Incomplete approaches to the incompleteness theorems

The incompleteness theorems appear mysterious to many people, from sheer confusion of the statements themselves, to wrongfully applying the theorems to scenarios way out proportion, such as (dis)proving the existence of god. It doesn’t help that when actually learning about the theorems in a logic course, most details are usually admitted. This is probably not the case at all universities, of course, but I have now personally experienced two different approaches to Gödel’s theorems:

1. Spend most of the time on the recursion theory prerequisites to the theorems, without actually covering the theorems themselves, save for the statements;
2. Skip the recursion theory and only give an informal argument of the incompleteness theorems without really showing why we should care about recursiveness.

The reason for not giving a full account of the theorems is of course the perennial enemy of lecturers: time. What I’ll try to do in this post is still not to give a complete account of the proofs, but try to explain how it all fits together. Fill in the gaps that I’ve at least encountered during my studies, which can then hopefully help others stitch together whichever parts they might have learned throughout their studies. Here we go.

I’m going to focus on the first incompleteness theorem, since the second one is really a corollary to the first one. Let’s start with the statement itself.

First Incompleteness Theorem (Gödel ’31). Let $\mathcal L\supseteq\mathcal L_{\mathbb N}$ be a language. Then any consistent recursive set of $\mathcal L$-sentences $\Gamma$ such that $\Gamma\vdash A_E$, is incomplete.

Let’s start off by examining the statement. We have to figure out what $\mathcal L_{\mathbb N}$ is, what we mean by recursive and what $A_E$ is. Firstly, $\mathcal L_{\mathbb N}:=\{0,S,+,\cdot,E,<\}$ is simply the language of arithmetic, where $S$ is the symbol for the successor function, and $E$ is the symbol for exponentiation. So far so good.

To define recursiveness we first define recursive functions. Before we do that, the intuitive idea behind any recursive set is that we (in principle) can sit down and write a computer program, which takes an object $x$ as input and after a finite amount of time will answer whether $x$ is in our set or not. Now, the simplest way to define what recursive functions are is probably that it’s the smallest set of functions $f:\mathbb N^m\to\mathbb N$ containing the functions:

• Successor $S:\mathbb N\to\mathbb N$;
• Addition $+:\mathbb N^2\to\mathbb N$;
• Multiplication $\cdot:\mathbb N^2\to\mathbb N$;
• Exponentiation $E:\mathbb N^2\to\mathbb N$;
• The characteristic functions $\chi_=:\mathbb N^2\to\mathbb N$ and $\chi_<:\mathbb N^2\to\mathbb N$ of equality = and less-than <;
• Constant functions $c_m:\mathbb N\to\mathbb N$ for all $m\in\mathbb N$;
• Projections $I^m_i:\mathbb N^m\to\mathbb N$, given as $I^m_i(n_1,\hdots,n_m):=n_i$;

and furthermore closed under composition and minimalisation, where the latter means that if $f:\mathbb N^{m+1}\to\mathbb N$ is recursive such that whenever $\vec n\in\mathbb N^m$ there exists $k\in\mathbb N$ satisfying $f(\vec n,k)=0$, then the function $\mu[f]:\mathbb N^m\to\mathbb N$ given as

$\mu[f](\vec n):=\min\{k\in\mathbb N\mid f(\vec n,k=0\}$

is recursive as well. We can also talk about recursive relations $R\subseteq\mathbb N^m$, which is simply when the characteristic function of the relation is recursive.

But we’re not really talking about sets of natural numbers in the statement of the theorem; we’re talking about sets of $\mathcal L$-sentences. To fix this, we encode all symbols in our language $L$, and all sentences of these symbols as numbers. This can be done in many different ways, and any such encoding is called a Gödel encoding. Writing $\sharp\Gamma\subseteq\mathbb N$ for the encoded version of $\Gamma$, we then say $\Gamma$ is recursive if $\sharp\Gamma$ is.

Now, why do we care about recursiveness in the incompleteness context? I’ll get to that, hang on. First, let’s clarify what $A_E$ is. It’s the following (weak) system of axioms in the language of arithmetic, from Enderton’s book:

• (S1) $\forall x Sx\neq 0$;
• (S2) $\forall x\forall y(Sx=Sy\to x=y)$;
• (L1) $\forall x\forall y(x;
• (L2) $\forall x x\not <0$;
• (L3) $\forall x\forall y(x;
• (A1) $\forall x x+0=x$;
• (A2) $\forall x\forall y x+Sy=S(x+y)$;
• (M1) $\forall x x\cdot 0=0$;
• (M2) $\forall x\forall y x\cdot Sy=(x\cdot y)+x$;
• (E1) $\forall x E(x,0)=S0$;
• (E2) $\forall x\forall y E(x,Sy)=E(x,y)\cdot x$.

It’s a fair amount of axioms, but they should all be “intuitively true” when thinking about the standard interpretation of the symbols in $\mathcal L_{\mathbb N}$. The specific choice of system is not important: all we need is a reasonably strong system which is finite.

Okay, so now we at least know what we’re working with. The next (tedious) part, which is the one I’ll skip here, is proving that basically everything in our formula-toolbox is recursive: the set of terms, set of formulas, set of variables, substitution and provability. This last one gives us a recursive relation $\text{Proof}_\Gamma\subseteq\mathbb N^2$, where $\text{Proof}_\Gamma(x,y)$ expresses the fact that $x$ encodes a proof from $\Gamma$ of an $\mathcal L$-formula $\varphi$, and $y$ encodes $\varphi$.

Having happily skipped that part, the next part is explaining why we care about recursiveness. For a relation $R\subseteq\mathbb N^m$ we say that an $\mathcal L_{\mathbb N}$-formula $\varphi(v_1,\hdots,v_m)$ represents $R$ in $A_E$ if, letting $v_1,\hdots,v_m$ be the free variables of $\varphi$ and letting $n_1,\hdots,n_m\in\mathbb N$,

• If $R(n_1,\hdots,n_m)$ holds then $A_E\vdash\varphi(\bar n_1/v_1,\hdots,\bar n_m/v_m)$, and
• If $R(n_1,\hdots,n_m)$ fails then $A_E\vdash\lnot\varphi(\bar n_1/v_1,\hdots,\bar n_m/v_m)$.

Here $\bar n$ is the $\mathcal L_{\mathbb N}$-term $SS\cdots S0$ with $n$ many $S$‘s. That $\varphi$ represents $R$ can be seen as a syntactic strengthening of $R$ being definable over the standard model of arithmetic. Indeed, if $A_E$ was replaced by the theory of that model, then representability would coincide with definability. Now, here’s the reason why we care about recursiveness.

Lemma (aka why recursiveness is useful). Every relation $R\subseteq\mathbb N^m$ is recursive iff it’s representable in $A_E$.

Now we can move on to the (in my mind) fun bits of the argument.

There are (at least) two proofs of the first incompleteness theorem floating about. There’s Gödel’s original method, consisting of constructing an explicit sentence $G$ which witnesses the incompleteness of $\Gamma$, i.e. that $\Gamma\not\vdash G$ and $\Gamma\not\vdash\lnot G$, and the abstract method which only shows that such a sentence exists. The abstract method, which doesn’t construct an explicit witness to the incompleteness, does a detour through Tarski’s Undefinability Theorem and is quite slick, but I’ll stick with Gödel’s constructive approach here instead, to emphasise that we’re really just formalising the liar paradox.

The crucial lemma is then the following, from which the incompleteness theorem is a simple corollary. The lemma essentially says that we can construct self-referential sentences in our weak system $A_E$.

Fixed-point lemma. Let $\varphi(x)$ be an $\mathcal L_{\mathbb N}$-formula. Then there exists an $\mathcal L_{\mathbb N}$-sentence $\sigma$ such that

$A_E\vdash\varphi(\overline{\sharp\sigma}/x)\leftrightarrow\sigma$.

Proof. Let $y\neq x$ be a variable and define a function $F:\mathbb N\to\mathbb N$ as $F(m):=\sharp(\psi(\overline{\sharp\psi}/y))$ whenever $m=\sharp\psi$ for an $\mathcal L_{\mathbb N}$-formula $\psi(y)$, and $F(m):=0$ otherwise.

By the above-mentioned results, this function is recursive and hence also representable in $A_E$, meaning that we get an $\mathcal L_{\mathbb N}$-formula $\theta(w,z)$ such that for any $\mathcal L_{\mathbb N}$-formula $\psi(y)$ we get that

$A_E\vdash\forall z(\theta(\overline{\sharp\psi}/w)\leftrightarrow z=\overline{\sharp(\psi(\overline{\sharp\psi}/y))})$,

just by definition of representability (where we identify $F$ with its graph). We may assume that $w=y$ and $z=x$, so that we wind up with

$A_E\vdash\forall x(\theta(\overline{\sharp\psi}/y)\leftrightarrow x=\overline{\sharp(\psi(\overline{\sharp\psi}/y))})\qquad (1)$

for every $\psi(y)$. Define now $\chi:\equiv\forall x(\theta\to\varphi)$ and note that $\chi$ only has $y$ free. Our desired formula is then

$\sigma:\equiv\chi(\overline{\sharp\chi}/y)\quad (\equiv\forall x(\theta(\overline{\sharp\chi}/y)\to\varphi)\qquad (2)$.

Note here that $\sigma$ says “if $x$ is the Gödel number of the sentence ‘$\chi$ is true of itself’ then $\varphi$ is true of $x$“. But since $\sigma$ also says that “$\chi$ is true of itself”, what $\sigma$ is really saying is that $\varphi$ holds of $\sharp\sigma$! Now, putting together (1) and (2) we get that

$A_E\vdash\forall x(\theta(\overline{\sharp\chi}/y)\leftrightarrow x=\overline{\sharp\sigma})\qquad (3)$.

It remains to show that $\sigma$ actually works, which is shown in the following technical claim.

Claim. $A_E\vdash\varphi(\overline{\sharp\sigma}/x)\leftrightarrow\sigma$.

Proof of claim. $(\leftarrow)$:

• $A_E,\sigma\vdash\theta(\overline{\sharp\chi}/y)(\overline{\sharp\sigma}/x)\to\varphi(\overline{\sharp\sigma}/x)\qquad (\forall{-}E)+(2)$
• $A_E\vdash\theta(\overline{\sharp\chi}/y)(\overline{\sharp\sigma}/x)\qquad (3)$
• $A_E,\sigma\vdash\varphi(\overline{\sharp\sigma}/x)\qquad (MP)$
• $A_E\vdash\sigma\to\varphi(\overline{\sharp\sigma}/x)\qquad (Ded)$.

$(\rightarrow)$:

• $A_E\vdash\theta(\overline{\sharp\chi}/y)\to x=\overline{\sharp\sigma}\qquad (3)$
• $A_E,\theta(\overline{\sharp\chi}/y)\vdash x=\overline{\sharp\sigma}\qquad (Ded)$
• $\vdash x=\overline{\sharp\sigma}\to(\varphi(\overline{\sharp\sigma}/x)\to\varphi(x/x))\qquad (Axiom)$
• $A,\theta(\overline{\sharp\chi}/y)\vdash\varphi(\overline{\sharp\sigma}/x)\to\varphi\qquad (MP)$
• $A_E,\theta(\overline{\sharp\chi}/y),\varphi(\overline{\sharp\sigma}/x)\vdash\varphi\qquad (Ded)$
• $A,\varphi(\overline{\sharp\sigma}/x)\vdash\theta(\overline{\sharp\chi}/y)\to\varphi\qquad (Ded)$
• $A_E,\varphi(\overline{\sharp\sigma}/x)\vdash\forall x(\theta(\overline{\sharp\chi}/y)\to\varphi)\qquad (Gen)$
• $A_E\vdash \varphi(\overline{\sharp\sigma}/x)\to\sigma\qquad (Ded)$

This finishes the proof of the fixed-point lemma. QED

Using the lemma, we can now prove the (first) incompleteness theorem.

Proof of the first incompleteness theorem. Recall that we have a recursive relation $\text{Proof}_\Gamma\subseteq\mathbb N^2$, where $\text{Proof}_\Gamma(x,y)$ expresses the fact that $x$ encodes a proof from $\Gamma$ of an $\mathcal L$-formula $\varphi$, and $y$ encodes $\varphi$. This is then also representable in $A_E$, so we get an $\mathcal L_{\mathbb N}$-formula $\varphi(x,y)$ representing it in $A_E$. Define then

$\textsf{Prv}_\Gamma(x):\equiv\exists z\varphi(z,x)$,

intuitively meaning “$x$ is provable”. More precisely, $\Gamma\vdash\psi$ implies that $A_E\vdash\textsf{Prv}_\Gamma(\overline{\sharp\psi})$, by definition of representability. The fixed-point lemma grants us an $\mathcal L_{\mathbb N}$-sentence $G_\Gamma$ such that

$A_E\vdash\lnot\textsf{Prv}_\Gamma(\overline{\sharp G_\Gamma})\leftrightarrow G_\Gamma\qquad (1)$,

i.e. $G_\Gamma$ says that “I am not provable from $\Gamma$“. Now, if $\Gamma\vdash G_\Gamma$ then $A_E\vdash\textsf{Prv}_\Gamma(\overline{\sharp G_\Gamma})$ by the above, but we also get that $\Gamma\cup A_E\vdash\lnot\textsf{Prv}_\Gamma(\overline{\sharp G_\Gamma})$ by (1), contradicting the consistency of $\Gamma\cup A_E$.

Next, assuming $\Gamma\vdash\lnot G_\Gamma$ we get that $\Gamma\cup A_E\vdash\textsf{Prv}_\Gamma(\overline{\sharp G_\Gamma})$ by (1). Using the consistency of $\Gamma\cup A_E$ we get an $\mathcal L_{\mathbb N}$-model $\mathcal M\vDash\Gamma\cup A_E$. By definition of $\textsf{Prv}_\Gamma$ it holds that $\mathcal M\models\varphi(\overline n,\overline{\sharp G_\Gamma})$ for some $n\in\mathbb N$. The soundness theorem then implies that $A_E\not\vdash\lnot\varphi(\overline n,\overline{\sharp G_\Gamma})$ and representability then yields that $n=\sharp\mathcal D$ for some deduction $\mathcal D$ of $G_\Gamma$ from $\Gamma$. This means that $\Gamma\vdash G_\Gamma$, contradicting the consistency of $\Gamma$ again. QED

Aaaand that finishes this (quite long) post. Hope this is useful to some of you!