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1 
\chapter{Functional Programming in HOL}


2 


3 
Although on the surface this chapter is mainly concerned with how to write


4 
functional programs in HOL and how to verify them, most of the


5 
constructs and proof procedures introduced are general purpose and recur in


6 
any specification or verification task.


7 


8 
The dedicated functional programmer should be warned: HOL offers only what


9 
could be called {\em total functional programming}  all functions in HOL


10 
must be total; lazy data structures are not directly available. On the


11 
positive side, functions in HOL need not be computable: HOL is a


12 
specification language that goes well beyond what can be expressed as a


13 
program. However, for the time being we concentrate on the computable.


14 


15 
\section{An introductory theory}


16 
\label{sec:introtheory}


17 


18 
Functional programming needs datatypes and functions. Both of them can be


19 
defined in a theory with a syntax reminiscent of languages like ML or


20 
Haskell. As an example consider the theory in Fig.~\ref{fig:ToyList}.


21 


22 
\begin{figure}[htbp]


23 
\begin{ttbox}\makeatother


24 
\input{ToyList/ToyList.thy}\end{ttbox}


25 
\caption{A theory of lists}


26 
\label{fig:ToyList}


27 
\end{figure}


28 


29 
HOL already has a predefined theory of lists called \texttt{List} 


30 
\texttt{ToyList} is merely a small fragment of it chosen as an example. In


31 
contrast to what is recommended in \S\ref{sec:Basic:Theories},


32 
\texttt{ToyList} is not based on \texttt{Main} but on \texttt{Datatype}, a


33 
theory that contains everything required for datatype definitions but does


34 
not have \texttt{List} as a parent, thus avoiding ambiguities caused by


35 
defining lists twice.


36 


37 
The \ttindexbold{datatype} \texttt{list} introduces two constructors


38 
\texttt{Nil} and \texttt{Cons}, the empty list and the operator that adds an


39 
element to the front of a list. For example, the term \texttt{Cons True (Cons


40 
False Nil)} is a value of type \texttt{bool~list}, namely the list with the


41 
elements \texttt{True} and \texttt{False}. Because this notation becomes


42 
unwieldy very quickly, the datatype declaration is annotated with an


43 
alternative syntax: instead of \texttt{Nil} and \texttt{Cons}~$x$~$xs$ we can


44 
write \index{#@{\tt[]}bold}\texttt{[]} and


45 
\texttt{$x$~\#~$xs$}\index{#@{\tt\#}bold}. In fact, this alternative syntax


46 
is the standard syntax. Thus the list \texttt{Cons True (Cons False Nil)}


47 
becomes \texttt{True \# False \# []}. The annotation \ttindexbold{infixr}


48 
means that \texttt{\#} associates to the right, i.e.\ the term \texttt{$x$ \#


49 
$y$ \# $z$} is read as \texttt{$x$ \# ($y$ \# $z$)} and not as \texttt{($x$


50 
\# $y$) \# $z$}.


51 


52 
\begin{warn}


53 
Syntax annotations are a powerful but completely optional feature. You


54 
could drop them from theory \texttt{ToyList} and go back to the identifiers


55 
\texttt{Nil} and \texttt{Cons}. However, lists are such a central datatype


56 
that their syntax is highly customized. We recommend that novices should


57 
not use syntax annotations in their own theories.


58 
\end{warn}


59 


60 
Next, the functions \texttt{app} and \texttt{rev} are declared. In contrast


61 
to ML, Isabelle insists on explicit declarations of all functions (keyword


62 
\ttindexbold{consts}). (Apart from the declarationbeforeuse restriction,


63 
the order of items in a theory file is unconstrained.) Function \texttt{app}


64 
is annotated with concrete syntax too. Instead of the prefix syntax


65 
\texttt{app}~$xs$~$ys$ the infix $xs$~\texttt{\at}~$ys$ becomes the preferred


66 
form.


67 


68 
Both functions are defined recursively. The equations for \texttt{app} and


69 
\texttt{rev} hardly need comments: \texttt{app} appends two lists and

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\texttt{rev} reverses a list. The keyword \ttindex{primrec} indicates that

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the recursion is of a particularly primitive kind where each recursive call


72 
peels off a datatype constructor from one of the arguments (see


73 
\S\ref{sec:datatype}). Thus the recursion always terminates, i.e.\ the


74 
function is \bfindex{total}.


75 


76 
The termination requirement is absolutely essential in HOL, a logic of total


77 
functions. If we were to drop it, inconsistencies could quickly arise: the


78 
``definition'' $f(n) = f(n)+1$ immediately leads to $0 = 1$ by subtracting


79 
$f(n)$ on both sides.


80 
% However, this is a subtle issue that we cannot discuss here further.


81 


82 
\begin{warn}


83 
As we have indicated, the desire for total functions is not a gratuitously


84 
imposed restriction but an essential characteristic of HOL. It is only


85 
because of totality that reasoning in HOL is comparatively easy. More


86 
generally, the philosophy in HOL is not to allow arbitrary axioms (such as


87 
function definitions whose totality has not been proved) because they


88 
quickly lead to inconsistencies. Instead, fixed constructs for introducing


89 
types and functions are offered (such as \texttt{datatype} and


90 
\texttt{primrec}) which are guaranteed to preserve consistency.


91 
\end{warn}


92 


93 
A remark about syntax. The textual definition of a theory follows a fixed


94 
syntax with keywords like \texttt{datatype} and \texttt{end} (see


95 
Fig.~\ref{fig:keywords} in Appendix~\ref{sec:Appendix} for a full list).


96 
Embedded in this syntax are the types and formulae of HOL, whose syntax is


97 
extensible, e.g.\ by new userdefined infix operators


98 
(see~\ref{sec:infixsyntax}). To distinguish the two levels, everything


99 
HOLspecific should be enclosed in \texttt{"}\dots\texttt{"}. The same holds


100 
for identifiers that happen to be keywords, as in


101 
\begin{ttbox}


102 
consts "end" :: 'a list => 'a


103 
\end{ttbox}


104 
To lessen this burden, quotation marks around types can be dropped,


105 
provided their syntax does not go beyond what is described in


106 
\S\ref{sec:TypesTermsForms}. Types containing further operators, e.g.\

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\label{startype} \texttt{*} for Cartesian products, need quotation marks.

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108 


109 
When Isabelle prints a syntax error message, it refers to the HOL syntax as


110 
the \bfindex{inner syntax}.


111 


112 
\section{An introductory proof}


113 
\label{sec:introproof}


114 


115 
Having defined \texttt{ToyList}, we load it with the ML command


116 
\begin{ttbox}


117 
use_thy "ToyList";


118 
\end{ttbox}


119 
and are ready to prove a few simple theorems. This will illustrate not just


120 
the basic proof commands but also the typical proof process.


121 


122 
\subsubsection*{Main goal: \texttt{rev(rev xs) = xs}}


123 


124 
Our goal is to show that reversing a list twice produces the original


125 
list. Typing


126 
\begin{ttbox}


127 
\input{ToyList/thm}\end{ttbox}


128 
establishes a new goal to be proved in the context of the current theory,


129 
which is the one we just loaded. Isabelle's response is to print the current proof state:


130 
\begin{ttbox}


131 
{\out Level 0}


132 
{\out rev (rev xs) = xs}


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{\out 1. rev (rev xs) = xs}


134 
\end{ttbox}


135 
Until we have finished a proof, the proof state always looks like this:


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\begin{ttbox}


137 
{\out Level \(i\)}


138 
{\out \(G\)}


139 
{\out 1. \(G@1\)}


140 
{\out \(\vdots\)}


141 
{\out \(n\). \(G@n\)}


142 
\end{ttbox}


143 
where \texttt{Level}~$i$ indicates that we are $i$ steps into the proof, $G$


144 
is the overall goal that we are trying to prove, and the numbered lines


145 
contain the subgoals $G@1$, \dots, $G@n$ that we need to prove to establish


146 
$G$. At \texttt{Level 0} there is only one subgoal, which is identical with


147 
the overall goal. Normally $G$ is constant and only serves as a reminder.


148 
Hence we rarely show it in this tutorial.


149 


150 
Let us now get back to \texttt{rev(rev xs) = xs}. Properties of recursively


151 
defined functions are best established by induction. In this case there is


152 
not much choice except to induct on \texttt{xs}:


153 
\begin{ttbox}


154 
\input{ToyList/inductxs}\end{ttbox}


155 
This tells Isabelle to perform induction on variable \texttt{xs} in subgoal


156 
1. The new proof state contains two subgoals, namely the base case


157 
(\texttt{Nil}) and the induction step (\texttt{Cons}):


158 
\begin{ttbox}


159 
{\out 1. rev (rev []) = []}


160 
{\out 2. !!a list. rev (rev list) = list ==> rev (rev (a # list)) = a # list}


161 
\end{ttbox}


162 
The induction step is an example of the general format of a subgoal:


163 
\begin{ttbox}


164 
{\out \(i\). !!\(x@1 \dots x@n\). {\it assumptions} ==> {\it conclusion}}


165 
\end{ttbox}\index{==>@{\tt==>}bold}


166 
The prefix of bound variables \texttt{!!\(x@1 \dots x@n\)} can be ignored


167 
most of the time, or simply treated as a list of variables local to this


168 
subgoal. Their deeper significance is explained in \S\ref{sec:PCproofs}. The


169 
{\it assumptions} are the local assumptions for this subgoal and {\it


170 
conclusion} is the actual proposition to be proved. Typical proof steps


171 
that add new assumptions are induction or case distinction. In our example


172 
the only assumption is the induction hypothesis \texttt{rev (rev list) =


173 
list}, where \texttt{list} is a variable name chosen by Isabelle. If there


174 
are multiple assumptions, they are enclosed in the bracket pair


175 
\texttt{[}\index{==>@\ttlbrbold} and \texttt{]}\index{==>@\ttrbrbold}


176 
and separated by semicolons.


177 


178 
Let us try to solve both goals automatically:


179 
\begin{ttbox}


180 
\input{ToyList/autotac}\end{ttbox}


181 
This command tells Isabelle to apply a proof strategy called


182 
\texttt{Auto_tac} to all subgoals. Essentially, \texttt{Auto_tac} tries to


183 
`simplify' the subgoals. In our case, subgoal~1 is solved completely (thanks


184 
to the equation \texttt{rev [] = []}) and disappears; the simplified version


185 
of subgoal~2 becomes the new subgoal~1:


186 
\begin{ttbox}\makeatother


187 
{\out 1. !!a list. rev(rev list) = list ==> rev(rev list @ a # []) = a # list}


188 
\end{ttbox}


189 
In order to simplify this subgoal further, a lemma suggests itself.


190 


191 
\subsubsection*{First lemma: \texttt{rev(xs \at~ys) = (rev ys) \at~(rev xs)}}


192 


193 
We start the proof as usual:


194 
\begin{ttbox}\makeatother


195 
\input{ToyList/lemma1}\end{ttbox}


196 
There are two variables that we could induct on: \texttt{xs} and


197 
\texttt{ys}. Because \texttt{\at} is defined by recursion on


198 
the first argument, \texttt{xs} is the correct one:


199 
\begin{ttbox}


200 
\input{ToyList/inductxs}\end{ttbox}


201 
This time not even the base case is solved automatically:


202 
\begin{ttbox}\makeatother


203 
by(Auto_tac);


204 
{\out 1. rev ys = rev ys @ []}


205 
{\out 2. \dots}


206 
\end{ttbox}


207 
We need another lemma.


208 


209 
\subsubsection*{Second lemma: \texttt{xs \at~[] = xs}}


210 


211 
This time the canonical proof procedure


212 
\begin{ttbox}\makeatother


213 
\input{ToyList/lemma2}\input{ToyList/inductxs}\input{ToyList/autotac}\end{ttbox}


214 
leads to the desired message \texttt{No subgoals!}:


215 
\begin{ttbox}\makeatother


216 
{\out Level 2}


217 
{\out xs @ [] = xs}


218 
{\out No subgoals!}


219 
\end{ttbox}


220 
Now we can give the lemma just proved a suitable name


221 
\begin{ttbox}

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\input{ToyList/qed2}\end{ttbox}\index{*qed}

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and tell Isabelle to use this lemma in all future proofs by simplification:


224 
\begin{ttbox}


225 
\input{ToyList/addsimps2}\end{ttbox}


226 
Note that in the theorem \texttt{app_Nil2} the free variable \texttt{xs} has


227 
been replaced by the unknown \texttt{?xs}, just as explained in


228 
\S\ref{sec:variables}.


229 


230 
Going back to the proof of the first lemma


231 
\begin{ttbox}\makeatother


232 
\input{ToyList/lemma1}\input{ToyList/inductxs}\input{ToyList/autotac}\end{ttbox}


233 
we find that this time \texttt{Auto_tac} solves the base case, but the


234 
induction step merely simplifies to


235 
\begin{ttbox}\makeatother


236 
{\out 1. !!a list.}


237 
{\out rev (list @ ys) = rev ys @ rev list}


238 
{\out ==> (rev ys @ rev list) @ a # [] = rev ys @ rev list @ a # []}


239 
\end{ttbox}


240 
Now we need to remember that \texttt{\at} associates to the right, and that


241 
\texttt{\#} and \texttt{\at} have the same priority (namely the \texttt{65}


242 
in the definition of \texttt{ToyList}). Thus the conclusion really is


243 
\begin{ttbox}\makeatother


244 
{\out ==> (rev ys @ rev list) @ (a # []) = rev ys @ (rev list @ (a # []))}


245 
\end{ttbox}


246 
and the missing lemma is associativity of \texttt{\at}.


247 


248 
\subsubsection*{Third lemma: \texttt{(xs \at~ys) \at~zs = xs \at~(ys \at~zs)}}


249 


250 
This time the canonical proof procedure


251 
\begin{ttbox}\makeatother


252 
\input{ToyList/lemma3}\end{ttbox}


253 
succeeds without further ado. Again we name the lemma and add it to


254 
the set of lemmas used during simplification:


255 
\begin{ttbox}


256 
\input{ToyList/qed3}\end{ttbox}


257 
Now we can go back and prove the first lemma


258 
\begin{ttbox}\makeatother


259 
\input{ToyList/lemma1}\input{ToyList/inductxs}\input{ToyList/autotac}\end{ttbox}


260 
add it to the simplification lemmas


261 
\begin{ttbox}


262 
\input{ToyList/qed1}\end{ttbox}


263 
and then solve our main theorem:


264 
\begin{ttbox}\makeatother


265 
\input{ToyList/thm}\input{ToyList/inductxs}\input{ToyList/autotac}\end{ttbox}


266 


267 
\subsubsection*{Review}


268 


269 
This is the end of our toy proof. It should have familiarized you with


270 
\begin{itemize}


271 
\item the standard theorem proving procedure:


272 
state a goal; proceed with proof until a new lemma is required; prove that


273 
lemma; come back to the original goal.


274 
\item a specific procedure that works well for functional programs:


275 
induction followed by allout simplification via \texttt{Auto_tac}.


276 
\item a basic repertoire of proof commands.


277 
\end{itemize}


278 


279 


280 
\section{Some helpful commands}


281 
\label{sec:commandsandhints}


282 


283 
This section discusses a few basic commands for manipulating the proof state


284 
and can be skipped by casual readers.


285 


286 
There are two kinds of commands used during a proof: the actual proof


287 
commands and auxiliary commands for examining the proof state and controlling


288 
the display. Proof commands are always of the form


289 
\texttt{by(\textit{tactic});}\indexbold{tactic} where \textbf{tactic} is a


290 
synonym for ``theorem proving function''. Typical examples are


291 
\texttt{induct_tac} and \texttt{Auto_tac}  the suffix \texttt{_tac} is


292 
merely a mnemonic. Further tactics are introduced throughout the tutorial.


293 


294 
%Tactics can also be modified. For example,


295 
%\begin{ttbox}


296 
%by(ALLGOALS Asm_simp_tac);


297 
%\end{ttbox}


298 
%tells Isabelle to apply \texttt{Asm_simp_tac} to all subgoals. For more on


299 
%tactics and how to combine them see~\S\ref{sec:Tactics}.


300 


301 
The most useful auxiliary commands are:


302 
\begin{description}


303 
\item[Printing the current state]


304 
Type \texttt{pr();} to redisplay the current proof state, for example when it


305 
has disappeared off the screen.


306 
\item[Limiting the number of subgoals]


307 
Typing \texttt{prlim $k$;} tells Isabelle to print only the first $k$


308 
subgoals from now on and redisplays the current proof state. This is helpful


309 
when there are many subgoals.


310 
\item[Undoing] Typing \texttt{undo();} undoes the effect of the last


311 
tactic.


312 
\item[Context switch] Every proof happens in the context of a


313 
\bfindex{current theory}. By default, this is the last theory loaded. If


314 
you want to prove a theorem in the context of a different theory


315 
\texttt{T}, you need to type \texttt{context T.thy;}\index{*contextbold}


316 
first. Of course you need to change the context again if you want to go


317 
back to your original theory.


318 
\item[Displaying types] We have already mentioned the flag


319 
\ttindex{show_types} above. It can also be useful for detecting typos in


320 
formulae early on. For example, if \texttt{show_types} is set and the goal


321 
\texttt{rev(rev xs) = xs} is started, Isabelle prints the additional output


322 
\begin{ttbox}


323 
{\out Variables:}


324 
{\out xs :: 'a list}


325 
\end{ttbox}


326 
which tells us that Isabelle has correctly inferred that


327 
\texttt{xs} is a variable of list type. On the other hand, had we


328 
made a typo as in \texttt{rev(re xs) = xs}, the response


329 
\begin{ttbox}


330 
Variables:


331 
re :: 'a list => 'a list


332 
xs :: 'a list


333 
\end{ttbox}


334 
would have alerted us because of the unexpected variable \texttt{re}.


335 
\item[(Re)loading theories]\index{loading theories}\index{reloading theories}


336 
Initially you load theory \texttt{T} by typing \ttindex{use_thy}~\texttt{"T";},


337 
which loads all parent theories of \texttt{T} automatically, if they are not


338 
loaded already. If you modify \texttt{T.thy} or \texttt{T.ML}, you can


339 
reload it by typing \texttt{use_thy~"T";} again. This time, however, only


340 
\texttt{T} is reloaded. If some of \texttt{T}'s parents have changed as well,

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type \ttindexbold{update_thy}~\texttt{"T";} to reload \texttt{T} and all of


342 
its parents that have changed (or have changed parents).

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\end{description}


344 
Further commands are found in the Reference Manual.


345 


346 


347 
\section{Datatypes}


348 
\label{sec:datatype}


349 


350 
Inductive datatypes are part of almost every nontrivial application of HOL.


351 
First we take another look at a very important example, the datatype of


352 
lists, before we turn to datatypes in general. The section closes with a


353 
case study.


354 


355 


356 
\subsection{Lists}


357 

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Lists are one of the essential datatypes in computing. Readers of this


359 
tutorial and users of HOL need to be familiar with their basic operations.


360 
Theory \texttt{ToyList} is only a small fragment of HOL's predefined theory

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\texttt{List}\footnote{\url{http://isabelle.in.tum.de/library/HOL/List.html}}.

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The latter contains many further operations. For example, the functions


363 
\ttindexbold{hd} (`head') and \ttindexbold{tl} (`tail') return the first


364 
element and the remainder of a list. (However, patternmatching is usually

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preferable to \texttt{hd} and \texttt{tl}.) Theory \texttt{List} also


366 
contains more syntactic sugar:

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\texttt{[}$x@1$\texttt{,}\dots\texttt{,}$x@n$\texttt{]} abbreviates

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$x@1$\texttt{\#}\dots\texttt{\#}$x@n$\texttt{\#[]}. In the rest of the


369 
tutorial we always use HOL's predefined lists.

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371 


372 
\subsection{The general format}


373 
\label{sec:generaldatatype}


374 


375 
The general HOL \texttt{datatype} definition is of the form


376 
\[


377 
\mathtt{datatype}~(\alpha@1, \dots, \alpha@n) \, t ~=~


378 
C@1~\tau@{11}~\dots~\tau@{1k@1} ~\mid~ \dots ~\mid~


379 
C@m~\tau@{m1}~\dots~\tau@{mk@m}


380 
\]


381 
where $\alpha@i$ are type variables (the parameters), $C@i$ are distinct


382 
constructor names and $\tau@{ij}$ are types; it is customary to capitalize


383 
the first letter in constructor names. There are a number of


384 
restrictions (such as the type should not be empty) detailed

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elsewhere~\cite{isabelleHOL}. Isabelle notifies you if you violate them.

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387 
Laws about datatypes, such as \verb$[] ~= x#xs$ and \texttt{(x\#xs = y\#ys) =


388 
(x=y \& xs=ys)}, are used automatically during proofs by simplification.


389 
The same is true for the equations in primitive recursive function


390 
definitions.


391 


392 
\subsection{Primitive recursion}


393 


394 
Functions on datatypes are usually defined by recursion. In fact, most of the


395 
time they are defined by what is called \bfindex{primitive recursion}.

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The keyword \ttindexbold{primrec} is followed by a list of equations

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\[ f \, x@1 \, \dots \, (C \, y@1 \, \dots \, y@k)\, \dots \, x@n = r \]


398 
such that $C$ is a constructor of the datatype $t$ and all recursive calls of


399 
$f$ in $r$ are of the form $f \, \dots \, y@i \, \dots$ for some $i$. Thus


400 
Isabelle immediately sees that $f$ terminates because one (fixed!) argument


401 
becomes smaller with every recursive call. There must be exactly one equation


402 
for each constructor. Their order is immaterial.


403 
A more general method for defining total recursive functions is explained in


404 
\S\ref{sec:recdef}.


405 


406 
\begin{exercise}


407 
Given the datatype of binary trees


408 
\begin{ttbox}


409 
\input{Misc/tree}\end{ttbox}


410 
define a function \texttt{mirror} that mirrors the structure of a binary tree


411 
by swapping subtrees (recursively). Prove \texttt{mirror(mirror(t)) = t}.


412 
\end{exercise}


413 


414 
\subsection{\texttt{case}expressions}


415 
\label{sec:caseexpressions}


416 


417 
HOL also features \ttindexbold{case}expressions for analyzing elements of a


418 
datatype. For example,


419 
\begin{ttbox}


420 
case xs of [] => 0  y#ys => y


421 
\end{ttbox}


422 
evaluates to \texttt{0} if \texttt{xs} is \texttt{[]} and to \texttt{y} if


423 
\texttt{xs} is \texttt{y\#ys}. (Since the result in both branches must be of


424 
the same type, it follows that \texttt{y::nat} and hence


425 
\texttt{xs::(nat)list}.)


426 


427 
In general, if $e$ is a term of the datatype $t$ defined in


428 
\S\ref{sec:generaldatatype} above, the corresponding


429 
\texttt{case}expression analyzing $e$ is


430 
\[


431 
\begin{array}{rrcl}


432 
\mbox{\tt case}~e~\mbox{\tt of} & C@1~x@{11}~\dots~x@{1k@1} & \To & e@1 \\


433 
\vdots \\


434 
\mid & C@m~x@{m1}~\dots~x@{mk@m} & \To & e@m


435 
\end{array}


436 
\]


437 


438 
\begin{warn}


439 
{\em All} constructors must be present, their order is fixed, and nested


440 
patterns are not supported. Violating these restrictions results in strange


441 
error messages.


442 
\end{warn}


443 
\noindent


444 
Nested patterns can be simulated by nested \texttt{case}expressions: instead


445 
of


446 
\begin{ttbox}


447 
case xs of [] => 0  [x] => x  x#(y#zs) => y


448 
\end{ttbox}


449 
write


450 
\begin{ttbox}


451 
case xs of [] => 0  x#ys => (case ys of [] => x  y#zs => y)


452 
\end{ttbox}


453 
Note that \texttt{case}expressions should be enclosed in parentheses to


454 
indicate their scope.


455 


456 
\subsection{Structural induction}


457 


458 
Almost all the basic laws about a datatype are applied automatically during


459 
simplification. Only induction is invoked by hand via \texttt{induct_tac},


460 
which works for any datatype. In some cases, induction is overkill and a case


461 
distinction over all constructors of the datatype suffices. This is performed


462 
by \ttindexbold{exhaust_tac}. A trivial example:


463 
\begin{ttbox}


464 
\input{Misc/exhaust.ML}{\out1. xs = [] ==> (case xs of [] => []  y # ys => xs) = xs}


465 
{\out2. !!a list. xs = a # list ==> (case xs of [] => []  y # ys => xs) = xs}


466 
\input{Misc/autotac.ML}\end{ttbox}


467 
Note that this particular case distinction could have been automated


468 
completely. See~\S\ref{sec:SimpFeatures}.


469 


470 
\begin{warn}


471 
Induction is only allowed on a free variable that should not occur among


472 
the assumptions of the subgoal. Exhaustion works for arbitrary terms.


473 
\end{warn}


474 


475 
\subsection{Case study: boolean expressions}


476 
\label{sec:boolex}


477 


478 
The aim of this case study is twofold: it shows how to model boolean


479 
expressions and some algorithms for manipulating them, and it demonstrates


480 
the constructs introduced above.


481 


482 
\subsubsection{How can we model boolean expressions?}


483 


484 
We want to represent boolean expressions built up from variables and


485 
constants by negation and conjunction. The following datatype serves exactly


486 
that purpose:


487 
\begin{ttbox}


488 
\input{Ifexpr/boolex}\end{ttbox}


489 
The two constants are represented by the terms \texttt{Const~True} and


490 
\texttt{Const~False}. Variables are represented by terms of the form


491 
\texttt{Var}~$n$, where $n$ is a natural number (type \texttt{nat}).


492 
For example, the formula $P@0 \land \neg P@1$ is represented by the term


493 
\texttt{And~(Var~0)~(Neg(Var~1))}.


494 


495 
\subsubsection{What is the value of boolean expressions?}


496 


497 
The value of a boolean expressions depends on the value of its variables.


498 
Hence the function \texttt{value} takes an additional parameter, an {\em


499 
environment} of type \texttt{nat~=>~bool}, which maps variables to their


500 
values:


501 
\begin{ttbox}


502 
\input{Ifexpr/value}\end{ttbox}


503 


504 
\subsubsection{Ifexpressions}


505 


506 
An alternative and often more efficient (because in a certain sense


507 
canonical) representation are socalled \textit{Ifexpressions\/} built up


508 
from constants (\texttt{CIF}), variables (\texttt{VIF}) and conditionals


509 
(\texttt{IF}):


510 
\begin{ttbox}


511 
\input{Ifexpr/ifex}\end{ttbox}


512 
The evaluation if Ifexpressions proceeds as for \texttt{boolex}:


513 
\begin{ttbox}


514 
\input{Ifexpr/valif}\end{ttbox}


515 


516 
\subsubsection{Transformation into and of Ifexpressions}


517 


518 
The type \texttt{boolex} is close to the customary representation of logical


519 
formulae, whereas \texttt{ifex} is designed for efficiency. Thus we need to


520 
translate from \texttt{boolex} into \texttt{ifex}:


521 
\begin{ttbox}


522 
\input{Ifexpr/bool2if}\end{ttbox}


523 
At last, we have something we can verify: that \texttt{bool2if} preserves the


524 
value of its argument.


525 
\begin{ttbox}


526 
\input{Ifexpr/bool2if.ML}\end{ttbox}


527 
The proof is canonical:


528 
\begin{ttbox}


529 
\input{Ifexpr/proof.ML}\end{ttbox}


530 
In fact, all proofs in this case study look exactly like this. Hence we do


531 
not show them below.


532 


533 
More interesting is the transformation of Ifexpressions into a normal form


534 
where the first argument of \texttt{IF} cannot be another \texttt{IF} but


535 
must be a constant or variable. Such a normal form can be computed by


536 
repeatedly replacing a subterm of the form \texttt{IF~(IF~b~x~y)~z~u} by


537 
\texttt{IF b (IF x z u) (IF y z u)}, which has the same value. The following


538 
primitive recursive functions perform this task:


539 
\begin{ttbox}


540 
\input{Ifexpr/normif}


541 
\input{Ifexpr/norm}\end{ttbox}


542 
Their interplay is a bit tricky, and we leave it to the reader to develop an


543 
intuitive understanding. Fortunately, Isabelle can help us to verify that the


544 
transformation preserves the value of the expression:


545 
\begin{ttbox}


546 
\input{Ifexpr/norm.ML}\end{ttbox}


547 
The proof is canonical, provided we first show the following lemma (which


548 
also helps to understand what \texttt{normif} does) and make it available


549 
for simplification via \texttt{Addsimps}:


550 
\begin{ttbox}


551 
\input{Ifexpr/normif.ML}\end{ttbox}


552 


553 
But how can we be sure that \texttt{norm} really produces a normal form in


554 
the above sense? We have to prove


555 
\begin{ttbox}


556 
\input{Ifexpr/normal_norm.ML}\end{ttbox}


557 
where \texttt{normal} expresses that an Ifexpression is in normal form:


558 
\begin{ttbox}


559 
\input{Ifexpr/normal}\end{ttbox}


560 
Of course, this requires a lemma about normality of \texttt{normif}


561 
\begin{ttbox}


562 
\input{Ifexpr/normal_normif.ML}\end{ttbox}


563 
that has to be made available for simplification via \texttt{Addsimps}.


564 


565 
How does one come up with the required lemmas? Try to prove the main theorems


566 
without them and study carefully what \texttt{Auto_tac} leaves unproved. This


567 
has to provide the clue.


568 
The necessity of universal quantification (\texttt{!t e}) in the two lemmas


569 
is explained in \S\ref{sec:InductionHeuristics}


570 


571 
\begin{exercise}

7569

572 
We strengthen the definition of a \texttt{normal} Ifexpression as follows:

5375

573 
the first argument of all \texttt{IF}s must be a variable. Adapt the above


574 
development to this changed requirement. (Hint: you may need to formulate


575 
some of the goals as implications (\texttt{>}) rather than equalities


576 
(\texttt{=}).)


577 
\end{exercise}


578 


579 
\section{Some basic types}


580 

6577

581 

5375

582 
\subsection{Natural numbers}

6577

583 
\index{arithmetic(}

5375

584 


585 
The type \ttindexbold{nat} of natural numbers is predefined and behaves like


586 
\begin{ttbox}


587 
datatype nat = 0  Suc nat


588 
\end{ttbox}


589 
In particular, there are \texttt{case}expressions, for example


590 
\begin{ttbox}


591 
case n of 0 => 0  Suc m => m


592 
\end{ttbox}


593 
primitive recursion, for example


594 
\begin{ttbox}


595 
\input{Misc/natsum}\end{ttbox}


596 
and induction, for example


597 
\begin{ttbox}


598 
\input{Misc/NatSum.ML}\ttbreak


599 
{\out sum n + sum n = n * Suc n}


600 
{\out No subgoals!}


601 
\end{ttbox}


602 


603 
The usual arithmetic operations \ttindexbold{+}, \ttindexbold{},

6577

604 
\ttindexbold{*}, \ttindexbold{div}, \ttindexbold{mod}, \ttindexbold{min} and


605 
\ttindexbold{max} are predefined, as are the relations \ttindexbold{<=} and


606 
\ttindexbold{<}. There is even a least number operation \ttindexbold{LEAST}.


607 
For example, \texttt{(LEAST n.$\,$1 < n) = 2} (HOL does not prove this


608 
completely automatically).

5375

609 


610 
\begin{warn}

6577

611 
The operations \ttindexbold{+}, \ttindexbold{}, \ttindexbold{*},


612 
\ttindexbold{min}, \ttindexbold{max}, \ttindexbold{<=} and \ttindexbold{<}


613 
are overloaded, i.e.\ they are available not just for natural numbers but


614 
at other types as well (see \S\ref{sec:TypeClasses}). For example, given

5375

615 
the goal \texttt{x+y = y+x}, there is nothing to indicate that you are

6577

616 
talking about natural numbers. Hence Isabelle can only infer that

5375

617 
\texttt{x} and \texttt{y} are of some arbitrary type where \texttt{+} is


618 
declared. As a consequence, you will be unable to prove the goal (although


619 
it may take you some time to realize what has happened if


620 
\texttt{show_types} is not set). In this particular example, you need to


621 
include an explicit type constraint, for example \texttt{x+y = y+(x::nat)}.


622 
If there is enough contextual information this may not be necessary:


623 
\texttt{x+0 = x} automatically implies \texttt{x::nat}.


624 
\end{warn}


625 

6577

626 
Simple arithmetic goals are proved automatically by both \texttt{Auto_tac}


627 
and the simplification tactics introduced in \S\ref{sec:Simplification}. For


628 
example, the goal


629 
\begin{ttbox}


630 
\input{Misc/arith1.ML}\end{ttbox}


631 
is proved automatically. The main restriction is that only addition is taken


632 
into account; other arithmetic operations and quantified formulae are ignored.


633 


634 
For more complex goals, there is the special tactic \ttindexbold{arith_tac}. It


635 
proves arithmetic goals involving the usual logical connectives (\verb$~$,


636 
\verb$&$, \verb$$, \verb$>$), the relations \texttt{<=} and \texttt{<},


637 
and the operations \ttindexbold{+}, \ttindexbold{}, \ttindexbold{min} and


638 
\ttindexbold{max}. For example, it can prove


639 
\begin{ttbox}


640 
\input{Misc/arith2.ML}\end{ttbox}


641 
because \texttt{k*k} can be treated as atomic.


642 
In contrast, $n*n = n \Longrightarrow n=0 \lor n=1$ is not


643 
even proved by \texttt{arith_tac} because the proof relies essentially on


644 
properties of multiplication.


645 


646 
\begin{warn}


647 
The running time of \texttt{arith_tac} is exponential in the number of


648 
occurrences of \ttindexbold{}, \ttindexbold{min} and \ttindexbold{max}


649 
because they are first eliminated by case distinctions.


650 


651 
\texttt{arith_tac} is incomplete even for the restricted class of formulae


652 
described above (known as ``linear arithmetic''). If divisibility plays a


653 
role, it may fail to prove a valid formula, for example $m+m \neq n+n+1$.


654 
Fortunately, such examples are rare in practice.


655 
\end{warn}


656 


657 
\index{arithmetic)}


658 

5375

659 


660 
\subsection{Products}


661 


662 
HOL also has pairs: \texttt{($a@1$,$a@2$)} is of type \texttt{$\tau@1$ *


663 
$\tau@2$} provided each $a@i$ is of type $\tau@i$. The components of a pair


664 
are extracted by \texttt{fst} and \texttt{snd}:


665 
\texttt{fst($x$,$y$) = $x$} and \texttt{snd($x$,$y$) = $y$}. Tuples


666 
are simulated by pairs nested to the right:


667 
\texttt{($a@1$,$a@2$,$a@3$)} and \texttt{$\tau@1$ * $\tau@2$ * $\tau@3$}


668 
stand for \texttt{($a@1$,($a@2$,$a@3$))} and \texttt{$\tau@1$ * ($\tau@2$ *


669 
$\tau@3$)}. Therefore \texttt{fst(snd($a@1$,$a@2$,$a@3$)) = $a@2$}.


670 


671 
It is possible to use (nested) tuples as patterns in abstractions, for


672 
example \texttt{\%(x,y,z).x+y+z} and \texttt{\%((x,y),z).x+y+z}.


673 


674 
In addition to explicit $\lambda$abstractions, tuple patterns can be used in


675 
most variable binding constructs. Typical examples are


676 
\begin{ttbox}


677 
let (x,y) = f z in (y,x)


678 


679 
case xs of [] => 0  (x,y)\#zs => x+y


680 
\end{ttbox}


681 
Further important examples are quantifiers and sets.


682 


683 
\begin{warn}


684 
Abstraction over pairs and tuples is merely a convenient shorthand for a more


685 
complex internal representation. Thus the internal and external form of a


686 
term may differ, which can affect proofs. If you want to avoid this


687 
complication, use \texttt{fst} and \texttt{snd}, i.e.\ write


688 
\texttt{\%p.~fst p + snd p} instead of \texttt{\%(x,y).~x + y}.


689 
See~\S\ref{} for theorem proving with tuple patterns.


690 
\end{warn}


691 


692 


693 
\section{Definitions}


694 
\label{sec:Definitions}


695 


696 
A definition is simply an abbreviation, i.e.\ a new name for an existing


697 
construction. In particular, definitions cannot be recursive. Isabelle offers


698 
definitions on the level of types and terms. Those on the type level are


699 
called type synonyms, those on the term level are called (constant)


700 
definitions.


701 


702 


703 
\subsection{Type synonyms}


704 
\indexbold{type synonyms}


705 


706 
Type synonyms are similar to those found in ML. Their syntax is fairly self


707 
explanatory:


708 
\begin{ttbox}


709 
\input{Misc/types}\end{ttbox}\indexbold{*types}


710 
The synonym \texttt{alist} shows that in general the type on the righthand


711 
side needs to be enclosed in double quotation marks


712 
(see the end of~\S\ref{sec:introtheory}).


713 


714 
Internally all synonyms are fully expanded. As a consequence Isabelle's


715 
output never contains synonyms. Their main purpose is to improve the


716 
readability of theory definitions. Synonyms can be used just like any other


717 
type:


718 
\begin{ttbox}


719 
\input{Misc/consts}\end{ttbox}


720 


721 
\subsection{Constant definitions}


722 
\label{sec:ConstDefinitions}


723 


724 
The above constants \texttt{nand} and \texttt{exor} are nonrecursive and can


725 
therefore be defined directly by


726 
\begin{ttbox}


727 
\input{Misc/defs}\end{ttbox}\indexbold{*defs}


728 
where \texttt{defs} is a keyword and \texttt{nand_def} and \texttt{exor_def}


729 
are arbitrary usersupplied names.


730 
The symbol \texttt{==}\index{==>@{\tt==}bold} is a special form of equality


731 
that should only be used in constant definitions.


732 
Declarations and definitions can also be merged


733 
\begin{ttbox}


734 
\input{Misc/constdefs}\end{ttbox}\indexbold{*constdefs}


735 
in which case the default name of each definition is $f$\texttt{_def}, where


736 
$f$ is the name of the defined constant.


737 


738 
Note that patternmatching is not allowed, i.e.\ each definition must be of


739 
the form $f\,x@1\,\dots\,x@n$\texttt{~==~}$t$.


740 


741 
Section~\S\ref{sec:Simplification} explains how definitions are used


742 
in proofs.


743 


744 
\begin{warn}


745 
A common mistake when writing definitions is to introduce extra free variables


746 
on the righthand side as in the following fictitious definition:


747 
\begin{ttbox}


748 
defs prime_def "prime(p) == (m divides p) > (m=1  m=p)"


749 
\end{ttbox}


750 
Isabelle rejects this `definition' because of the extra {\tt m} on the


751 
righthand side, which would introduce an inconsistency. What you should have


752 
written is


753 
\begin{ttbox}


754 
defs prime_def "prime(p) == !m. (m divides p) > (m=1  m=p)"


755 
\end{ttbox}


756 
\end{warn}


757 


758 


759 


760 


761 
\chapter{More Functional Programming}


762 


763 
The purpose of this chapter is to deepen the reader's understanding of the


764 
concepts encountered so far and to introduce an advanced method for defining


765 
recursive functions. The first two sections give a structured presentation of


766 
theorem proving by simplification (\S\ref{sec:Simplification}) and


767 
discuss important heuristics for induction (\S\ref{sec:InductionHeuristics}). They


768 
can be skipped by readers less interested in proofs. They are followed by a


769 
case study, a compiler for expressions (\S\ref{sec:ExprCompiler}).


770 
Finally we present a very general method for defining recursive functions


771 
that goes well beyond what \texttt{primrec} allows (\S\ref{sec:recdef}).


772 


773 


774 
\section{Simplification}


775 
\label{sec:Simplification}


776 


777 
So far we have proved our theorems by \texttt{Auto_tac}, which


778 
`simplifies' all subgoals. In fact, \texttt{Auto_tac} can do much more than


779 
that, except that it did not need to so far. However, when you go beyond toy


780 
examples, you need to understand the ingredients of \texttt{Auto_tac}.


781 
This section covers the tactic that \texttt{Auto_tac} always applies first,


782 
namely simplification.


783 


784 
Simplification is one of the central theorem proving tools in Isabelle and


785 
many other systems. The tool itself is called the \bfindex{simplifier}. The


786 
purpose of this section is twofold: to introduce the many features of the


787 
simplifier (\S\ref{sec:SimpFeatures}) and to explain a little bit how the


788 
simplifier works (\S\ref{sec:SimpHow}). Anybody intending to use HOL should


789 
read \S\ref{sec:SimpFeatures}, and the serious student should read


790 
\S\ref{sec:SimpHow} as well in order to understand what happened in case


791 
things do not simplify as expected.


792 


793 


794 
\subsection{Using the simplifier}


795 
\label{sec:SimpFeatures}


796 


797 
In its most basic form, simplification means repeated application of


798 
equations from left to right. For example, taking the rules for \texttt{\at}


799 
and applying them to the term \texttt{[0,1] \at\ []} results in a sequence of


800 
simplification steps:


801 
\begin{ttbox}\makeatother


802 
(0#1#[]) @ [] \(\leadsto\) 0#((1#[]) @ []) \(\leadsto\) 0#(1#([] @ [])) \(\leadsto\) 0#1#[]


803 
\end{ttbox}


804 
This is also known as {\em term rewriting} and the equations are referred


805 
to as {\em rewrite rules}. This is more honest than `simplification' because


806 
the terms do not necessarily become simpler in the process.


807 


808 
\subsubsection{Simpsets}


809 


810 
To facilitate simplification, each theory has an associated set of


811 
simplification rules, known as a \bfindex{simpset}. Within a theory,


812 
proofs by simplification refer to the associated simpset by default.


813 
The simpset of a theory is built up as follows: starting with the union of


814 
the simpsets of the parent theories, each occurrence of a \texttt{datatype}


815 
or \texttt{primrec} construct augments the simpset. Explicit definitions are


816 
not added automatically. Users can add new theorems via \texttt{Addsimps} and


817 
delete them again later by \texttt{Delsimps}.


818 


819 
You may augment a simpset not just by equations but by pretty much any


820 
theorem. The simplifier will try to make sense of it. For example, a theorem


821 
\verb$~$$P$ is automatically turned into \texttt{$P$ = False}. The details are


822 
explained in \S\ref{sec:SimpHow}.


823 


824 
As a rule of thumb, rewrite rules that really simplify a term (like


825 
\texttt{xs \at\ [] = xs} and \texttt{rev(rev xs) = xs}) should be added to the


826 
current simpset right after they have been proved. Those of a more specific


827 
nature (e.g.\ distributivity laws, which alter the structure of terms


828 
considerably) should only be added for specific proofs and deleted again


829 
afterwards. Conversely, it may also happen that a generally useful rule


830 
needs to be removed for a certain proof and is added again afterwards. The


831 
need of frequent temporary additions or deletions may indicate a badly


832 
designed simpset.


833 
\begin{warn}


834 
Simplification may not terminate, for example if both $f(x) = g(x)$ and


835 
$g(x) = f(x)$ are in the simpset. It is the user's responsibility not to


836 
include rules that can lead to nontermination, either on their own or in


837 
combination with other rules.


838 
\end{warn}


839 


840 
\subsubsection{Simplification tactics}


841 


842 
There are four main simplification tactics:


843 
\begin{ttdescription}


844 
\item[\ttindexbold{Simp_tac} $i$] simplifies the conclusion of subgoal~$i$


845 
using the theory's simpset. It may solve the subgoal completely if it has


846 
become trivial. For example:


847 
\begin{ttbox}\makeatother


848 
{\out 1. [] @ [] = []}


849 
by(Simp_tac 1);


850 
{\out No subgoals!}


851 
\end{ttbox}


852 


853 
\item[\ttindexbold{Asm_simp_tac}]


854 
is like \verb$Simp_tac$, but extracts additional rewrite rules from


855 
the assumptions of the subgoal. For example, it solves


856 
\begin{ttbox}\makeatother

6691

857 
{\out 1. xs = [] ==> xs @ ys = ys @ xs}

5375

858 
\end{ttbox}


859 
which \texttt{Simp_tac} does not do.


860 


861 
\item[\ttindexbold{Full_simp_tac}] is like \verb$Simp_tac$, but also


862 
simplifies the assumptions (without using the assumptions to


863 
simplify each other or the actual goal).


864 


865 
\item[\ttindexbold{Asm_full_simp_tac}] is like \verb$Asm_simp_tac$,


866 
but also simplifies the assumptions. In particular, assumptions can


867 
simplify each other. For example:


868 
\begin{ttbox}\makeatother


869 
\out{ 1. [ xs @ zs = ys @ xs; [] @ xs = [] @ [] ] ==> ys = zs}


870 
by(Asm_full_simp_tac 1);


871 
{\out No subgoals!}


872 
\end{ttbox}


873 
The second assumption simplifies to \texttt{xs = []}, which in turn


874 
simplifies the first assumption to \texttt{zs = ys}, thus reducing the


875 
conclusion to \texttt{ys = ys} and hence to \texttt{True}.


876 
(See also the paragraph on tracing below.)


877 
\end{ttdescription}


878 
\texttt{Asm_full_simp_tac} is the most powerful of this quartet of


879 
tactics. In fact, \texttt{Auto_tac} starts by applying


880 
\texttt{Asm_full_simp_tac} to all subgoals. The only reason for the existence


881 
of the other three tactics is that sometimes one wants to limit the amount of


882 
simplification, for example to avoid nontermination:


883 
\begin{ttbox}\makeatother


884 
{\out 1. ! x. f x = g (f (g x)) ==> f [] = f [] @ []}


885 
\end{ttbox}


886 
is solved by \texttt{Simp_tac}, but \texttt{Asm_simp_tac} and


887 
\texttt{Asm_full_simp_tac} loop because the rewrite rule \texttt{f x = g(f(g


888 
x))} extracted from the assumption does not terminate. Isabelle notices


889 
certain simple forms of nontermination, but not this one.


890 


891 
\subsubsection{Modifying simpsets locally}


892 


893 
If a certain theorem is merely needed in one proof by simplification, the


894 
pattern


895 
\begin{ttbox}


896 
Addsimps [\(rare_theorem\)];


897 
by(Simp_tac 1);


898 
Delsimps [\(rare_theorem\)];


899 
\end{ttbox}


900 
is awkward. Therefore there are lowercase versions of the simplification


901 
tactics (\ttindexbold{simp_tac}, \ttindexbold{asm_simp_tac},


902 
\ttindexbold{full_simp_tac}, \ttindexbold{asm_full_simp_tac}) and of the


903 
simpset modifiers (\ttindexbold{addsimps}, \ttindexbold{delsimps})


904 
that do not access or modify the implicit simpset but explicitly take a


905 
simpset as an argument. For example, the above three lines become


906 
\begin{ttbox}


907 
by(simp_tac (simpset() addsimps [\(rare_theorem\)]) 1);


908 
\end{ttbox}


909 
where the result of the function call \texttt{simpset()} is the simpset of


910 
the current theory and \texttt{addsimps} is an infix function. The implicit


911 
simpset is read once but not modified.


912 
This is far preferable to pairs of \texttt{Addsimps} and \texttt{Delsimps}.


913 
Local modifications can be stacked as in


914 
\begin{ttbox}


915 
by(simp_tac (simpset() addsimps [\(rare_theorem\)] delsimps [\(some_thm\)]) 1);


916 
\end{ttbox}


917 


918 
\subsubsection{Rewriting with definitions}


919 


920 
Constant definitions (\S\ref{sec:ConstDefinitions}) are not automatically


921 
included in the simpset of a theory. Hence such definitions are not expanded


922 
automatically either, just as it should be: definitions are introduced for


923 
the purpose of abbreviating complex concepts. Of course we need to expand the


924 
definitions initially to derive enough lemmas that characterize the concept


925 
sufficiently for us to forget the original definition completely. For


926 
example, given the theory


927 
\begin{ttbox}


928 
\input{Misc/Exor.thy}\end{ttbox}


929 
we may want to prove \verb$exor A (~A)$. Instead of \texttt{Goal} we use


930 
\begin{ttbox}


931 
\input{Misc/exorgoal.ML}\end{ttbox}


932 
which tells Isabelle to expand the definition of \texttt{exor}the first


933 
argument of \texttt{Goalw} can be a list of definitionsin the initial goal:


934 
\begin{ttbox}


935 
{\out exor A (~ A)}


936 
{\out 1. A & ~ ~ A  ~ A & ~ A}


937 
\end{ttbox}


938 
In this simple example, the goal is proved by \texttt{Simp_tac}.


939 
Of course the resulting theorem is insufficient to characterize \texttt{exor}


940 
completely.


941 


942 
In case we want to expand a definition in the middle of a proof, we can


943 
simply add the definition locally to the simpset:


944 
\begin{ttbox}


945 
\input{Misc/exorproof.ML}\end{ttbox}


946 
You should normally not add the definition permanently to the simpset


947 
using \texttt{Addsimps} because this defeats the whole purpose of an


948 
abbreviation.


949 


950 
\begin{warn}


951 
If you have defined $f\,x\,y$\texttt{~==~}$t$ then you can only expand


952 
occurrences of $f$ with at least two arguments. Thus it is safer to define


953 
$f$\texttt{~==~\%$x\,y$.}$\;t$.


954 
\end{warn}


955 


956 
\subsubsection{Simplifying \texttt{let}expressions}


957 


958 
Proving a goal containing \ttindex{let}expressions invariably requires the


959 
\texttt{let}constructs to be expanded at some point. Since


960 
\texttt{let}\texttt{in} is just syntactic sugar for a defined constant


961 
(called \texttt{Let}), expanding \texttt{let}constructs means rewriting with


962 
\texttt{Let_def}:


963 
%context List.thy;


964 
%Goal "(let xs = [] in xs@xs) = ys";


965 
\begin{ttbox}\makeatother


966 
{\out 1. (let xs = [] in xs @ xs) = ys}


967 
by(simp_tac (simpset() addsimps [Let_def]) 1);


968 
{\out 1. [] = ys}


969 
\end{ttbox}


970 
If, in a particular context, there is no danger of a combinatorial explosion


971 
of nested \texttt{let}s one could even add \texttt{Let_def} permanently via


972 
\texttt{Addsimps}.


973 


974 
\subsubsection{Conditional equations}


975 


976 
So far all examples of rewrite rules were equations. The simplifier also


977 
accepts {\em conditional\/} equations, for example


978 
\begin{ttbox}


979 
xs ~= [] ==> hd xs # tl xs = xs \hfill \((*)\)


980 
\end{ttbox}


981 
(which is proved by \texttt{exhaust_tac} on \texttt{xs} followed by


982 
\texttt{Asm_full_simp_tac} twice). Assuming that this theorem together with


983 
%\begin{ttbox}\makeatother


984 
\texttt{(rev xs = []) = (xs = [])}


985 
%\end{ttbox}


986 
are part of the simpset, the subgoal


987 
\begin{ttbox}\makeatother


988 
{\out 1. xs ~= [] ==> hd(rev xs) # tl(rev xs) = rev xs}


989 
\end{ttbox}


990 
is proved by simplification:


991 
the conditional equation $(*)$ above


992 
can simplify \texttt{hd(rev~xs)~\#~tl(rev~xs)} to \texttt{rev xs}


993 
because the corresponding precondition \verb$rev xs ~= []$ simplifies to


994 
\verb$xs ~= []$, which is exactly the local assumption of the subgoal.


995 


996 


997 
\subsubsection{Automatic case splits}


998 


999 
Goals containing \ttindex{if}expressions are usually proved by case


1000 
distinction on the condition of the \texttt{if}. For example the goal


1001 
\begin{ttbox}


1002 
{\out 1. ! xs. if xs = [] then rev xs = [] else rev xs ~= []}


1003 
\end{ttbox}


1004 
can be split into


1005 
\begin{ttbox}


1006 
{\out 1. ! xs. (xs = [] > rev xs = []) \& (xs ~= [] > rev xs ~= [])}


1007 
\end{ttbox}


1008 
by typing


1009 
\begin{ttbox}


1010 
\input{Misc/splitif.ML}\end{ttbox}


1011 
Because this is almost always the right proof strategy, the simplifier


1012 
performs casesplitting on \texttt{if}s automatically. Try \texttt{Simp_tac}


1013 
on the initial goal above.


1014 


1015 
This splitting idea generalizes from \texttt{if} to \ttindex{case}:


1016 
\begin{ttbox}\makeatother


1017 
{\out 1. (case xs of [] => zs  y#ys => y#(ys@zs)) = xs@zs}


1018 
\end{ttbox}


1019 
becomes


1020 
\begin{ttbox}\makeatother


1021 
{\out 1. (xs = [] > zs = xs @ zs) &}


1022 
{\out (! a list. xs = a # list > a # list @ zs = xs @ zs)}


1023 
\end{ttbox}


1024 
by typing


1025 
\begin{ttbox}


1026 
\input{Misc/splitlist.ML}\end{ttbox}


1027 
In contrast to \texttt{if}expressions, the simplifier does not split


1028 
\texttt{case}expressions by default because this can lead to nontermination


1029 
in case of recursive datatypes.


1030 
Nevertheless the simplifier can be instructed to perform \texttt{case}splits


1031 
by adding the appropriate rule to the simpset:


1032 
\begin{ttbox}


1033 
by(simp_tac (simpset() addsplits [split_list_case]) 1);


1034 
\end{ttbox}\indexbold{*addsplits}


1035 
solves the initial goal outright, which \texttt{Simp_tac} alone will not do.


1036 


1037 
In general, every datatype $t$ comes with a rule


1038 
\texttt{$t$.split} that can be added to the simpset either


1039 
locally via \texttt{addsplits} (see above), or permanently via


1040 
\begin{ttbox}


1041 
Addsplits [\(t\).split];


1042 
\end{ttbox}\indexbold{*Addsplits}


1043 
Splitrules can be removed globally via \ttindexbold{Delsplits} and locally


1044 
via \ttindexbold{delsplits} as, for example, in


1045 
\begin{ttbox}


1046 
by(simp_tac (simpset() addsimps [\(\dots\)] delsplits [split_if]) 1);


1047 
\end{ttbox}


1048 


1049 

6577

1050 
\subsubsection{Arithmetic}


1051 
\index{arithmetic}


1052 


1053 
The simplifier routinely solves a small class of linear arithmetic formulae


1054 
(over types \texttt{nat} and \texttt{int}): it only takes into account


1055 
assumptions and conclusions that are (possibly negated) (in)equalities


1056 
(\texttt{=}, \texttt{<=}, \texttt{<}) and it only knows about addition. Thus


1057 
\begin{ttbox}


1058 
\input{Misc/arith1.ML}\end{ttbox}


1059 
is proved by simplification, whereas the only slightly more complex


1060 
\begin{ttbox}


1061 
\input{Misc/arith3.ML}\end{ttbox}


1062 
is not proved by simplification and requires \texttt{arith_tac}.


1063 

5375

1064 
\subsubsection{Permutative rewrite rules}


1065 


1066 
A rewrite rule is {\bf permutative} if the lefthand side and righthand side


1067 
are the same up to renaming of variables. The most common permutative rule


1068 
is commutativity: $x+y = y+x$. Another example is $(xy)z = (xz)y$. Such


1069 
rules are problematic because once they apply, they can be used forever.


1070 
The simplifier is aware of this danger and treats permutative rules

6606

1071 
separately. For details see~\cite{isabelleref}.

5375

1072 


1073 
\subsubsection{Tracing}


1074 
\indexbold{tracing the simplifier}


1075 


1076 
Using the simplifier effectively may take a bit of experimentation. Set the


1077 
\verb$trace_simp$ flag to get a better idea of what is going on:


1078 
\begin{ttbox}\makeatother


1079 
{\out 1. rev [x] = []}


1080 
\ttbreak


1081 
set trace_simp;


1082 
by(Simp_tac 1);


1083 
\ttbreak\makeatother


1084 
{\out Applying instance of rewrite rule:}


1085 
{\out rev (?x # ?xs) == rev ?xs @ [?x]}


1086 
{\out Rewriting:}


1087 
{\out rev [x] == rev [] @ [x]}


1088 
\ttbreak


1089 
{\out Applying instance of rewrite rule:}


1090 
{\out rev [] == []}


1091 
{\out Rewriting:}


1092 
{\out rev [] == []}


1093 
\ttbreak\makeatother


1094 
{\out Applying instance of rewrite rule:}


1095 
{\out [] @ ?y == ?y}


1096 
{\out Rewriting:}


1097 
{\out [] @ [x] == [x]}


1098 
\ttbreak


1099 
{\out Applying instance of rewrite rule:}


1100 
{\out ?x # ?t = ?t == False}


1101 
{\out Rewriting:}


1102 
{\out [x] = [] == False}


1103 
\ttbreak


1104 
{\out Level 1}


1105 
{\out rev [x] = []}


1106 
{\out 1. False}


1107 
\end{ttbox}


1108 
In more complicated cases, the trace can be enormous, especially since


1109 
invocations of the simplifier are often nested (e.g.\ when solving conditions


1110 
of rewrite rules).


1111 


1112 
\subsection{How it works}


1113 
\label{sec:SimpHow}


1114 


1115 
\subsubsection{Higherorder patterns}


1116 


1117 
\subsubsection{Local assumptions}


1118 


1119 
\subsubsection{The preprocessor}


1120 


1121 
\section{Induction heuristics}


1122 
\label{sec:InductionHeuristics}


1123 


1124 
The purpose of this section is to illustrate some simple heuristics for


1125 
inductive proofs. The first one we have already mentioned in our initial


1126 
example:


1127 
\begin{quote}


1128 
{\em 1. Theorems about recursive functions are proved by induction.}


1129 
\end{quote}


1130 
In case the function has more than one argument


1131 
\begin{quote}


1132 
{\em 2. Do induction on argument number $i$ if the function is defined by


1133 
recursion in argument number $i$.}


1134 
\end{quote}


1135 
When we look at the proof of


1136 
\begin{ttbox}\makeatother


1137 
(xs @ ys) @ zs = xs @ (ys @ zs)


1138 
\end{ttbox}


1139 
in \S\ref{sec:introproof} we find (a) \texttt{\at} is recursive in the first


1140 
argument, (b) \texttt{xs} occurs only as the first argument of \texttt{\at},


1141 
and (c) both \texttt{ys} and \texttt{zs} occur at least once as the second


1142 
argument of \texttt{\at}. Hence it is natural to perform induction on


1143 
\texttt{xs}.


1144 


1145 
The key heuristic, and the main point of this section, is to


1146 
generalize the goal before induction. The reason is simple: if the goal is


1147 
too specific, the induction hypothesis is too weak to allow the induction


1148 
step to go through. Let us now illustrate the idea with an example.


1149 


1150 
We define a tailrecursive version of listreversal,


1151 
i.e.\ one that can be compiled into a loop:


1152 
\begin{ttbox}


1153 
\input{Misc/Itrev.thy}\end{ttbox}


1154 
The behaviour of \texttt{itrev} is simple: it reverses its first argument by


1155 
stacking its elements onto the second argument, and returning that second


1156 
argument when the first one becomes empty.


1157 
We need to show that \texttt{itrev} does indeed reverse its first argument


1158 
provided the second one is empty:


1159 
\begin{ttbox}


1160 
\input{Misc/itrev1.ML}\end{ttbox}


1161 
There is no choice as to the induction variable, and we immediately simplify:


1162 
\begin{ttbox}


1163 
\input{Misc/induct_auto.ML}\ttbreak\makeatother


1164 
{\out1. !!a list. itrev list [] = rev list\(\;\)==> itrev list [a] = rev list @ [a]}


1165 
\end{ttbox}


1166 
Just as predicted above, the overall goal, and hence the induction


1167 
hypothesis, is too weak to solve the induction step because of the fixed


1168 
\texttt{[]}. The corresponding heuristic:


1169 
\begin{quote}


1170 
{\em 3. Generalize goals for induction by replacing constants by variables.}


1171 
\end{quote}


1172 
Of course one cannot do this na\"{\i}vely: \texttt{itrev xs ys = rev xs} is


1173 
just not true  the correct generalization is


1174 
\begin{ttbox}\makeatother


1175 
\input{Misc/itrev2.ML}\end{ttbox}


1176 
If \texttt{ys} is replaced by \texttt{[]}, the righthand side simplifies to


1177 
\texttt{rev xs}, just as required.


1178 


1179 
In this particular instance it is easy to guess the right generalization,


1180 
but in more complex situations a good deal of creativity is needed. This is


1181 
the main source of complications in inductive proofs.


1182 


1183 
Although we now have two variables, only \texttt{xs} is suitable for


1184 
induction, and we repeat our above proof attempt. Unfortunately, we are still


1185 
not there:


1186 
\begin{ttbox}\makeatother


1187 
{\out 1. !!a list.}


1188 
{\out itrev list ys = rev list @ ys}


1189 
{\out ==> itrev list (a # ys) = rev list @ a # ys}


1190 
\end{ttbox}


1191 
The induction hypothesis is still too weak, but this time it takes no


1192 
intuition to generalize: the problem is that \texttt{ys} is fixed throughout


1193 
the subgoal, but the induction hypothesis needs to be applied with


1194 
\texttt{a \# ys} instead of \texttt{ys}. Hence we prove the theorem


1195 
for all \texttt{ys} instead of a fixed one:


1196 
\begin{ttbox}\makeatother


1197 
\input{Misc/itrev3.ML}\end{ttbox}


1198 
This time induction on \texttt{xs} followed by simplification succeeds. This


1199 
leads to another heuristic for generalization:


1200 
\begin{quote}


1201 
{\em 4. Generalize goals for induction by universally quantifying all free


1202 
variables {\em(except the induction variable itself!)}.}


1203 
\end{quote}


1204 
This prevents trivial failures like the above and does not change the


1205 
provability of the goal. Because it is not always required, and may even


1206 
complicate matters in some cases, this heuristic is often not


1207 
applied blindly.


1208 


1209 
A final point worth mentioning is the orientation of the equation we just


1210 
proved: the more complex notion (\texttt{itrev}) is on the lefthand


1211 
side, the simpler \texttt{rev} on the righthand side. This constitutes


1212 
another, albeit weak heuristic that is not restricted to induction:


1213 
\begin{quote}


1214 
{\em 5. The righthand side of an equation should (in some sense) be


1215 
simpler than the lefthand side.}


1216 
\end{quote}


1217 
The heuristic is tricky to apply because it is not obvious that


1218 
\texttt{rev xs \at\ ys} is simpler than \texttt{itrev xs ys}. But see what


1219 
happens if you try to prove the symmetric equation!


1220 


1221 


1222 
\section{Case study: compiling expressions}


1223 
\label{sec:ExprCompiler}


1224 


1225 
The task is to develop a compiler from a generic type of expressions (built


1226 
up from variables, constants and binary operations) to a stack machine. This


1227 
generic type of expressions is a generalization of the boolean expressions in


1228 
\S\ref{sec:boolex}. This time we do not commit ourselves to a particular


1229 
type of variables or values but make them type parameters. Neither is there


1230 
a fixed set of binary operations: instead the expression contains the


1231 
appropriate function itself.


1232 
\begin{ttbox}


1233 
\input{CodeGen/expr}\end{ttbox}


1234 
The three constructors represent constants, variables and the combination of


1235 
two subexpressions with a binary operation.


1236 


1237 
The value of an expression w.r.t.\ an environment that maps variables to


1238 
values is easily defined:


1239 
\begin{ttbox}


1240 
\input{CodeGen/value}\end{ttbox}


1241 


1242 
The stack machine has three instructions: load a constant value onto the


1243 
stack, load the contents of a certain address onto the stack, and apply a


1244 
binary operation to the two topmost elements of the stack, replacing them by


1245 
the result. As for \texttt{expr}, addresses and values are type parameters:


1246 
\begin{ttbox}


1247 
\input{CodeGen/instr}\end{ttbox}


1248 


1249 
The execution of the stack machine is modelled by a function \texttt{exec}


1250 
that takes a store (modelled as a function from addresses to values, just


1251 
like the environment for evaluating expressions), a stack (modelled as a


1252 
list) of values and a list of instructions and returns the stack at the end


1253 
of the execution  the store remains unchanged:


1254 
\begin{ttbox}


1255 
\input{CodeGen/exec}\end{ttbox}


1256 
Recall that \texttt{hd} and \texttt{tl}


1257 
return the first element and the remainder of a list.


1258 


1259 
Because all functions are total, \texttt{hd} is defined even for the empty


1260 
list, although we do not know what the result is. Thus our model of the


1261 
machine always terminates properly, although the above definition does not


1262 
tell us much about the result in situations where \texttt{Apply} was executed


1263 
with fewer than two elements on the stack.


1264 


1265 
The compiler is a function from expressions to a list of instructions. Its


1266 
definition is pretty much obvious:


1267 
\begin{ttbox}\makeatother


1268 
\input{CodeGen/comp}\end{ttbox}


1269 


1270 
Now we have to prove the correctness of the compiler, i.e.\ that the


1271 
execution of a compiled expression results in the value of the expression:


1272 
\begin{ttbox}


1273 
exec s [] (comp e) = [value s e]


1274 
\end{ttbox}


1275 
This is generalized to


1276 
\begin{ttbox}


1277 
\input{CodeGen/goal2.ML}\end{ttbox}


1278 
and proved by induction on \texttt{e} followed by simplification, once we


1279 
have the following lemma about executing the concatenation of two instruction


1280 
sequences:


1281 
\begin{ttbox}\makeatother

6099

1282 
\input{CodeGen/goal1.ML}\end{ttbox}

5375

1283 
This requires induction on \texttt{xs} and ordinary simplification for the


1284 
base cases. In the induction step, simplification leaves us with a formula


1285 
that contains two \texttt{case}expressions over instructions. Thus we add


1286 
automatic case splitting as well, which finishes the proof:


1287 
\begin{ttbox}


1288 
\input{CodeGen/simpsplit.ML}\end{ttbox}


1289 


1290 
We could now go back and prove \texttt{exec s [] (comp e) = [value s e]}


1291 
merely by simplification with the generalized version we just proved.


1292 
However, this is unnecessary because the generalized version fully subsumes


1293 
its instance.


1294 

5850

1295 


1296 
\section{Advanced datatypes}


1297 
\index{*datatype(}


1298 
\index{*primrec(}


1299 

6099

1300 
This section presents advanced forms of \texttt{datatype}s and (in the near


1301 
future!) records.


1302 

5850

1303 
\subsection{Mutual recursion}


1304 


1305 
Sometimes it is necessary to define two datatypes that depend on each


1306 
other. This is called \textbf{mutual recursion}. As an example consider a


1307 
language of arithmetic and boolean expressions where


1308 
\begin{itemize}


1309 
\item arithmetic expressions contain boolean expressions because there are


1310 
conditional arithmetic expressions like ``if $m<n$ then $nm$ else $mn$'',


1311 
and


1312 
\item boolean expressions contain arithmetic expressions because of


1313 
comparisons like ``$m<n$''.


1314 
\end{itemize}


1315 
In Isabelle this becomes


1316 
\begin{ttbox}


1317 
\input{Datatype/abdata}\end{ttbox}\indexbold{*and}


1318 
Type \texttt{aexp} is similar to \texttt{expr} in \S\ref{sec:ExprCompiler},


1319 
except that we have fixed the values to be of type \texttt{nat} and that we


1320 
have fixed the two binary operations \texttt{Sum} and \texttt{Diff}. Boolean


1321 
expressions can be arithmetic comparisons, conjunctions and negations.


1322 
The semantics is fixed via two evaluation functions


1323 
\begin{ttbox}


1324 
\input{Datatype/abconstseval}\end{ttbox}


1325 
that take an environment (a mapping from variables \texttt{'a} to values


1326 
\texttt{nat}) and an expression and return its arithmetic/boolean


1327 
value. Since the datatypes are mutually recursive, so are functions that


1328 
operate on them. Hence they need to be defined in a single \texttt{primrec}


1329 
section:


1330 
\begin{ttbox}


1331 
\input{Datatype/abevala}


1332 
\input{Datatype/abevalb}\end{ttbox}


1333 


1334 
%In general, given $n$ mutually recursive datatypes, whenever you define a


1335 
%\texttt{primrec} function on one of them, Isabelle expects you to define $n$


1336 
%(possibly mutually recursive) functions simultaneously.


1337 


1338 
In the same fashion we also define two functions that perform substitution:


1339 
\begin{ttbox}


1340 
\input{Datatype/abconstssubst}\end{ttbox}


1341 
The first argument is a function mapping variables to expressions, the


1342 
substitution. It is applied to all variables in the second argument. As a


1343 
result, the type of variables in the expression may change from \texttt{'a}


1344 
to \texttt{'b}. Note that there are only arithmetic and no boolean variables.


1345 
\begin{ttbox}


1346 
\input{Datatype/absubsta}


1347 
\input{Datatype/absubstb}\end{ttbox}


1348 


1349 
Now we can prove a fundamental theorem about the interaction between


1350 
evaluation and substitution: applying a substitution $s$ to an expression $a$


1351 
and evaluating the result in an environment $env$ yields the same result as


1352 
evaluation $a$ in the environment that maps every variable $x$ to the value


1353 
of $s(x)$ under $env$. If you try to prove this separately for arithmetic or


1354 
boolean expressions (by induction), you find that you always need the other


1355 
theorem in the induction step. Therefore you need to state and prove both


1356 
theorems simultaneously:


1357 
\begin{quote}\small


1358 
\verbatiminput{Datatype/abgoalind.ML}

7848

1359 
\end{quote}

5850

1360 
The resulting 8 goals (one for each constructor) are proved in one fell swoop


1361 
\texttt{by Auto_tac;}.


1362 


1363 
In general, given $n$ mutually recursive datatypes $\tau@1$, \dots, $\tau@n$,


1364 
Isabelle expects an inductive proof to start with a goal of the form


1365 
\[ P@1(x@1)\ \texttt{\&}\ \dots\ \texttt{\&}\ P@n(x@n) \]


1366 
where each variable $x@i$ is of type $\tau@i$. Induction is started by


1367 
\begin{ttbox}

7848

1368 
by(induct_tac "\(x@1\) \(\dots\) \(x@n\)" \(k\));

5850

1369 
\end{ttbox}


1370 


1371 
\begin{exercise}


1372 
Define a function \texttt{norma} of type \texttt{'a aexp => 'a aexp} that


1373 
replaces \texttt{IF}s with complex boolean conditions by nested


1374 
\texttt{IF}s where each condition is a \texttt{Less}  \texttt{And} and


1375 
\texttt{Neg} should be eliminated completely. Prove that \texttt{norma}


1376 
preserves the value of an expression and that the result of \texttt{norma}


1377 
is really normal, i.e.\ no more \texttt{And}s and \texttt{Neg}s occur in


1378 
it. ({\em Hint:} proceed as in \S\ref{sec:boolex}).


1379 
\end{exercise}


1380 


1381 
\subsection{Nested recursion}


1382 


1383 
So far, all datatypes had the property that on the righthand side of their


1384 
definition they occurred only at the toplevel, i.e.\ directly below a


1385 
constructor. This is not the case any longer for the following model of terms


1386 
where function symbols can be applied to a list of arguments:


1387 
\begin{ttbox}


1388 
\input{Datatype/tdata}\end{ttbox}


1389 
Parameter \texttt{'a} is the type of variables and \texttt{'b} the type of


1390 
function symbols.


1391 
A mathematical term like $f(x,g(y))$ becomes \texttt{App f [Var x, App g


1392 
[Var y]]}, where \texttt{f}, \texttt{g}, \texttt{x}, \texttt{y} are


1393 
suitable values, e.g.\ numbers or strings.


1394 


1395 
What complicates the definition of \texttt{term} is the nested occurrence of


1396 
\texttt{term} inside \texttt{list} on the righthand side. In principle,


1397 
nested recursion can be eliminated in favour of mutual recursion by unfolding


1398 
the offending datatypes, here \texttt{list}. The result for \texttt{term}


1399 
would be something like


1400 
\begin{ttbox}


1401 
\input{Datatype/tunfoldeddata}\end{ttbox}

7569

1402 
Although we do not recommend this unfolding to the user, it shows how to


1403 
simulate nested recursion by mutual recursion.


1404 
Now we return to the initial definition of \texttt{term} using

5850

1405 
nested recursion.


1406 


1407 
Let us define a substitution function on terms. Because terms involve term


1408 
lists, we need to define two substitution functions simultaneously:


1409 
\begin{ttbox}


1410 
\input{Datatype/tconstssubst}


1411 
\input{Datatype/tsubst}


1412 
\input{Datatype/tsubsts}\end{ttbox}


1413 
Similarly, when proving a statement about terms inductively, we need


1414 
to prove a related statement about term lists simultaneously. For example,


1415 
the fact that the identity substitution does not change a term needs to be


1416 
strengthened and proved as follows:


1417 
\begin{quote}\small


1418 
\verbatiminput{Datatype/tidproof.ML}


1419 
\end{quote}


1420 
Note that \texttt{Var} is the identity substitution because by definition it


1421 
leaves variables unchanged: \texttt{subst Var (Var $x$) = Var $x$}. Note also


1422 
that the type annotations are necessary because otherwise there is nothing in

6099

1423 
the goal to enforce that both halves of the goal talk about the same type

5850

1424 
parameters \texttt{('a,'b)}. As a result, induction would fail

6099

1425 
because the two halves of the goal would be unrelated.

5850

1426 


1427 
\begin{exercise}

6099

1428 
The fact that substitution distributes over composition can be expressed


1429 
roughly as follows:

5850

1430 
\begin{ttbox}


1431 
subst (f o g) t = subst f (subst g t)


1432 
\end{ttbox}


1433 
Correct this statement (you will find that it does not typecheck),


1434 
strengthen it and prove it. (Note: \texttt{o} is function composition; its


1435 
definition is found in the theorem \texttt{o_def}).


1436 
\end{exercise}


1437 

7569

1438 
Returning to the definition of \texttt{subst}, we have to admit that it does


1439 
not really need the auxiliary \texttt{substs} function. The \texttt{App}case


1440 
can directly be expressed as

5850

1441 
\begin{ttbox}


1442 
\input{Datatype/appmap}\end{ttbox}

7569

1443 
Although Isabelle insists on the more verbose version, we can easily {\em


1444 
prove} that the \texttt{map}equation holds: simply by induction on


1445 
\texttt{ts} followed by \texttt{Auto_tac}.

5850

1446 


1447 
%FIXME: forward pointer to section where better induction principles are


1448 
%derived?


1449 


1450 
\begin{exercise}


1451 
Define a function \texttt{rev_term} of type \texttt{term > term} that


1452 
recursively reverses the order of arguments of all function symbols in a


1453 
term. Prove that \texttt{rev_term(rev_term t) = t}.


1454 
\end{exercise}


1455 


1456 
Of course, you may also combine mutual and nested recursion as in the


1457 
following example


1458 
\begin{ttbox}


1459 
\input{Datatype/mutnested}\end{ttbox}


1460 
taken from an operational semantics of applied lambda terms. Note that double


1461 
quotes are required around the type involving \texttt{*}, as explained on


1462 
page~\pageref{startype}.


1463 


1464 
\subsection{The limits of nested recursion}


1465 


1466 
How far can we push nested recursion? By the unfolding argument above, we can


1467 
reduce nested to mutual recursion provided the nested recursion only involves

7569

1468 
previously defined datatypes. The \texttt{data} example above involves


1469 
products, which is fine because products are also datatypes.

5850

1470 
However, functions are most emphatically not allowed:


1471 
\begin{ttbox}


1472 
datatype t = C (t => bool)


1473 
\end{ttbox}


1474 
is a real can of worms: in HOL it must be ruled out because it requires a


1475 
type \texttt{t} such that \texttt{t} and its power set \texttt{t => bool}


1476 
have the same cardinalityan impossibility.


1477 
In theory, we could allow limited forms of recursion involving function


1478 
spaces. For example


1479 
\begin{ttbox}


1480 
datatype t = C (nat => t)  D


1481 
\end{ttbox}


1482 
is unproblematic. However, Isabelle does not support recursion involving


1483 
\texttt{=>} at all at the moment.


1484 


1485 
For a theoretical analysis of what kinds of datatypes are feasible in HOL


1486 
see, for example,~\cite{GunterHOL92}. There are alternatives to pure HOL:

6606

1487 
LCF~\cite{paulson87} is a logic where types like

5850

1488 
\begin{ttbox}


1489 
datatype t = C (t > t)


1490 
\end{ttbox}

6099

1491 
do indeed make sense (note the intentionally different arrow \texttt{>}!).

5850

1492 
There is even a version of LCF on top of HOL, called

6606

1493 
HOLCF~\cite{MuellerNvOS99}.

5850

1494 


1495 
\index{*primrec)}


1496 
\index{*datatype)}


1497 


1498 
\subsection{Case study: Tries}


1499 


1500 
Tries are a classic search tree data structure~\cite{Knuth375} for fast


1501 
indexing with strings. Figure~\ref{fig:trie} gives a graphical example of a


1502 
trie containing the words ``all'', ``an'', ``ape'', ``can'', ``car'' and


1503 
``cat''. When searching a string in a trie, the letters of the string are


1504 
examined sequentially. Each letter determines which subtrie to search next.


1505 
In this case study we model tries as a datatype, define a lookup and an


1506 
update function, and prove that they behave as expected.


1507 


1508 
\begin{figure}[htbp]


1509 
\begin{center}


1510 
\unitlength1mm


1511 
\begin{picture}(60,30)


1512 
\put( 5, 0){\makebox(0,0)[b]{l}}


1513 
\put(25, 0){\makebox(0,0)[b]{e}}


1514 
\put(35, 0){\makebox(0,0)[b]{n}}


1515 
\put(45, 0){\makebox(0,0)[b]{r}}
