a function for generating simplified instances of the case analysis rule. It works for datatypes and for inductive definitions involving datatypes, such as an inductively defined relation between lists. It instantiates {\tt elim} with a user-supplied term then simplifies the cases using freeness of the underlying datatype. The simplified rules perform `rule inversion' on the inductive definition. Section~\S\ref{mkcases} presents an example. \section{Induction and coinduction rules} Here we must consider inductive and coinductive definitions separately. For an inductive definition, the package returns an induction rule derived directly from the properties of least fixedpoints, as well as a modified rule for mutual recursion and inductively defined relations. For a coinductive definition, the package returns a basic coinduction rule. \subsection{The basic induction rule}\label{basic-ind-sec} The basic rule, called {\tt induct}, is appropriate in most situations. For inductive definitions, it is strong rule induction~\cite{camilleri92}; for datatype definitions (see below), it is just structural induction. The induction rule for an inductively defined set~$R$ has the following form. The major premise is $x\in R$. There is a minor premise for each introduction rule: \begin{itemize} \item If the introduction rule concludes $t\in R_i$, then the minor premise is~$P(t)$. \item The minor premise's eigenvariables are precisely the introduction rule's free variables that are not parameters of~$R$. For instance, the eigenvariables in the $\Fin(A)$ rule below are $a$ and $b$, but not~$A$. \item If the introduction rule has a premise $t\in R_i$, then the minor premise discharges the assumption $t\in R_i$ and the induction hypothesis~$P(t)$. If the introduction rule has a premise $t\in M(R_i)$ then the minor premise discharges the single assumption \[ t\in M(\{z\in R_i. P(z)\}). \] Because $M$ is monotonic, this assumption implies $t\in M(R_i)$. The occurrence of $P$ gives the effect of an induction hypothesis, which may be exploited by appealing to properties of~$M$. \end{itemize} The induction rule for $\Fin(A)$ resembles the elimination rule shown above, but includes an induction hypothesis: \[ \infer{P(x)}{x\in\Fin(A) & P(\emptyset) & \infer*{P(\{a\}\un b)}{[a\in A & b\in\Fin(A) & P(b)]_{a,b}} } \] Stronger induction rules often suggest themselves. We can derive a rule for $\Fin(A)$ whose third premise discharges the extra assumption $a\not\in b$. The Isabelle/ZF theory defines the {\bf rank} of a set~\cite[\S3.4]{paulson-set-II}, which supports well-founded induction and recursion over datatypes. The package proves a rule for mutual induction and inductive relations. \subsection{Mutual induction} The mutual induction rule is called {\tt mutual\_induct}. It differs from the basic rule in several respects: \begin{itemize} \item Instead of a single predicate~$P$, it uses $n$ predicates $P_1$, \ldots,~$P_n$: one for each recursive set. \item There is no major premise such as $x\in R_i$. Instead, the conclusion refers to all the recursive sets: \[ (\forall z.z\in R_1\imp P_1(z))\conj\cdots\conj (\forall z.z\in R_n\imp P_n(z)) \] Proving the premises establishes $P_i(z)$ for $z\in R_i$ and $i=1$, \ldots,~$n$. \item If the domain of some $R_i$ is the Cartesian product $A_1\times\cdots\times A_m$ (associated to the right), then the corresponding predicate $P_i$ takes $m$ arguments and the corresponding conjunct of the conclusion is \[ (\forall z_1\ldots z_m.\pair{z_1,\ldots,z_m}\in R_i\imp P_i(z_1,\ldots,z_m)) \] \end{itemize} The last point above simplifies reasoning about inductively defined relations. It eliminates the need to express properties of $z_1$, \ldots,~$z_m$ as properties of the tuple $\pair{z_1,\ldots,z_m}$. \subsection{Coinduction}\label{coind-sec} A coinductive definition yields a primitive coinduction rule, with no refinements such as those for the induction rules. (Experience may suggest refinements later.) Consider the codatatype of lazy lists as an example. For suitable definitions of $\LNil$ and $\LCons$, lazy lists may be defined as the greatest fixedpoint satisfying the rules \[ \LNil\in\llist(A) \qquad \infer[(-)]{\LCons(a,l)\in\llist(A)}{a\in A & l\in\llist(A)} \] The $(-)$ tag stresses that this is a coinductive definition. A suitable domain for $\llist(A)$ is $\quniv(A)$, a set closed under variant forms of sum and product for representing infinite data structures (see~\S\ref{univ-sec}). Coinductive definitions use these variant sums and products. The package derives an {\tt unfold} theorem similar to that for $\Fin(A)$. Then it proves the theorem {\tt coinduct}, which expresses that $\llist(A)$ is the greatest solution to this equation contained in $\quniv(A)$: \[ \infer{x\in\llist(A)}{x\in X & X\sbs \quniv(A) & \infer*{z=\LNil\disj \bigl(\exists a\,l.\, z=\LCons(a,l) \conj a\in A \conj l\in X\un\llist(A) \bigr)} {[z\in X]_z}} % \begin{array}[t]{@{}l} % z=\LCons(a,l) \conj a\in A \conj{}\\ % l\in X\un\llist(A) \bigr) % \end{array} }{[z\in X]_z}} \] This rule complements the introduction rules; it provides a means of showing $x\in\llist(A)$ when $x$ is infinite. For instance, if $x=\LCons(0,x)$ then applying the rule with $X=\{x\}$ proves $x\in\llist(\nat)$. (Here $\nat$ is the set of natural numbers.) Having $X\un\llist(A)$ instead of simply $X$ in the third premise above represents a slight strengthening of the greatest fixedpoint property. I discuss several forms of coinduction rules elsewhere~\cite{paulson-coind}. \section{Examples of inductive and coinductive definitions}\label{ind-eg-sec} This section presents several examples: the finite powerset operator, lists of $n$ elements, bisimulations on lazy lists, the well-founded part of a relation, and the primitive recursive functions. \subsection{The finite powerset operator} This operator has been discussed extensively above. Here is the corresponding invocation in an Isabelle theory file. Note that $\cons(a,b)$ abbreviates $\{a\}\un b$ in Isabelle/ZF. \begin{ttbox} Finite = Arith + consts Fin :: "i=>i" inductive domains "Fin(A)" <= "Pow(A)" intrs emptyI "0 : Fin(A)" consI "[| a: A; b: Fin(A) |] ==> cons(a,b) : Fin(A)" type_intrs "[empty_subsetI, cons_subsetI, PowI]" type_elims "[make_elim PowD]" end \end{ttbox} Theory {\tt Finite} extends the parent theory {\tt Arith} by declaring the unary function symbol~$\Fin$, which is defined inductively. Its domain is specified as $\pow(A)$, where $A$ is the parameter appearing in the introduction rules. For type-checking, we supply two introduction rules: \[ \emptyset\sbs A \qquad \infer{\{a\}\un B\sbs C}{a\in C & B\sbs C} \] A further introduction rule and an elimination rule express the two directions of the equivalence $A\in\pow(B)\bimp A\sbs B$. Type-checking involves mostly introduction rules. Like all Isabelle theory files, this one yields a structure containing the new theory as an \ML{} value. Structure {\tt Finite} also has a substructure, called~{\tt Fin}. After declaring \hbox{\tt open Finite;} we can refer to the $\Fin(A)$ introduction rules as the list {\tt Fin.intrs} or individually as {\tt Fin.emptyI} and {\tt Fin.consI}. The induction rule is {\tt Fin.induct}. \subsection{Lists of $n$ elements}\label{listn-sec} This has become a standard example of an inductive definition. Following Paulin-Mohring~\cite{paulin92}, we could attempt to define a new datatype $\listn(A,n)$, for lists of length~$n$, as an $n$-indexed family of sets. But her introduction rules \[ \hbox{\tt Niln}\in\listn(A,0) \qquad \infer{\hbox{\tt Consn}(n,a,l)\in\listn(A,\succ(n))} {n\in\nat & a\in A & l\in\listn(A,n)} \] are not acceptable to the inductive definition package: $\listn$ occurs with three different parameter lists in the definition. The Isabelle version of this example suggests a general treatment of varying parameters. Here, we use the existing datatype definition of $\lst(A)$, with constructors $\Nil$ and~$\Cons$. Then incorporate the parameter~$n$ into the inductive set itself, defining $\listn(A)$ as a relation. It consists of pairs $\pair{n,l}$ such that $n\in\nat$ and~$l\in\lst(A)$ and $l$ has length~$n$. In fact, $\listn(A)$ is the converse of the length function on~$\lst(A)$. The Isabelle/ZF introduction rules are \[ \pair{0,\Nil}\in\listn(A) \qquad \infer{\pair{\succ(n),\Cons(a,l)}\in\listn(A)} {a\in A & \pair{n,l}\in\listn(A)} \] The Isabelle theory file takes, as parent, the theory~{\tt List} of lists. We declare the constant~$\listn$ and supply an inductive definition, specifying the domain as $\nat\times\lst(A)$: \begin{ttbox} ListN = List + consts listn ::"i=>i" inductive domains "listn(A)" <= "nat*list(A)" intrs NilI "<0,Nil> : listn(A)" ConsI "[| a: A; <n,l> : listn(A) |] ==> <succ(n), Cons(a,l)> : listn(A)" type_intrs "nat_typechecks @ list.intrs" end \end{ttbox} The type-checking rules include those for 0, $\succ$, $\Nil$ and $\Cons$. Because $\listn(A)$ is a set of pairs, type-checking requires the equivalence $\pair{a,b}\in A\times B \bimp a\in A \conj b\in B$; the package always includes the necessary rules. The package returns introduction, elimination and induction rules for $\listn$. The basic induction rule, {\tt ListN.induct}, is \[ \infer{P(x)}{x\in\listn(A) & P(\pair{0,\Nil}) & \infer*{P(\pair{\succ(n),\Cons(a,l)})} {[a\in A & \pair{n,l}\in\listn(A) & P(\pair{n,l})]_{a,l,n}}} \] This rule requires the induction formula to be a unary property of pairs,~$P(\pair{n,l})$. The alternative rule, {\tt ListN.mutual\_induct}, uses a binary property instead: \[ \infer{\forall n\,l. \pair{n,l}\in\listn(A) \imp P(n,l)} {P(0,\Nil) & \infer*{P(\succ(n),\Cons(a,l))} {[a\in A & \pair{n,l}\in\listn(A) & P(n,l)]_{a,l,n}}} \] It is now a simple matter to prove theorems about $\listn(A)$, such as \[ \forall l\in\lst(A). \pair{\length(l),\, l}\in\listn(A) \] \[ \listn(A)``\{n\} = \{l\in\lst(A). \length(l)=n\} \] This latter result --- here $r``X$ denotes the image of $X$ under $r$ --- asserts that the inductive definition agrees with the obvious notion of $n$-element list. Unlike in Coq, the definition does not declare a new datatype. A `list of $n$ elements' really is a list and is subject to list operators such as append (concatenation). For example, a trivial induction on $\pair{m,l}\in\listn(A)$ yields \[ \infer{\pair{m\mathbin{+} m,\, l@l'}\in\listn(A)} {\pair{m,l}\in\listn(A) & \pair{m',l'}\in\listn(A)} \] where $+$ here denotes addition on the natural numbers and @ denotes append. \subsection{A demonstration of rule inversion}\label{mkcases} The elimination rule, {\tt ListN.elim}, is cumbersome: \[ \infer{Q}{x\in\listn(A) & \infer*{Q}{[x = \pair{0,\Nil}]} & \infer*{Q} {\left[\begin{array}{l} x = \pair{\succ(n),\Cons(a,l)} \\ a\in A \\ \pair{n,l}\in\listn(A) \end{array} \right]_{a,l,n}}} \] The ML function {\tt ListN.mk\_cases} generates simplified instances of this rule. It works by freeness reasoning on the list constructors: $\Cons(a,l)$ is injective in its two arguments and differs from~$\Nil$. If $x$ is $\pair{i,\Nil}$ or $\pair{i,\Cons(a,l)}$ then {\tt ListN.mk\_cases} deduces the corresponding form of~$i$; this is called rule inversion. For example, \begin{ttbox} ListN.mk_cases List.con_defs "<i,Cons(a,l)> : listn(A)" \end{ttbox} yields a rule with only two premises: \[ \infer{Q}{\pair{i, \Cons(a,l)}\in\listn(A) & \infer*{Q} {\left[\begin{array}{l} i = \succ(n) \\ a\in A \\ \pair{n,l}\in\listn(A) \end{array} \right]_{n}}} \] The package also has built-in rules for freeness reasoning about $0$ and~$\succ$. So if $x$ is $\pair{0,l}$ or $\pair{\succ(i),l}$, then {\tt ListN.mk\_cases} can similarly deduce the corresponding form of~$l$. The function {\tt mk\_cases} is also useful with datatype definitions. The instance from the definition of lists, namely {\tt List.mk\_cases}, can prove the rule \[ \infer{Q}{\Cons(a,l)\in\lst(A) & & \infer*{Q}{[a\in A &l\in\lst(A)]} } \] A typical use of {\tt mk\_cases} concerns inductive definitions of evaluation relations. Then rule inversion yields case analysis on possible evaluations. For example, the Isabelle/ZF theory includes a short proof of the diamond property for parallel contraction on combinators. \subsection{A coinductive definition: bisimulations on lazy lists} This example anticipates the definition of the codatatype $\llist(A)$, which consists of finite and infinite lists over~$A$. Its constructors are $\LNil$ and $\LCons$, satisfying the introduction rules shown in~\S\ref{coind-sec}. Because $\llist(A)$ is defined as a greatest fixedpoint and uses the variant pairing and injection operators, it contains non-well-founded elements such as solutions to $\LCons(a,l)=l$. The next step in the development of lazy lists is to define a coinduction principle for proving equalities. This is done by showing that the equality relation on lazy lists is the greatest fixedpoint of some monotonic operation. The usual approach~\cite{pitts94} is to define some notion of bisimulation for lazy lists, define equivalence to be the greatest bisimulation, and finally to prove that two lazy lists are equivalent if and only if they are equal. The coinduction rule for equivalence then yields a coinduction principle for equalities. A binary relation $R$ on lazy lists is a {\bf bisimulation} provided $R\sbs R^+$, where $R^+$ is the relation \[ \{\pair{\LNil,\LNil}\} \un \{\pair{\LCons(a,l),\LCons(a,l')} . a\in A \conj \pair{l,l'}\in R\}. \] A pair of lazy lists are {\bf equivalent} if they belong to some bisimulation. Equivalence can be coinductively defined as the greatest fixedpoint for the introduction rules \[ \pair{\LNil,\LNil} \in\lleq(A) \qquad \infer[(-)]{\pair{\LCons(a,l),\LCons(a,l')} \in\lleq(A)} {a\in A & \pair{l,l'}\in \lleq(A)} \] To make this coinductive definition, the theory file includes (after the declaration of $\llist(A)$) the following lines: \begin{ttbox} consts lleq :: "i=>i" coinductive domains "lleq(A)" <= "llist(A) * llist(A)" intrs LNil "<LNil, LNil> : lleq(A)" LCons "[| a:A; <l,l'>: lleq(A) |] ==> <LCons(a,l), LCons(a,l')>: lleq(A)" type_intrs "llist.intrs" \end{ttbox} Again, {\tt addconsts} declares a constant for $\lleq$ in the parent theory. The domain of $\lleq(A)$ is $\llist(A)\times\llist(A)$. The type-checking rules include the introduction rules for $\llist(A)$, whose declaration is discussed below (\S\ref{lists-sec}). The package returns the introduction rules and the elimination rule, as usual. But instead of induction rules, it returns a coinduction rule. The rule is too big to display in the usual notation; its conclusion is $x\in\lleq(A)$ and its premises are $x\in X$, ${X\sbs\llist(A)\times\llist(A)}$ and \[ \infer*{z=\pair{\LNil,\LNil}\disj \bigl(\exists a\,l\,l'.\, z=\pair{\LCons(a,l),\LCons(a,l')} \conj a\in A \conj\pair{l,l'}\in X\un\lleq(A) \bigr) % \begin{array}[t]{@{}l} % z=\pair{\LCons(a,l),\LCons(a,l')} \conj a\in A \conj{}\\ % \pair{l,l'}\in X\un\lleq(A) \bigr) % \end{array} }{[z\in X]_z} \] Thus if $x\in X$, where $X$ is a bisimulation contained in the domain of $\lleq(A)$, then $x\in\lleq(A)$. It is easy to show that $\lleq(A)$ is reflexive: the equality relation is a bisimulation. And $\lleq(A)$ is symmetric: its converse is a bisimulation. But showing that $\lleq(A)$ coincides with the equality relation takes some work. \subsection{The accessible part of a relation}\label{acc-sec} Let $\prec$ be a binary relation on~$D$; in short, $(\prec)\sbs D\times D$. The {\bf accessible} or {\bf well-founded} part of~$\prec$, written $\acc(\prec)$, is essentially that subset of~$D$ for which $\prec$ admits no infinite decreasing chains~\cite{aczel77}. Formally, $\acc(\prec)$ is inductively defined to be the least set that contains $a$ if it contains all $\prec$-predecessors of~$a$, for $a\in D$. Thus we need an introduction rule of the form \[ \infer{a\in\acc(\prec)}{\forall y.y\prec a\imp y\in\acc(\prec)} \] Paulin-Mohring treats this example in Coq~\cite{paulin92}, but it causes difficulties for other systems. Its premise is not acceptable to the inductive definition package of the Cambridge HOL system~\cite{camilleri92}. It is also unacceptable to Isabelle package (recall \S\ref{intro-sec}), but fortunately can be transformed into the acceptable form $t\in M(R)$. The powerset operator is monotonic, and $t\in\pow(R)$ is equivalent to $t\sbs R$. This in turn is equivalent to $\forall y\in t. y\in R$. To express $\forall y.y\prec a\imp y\in\acc(\prec)$ we need only find a term~$t$ such that $y\in t$ if and only if $y\prec a$. A suitable $t$ is the inverse image of~$\{a\}$ under~$\prec$. The theory file below follows this approach. Here $r$ is~$\prec$ and $\field(r)$ refers to~$D$, the domain of $\acc(r)$. (The field of a relation is the union of its domain and range.) Finally $r^{-}``\{a\}$ denotes the inverse image of~$\{a\}$ under~$r$. We supply the theorem {\tt Pow\_mono}, which asserts that $\pow$ is monotonic. \begin{ttbox} Acc = WF + consts acc :: "i=>i" inductive domains "acc(r)" <= "field(r)" intrs vimage "[| r-``\{a\}: Pow(acc(r)); a: field(r) |] ==> a: acc(r)" monos "[Pow_mono]" end \end{ttbox} The Isabelle theory proceeds to prove facts about $\acc(\prec)$. For instance, $\prec$ is well-founded if and only if its field is contained in $\acc(\prec)$. As mentioned in~\S\ref{basic-ind-sec}, a premise of the form $t\in M(R)$ gives rise to an unusual induction hypothesis. Let us examine the induction rule, {\tt Acc.induct}: \[ \infer{P(x)}{x\in\acc(r) & \infer*{P(a)}{[r^{-}``\{a\}\in\pow(\{z\in\acc(r).P(z)\}) & a\in\field(r)]_a}} \] The strange induction hypothesis is equivalent to $\forall y. \pair{y,a}\in r\imp y\in\acc(r)\conj P(y)$. Therefore the rule expresses well-founded induction on the accessible part of~$\prec$. The use of inverse image is not essential. The Isabelle package can accept introduction rules with arbitrary premises of the form $\forall \vec{y}.P(\vec{y})\imp f(\vec{y})\in R$. The premise can be expressed equivalently as \[ \{z\in D. P(\vec{y}) \conj z=f(\vec{y})\} \in \pow(R) \] provided $f(\vec{y})\in D$ for all $\vec{y}$ such that~$P(\vec{y})$. The following section demonstrates another use of the premise $t\in M(R)$, where $M=\lst$. \subsection{The primitive recursive functions}\label{primrec-sec} The primitive recursive functions are traditionally defined inductively, as a subset of the functions over the natural numbers. One difficulty is that functions of all arities are taken together, but this is easily circumvented by regarding them as functions on lists. Another difficulty, the notion of composition, is less easily circumvented. Here is a more precise definition. Letting $\vec{x}$ abbreviate $x_0,\ldots,x_{n-1}$, we can write lists such as $[\vec{x}]$, $[y+1,\vec{x}]$, etc. A function is {\bf primitive recursive} if it belongs to the least set of functions in $\lst(\nat)\to\nat$ containing \begin{itemize} \item The {\bf successor} function $\SC$, such that $\SC[y,\vec{x}]=y+1$. \item All {\bf constant} functions $\CONST(k)$, such that $\CONST(k)[\vec{x}]=k$. \item All {\bf projection} functions $\PROJ(i)$, such that $\PROJ(i)[\vec{x}]=x_i$ if $0\leq i<n$. \item All {\bf compositions} $\COMP(g,[f_0,\ldots,f_{m-1}])$, where $g$ and $f_0$, \ldots, $f_{m-1}$ are primitive recursive, such that \begin{eqnarray*} \COMP(g,[f_0,\ldots,f_{m-1}])[\vec{x}] & = & g[f_0[\vec{x}],\ldots,f_{m-1}[\vec{x}]]. \end{eqnarray*} \item All {\bf recursions} $\PREC(f,g)$, where $f$ and $g$ are primitive recursive, such that \begin{eqnarray*} \PREC(f,g)[0,\vec{x}] & = & f[\vec{x}] \\ \PREC(f,g)[y+1,\vec{x}] & = & g[\PREC(f,g)[y,\vec{x}],\, y,\, \vec{x}]. \end{eqnarray*} \end{itemize} Composition is awkward because it combines not two functions, as is usual, but $m+1$ functions. In her proof that Ackermann's function is not primitive recursive, Nora Szasz was unable to formalize this definition directly~\cite{szasz93}. So she generalized primitive recursion to tuple-valued functions. This modified the inductive definition such that each operation on primitive recursive functions combined just two functions. \begin{figure} \begin{ttbox} Primrec = List + consts primrec :: "i" SC :: "i" \(\vdots\) defs SC_def "SC == lam l:list(nat).list_case(0, %x xs.succ(x), l)" \(\vdots\) inductive domains "primrec" <= "list(nat)->nat" intrs SC "SC : primrec" CONST "k: nat ==> CONST(k) : primrec" PROJ "i: nat ==> PROJ(i) : primrec" COMP "[| g: primrec; fs: list(primrec) |] ==> COMP(g,fs): primrec" PREC "[| f: primrec; g: primrec |] ==> PREC(f,g): primrec" monos "[list_mono]" con_defs "[SC_def,CONST_def,PROJ_def,COMP_def,PREC_def]" type_intrs "nat_typechecks @ list.intrs @ \ttback \ttback [lam_type, list_case_type, drop_type, map_type, \ttback \ttback apply_type, rec_type]" end \end{ttbox} \hrule \caption{Inductive definition of the primitive recursive functions} \label{primrec-fig} \end{figure} \def\fs{{\it fs}} Szasz was using ALF, but Coq and HOL would also have problems accepting this definition. Isabelle's package accepts it easily since $[f_0,\ldots,f_{m-1}]$ is a list of primitive recursive functions and $\lst$ is monotonic. There are five introduction rules, one for each of the five forms of primitive recursive function. Let us examine the one for $\COMP$: \[ \infer{\COMP(g,\fs)\in\primrec}{g\in\primrec & \fs\in\lst(\primrec)} \] The induction rule for $\primrec$ has one case for each introduction rule. Due to the use of $\lst$ as a monotone operator, the composition case has an unusual induction hypothesis: \[ \infer*{P(\COMP(g,\fs))} {[g\in\primrec & \fs\in\lst(\{z\in\primrec.P(z)\})]_{\fs,g}} \] The hypothesis states that $\fs$ is a list of primitive recursive functions satisfying the induction formula. Proving the $\COMP$ case typically requires structural induction on lists, yielding two subcases: either $\fs=\Nil$ or else $\fs=\Cons(f,\fs')$, where $f\in\primrec$, $P(f)$, and $\fs'$ is another list of primitive recursive functions satisfying~$P$. Figure~\ref{primrec-fig} presents the theory file. Theory {\tt Primrec} defines the constants $\SC$, $\CONST$, etc. These are not constructors of a new datatype, but functions over lists of numbers. Their definitions, most of which are omitted, consist of routine list programming. In Isabelle/ZF, the primitive recursive functions are defined as a subset of the function set $\lst(\nat)\to\nat$. The Isabelle theory goes on to formalize Ackermann's function and prove that it is not primitive recursive, using the induction rule {\tt Primrec.induct}. The proof follows Szasz's excellent account. \section{Datatypes and codatatypes}\label{data-sec} A (co)datatype definition is a (co)inductive definition with automatically defined constructors and a case analysis operator. The package proves that the case operator inverts the constructors and can prove freeness theorems involving any pair of constructors. \subsection{Constructors and their domain}\label{univ-sec} Conceptually, our two forms of definition are distinct. A (co)inductive definition selects a subset of an existing set; a (co)datatype definition creates a new set. But the package reduces the latter to the former. A set having strong closure properties must serve as the domain of the (co)inductive definition. Constructing this set requires some theoretical effort, which must be done anyway to show that (co)datatypes exist. It is not obvious that standard set theory is suitable for defining codatatypes. Isabelle/ZF defines the standard notion of Cartesian product $A\times B$, containing ordered pairs $\pair{a,b}$. Now the $m$-tuple $\pair{x_1,\ldots,x_m}$ is the empty set~$\emptyset$ if $m=0$, simply $x_1$ if $m=1$ and $\pair{x_1,\pair{x_2,\ldots,x_m}}$ if $m\geq2$. Isabelle/ZF also defines the disjoint sum $A+B$, containing injections $\Inl(a)\equiv\pair{0,a}$ and $\Inr(b)\equiv\pair{1,b}$. A datatype constructor $\Con(x_1,\ldots,x_m)$ is defined to be $h(\pair{x_1,\ldots,x_m})$, where $h$ is composed of $\Inl$ and~$\Inr$. In a mutually recursive definition, all constructors for the set~$R_i$ have the outer form~$h_{in}$, where $h_{in}$ is the injection described in~\S\ref{mutual-sec}. Further nested injections ensure that the constructors for~$R_i$ are pairwise distinct. Isabelle/ZF defines the set $\univ(A)$, which contains~$A$ and furthermore contains $\pair{a,b}$, $\Inl(a)$ and $\Inr(b)$ for $a$, $b\in\univ(A)$. In a typical datatype definition with set parameters $A_1$, \ldots, $A_k$, a suitable domain for all the recursive sets is $\univ(A_1\un\cdots\un A_k)$. This solves the problem for datatypes~\cite[\S4.2]{paulson-set-II}. The standard pairs and injections can only yield well-founded constructions. This eases the (manual!) definition of recursive functions over datatypes. But they are unsuitable for codatatypes, which typically contain non-well-founded objects. To support codatatypes, Isabelle/ZF defines a variant notion of ordered pair, written~$\pair{a;b}$. It also defines the corresponding variant notion of Cartesian product $A\otimes B$, variant injections $\QInl(a)$ and~$\QInr(b)$ and variant disjoint sum $A\oplus B$. Finally it defines the set $\quniv(A)$, which contains~$A$ and furthermore contains $\pair{a;b}$, $\QInl(a)$ and $\QInr(b)$ for $a$, $b\in\quniv(A)$. In a typical codatatype definition with set parameters $A_1$, \ldots, $A_k$, a suitable domain is $\quniv(A_1\un\cdots\un A_k)$. This approach using standard ZF set theory~\cite{paulson-final} is an alternative to adopting Aczel's Anti-Foundation Axiom~\cite{aczel88}. \subsection{The case analysis operator} The (co)datatype package automatically defines a case analysis operator, called {\tt$R$\_case}. A mutually recursive definition still has only one operator, whose name combines those of the recursive sets: it is called {\tt$R_1$\_\ldots\_$R_n$\_case}. The case operator is analogous to those for products and sums. Datatype definitions employ standard products and sums, whose operators are $\split$ and $\case$ and satisfy the equations \begin{eqnarray*} \split(f,\pair{x,y}) & = & f(x,y) \\ \case(f,g,\Inl(x)) & = & f(x) \\ \case(f,g,\Inr(y)) & = & g(y) \end{eqnarray*} Suppose the datatype has $k$ constructors $\Con_1$, \ldots,~$\Con_k$. Then its case operator takes $k+1$ arguments and satisfies an equation for each constructor: \begin{eqnarray*} R\hbox{\_case}(f_1,\ldots,f_k, {\tt Con}_i(\vec{x})) & = & f_i(\vec{x}), \qquad i = 1, \ldots, k \end{eqnarray*} The case operator's definition takes advantage of Isabelle's representation of syntax in the typed $\lambda$-calculus; it could readily be adapted to a theorem prover for higher-order logic. If $f$ and~$g$ have meta-type $i\To i$ then so do $\split(f)$ and $\case(f,g)$. This works because $\split$ and $\case$ operate on their last argument. They are easily combined to make complex case analysis operators. Here are two examples: \begin{itemize} \item $\split(\lambda x.\split(f(x)))$ performs case analysis for $A\times (B\times C)$, as is easily verified: \begin{eqnarray*} \split(\lambda x.\split(f(x)), \pair{a,b,c}) & = & (\lambda x.\split(f(x))(a,\pair{b,c}) \\ & = & \split(f(a), \pair{b,c}) \\ & = & f(a,b,c) \end{eqnarray*} \item $\case(f,\case(g,h))$ performs case analysis for $A+(B+C)$; let us verify one of the three equations: \begin{eqnarray*} \case(f,\case(g,h), \Inr(\Inl(b))) & = & \case(g,h,\Inl(b)) \\ & = & g(b) \end{eqnarray*} \end{itemize} Codatatype definitions are treated in precisely the same way. They express case operators using those for the variant products and sums, namely $\qsplit$ and~$\qcase$. \medskip \ifCADE The package has processed all the datatypes discussed in my earlier paper~\cite{paulson-set-II} and the codatatype of lazy lists. Space limitations preclude discussing these examples here, but they are distributed with Isabelle. \typeout{****Omitting datatype examples from CADE version!} \else To see how constructors and the case analysis operator are defined, let us examine some examples. These include lists and trees/forests, which I have discussed extensively in another paper~\cite{paulson-set-II}. \subsection{Example: lists and lazy lists}\label{lists-sec} Here is a theory file that declares the datatype of lists: \begin{ttbox} List = Univ + consts list :: "i=>i" datatype "list(A)" = Nil | Cons ("a:A", "l: list(A)") end \end{ttbox} And here is the theory file that declares the codatatype of lazy lists: \begin{ttbox} LList = QUniv + consts llist :: "i=>i" codatatype "llist(A)" = LNil | LCons ("a: A", "l: llist(A)") end \end{ttbox} They highlight the (many) similarities and (few) differences between datatype and codatatype definitions.\footnote{The real theory files contain many more declarations, mainly of functions over lists; the declaration of lazy lists is followed by the coinductive definition of lazy list equality.} Each form of list has two constructors, one for the empty list and one for adding an element to a list. Each takes a parameter, defining the set of lists over a given set~$A$. Each uses the appropriate domain from a Isabelle/ZF theory: \begin{itemize} \item $\lst(A)$ requires the parent theory {\tt Univ}. The package automatically uses the domain $\univ(A)$ (the default choice can be overridden). \item $\llist(A)$ requires the parent theory {\tt QUniv}. The package automatically uses the domain $\quniv(A)$. \end{itemize} Since $\lst(A)$ is a datatype, it enjoys a structural induction rule, {\tt List.induct}: \[ \infer{P(x)}{x\in\lst(A) & P(\Nil) & \infer*{P(\Cons(a,l))}{[a\in A & l\in\lst(A) & P(l)]_{a,l}} } \] Induction and freeness yield the law $l\not=\Cons(a,l)$. To strengthen this, Isabelle/ZF defines the rank of a set and proves that the standard pairs and injections have greater rank than their components. An immediate consequence, which justifies structural recursion on lists \cite[\S4.3]{paulson-set-II}, is \[ \rank(l) < \rank(\Cons(a,l)). \] Since $\llist(A)$ is a codatatype, it has no induction rule. Instead it has the coinduction rule shown in \S\ref{coind-sec}. Since variant pairs and injections are monotonic and need not have greater rank than their components, fixedpoint operators can create cyclic constructions. For example, the definition \begin{eqnarray*} \lconst(a) & \equiv & \lfp(\univ(a), \lambda l. \LCons(a,l)) \end{eqnarray*} yields $\lconst(a) = \LCons(a,\lconst(a))$. \medskip It may be instructive to examine the definitions of the constructors and case operator for $\lst(A)$. The definitions for $\llist(A)$ are similar. The list constructors are defined as follows: \begin{eqnarray*} \Nil & = & \Inl(\emptyset) \\ \Cons(a,l) & = & \Inr(\pair{a,l}) \end{eqnarray*} The operator $\lstcase$ performs case analysis on these two alternatives: \begin{eqnarray*} \lstcase(c,h) & \equiv & \case(\lambda u.c, \split(h)) \end{eqnarray*} Let us verify the two equations: \begin{eqnarray*} \lstcase(c, h, \Nil) & = & \case(\lambda u.c, \split(h), \Inl(\emptyset)) \\ & = & (\lambda u.c)(\emptyset) \\ & = & c\\[1ex] \lstcase(c, h, \Cons(x,y)) & = & \case(\lambda u.c, \split(h), \Inr(\pair{x,y})) \\ & = & \split(h, \pair{x,y}) \\ & = & h(x,y) \end{eqnarray*} \subsection{Example: mutual recursion} In mutually recursive trees and forests~\cite[\S4.5]{paulson-set-II}, trees have the one constructor $\Tcons$, while forests have the two constructors $\Fnil$ and~$\Fcons$: \begin{ttbox} TF = List + consts tree, forest, tree_forest :: "i=>i" datatype "tree(A)" = Tcons ("a: A", "f: forest(A)") and "forest(A)" = Fnil | Fcons ("t: tree(A)", "f: forest(A)") end \end{ttbox} The three introduction rules define the mutual recursion. The distinguishing feature of this example is its two induction rules. The basic induction rule is called {\tt TF.induct}: \[ \infer{P(x)}{x\in\TF(A) & \infer*{P(\Tcons(a,f))} {\left[\begin{array}{l} a\in A \\ f\in\forest(A) \\ P(f) \end{array} \right]_{a,f}} & P(\Fnil) & \infer*{P(\Fcons(t,f))} {\left[\begin{array}{l} t\in\tree(A) \\ P(t) \\ f\in\forest(A) \\ P(f) \end{array} \right]_{t,f}} } \] This rule establishes a single predicate for $\TF(A)$, the union of the recursive sets. Although such reasoning is sometimes useful \cite[\S4.5]{paulson-set-II}, a proper mutual induction rule should establish separate predicates for $\tree(A)$ and $\forest(A)$. The package calls this rule {\tt TF.mutual\_induct}. Observe the usage of $P$ and $Q$ in the induction hypotheses: \[ \infer{(\forall z. z\in\tree(A)\imp P(z)) \conj (\forall z. z\in\forest(A)\imp Q(z))} {\infer*{P(\Tcons(a,f))} {\left[\begin{array}{l} a\in A \\ f\in\forest(A) \\ Q(f) \end{array} \right]_{a,f}} & Q(\Fnil) & \infer*{Q(\Fcons(t,f))} {\left[\begin{array}{l} t\in\tree(A) \\ P(t) \\ f\in\forest(A) \\ Q(f) \end{array} \right]_{t,f}} } \] As mentioned above, the package does not define a structural recursion operator. I have described elsewhere how this is done \cite[\S4.5]{paulson-set-II}. Both forest constructors have the form $\Inr(\cdots)$, while the tree constructor has the form $\Inl(\cdots)$. This pattern would hold regardless of how many tree or forest constructors there were. \begin{eqnarray*} \Tcons(a,l) & = & \Inl(\pair{a,l}) \\ \Fnil & = & \Inr(\Inl(\emptyset)) \\ \Fcons(a,l) & = & \Inr(\Inr(\pair{a,l})) \end{eqnarray*} There is only one case operator; it works on the union of the trees and forests: \begin{eqnarray*} {\tt tree\_forest\_case}(f,c,g) & \equiv & \case(\split(f),\, \case(\lambda u.c, \split(g))) \end{eqnarray*} \subsection{A four-constructor datatype} Finally let us consider a fairly general datatype. It has four constructors $\Con_0$, \ldots, $\Con_3$, with the corresponding arities. \begin{ttbox} Data = Univ + consts data :: "[i,i] => i" datatype "data(A,B)" = Con0 | Con1 ("a: A") | Con2 ("a: A", "b: B") | Con3 ("a: A", "b: B", "d: data(A,B)") end \end{ttbox} Because this datatype has two set parameters, $A$ and~$B$, the package automatically supplies $\univ(A\un B)$ as its domain. The structural induction rule has four minor premises, one per constructor: \[ \infer{P(x)}{x\in\data(A,B) & P(\Con_0) & \infer*{P(\Con_1(a))}{[a\in A]_a} & \infer*{P(\Con_2(a,b))} {\left[\begin{array}{l} a\in A \\ b\in B \end{array} \right]_{a,b}} & \infer*{P(\Con_3(a,b,d))} {\left[\begin{array}{l} a\in A \\ b\in B \\ d\in\data(A,B) \\ P(d) \end{array} \right]_{a,b,d}} } \] The constructor definitions are \begin{eqnarray*} \Con_0 & = & \Inl(\Inl(\emptyset)) \\ \Con_1(a) & = & \Inl(\Inr(a)) \\ \Con_2(a,b) & = & \Inr(\Inl(\pair{a,b})) \\ \Con_3(a,b,c) & = & \Inr(\Inr(\pair{a,b,c})). \end{eqnarray*} The case operator is \begin{eqnarray*} {\tt data\_case}(f_0,f_1,f_2,f_3) & \equiv & \case(\begin{array}[t]{@{}l} \case(\lambda u.f_0,\; f_1),\, \\ \case(\split(f_2),\; \split(\lambda v.\split(f_3(v)))) ) \end{array} \end{eqnarray*} This may look cryptic, but the case equations are trivial to verify. In the constructor definitions, the injections are balanced. A more naive approach is to define $\Con_3(a,b,c)$ as $\Inr(\Inr(\Inr(\pair{a,b,c})))$; instead, each constructor has two injections. The difference here is small. But the ZF examples include a 60-element enumeration type, where each constructor has 5 or~6 injections. The naive approach would require 1 to~59 injections; the definitions would be quadratic in size. It is like the difference between the binary and unary numeral systems. The result structure contains the case operator and constructor definitions as the theorem list \verb|con_defs|. It contains the case equations, such as \begin{eqnarray*} {\tt data\_case}(f_0,f_1,f_2,f_3,\Con_3(a,b,c)) & = & f_3(a,b,c), \end{eqnarray*} as the theorem list \verb|case_eqns|. There is one equation per constructor. \subsection{Proving freeness theorems} There are two kinds of freeness theorems: \begin{itemize} \item {\bf injectiveness} theorems, such as \[ \Con_2(a,b) = \Con_2(a',b') \bimp a=a' \conj b=b' \] \item {\bf distinctness} theorems, such as \[ \Con_1(a) \not= \Con_2(a',b') \] \end{itemize} Since the number of such theorems is quadratic in the number of constructors, the package does not attempt to prove them all. Instead it returns tools for proving desired theorems --- either explicitly or `on the fly' during simplification or classical reasoning. The theorem list \verb|free_iffs| enables the simplifier to perform freeness reasoning. This works by incremental unfolding of constructors that appear in equations. The theorem list contains logical equivalences such as \begin{eqnarray*} \Con_0=c & \bimp & c=\Inl(\Inl(\emptyset)) \\ \Con_1(a)=c & \bimp & c=\Inl(\Inr(a)) \\ & \vdots & \\ \Inl(a)=\Inl(b) & \bimp & a=b \\ \Inl(a)=\Inr(b) & \bimp & {\tt False} \\ \pair{a,b} = \pair{a',b'} & \bimp & a=a' \conj b=b' \end{eqnarray*} For example, these rewrite $\Con_1(a)=\Con_1(b)$ to $a=b$ in four steps. The theorem list \verb|free_SEs| enables the classical reasoner to perform similar replacements. It consists of elimination rules to replace $\Con_0=c$ by $c=\Inl(\Inl(\emptyset))$ and so forth, in the assumptions. Such incremental unfolding combines freeness reasoning with other proof steps. It has the unfortunate side-effect of unfolding definitions of constructors in contexts such as $\exists x.\Con_1(a)=x$, where they should be left alone. Calling the Isabelle tactic {\tt fold\_tac con\_defs} restores the defined constants. \fi %CADE \section{Related work}\label{related} The use of least fixedpoints to express inductive definitions seems obvious. Why, then, has this technique so seldom been implemented? Most automated logics can only express inductive definitions by asserting new axioms. Little would be left of Boyer and Moore's logic~\cite{bm79} if their shell principle were removed. With ALF the situation is more complex; earlier versions of Martin-L\"of's type theory could (using wellordering types) express datatype definitions, but the version underlying ALF requires new rules for each definition~\cite{dybjer91}. With Coq the situation is subtler still; its underlying Calculus of Constructions can express inductive definitions~\cite{huet88}, but cannot quite handle datatype definitions~\cite{paulin92}. It seems that researchers tried hard to circumvent these problems before finally extending the Calculus with rule schemes for strictly positive operators. Higher-order logic can express inductive definitions through quantification over unary predicates. The following formula expresses that~$i$ belongs to the least set containing~0 and closed under~$\succ$: \[ \forall P. P(0)\conj (\forall x.P(x)\imp P(\succ(x))) \imp P(i) \] This technique can be used to prove the Knaster-Tarski Theorem, but it is little used in the Cambridge HOL system. Melham~\cite{melham89} clearly describes the development. The natural numbers are defined as shown above, but lists are defined as functions over the natural numbers. Unlabelled trees are defined using G\"odel numbering; a labelled tree consists of an unlabelled tree paired with a list of labels. Melham's datatype package expresses the user's datatypes in terms of labelled trees. It has been highly successful, but a fixedpoint approach might have yielded greater functionality with less effort. Melham's inductive definition package~\cite{camilleri92} uses quantification over predicates, which is implicitly a fixedpoint approach. Instead of formalizing the notion of monotone function, it requires definitions to consist of finitary rules, a syntactic form that excludes many monotone inductive definitions. The earliest use of least fixedpoints is probably Robin Milner's datatype package for Edinburgh LCF~\cite{milner-ind}. Brian Monahan extended this package considerably~\cite{monahan84}, as did I in unpublished work.\footnote{The datatype package described in my LCF book~\cite{paulson87} does {\it not\/} make definitions, but merely asserts axioms.} LCF is a first-order logic of domain theory; the relevant fixedpoint theorem is not Knaster-Tarski but concerns fixedpoints of continuous functions over domains. LCF is too weak to express recursive predicates. Thus it would appear that the Isabelle package is the first to be based on the Knaster-Tarski Theorem. \section{Conclusions and future work} Higher-order logic and set theory are both powerful enough to express inductive definitions. A growing number of theorem provers implement one of these~\cite{IMPS,saaltink-fme}. The easiest sort of inductive definition package to write is one that asserts new axioms, not one that makes definitions and proves theorems about them. But asserting axioms could introduce unsoundness. The fixedpoint approach makes it fairly easy to implement a package for (co)inductive definitions that does not assert axioms. It is efficient: it processes most definitions in seconds and even a 60-constructor datatype requires only two minutes. It is also simple: the package consists of under 1100 lines (35K bytes) of Standard ML code. The first working version took under a week to code. In set theory, care is required to ensure that the inductive definition yields a set (rather than a proper class). This problem is inherent to set theory, whether or not the Knaster-Tarski Theorem is employed. We must exhibit a bounding set (called a domain above). For inductive definitions, this is often trivial. For datatype definitions, I have had to formalize much set theory. To justify infinitely branching datatype definitions, I have had to develop a theory of cardinal arithmetic, such as the theorem that if $\kappa$ is an infinite cardinal and $|X(\alpha)| \le \kappa$ for all $\alpha<\kappa$ then $|\union\sb{\alpha<\kappa} X(\alpha)| \le \kappa$. The need for such efforts is not a drawback of the fixedpoint approach, for the alternative is to take such definitions on faith. The approach is not restricted to set theory. It should be suitable for any logic that has some notion of set and the Knaster-Tarski Theorem. I have ported the (co)inductive definition package from Isabelle/ZF to Isabelle/HOL (higher-order logic). I hope to port the (co)datatype package later. HOL represents sets by unary predicates; defining the corresponding types may cause complications. \bibliographystyle{springer} \bibliography{string-abbrv,atp,theory,funprog,isabelle} %%%%%\doendnotes \ifCADE\typeout{****Omitting appendices from CADE version!} \else \newpage \appendix \section{Inductive and coinductive definitions: users guide} A theory file may contain any number of inductive and coinductive definitions. They may be intermixed with other declarations; in particular, the (co)inductive sets {\bf must} be declared separately as constants, and may have mixfix syntax or be subject to syntax translations. Each (co)inductive definition adds definitions to the theory and also proves some theorems. Each definition creates an ML structure, which is a substructure of the main theory structure. \subsection{The result structure} Many of the result structure's components have been discussed in~\S\ref{basic-sec}; others are self-explanatory. \begin{description} \item[\tt thy] is the new theory containing the recursive sets. \item[\tt defs] is the list of definitions of the recursive sets. \item[\tt bnd\_mono] is a monotonicity theorem for the fixedpoint operator. \item[\tt unfold] is a fixedpoint equation for the recursive set (the union of the recursive sets, in the case of mutual recursion). \item[\tt dom\_subset] is a theorem stating inclusion in the domain. \item[\tt intrs] is the list of introduction rules, now proved as theorems, for the recursive sets. The rules are also available individually, using the names given them in the theory file. \item[\tt elim] is the elimination rule. \item[\tt mk\_cases] is a function to create simplified instances of {\tt elim}, using freeness reasoning on some underlying datatype. \end{description} For an inductive definition, the result structure contains two induction rules, {\tt induct} and \verb|mutual_induct|. For a coinductive definition, it contains the rule \verb|coinduct|. Figure~\ref{def-result-fig} summarizes the two result signatures, specifying the types of all these components. \begin{figure} \begin{ttbox} sig val thy : theory val defs : thm list val bnd_mono : thm val unfold : thm val dom_subset : thm val intrs : thm list val elim : thm val mk_cases : thm list -> string -> thm {\it(Inductive definitions only)} val induct : thm val mutual_induct: thm {\it(Coinductive definitions only)} val coinduct : thm end \end{ttbox} \hrule \caption{The result of a (co)inductive definition} \label{def-result-fig} \end{figure} \subsection{The syntax of a (co)inductive definition} An inductive definition has the form \begin{ttbox} inductive domains {\it domain declarations} intrs {\it introduction rules} monos {\it monotonicity theorems} con_defs {\it constructor definitions} type_intrs {\it introduction rules for type-checking} type_elims {\it elimination rules for type-checking} \end{ttbox} A coinductive definition is identical save that it starts with the keyword {\tt coinductive}. The {\tt monos}, {\tt con\_defs}, {\tt type\_intrs} and {\tt type\_elims} sections are optional. If present, each is specified as a string, which must be a valid ML expression of type {\tt thm list}. It is simply inserted into the {\tt .thy.ML} file; if it is ill-formed, it will trigger ML error messages. You can then inspect the file on your directory. \begin{description} \item[\it domain declarations] consist of one or more items of the form {\it string\/}~{\tt <=}~{\it string}, associating each recursive set with its domain. \item[\it introduction rules] specify one or more introduction rules in the form {\it ident\/}~{\it string}, where the identifier gives the name of the rule in the result structure. \item[\it monotonicity theorems] are required for each operator applied to a recursive set in the introduction rules. There {\bf must} be a theorem of the form $A\sbs B\Imp M(A)\sbs M(B)$, for each premise $t\in M(R_i)$ in an introduction rule! \item[\it constructor definitions] contain definitions of constants appearing in the introduction rules. The (co)datatype package supplies the constructors' definitions here. Most (co)inductive definitions omit this section; one exception is the primitive recursive functions example (\S\ref{primrec-sec}). \item[\it type\_intrs] consists of introduction rules for type-checking the definition, as discussed in~\S\ref{basic-sec}. They are applied using depth-first search; you can trace the proof by setting \verb|trace_DEPTH_FIRST := true|. \item[\it type\_elims] consists of elimination rules for type-checking the definition. They are presumed to be `safe' and are applied as much as possible, prior to the {\tt type\_intrs} search. \end{description} The package has a few notable restrictions: \begin{itemize} \item The theory must separately declare the recursive sets as constants. \item The names of the recursive sets must be identifiers, not infix operators. \item Side-conditions must not be conjunctions. However, an introduction rule may contain any number of side-conditions. \item Side-conditions of the form $x=t$, where the variable~$x$ does not occur in~$t$, will be substituted through the rule \verb|mutual_induct|. \end{itemize} \section{Datatype and codatatype definitions: users guide} This section explains how to include (co)datatype declarations in a theory file. \subsection{The result structure} The result structure extends that of (co)inductive definitions (Figure~\ref{def-result-fig}) with several additional items: \begin{ttbox} val con_defs : thm list val case_eqns : thm list val free_iffs : thm list val free_SEs : thm list val mk_free : string -> thm \end{ttbox} Most of these have been discussed in~\S\ref{data-sec}. Here is a summary: \begin{description} \item[\tt con\_defs] is a list of definitions: the case operator followed by the constructors. This theorem list can be supplied to \verb|mk_cases|, for example. \item[\tt case\_eqns] is a list of equations, stating that the case operator inverts each constructor. \item[\tt free\_iffs] is a list of logical equivalences to perform freeness reasoning by rewriting. A typical application has the form \begin{ttbox} by (asm_simp_tac (ZF_ss addsimps free_iffs) 1); \end{ttbox} \item[\tt free\_SEs] is a list of `safe' elimination rules to perform freeness reasoning. It can be supplied to \verb|eresolve_tac| or to the classical reasoner: \begin{ttbox} by (fast_tac (ZF_cs addSEs free_SEs) 1); \end{ttbox} \item[\tt mk\_free] is a function to prove freeness properties, specified as strings. The theorems can be expressed in various forms, such as logical equivalences or elimination rules. \end{description} The result structure also inherits everything from the underlying (co)inductive definition, such as the introduction rules, elimination rule, and (co)induction rule. \subsection{The syntax of a (co)datatype definition} A datatype definition has the form \begin{ttbox} datatype <={\it domain} {\it datatype declaration} and {\it datatype declaration} and \ldots monos {\it monotonicity theorems} type_intrs {\it introduction rules for type-checking} type_elims {\it elimination rules for type-checking} \end{ttbox} A codatatype definition is identical save that it starts with the keyword {\tt codatatype}. The syntax is rather complicated; please consult the examples above (\S\ref{lists-sec}) and the theory files on the ZF source directory. The {\tt monos}, {\tt type\_intrs} and {\tt type\_elims} sections are optional. They are treated like their counterparts in a (co)inductive definition, as described above. The package supplements your type-checking rules (if any) with additional ones that should cope with any finitely-branching (co)datatype definition. \begin{description} \item[\it domain] specifies a single domain to use for all the mutually recursive (co)datatypes. If it (and the preceeding~{\tt <=}) are omitted, the package supplies a domain automatically. Suppose the definition involves the set parameters $A_1$, \ldots, $A_k$. Then $\univ(A_1\un\cdots\un A_k)$ is used for a datatype definition and $\quniv(A_1\un\cdots\un A_k)$ is used for a codatatype definition. These choices should work for all finitely-branching (co)datatype definitions. For examples of infinitely-branching datatype definitions, see the file {\tt ZF/ex/Brouwer.thy}. \item[\it datatype declaration] has the form \begin{quote} {\it string\/} {\tt =} {\it constructor} {\tt|} {\it constructor} {\tt|} \ldots \end{quote} The {\it string\/} is the datatype, say {\tt"list(A)"}. Each {\it constructor\/} has the form \begin{quote} {\it name\/} {\tt(} {\it premise} {\tt,} {\it premise} {\tt,} \ldots {\tt)} {\it mixfix\/} \end{quote} The {\it name\/} specifies a new constructor while the {\it premises\/} its typing conditions. The optional {\it mixfix\/} phrase may give the constructor infix, for example. Mutually recursive {\it datatype declarations\/} are separated by the keyword~{\tt and}. \end{description} \paragraph*{Note.} In the definitions of the constructors, the right-hand sides may overlap. For instance, the datatype of combinators has constructors defined by \begin{eqnarray*} {\tt K} & \equiv & \Inl(\emptyset) \\ {\tt S} & \equiv & \Inr(\Inl(\emptyset)) \\ p{\tt\#}q & \equiv & \Inr(\Inl(\pair{p,q})) \end{eqnarray*} Unlike in previous versions of Isabelle, \verb|fold_tac| now ensures that the longest right-hand sides are folded first. \fi \end{document}